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Diffstat (limited to 'Documentation/filesystems')
94 files changed, 25077 insertions, 0 deletions
diff --git a/Documentation/filesystems/00-INDEX b/Documentation/filesystems/00-INDEX new file mode 100644 index 00000000..8c624a18 --- /dev/null +++ b/Documentation/filesystems/00-INDEX @@ -0,0 +1,120 @@ +00-INDEX + - this file (info on some of the filesystems supported by linux). +Locking + - info on locking rules as they pertain to Linux VFS. +9p.txt + - 9p (v9fs) is an implementation of the Plan 9 remote fs protocol. +adfs.txt + - info and mount options for the Acorn Advanced Disc Filing System. +afs.txt + - info and examples for the distributed AFS (Andrew File System) fs. +affs.txt + - info and mount options for the Amiga Fast File System. +automount-support.txt + - information about filesystem automount support. +befs.txt + - information about the BeOS filesystem for Linux. +bfs.txt + - info for the SCO UnixWare Boot Filesystem (BFS). +ceph.txt + - info for the Ceph Distributed File System +cifs.txt + - description of the CIFS filesystem. +coda.txt + - description of the CODA filesystem. +configfs/ + - directory containing configfs documentation and example code. +cramfs.txt + - info on the cram filesystem for small storage (ROMs etc). +dentry-locking.txt + - info on the RCU-based dcache locking model. +directory-locking + - info about the locking scheme used for directory operations. +dlmfs.txt + - info on the userspace interface to the OCFS2 DLM. +dnotify.txt + - info about directory notification in Linux. +dnotify_test.c + - example program for dnotify +ecryptfs.txt + - docs on eCryptfs: stacked cryptographic filesystem for Linux. +exofs.txt + - info, usage, mount options, design about EXOFS. +ext2.txt + - info, mount options and specifications for the Ext2 filesystem. +ext3.txt + - info, mount options and specifications for the Ext3 filesystem. +ext4.txt + - info, mount options and specifications for the Ext4 filesystem. +files.txt + - info on file management in the Linux kernel. +fuse.txt + - info on the Filesystem in User SpacE including mount options. +gfs2.txt + - info on the Global File System 2. +hfs.txt + - info on the Macintosh HFS Filesystem for Linux. +hfsplus.txt + - info on the Macintosh HFSPlus Filesystem for Linux. +hpfs.txt + - info and mount options for the OS/2 HPFS. +inotify.txt + - info on the powerful yet simple file change notification system. +isofs.txt + - info and mount options for the ISO 9660 (CDROM) filesystem. +jfs.txt + - info and mount options for the JFS filesystem. +locks.txt + - info on file locking implementations, flock() vs. fcntl(), etc. +logfs.txt + - info on the LogFS flash filesystem. +mandatory-locking.txt + - info on the Linux implementation of Sys V mandatory file locking. +ncpfs.txt + - info on Novell Netware(tm) filesystem using NCP protocol. +nfs/ + - nfs-related documentation. +nilfs2.txt + - info and mount options for the NILFS2 filesystem. +ntfs.txt + - info and mount options for the NTFS filesystem (Windows NT). +ocfs2.txt + - info and mount options for the OCFS2 clustered filesystem. +porting + - various information on filesystem porting. +proc.txt + - info on Linux's /proc filesystem. +ramfs-rootfs-initramfs.txt + - info on the 'in memory' filesystems ramfs, rootfs and initramfs. +reiser4.txt + - info on the Reiser4 filesystem based on dancing tree algorithms. +relay.txt + - info on relay, for efficient streaming from kernel to user space. +romfs.txt + - description of the ROMFS filesystem. +seq_file.txt + - how to use the seq_file API +sharedsubtree.txt + - a description of shared subtrees for namespaces. +spufs.txt + - info and mount options for the SPU filesystem used on Cell. +sysfs-pci.txt + - info on accessing PCI device resources through sysfs. +sysfs.txt + - info on sysfs, a ram-based filesystem for exporting kernel objects. +sysv-fs.txt + - info on the SystemV/V7/Xenix/Coherent filesystem. +tmpfs.txt + - info on tmpfs, a filesystem that holds all files in virtual memory. +udf.txt + - info and mount options for the UDF filesystem. +ufs.txt + - info on the ufs filesystem. +vfat.txt + - info on using the VFAT filesystem used in Windows NT and Windows 95 +vfs.txt + - overview of the Virtual File System +xfs.txt + - info and mount options for the XFS filesystem. +xip.txt + - info on execute-in-place for file mappings. diff --git a/Documentation/filesystems/9p.txt b/Documentation/filesystems/9p.txt new file mode 100644 index 00000000..13de64c7 --- /dev/null +++ b/Documentation/filesystems/9p.txt @@ -0,0 +1,166 @@ + v9fs: Plan 9 Resource Sharing for Linux + ======================================= + +ABOUT +===== + +v9fs is a Unix implementation of the Plan 9 9p remote filesystem protocol. + +This software was originally developed by Ron Minnich <rminnich@sandia.gov> +and Maya Gokhale. Additional development by Greg Watson +<gwatson@lanl.gov> and most recently Eric Van Hensbergen +<ericvh@gmail.com>, Latchesar Ionkov <lucho@ionkov.net> and Russ Cox +<rsc@swtch.com>. + +The best detailed explanation of the Linux implementation and applications of +the 9p client is available in the form of a USENIX paper: + http://www.usenix.org/events/usenix05/tech/freenix/hensbergen.html + +Other applications are described in the following papers: + * XCPU & Clustering + http://xcpu.org/papers/xcpu-talk.pdf + * KVMFS: control file system for KVM + http://xcpu.org/papers/kvmfs.pdf + * CellFS: A New Programming Model for the Cell BE + http://xcpu.org/papers/cellfs-talk.pdf + * PROSE I/O: Using 9p to enable Application Partitions + http://plan9.escet.urjc.es/iwp9/cready/PROSE_iwp9_2006.pdf + * VirtFS: A Virtualization Aware File System pass-through + http://goo.gl/3WPDg + +USAGE +===== + +For remote file server: + + mount -t 9p 10.10.1.2 /mnt/9 + +For Plan 9 From User Space applications (http://swtch.com/plan9) + + mount -t 9p `namespace`/acme /mnt/9 -o trans=unix,uname=$USER + +For server running on QEMU host with virtio transport: + + mount -t 9p -o trans=virtio <mount_tag> /mnt/9 + +where mount_tag is the tag associated by the server to each of the exported +mount points. Each 9P export is seen by the client as a virtio device with an +associated "mount_tag" property. Available mount tags can be +seen by reading /sys/bus/virtio/drivers/9pnet_virtio/virtio<n>/mount_tag files. + +OPTIONS +======= + + trans=name select an alternative transport. Valid options are + currently: + unix - specifying a named pipe mount point + tcp - specifying a normal TCP/IP connection + fd - used passed file descriptors for connection + (see rfdno and wfdno) + virtio - connect to the next virtio channel available + (from QEMU with trans_virtio module) + rdma - connect to a specified RDMA channel + + uname=name user name to attempt mount as on the remote server. The + server may override or ignore this value. Certain user + names may require authentication. + + aname=name aname specifies the file tree to access when the server is + offering several exported file systems. + + cache=mode specifies a caching policy. By default, no caches are used. + loose = no attempts are made at consistency, + intended for exclusive, read-only mounts + fscache = use FS-Cache for a persistent, read-only + cache backend. + + debug=n specifies debug level. The debug level is a bitmask. + 0x01 = display verbose error messages + 0x02 = developer debug (DEBUG_CURRENT) + 0x04 = display 9p trace + 0x08 = display VFS trace + 0x10 = display Marshalling debug + 0x20 = display RPC debug + 0x40 = display transport debug + 0x80 = display allocation debug + 0x100 = display protocol message debug + 0x200 = display Fid debug + 0x400 = display packet debug + 0x800 = display fscache tracing debug + + rfdno=n the file descriptor for reading with trans=fd + + wfdno=n the file descriptor for writing with trans=fd + + maxdata=n the number of bytes to use for 9p packet payload (msize) + + port=n port to connect to on the remote server + + noextend force legacy mode (no 9p2000.u or 9p2000.L semantics) + + version=name Select 9P protocol version. Valid options are: + 9p2000 - Legacy mode (same as noextend) + 9p2000.u - Use 9P2000.u protocol + 9p2000.L - Use 9P2000.L protocol + + dfltuid attempt to mount as a particular uid + + dfltgid attempt to mount with a particular gid + + afid security channel - used by Plan 9 authentication protocols + + nodevmap do not map special files - represent them as normal files. + This can be used to share devices/named pipes/sockets between + hosts. This functionality will be expanded in later versions. + + access there are four access modes. + user = if a user tries to access a file on v9fs + filesystem for the first time, v9fs sends an + attach command (Tattach) for that user. + This is the default mode. + <uid> = allows only user with uid=<uid> to access + the files on the mounted filesystem + any = v9fs does single attach and performs all + operations as one user + client = ACL based access check on the 9p client + side for access validation + + cachetag cache tag to use the specified persistent cache. + cache tags for existing cache sessions can be listed at + /sys/fs/9p/caches. (applies only to cache=fscache) + +RESOURCES +========= + +Protocol specifications are maintained on github: +http://ericvh.github.com/9p-rfc/ + +9p client and server implementations are listed on +http://9p.cat-v.org/implementations + +A 9p2000.L server is being developed by LLNL and can be found +at http://code.google.com/p/diod/ + +There are user and developer mailing lists available through the v9fs project +on sourceforge (http://sourceforge.net/projects/v9fs). + +News and other information is maintained on a Wiki. +(http://sf.net/apps/mediawiki/v9fs/index.php). + +Bug reports may be issued through the kernel.org bugzilla +(http://bugzilla.kernel.org) + +For more information on the Plan 9 Operating System check out +http://plan9.bell-labs.com/plan9 + +For information on Plan 9 from User Space (Plan 9 applications and libraries +ported to Linux/BSD/OSX/etc) check out http://swtch.com/plan9 + + +STATUS +====== + +The 2.6 kernel support is working on PPC and x86. + +PLEASE USE THE KERNEL BUGZILLA TO REPORT PROBLEMS. (http://bugzilla.kernel.org) + diff --git a/Documentation/filesystems/Locking b/Documentation/filesystems/Locking new file mode 100644 index 00000000..57d827d6 --- /dev/null +++ b/Documentation/filesystems/Locking @@ -0,0 +1,534 @@ + The text below describes the locking rules for VFS-related methods. +It is (believed to be) up-to-date. *Please*, if you change anything in +prototypes or locking protocols - update this file. And update the relevant +instances in the tree, don't leave that to maintainers of filesystems/devices/ +etc. At the very least, put the list of dubious cases in the end of this file. +Don't turn it into log - maintainers of out-of-the-tree code are supposed to +be able to use diff(1). + Thing currently missing here: socket operations. Alexey? + +--------------------------- dentry_operations -------------------------- +prototypes: + int (*d_revalidate)(struct dentry *, struct nameidata *); + int (*d_hash)(const struct dentry *, const struct inode *, + struct qstr *); + int (*d_compare)(const struct dentry *, const struct inode *, + const struct dentry *, const struct inode *, + unsigned int, const char *, const struct qstr *); + int (*d_delete)(struct dentry *); + void (*d_release)(struct dentry *); + void (*d_iput)(struct dentry *, struct inode *); + char *(*d_dname)((struct dentry *dentry, char *buffer, int buflen); + struct vfsmount *(*d_automount)(struct path *path); + int (*d_manage)(struct dentry *, bool); + +locking rules: + rename_lock ->d_lock may block rcu-walk +d_revalidate: no no yes (ref-walk) maybe +d_hash no no no maybe +d_compare: yes no no maybe +d_delete: no yes no no +d_release: no no yes no +d_iput: no no yes no +d_dname: no no no no +d_automount: no no yes no +d_manage: no no yes (ref-walk) maybe + +--------------------------- inode_operations --------------------------- +prototypes: + int (*create) (struct inode *,struct dentry *,int, struct nameidata *); + struct dentry * (*lookup) (struct inode *,struct dentry *, struct nameid +ata *); + int (*link) (struct dentry *,struct inode *,struct dentry *); + int (*unlink) (struct inode *,struct dentry *); + int (*symlink) (struct inode *,struct dentry *,const char *); + int (*mkdir) (struct inode *,struct dentry *,int); + int (*rmdir) (struct inode *,struct dentry *); + int (*mknod) (struct inode *,struct dentry *,int,dev_t); + int (*rename) (struct inode *, struct dentry *, + struct inode *, struct dentry *); + int (*readlink) (struct dentry *, char __user *,int); + void * (*follow_link) (struct dentry *, struct nameidata *); + void (*put_link) (struct dentry *, struct nameidata *, void *); + void (*truncate) (struct inode *); + int (*permission) (struct inode *, int, unsigned int); + int (*check_acl)(struct inode *, int, unsigned int); + int (*setattr) (struct dentry *, struct iattr *); + int (*getattr) (struct vfsmount *, struct dentry *, struct kstat *); + int (*setxattr) (struct dentry *, const char *,const void *,size_t,int); + ssize_t (*getxattr) (struct dentry *, const char *, void *, size_t); + ssize_t (*listxattr) (struct dentry *, char *, size_t); + int (*removexattr) (struct dentry *, const char *); + void (*truncate_range)(struct inode *, loff_t, loff_t); + int (*fiemap)(struct inode *, struct fiemap_extent_info *, u64 start, u64 len); + +locking rules: + all may block + i_mutex(inode) +lookup: yes +create: yes +link: yes (both) +mknod: yes +symlink: yes +mkdir: yes +unlink: yes (both) +rmdir: yes (both) (see below) +rename: yes (all) (see below) +readlink: no +follow_link: no +put_link: no +truncate: yes (see below) +setattr: yes +permission: no (may not block if called in rcu-walk mode) +check_acl: no +getattr: no +setxattr: yes +getxattr: no +listxattr: no +removexattr: yes +truncate_range: yes +fiemap: no + Additionally, ->rmdir(), ->unlink() and ->rename() have ->i_mutex on +victim. + cross-directory ->rename() has (per-superblock) ->s_vfs_rename_sem. + ->truncate() is never called directly - it's a callback, not a +method. It's called by vmtruncate() - deprecated library function used by +->setattr(). Locking information above applies to that call (i.e. is +inherited from ->setattr() - vmtruncate() is used when ATTR_SIZE had been +passed). + +See Documentation/filesystems/directory-locking for more detailed discussion +of the locking scheme for directory operations. + +--------------------------- super_operations --------------------------- +prototypes: + struct inode *(*alloc_inode)(struct super_block *sb); + void (*destroy_inode)(struct inode *); + void (*dirty_inode) (struct inode *, int flags); + int (*write_inode) (struct inode *, struct writeback_control *wbc); + int (*drop_inode) (struct inode *); + void (*evict_inode) (struct inode *); + void (*put_super) (struct super_block *); + void (*write_super) (struct super_block *); + int (*sync_fs)(struct super_block *sb, int wait); + int (*freeze_fs) (struct super_block *); + int (*unfreeze_fs) (struct super_block *); + int (*statfs) (struct dentry *, struct kstatfs *); + int (*remount_fs) (struct super_block *, int *, char *); + void (*umount_begin) (struct super_block *); + int (*show_options)(struct seq_file *, struct vfsmount *); + ssize_t (*quota_read)(struct super_block *, int, char *, size_t, loff_t); + ssize_t (*quota_write)(struct super_block *, int, const char *, size_t, loff_t); + int (*bdev_try_to_free_page)(struct super_block*, struct page*, gfp_t); + +locking rules: + All may block [not true, see below] + s_umount +alloc_inode: +destroy_inode: +dirty_inode: +write_inode: +drop_inode: !!!inode->i_lock!!! +evict_inode: +put_super: write +write_super: read +sync_fs: read +freeze_fs: read +unfreeze_fs: read +statfs: maybe(read) (see below) +remount_fs: write +umount_begin: no +show_options: no (namespace_sem) +quota_read: no (see below) +quota_write: no (see below) +bdev_try_to_free_page: no (see below) + +->statfs() has s_umount (shared) when called by ustat(2) (native or +compat), but that's an accident of bad API; s_umount is used to pin +the superblock down when we only have dev_t given us by userland to +identify the superblock. Everything else (statfs(), fstatfs(), etc.) +doesn't hold it when calling ->statfs() - superblock is pinned down +by resolving the pathname passed to syscall. +->quota_read() and ->quota_write() functions are both guaranteed to +be the only ones operating on the quota file by the quota code (via +dqio_sem) (unless an admin really wants to screw up something and +writes to quota files with quotas on). For other details about locking +see also dquot_operations section. +->bdev_try_to_free_page is called from the ->releasepage handler of +the block device inode. See there for more details. + +--------------------------- file_system_type --------------------------- +prototypes: + int (*get_sb) (struct file_system_type *, int, + const char *, void *, struct vfsmount *); + struct dentry *(*mount) (struct file_system_type *, int, + const char *, void *); + void (*kill_sb) (struct super_block *); +locking rules: + may block +mount yes +kill_sb yes + +->mount() returns ERR_PTR or the root dentry; its superblock should be locked +on return. +->kill_sb() takes a write-locked superblock, does all shutdown work on it, +unlocks and drops the reference. + +--------------------------- address_space_operations -------------------------- +prototypes: + int (*writepage)(struct page *page, struct writeback_control *wbc); + int (*readpage)(struct file *, struct page *); + int (*sync_page)(struct page *); + int (*writepages)(struct address_space *, struct writeback_control *); + int (*set_page_dirty)(struct page *page); + int (*readpages)(struct file *filp, struct address_space *mapping, + struct list_head *pages, unsigned nr_pages); + int (*write_begin)(struct file *, struct address_space *mapping, + loff_t pos, unsigned len, unsigned flags, + struct page **pagep, void **fsdata); + int (*write_end)(struct file *, struct address_space *mapping, + loff_t pos, unsigned len, unsigned copied, + struct page *page, void *fsdata); + sector_t (*bmap)(struct address_space *, sector_t); + int (*invalidatepage) (struct page *, unsigned long); + int (*releasepage) (struct page *, int); + void (*freepage)(struct page *); + int (*direct_IO)(int, struct kiocb *, const struct iovec *iov, + loff_t offset, unsigned long nr_segs); + int (*get_xip_mem)(struct address_space *, pgoff_t, int, void **, + unsigned long *); + int (*migratepage)(struct address_space *, struct page *, struct page *); + int (*launder_page)(struct page *); + int (*is_partially_uptodate)(struct page *, read_descriptor_t *, unsigned long); + int (*error_remove_page)(struct address_space *, struct page *); + +locking rules: + All except set_page_dirty and freepage may block + + PageLocked(page) i_mutex +writepage: yes, unlocks (see below) +readpage: yes, unlocks +sync_page: maybe +writepages: +set_page_dirty no +readpages: +write_begin: locks the page yes +write_end: yes, unlocks yes +bmap: +invalidatepage: yes +releasepage: yes +freepage: yes +direct_IO: +get_xip_mem: maybe +migratepage: yes (both) +launder_page: yes +is_partially_uptodate: yes +error_remove_page: yes + + ->write_begin(), ->write_end(), ->sync_page() and ->readpage() +may be called from the request handler (/dev/loop). + + ->readpage() unlocks the page, either synchronously or via I/O +completion. + + ->readpages() populates the pagecache with the passed pages and starts +I/O against them. They come unlocked upon I/O completion. + + ->writepage() is used for two purposes: for "memory cleansing" and for +"sync". These are quite different operations and the behaviour may differ +depending upon the mode. + +If writepage is called for sync (wbc->sync_mode != WBC_SYNC_NONE) then +it *must* start I/O against the page, even if that would involve +blocking on in-progress I/O. + +If writepage is called for memory cleansing (sync_mode == +WBC_SYNC_NONE) then its role is to get as much writeout underway as +possible. So writepage should try to avoid blocking against +currently-in-progress I/O. + +If the filesystem is not called for "sync" and it determines that it +would need to block against in-progress I/O to be able to start new I/O +against the page the filesystem should redirty the page with +redirty_page_for_writepage(), then unlock the page and return zero. +This may also be done to avoid internal deadlocks, but rarely. + +If the filesystem is called for sync then it must wait on any +in-progress I/O and then start new I/O. + +The filesystem should unlock the page synchronously, before returning to the +caller, unless ->writepage() returns special WRITEPAGE_ACTIVATE +value. WRITEPAGE_ACTIVATE means that page cannot really be written out +currently, and VM should stop calling ->writepage() on this page for some +time. VM does this by moving page to the head of the active list, hence the +name. + +Unless the filesystem is going to redirty_page_for_writepage(), unlock the page +and return zero, writepage *must* run set_page_writeback() against the page, +followed by unlocking it. Once set_page_writeback() has been run against the +page, write I/O can be submitted and the write I/O completion handler must run +end_page_writeback() once the I/O is complete. If no I/O is submitted, the +filesystem must run end_page_writeback() against the page before returning from +writepage. + +That is: after 2.5.12, pages which are under writeout are *not* locked. Note, +if the filesystem needs the page to be locked during writeout, that is ok, too, +the page is allowed to be unlocked at any point in time between the calls to +set_page_writeback() and end_page_writeback(). + +Note, failure to run either redirty_page_for_writepage() or the combination of +set_page_writeback()/end_page_writeback() on a page submitted to writepage +will leave the page itself marked clean but it will be tagged as dirty in the +radix tree. This incoherency can lead to all sorts of hard-to-debug problems +in the filesystem like having dirty inodes at umount and losing written data. + + ->sync_page() locking rules are not well-defined - usually it is called +with lock on page, but that is not guaranteed. Considering the currently +existing instances of this method ->sync_page() itself doesn't look +well-defined... + + ->writepages() is used for periodic writeback and for syscall-initiated +sync operations. The address_space should start I/O against at least +*nr_to_write pages. *nr_to_write must be decremented for each page which is +written. The address_space implementation may write more (or less) pages +than *nr_to_write asks for, but it should try to be reasonably close. If +nr_to_write is NULL, all dirty pages must be written. + +writepages should _only_ write pages which are present on +mapping->io_pages. + + ->set_page_dirty() is called from various places in the kernel +when the target page is marked as needing writeback. It may be called +under spinlock (it cannot block) and is sometimes called with the page +not locked. + + ->bmap() is currently used by legacy ioctl() (FIBMAP) provided by some +filesystems and by the swapper. The latter will eventually go away. Please, +keep it that way and don't breed new callers. + + ->invalidatepage() is called when the filesystem must attempt to drop +some or all of the buffers from the page when it is being truncated. It +returns zero on success. If ->invalidatepage is zero, the kernel uses +block_invalidatepage() instead. + + ->releasepage() is called when the kernel is about to try to drop the +buffers from the page in preparation for freeing it. It returns zero to +indicate that the buffers are (or may be) freeable. If ->releasepage is zero, +the kernel assumes that the fs has no private interest in the buffers. + + ->freepage() is called when the kernel is done dropping the page +from the page cache. + + ->launder_page() may be called prior to releasing a page if +it is still found to be dirty. It returns zero if the page was successfully +cleaned, or an error value if not. Note that in order to prevent the page +getting mapped back in and redirtied, it needs to be kept locked +across the entire operation. + +----------------------- file_lock_operations ------------------------------ +prototypes: + void (*fl_copy_lock)(struct file_lock *, struct file_lock *); + void (*fl_release_private)(struct file_lock *); + + +locking rules: + file_lock_lock may block +fl_copy_lock: yes no +fl_release_private: maybe no + +----------------------- lock_manager_operations --------------------------- +prototypes: + int (*fl_compare_owner)(struct file_lock *, struct file_lock *); + void (*fl_notify)(struct file_lock *); /* unblock callback */ + int (*fl_grant)(struct file_lock *, struct file_lock *, int); + void (*fl_release_private)(struct file_lock *); + void (*fl_break)(struct file_lock *); /* break_lease callback */ + int (*fl_change)(struct file_lock **, int); + +locking rules: + file_lock_lock may block +fl_compare_owner: yes no +fl_notify: yes no +fl_grant: no no +fl_release_private: maybe no +fl_break: yes no +fl_change yes no + +--------------------------- buffer_head ----------------------------------- +prototypes: + void (*b_end_io)(struct buffer_head *bh, int uptodate); + +locking rules: + called from interrupts. In other words, extreme care is needed here. +bh is locked, but that's all warranties we have here. Currently only RAID1, +highmem, fs/buffer.c, and fs/ntfs/aops.c are providing these. Block devices +call this method upon the IO completion. + +--------------------------- block_device_operations ----------------------- +prototypes: + int (*open) (struct block_device *, fmode_t); + int (*release) (struct gendisk *, fmode_t); + int (*ioctl) (struct block_device *, fmode_t, unsigned, unsigned long); + int (*compat_ioctl) (struct block_device *, fmode_t, unsigned, unsigned long); + int (*direct_access) (struct block_device *, sector_t, void **, unsigned long *); + int (*media_changed) (struct gendisk *); + void (*unlock_native_capacity) (struct gendisk *); + int (*revalidate_disk) (struct gendisk *); + int (*getgeo)(struct block_device *, struct hd_geometry *); + void (*swap_slot_free_notify) (struct block_device *, unsigned long); + +locking rules: + bd_mutex +open: yes +release: yes +ioctl: no +compat_ioctl: no +direct_access: no +media_changed: no +unlock_native_capacity: no +revalidate_disk: no +getgeo: no +swap_slot_free_notify: no (see below) + +media_changed, unlock_native_capacity and revalidate_disk are called only from +check_disk_change(). + +swap_slot_free_notify is called with swap_lock and sometimes the page lock +held. + + +--------------------------- file_operations ------------------------------- +prototypes: + loff_t (*llseek) (struct file *, loff_t, int); + ssize_t (*read) (struct file *, char __user *, size_t, loff_t *); + ssize_t (*write) (struct file *, const char __user *, size_t, loff_t *); + ssize_t (*aio_read) (struct kiocb *, const struct iovec *, unsigned long, loff_t); + ssize_t (*aio_write) (struct kiocb *, const struct iovec *, unsigned long, loff_t); + int (*readdir) (struct file *, void *, filldir_t); + unsigned int (*poll) (struct file *, struct poll_table_struct *); + long (*unlocked_ioctl) (struct file *, unsigned int, unsigned long); + long (*compat_ioctl) (struct file *, unsigned int, unsigned long); + int (*mmap) (struct file *, struct vm_area_struct *); + int (*open) (struct inode *, struct file *); + int (*flush) (struct file *); + int (*release) (struct inode *, struct file *); + int (*fsync) (struct file *, int datasync); + int (*aio_fsync) (struct kiocb *, int datasync); + int (*fasync) (int, struct file *, int); + int (*lock) (struct file *, int, struct file_lock *); + ssize_t (*readv) (struct file *, const struct iovec *, unsigned long, + loff_t *); + ssize_t (*writev) (struct file *, const struct iovec *, unsigned long, + loff_t *); + ssize_t (*sendfile) (struct file *, loff_t *, size_t, read_actor_t, + void __user *); + ssize_t (*sendpage) (struct file *, struct page *, int, size_t, + loff_t *, int); + unsigned long (*get_unmapped_area)(struct file *, unsigned long, + unsigned long, unsigned long, unsigned long); + int (*check_flags)(int); + int (*flock) (struct file *, int, struct file_lock *); + ssize_t (*splice_write)(struct pipe_inode_info *, struct file *, loff_t *, + size_t, unsigned int); + ssize_t (*splice_read)(struct file *, loff_t *, struct pipe_inode_info *, + size_t, unsigned int); + int (*setlease)(struct file *, long, struct file_lock **); + long (*fallocate)(struct file *, int, loff_t, loff_t); +}; + +locking rules: + All may block except for ->setlease. + No VFS locks held on entry except for ->fsync and ->setlease. + +->fsync() has i_mutex on inode. + +->setlease has the file_list_lock held and must not sleep. + +->llseek() locking has moved from llseek to the individual llseek +implementations. If your fs is not using generic_file_llseek, you +need to acquire and release the appropriate locks in your ->llseek(). +For many filesystems, it is probably safe to acquire the inode +mutex or just to use i_size_read() instead. +Note: this does not protect the file->f_pos against concurrent modifications +since this is something the userspace has to take care about. + +->fasync() is responsible for maintaining the FASYNC bit in filp->f_flags. +Most instances call fasync_helper(), which does that maintenance, so it's +not normally something one needs to worry about. Return values > 0 will be +mapped to zero in the VFS layer. + +->readdir() and ->ioctl() on directories must be changed. Ideally we would +move ->readdir() to inode_operations and use a separate method for directory +->ioctl() or kill the latter completely. One of the problems is that for +anything that resembles union-mount we won't have a struct file for all +components. And there are other reasons why the current interface is a mess... + +->read on directories probably must go away - we should just enforce -EISDIR +in sys_read() and friends. + +--------------------------- dquot_operations ------------------------------- +prototypes: + int (*write_dquot) (struct dquot *); + int (*acquire_dquot) (struct dquot *); + int (*release_dquot) (struct dquot *); + int (*mark_dirty) (struct dquot *); + int (*write_info) (struct super_block *, int); + +These operations are intended to be more or less wrapping functions that ensure +a proper locking wrt the filesystem and call the generic quota operations. + +What filesystem should expect from the generic quota functions: + + FS recursion Held locks when called +write_dquot: yes dqonoff_sem or dqptr_sem +acquire_dquot: yes dqonoff_sem or dqptr_sem +release_dquot: yes dqonoff_sem or dqptr_sem +mark_dirty: no - +write_info: yes dqonoff_sem + +FS recursion means calling ->quota_read() and ->quota_write() from superblock +operations. + +More details about quota locking can be found in fs/dquot.c. + +--------------------------- vm_operations_struct ----------------------------- +prototypes: + void (*open)(struct vm_area_struct*); + void (*close)(struct vm_area_struct*); + int (*fault)(struct vm_area_struct*, struct vm_fault *); + int (*page_mkwrite)(struct vm_area_struct *, struct vm_fault *); + int (*access)(struct vm_area_struct *, unsigned long, void*, int, int); + +locking rules: + mmap_sem PageLocked(page) +open: yes +close: yes +fault: yes can return with page locked +page_mkwrite: yes can return with page locked +access: yes + + ->fault() is called when a previously not present pte is about +to be faulted in. The filesystem must find and return the page associated +with the passed in "pgoff" in the vm_fault structure. If it is possible that +the page may be truncated and/or invalidated, then the filesystem must lock +the page, then ensure it is not already truncated (the page lock will block +subsequent truncate), and then return with VM_FAULT_LOCKED, and the page +locked. The VM will unlock the page. + + ->page_mkwrite() is called when a previously read-only pte is +about to become writeable. The filesystem again must ensure that there are +no truncate/invalidate races, and then return with the page locked. If +the page has been truncated, the filesystem should not look up a new page +like the ->fault() handler, but simply return with VM_FAULT_NOPAGE, which +will cause the VM to retry the fault. + + ->access() is called when get_user_pages() fails in +acces_process_vm(), typically used to debug a process through +/proc/pid/mem or ptrace. This function is needed only for +VM_IO | VM_PFNMAP VMAs. + +================================================================================ + Dubious stuff + +(if you break something or notice that it is broken and do not fix it yourself +- at least put it here) diff --git a/Documentation/filesystems/Makefile b/Documentation/filesystems/Makefile new file mode 100644 index 00000000..a5dd114d --- /dev/null +++ b/Documentation/filesystems/Makefile @@ -0,0 +1,8 @@ +# kbuild trick to avoid linker error. Can be omitted if a module is built. +obj- := dummy.o + +# List of programs to build +hostprogs-y := dnotify_test + +# Tell kbuild to always build the programs +always := $(hostprogs-y) diff --git a/Documentation/filesystems/adfs.txt b/Documentation/filesystems/adfs.txt new file mode 100644 index 00000000..59497663 --- /dev/null +++ b/Documentation/filesystems/adfs.txt @@ -0,0 +1,75 @@ +Mount options for ADFS +---------------------- + + uid=nnn All files in the partition will be owned by + user id nnn. Default 0 (root). + gid=nnn All files in the partition will be in group + nnn. Default 0 (root). + ownmask=nnn The permission mask for ADFS 'owner' permissions + will be nnn. Default 0700. + othmask=nnn The permission mask for ADFS 'other' permissions + will be nnn. Default 0077. + ftsuffix=n When ftsuffix=0, no file type suffix will be applied. + When ftsuffix=1, a hexadecimal suffix corresponding to + the RISC OS file type will be added. Default 0. + +Mapping of ADFS permissions to Linux permissions +------------------------------------------------ + + ADFS permissions consist of the following: + + Owner read + Owner write + Other read + Other write + + (In older versions, an 'execute' permission did exist, but this + does not hold the same meaning as the Linux 'execute' permission + and is now obsolete). + + The mapping is performed as follows: + + Owner read -> -r--r--r-- + Owner write -> --w--w---w + Owner read and filetype UnixExec -> ---x--x--x + These are then masked by ownmask, eg 700 -> -rwx------ + Possible owner mode permissions -> -rwx------ + + Other read -> -r--r--r-- + Other write -> --w--w--w- + Other read and filetype UnixExec -> ---x--x--x + These are then masked by othmask, eg 077 -> ----rwxrwx + Possible other mode permissions -> ----rwxrwx + + Hence, with the default masks, if a file is owner read/write, and + not a UnixExec filetype, then the permissions will be: + + -rw------- + + However, if the masks were ownmask=0770,othmask=0007, then this would + be modified to: + -rw-rw---- + + There is no restriction on what you can do with these masks. You may + wish that either read bits give read access to the file for all, but + keep the default write protection (ownmask=0755,othmask=0577): + + -rw-r--r-- + + You can therefore tailor the permission translation to whatever you + desire the permissions should be under Linux. + +RISC OS file type suffix +------------------------ + + RISC OS file types are stored in bits 19..8 of the file load address. + + To enable non-RISC OS systems to be used to store files without losing + file type information, a file naming convention was devised (initially + for use with NFS) such that a hexadecimal suffix of the form ,xyz + denoted the file type: e.g. BasicFile,ffb is a BASIC (0xffb) file. This + naming convention is now also used by RISC OS emulators such as RPCEmu. + + Mounting an ADFS disc with option ftsuffix=1 will cause appropriate file + type suffixes to be appended to file names read from a directory. If the + ftsuffix option is zero or omitted, no file type suffixes will be added. diff --git a/Documentation/filesystems/affs.txt b/Documentation/filesystems/affs.txt new file mode 100644 index 00000000..81ac488e --- /dev/null +++ b/Documentation/filesystems/affs.txt @@ -0,0 +1,219 @@ +Overview of Amiga Filesystems +============================= + +Not all varieties of the Amiga filesystems are supported for reading and +writing. The Amiga currently knows six different filesystems: + +DOS\0 The old or original filesystem, not really suited for + hard disks and normally not used on them, either. + Supported read/write. + +DOS\1 The original Fast File System. Supported read/write. + +DOS\2 The old "international" filesystem. International means that + a bug has been fixed so that accented ("international") letters + in file names are case-insensitive, as they ought to be. + Supported read/write. + +DOS\3 The "international" Fast File System. Supported read/write. + +DOS\4 The original filesystem with directory cache. The directory + cache speeds up directory accesses on floppies considerably, + but slows down file creation/deletion. Doesn't make much + sense on hard disks. Supported read only. + +DOS\5 The Fast File System with directory cache. Supported read only. + +All of the above filesystems allow block sizes from 512 to 32K bytes. +Supported block sizes are: 512, 1024, 2048 and 4096 bytes. Larger blocks +speed up almost everything at the expense of wasted disk space. The speed +gain above 4K seems not really worth the price, so you don't lose too +much here, either. + +The muFS (multi user File System) equivalents of the above file systems +are supported, too. + +Mount options for the AFFS +========================== + +protect If this option is set, the protection bits cannot be altered. + +setuid[=uid] This sets the owner of all files and directories in the file + system to uid or the uid of the current user, respectively. + +setgid[=gid] Same as above, but for gid. + +mode=mode Sets the mode flags to the given (octal) value, regardless + of the original permissions. Directories will get an x + permission if the corresponding r bit is set. + This is useful since most of the plain AmigaOS files + will map to 600. + +reserved=num Sets the number of reserved blocks at the start of the + partition to num. You should never need this option. + Default is 2. + +root=block Sets the block number of the root block. This should never + be necessary. + +bs=blksize Sets the blocksize to blksize. Valid block sizes are 512, + 1024, 2048 and 4096. Like the root option, this should + never be necessary, as the affs can figure it out itself. + +quiet The file system will not return an error for disallowed + mode changes. + +verbose The volume name, file system type and block size will + be written to the syslog when the filesystem is mounted. + +mufs The filesystem is really a muFS, also it doesn't + identify itself as one. This option is necessary if + the filesystem wasn't formatted as muFS, but is used + as one. + +prefix=path Path will be prefixed to every absolute path name of + symbolic links on an AFFS partition. Default = "/". + (See below.) + +volume=name When symbolic links with an absolute path are created + on an AFFS partition, name will be prepended as the + volume name. Default = "" (empty string). + (See below.) + +Handling of the Users/Groups and protection flags +================================================= + +Amiga -> Linux: + +The Amiga protection flags RWEDRWEDHSPARWED are handled as follows: + + - R maps to r for user, group and others. On directories, R implies x. + + - If both W and D are allowed, w will be set. + + - E maps to x. + + - H and P are always retained and ignored under Linux. + + - A is always reset when a file is written to. + +User id and group id will be used unless set[gu]id are given as mount +options. Since most of the Amiga file systems are single user systems +they will be owned by root. The root directory (the mount point) of the +Amiga filesystem will be owned by the user who actually mounts the +filesystem (the root directory doesn't have uid/gid fields). + +Linux -> Amiga: + +The Linux rwxrwxrwx file mode is handled as follows: + + - r permission will set R for user, group and others. + + - w permission will set W and D for user, group and others. + + - x permission of the user will set E for plain files. + + - All other flags (suid, sgid, ...) are ignored and will + not be retained. + +Newly created files and directories will get the user and group ID +of the current user and a mode according to the umask. + +Symbolic links +============== + +Although the Amiga and Linux file systems resemble each other, there +are some, not always subtle, differences. One of them becomes apparent +with symbolic links. While Linux has a file system with exactly one +root directory, the Amiga has a separate root directory for each +file system (for example, partition, floppy disk, ...). With the Amiga, +these entities are called "volumes". They have symbolic names which +can be used to access them. Thus, symbolic links can point to a +different volume. AFFS turns the volume name into a directory name +and prepends the prefix path (see prefix option) to it. + +Example: +You mount all your Amiga partitions under /amiga/<volume> (where +<volume> is the name of the volume), and you give the option +"prefix=/amiga/" when mounting all your AFFS partitions. (They +might be "User", "WB" and "Graphics", the mount points /amiga/User, +/amiga/WB and /amiga/Graphics). A symbolic link referring to +"User:sc/include/dos/dos.h" will be followed to +"/amiga/User/sc/include/dos/dos.h". + +Examples +======== + +Command line: + mount Archive/Amiga/Workbench3.1.adf /mnt -t affs -o loop,verbose + mount /dev/sda3 /Amiga -t affs + +/etc/fstab entry: + /dev/sdb5 /amiga/Workbench affs noauto,user,exec,verbose 0 0 + +IMPORTANT NOTE +============== + +If you boot Windows 95 (don't know about 3.x, 98 and NT) while you +have an Amiga harddisk connected to your PC, it will overwrite +the bytes 0x00dc..0x00df of block 0 with garbage, thus invalidating +the Rigid Disk Block. Sheer luck has it that this is an unused +area of the RDB, so only the checksum doesn't match anymore. +Linux will ignore this garbage and recognize the RDB anyway, but +before you connect that drive to your Amiga again, you must +restore or repair your RDB. So please do make a backup copy of it +before booting Windows! + +If the damage is already done, the following should fix the RDB +(where <disk> is the device name). +DO AT YOUR OWN RISK: + + dd if=/dev/<disk> of=rdb.tmp count=1 + cp rdb.tmp rdb.fixed + dd if=/dev/zero of=rdb.fixed bs=1 seek=220 count=4 + dd if=rdb.fixed of=/dev/<disk> + +Bugs, Restrictions, Caveats +=========================== + +Quite a few things may not work as advertised. Not everything is +tested, though several hundred MB have been read and written using +this fs. For a most up-to-date list of bugs please consult +fs/affs/Changes. + +Filenames are truncated to 30 characters without warning (this +can be changed by setting the compile-time option AFFS_NO_TRUNCATE +in include/linux/amigaffs.h). + +Case is ignored by the affs in filename matching, but Linux shells +do care about the case. Example (with /wb being an affs mounted fs): + rm /wb/WRONGCASE +will remove /mnt/wrongcase, but + rm /wb/WR* +will not since the names are matched by the shell. + +The block allocation is designed for hard disk partitions. If more +than 1 process writes to a (small) diskette, the blocks are allocated +in an ugly way (but the real AFFS doesn't do much better). This +is also true when space gets tight. + +You cannot execute programs on an OFS (Old File System), since the +program files cannot be memory mapped due to the 488 byte blocks. +For the same reason you cannot mount an image on such a filesystem +via the loopback device. + +The bitmap valid flag in the root block may not be accurate when the +system crashes while an affs partition is mounted. There's currently +no way to fix a garbled filesystem without an Amiga (disk validator) +or manually (who would do this?). Maybe later. + +If you mount affs partitions on system startup, you may want to tell +fsck that the fs should not be checked (place a '0' in the sixth field +of /etc/fstab). + +It's not possible to read floppy disks with a normal PC or workstation +due to an incompatibility with the Amiga floppy controller. + +If you are interested in an Amiga Emulator for Linux, look at + +http://web.archive.org/web/*/http://www.freiburg.linux.de/~uae/ diff --git a/Documentation/filesystems/afs.txt b/Documentation/filesystems/afs.txt new file mode 100644 index 00000000..ffef91c4 --- /dev/null +++ b/Documentation/filesystems/afs.txt @@ -0,0 +1,247 @@ + ==================== + kAFS: AFS FILESYSTEM + ==================== + +Contents: + + - Overview. + - Usage. + - Mountpoints. + - Proc filesystem. + - The cell database. + - Security. + - Examples. + + +======== +OVERVIEW +======== + +This filesystem provides a fairly simple secure AFS filesystem driver. It is +under development and does not yet provide the full feature set. The features +it does support include: + + (*) Security (currently only AFS kaserver and KerberosIV tickets). + + (*) File reading and writing. + + (*) Automounting. + + (*) Local caching (via fscache). + +It does not yet support the following AFS features: + + (*) pioctl() system call. + + +=========== +COMPILATION +=========== + +The filesystem should be enabled by turning on the kernel configuration +options: + + CONFIG_AF_RXRPC - The RxRPC protocol transport + CONFIG_RXKAD - The RxRPC Kerberos security handler + CONFIG_AFS - The AFS filesystem + +Additionally, the following can be turned on to aid debugging: + + CONFIG_AF_RXRPC_DEBUG - Permit AF_RXRPC debugging to be enabled + CONFIG_AFS_DEBUG - Permit AFS debugging to be enabled + +They permit the debugging messages to be turned on dynamically by manipulating +the masks in the following files: + + /sys/module/af_rxrpc/parameters/debug + /sys/module/kafs/parameters/debug + + +===== +USAGE +===== + +When inserting the driver modules the root cell must be specified along with a +list of volume location server IP addresses: + + modprobe af_rxrpc + modprobe rxkad + modprobe kafs rootcell=cambridge.redhat.com:172.16.18.73:172.16.18.91 + +The first module is the AF_RXRPC network protocol driver. This provides the +RxRPC remote operation protocol and may also be accessed from userspace. See: + + Documentation/networking/rxrpc.txt + +The second module is the kerberos RxRPC security driver, and the third module +is the actual filesystem driver for the AFS filesystem. + +Once the module has been loaded, more modules can be added by the following +procedure: + + echo add grand.central.org 18.9.48.14:128.2.203.61:130.237.48.87 >/proc/fs/afs/cells + +Where the parameters to the "add" command are the name of a cell and a list of +volume location servers within that cell, with the latter separated by colons. + +Filesystems can be mounted anywhere by commands similar to the following: + + mount -t afs "%cambridge.redhat.com:root.afs." /afs + mount -t afs "#cambridge.redhat.com:root.cell." /afs/cambridge + mount -t afs "#root.afs." /afs + mount -t afs "#root.cell." /afs/cambridge + +Where the initial character is either a hash or a percent symbol depending on +whether you definitely want a R/W volume (hash) or whether you'd prefer a R/O +volume, but are willing to use a R/W volume instead (percent). + +The name of the volume can be suffixes with ".backup" or ".readonly" to +specify connection to only volumes of those types. + +The name of the cell is optional, and if not given during a mount, then the +named volume will be looked up in the cell specified during modprobe. + +Additional cells can be added through /proc (see later section). + + +=========== +MOUNTPOINTS +=========== + +AFS has a concept of mountpoints. In AFS terms, these are specially formatted +symbolic links (of the same form as the "device name" passed to mount). kAFS +presents these to the user as directories that have a follow-link capability +(ie: symbolic link semantics). If anyone attempts to access them, they will +automatically cause the target volume to be mounted (if possible) on that site. + +Automatically mounted filesystems will be automatically unmounted approximately +twenty minutes after they were last used. Alternatively they can be unmounted +directly with the umount() system call. + +Manually unmounting an AFS volume will cause any idle submounts upon it to be +culled first. If all are culled, then the requested volume will also be +unmounted, otherwise error EBUSY will be returned. + +This can be used by the administrator to attempt to unmount the whole AFS tree +mounted on /afs in one go by doing: + + umount /afs + + +=============== +PROC FILESYSTEM +=============== + +The AFS modules creates a "/proc/fs/afs/" directory and populates it: + + (*) A "cells" file that lists cells currently known to the afs module and + their usage counts: + + [root@andromeda ~]# cat /proc/fs/afs/cells + USE NAME + 3 cambridge.redhat.com + + (*) A directory per cell that contains files that list volume location + servers, volumes, and active servers known within that cell. + + [root@andromeda ~]# cat /proc/fs/afs/cambridge.redhat.com/servers + USE ADDR STATE + 4 172.16.18.91 0 + [root@andromeda ~]# cat /proc/fs/afs/cambridge.redhat.com/vlservers + ADDRESS + 172.16.18.91 + [root@andromeda ~]# cat /proc/fs/afs/cambridge.redhat.com/volumes + USE STT VLID[0] VLID[1] VLID[2] NAME + 1 Val 20000000 20000001 20000002 root.afs + + +================= +THE CELL DATABASE +================= + +The filesystem maintains an internal database of all the cells it knows and the +IP addresses of the volume location servers for those cells. The cell to which +the system belongs is added to the database when modprobe is performed by the +"rootcell=" argument or, if compiled in, using a "kafs.rootcell=" argument on +the kernel command line. + +Further cells can be added by commands similar to the following: + + echo add CELLNAME VLADDR[:VLADDR][:VLADDR]... >/proc/fs/afs/cells + echo add grand.central.org 18.9.48.14:128.2.203.61:130.237.48.87 >/proc/fs/afs/cells + +No other cell database operations are available at this time. + + +======== +SECURITY +======== + +Secure operations are initiated by acquiring a key using the klog program. A +very primitive klog program is available at: + + http://people.redhat.com/~dhowells/rxrpc/klog.c + +This should be compiled by: + + make klog LDLIBS="-lcrypto -lcrypt -lkrb4 -lkeyutils" + +And then run as: + + ./klog + +Assuming it's successful, this adds a key of type RxRPC, named for the service +and cell, eg: "afs@<cellname>". This can be viewed with the keyctl program or +by cat'ing /proc/keys: + + [root@andromeda ~]# keyctl show + Session Keyring + -3 --alswrv 0 0 keyring: _ses.3268 + 2 --alswrv 0 0 \_ keyring: _uid.0 + 111416553 --als--v 0 0 \_ rxrpc: afs@CAMBRIDGE.REDHAT.COM + +Currently the username, realm, password and proposed ticket lifetime are +compiled in to the program. + +It is not required to acquire a key before using AFS facilities, but if one is +not acquired then all operations will be governed by the anonymous user parts +of the ACLs. + +If a key is acquired, then all AFS operations, including mounts and automounts, +made by a possessor of that key will be secured with that key. + +If a file is opened with a particular key and then the file descriptor is +passed to a process that doesn't have that key (perhaps over an AF_UNIX +socket), then the operations on the file will be made with key that was used to +open the file. + + +======== +EXAMPLES +======== + +Here's what I use to test this. Some of the names and IP addresses are local +to my internal DNS. My "root.afs" partition has a mount point within it for +some public volumes volumes. + +insmod /tmp/rxrpc.o +insmod /tmp/rxkad.o +insmod /tmp/kafs.o rootcell=cambridge.redhat.com:172.16.18.91 + +mount -t afs \%root.afs. /afs +mount -t afs \%cambridge.redhat.com:root.cell. /afs/cambridge.redhat.com/ + +echo add grand.central.org 18.9.48.14:128.2.203.61:130.237.48.87 > /proc/fs/afs/cells +mount -t afs "#grand.central.org:root.cell." /afs/grand.central.org/ +mount -t afs "#grand.central.org:root.archive." /afs/grand.central.org/archive +mount -t afs "#grand.central.org:root.contrib." /afs/grand.central.org/contrib +mount -t afs "#grand.central.org:root.doc." /afs/grand.central.org/doc +mount -t afs "#grand.central.org:root.project." /afs/grand.central.org/project +mount -t afs "#grand.central.org:root.service." /afs/grand.central.org/service +mount -t afs "#grand.central.org:root.software." /afs/grand.central.org/software +mount -t afs "#grand.central.org:root.user." /afs/grand.central.org/user + +umount /afs +rmmod kafs +rmmod rxkad +rmmod rxrpc diff --git a/Documentation/filesystems/autofs4-mount-control.txt b/Documentation/filesystems/autofs4-mount-control.txt new file mode 100644 index 00000000..4c95935c --- /dev/null +++ b/Documentation/filesystems/autofs4-mount-control.txt @@ -0,0 +1,393 @@ + +Miscellaneous Device control operations for the autofs4 kernel module +==================================================================== + +The problem +=========== + +There is a problem with active restarts in autofs (that is to say +restarting autofs when there are busy mounts). + +During normal operation autofs uses a file descriptor opened on the +directory that is being managed in order to be able to issue control +operations. Using a file descriptor gives ioctl operations access to +autofs specific information stored in the super block. The operations +are things such as setting an autofs mount catatonic, setting the +expire timeout and requesting expire checks. As is explained below, +certain types of autofs triggered mounts can end up covering an autofs +mount itself which prevents us being able to use open(2) to obtain a +file descriptor for these operations if we don't already have one open. + +Currently autofs uses "umount -l" (lazy umount) to clear active mounts +at restart. While using lazy umount works for most cases, anything that +needs to walk back up the mount tree to construct a path, such as +getcwd(2) and the proc file system /proc/<pid>/cwd, no longer works +because the point from which the path is constructed has been detached +from the mount tree. + +The actual problem with autofs is that it can't reconnect to existing +mounts. Immediately one thinks of just adding the ability to remount +autofs file systems would solve it, but alas, that can't work. This is +because autofs direct mounts and the implementation of "on demand mount +and expire" of nested mount trees have the file system mounted directly +on top of the mount trigger directory dentry. + +For example, there are two types of automount maps, direct (in the kernel +module source you will see a third type called an offset, which is just +a direct mount in disguise) and indirect. + +Here is a master map with direct and indirect map entries: + +/- /etc/auto.direct +/test /etc/auto.indirect + +and the corresponding map files: + +/etc/auto.direct: + +/automount/dparse/g6 budgie:/autofs/export1 +/automount/dparse/g1 shark:/autofs/export1 +and so on. + +/etc/auto.indirect: + +g1 shark:/autofs/export1 +g6 budgie:/autofs/export1 +and so on. + +For the above indirect map an autofs file system is mounted on /test and +mounts are triggered for each sub-directory key by the inode lookup +operation. So we see a mount of shark:/autofs/export1 on /test/g1, for +example. + +The way that direct mounts are handled is by making an autofs mount on +each full path, such as /automount/dparse/g1, and using it as a mount +trigger. So when we walk on the path we mount shark:/autofs/export1 "on +top of this mount point". Since these are always directories we can +use the follow_link inode operation to trigger the mount. + +But, each entry in direct and indirect maps can have offsets (making +them multi-mount map entries). + +For example, an indirect mount map entry could also be: + +g1 \ + / shark:/autofs/export5/testing/test \ + /s1 shark:/autofs/export/testing/test/s1 \ + /s2 shark:/autofs/export5/testing/test/s2 \ + /s1/ss1 shark:/autofs/export1 \ + /s2/ss2 shark:/autofs/export2 + +and a similarly a direct mount map entry could also be: + +/automount/dparse/g1 \ + / shark:/autofs/export5/testing/test \ + /s1 shark:/autofs/export/testing/test/s1 \ + /s2 shark:/autofs/export5/testing/test/s2 \ + /s1/ss1 shark:/autofs/export2 \ + /s2/ss2 shark:/autofs/export2 + +One of the issues with version 4 of autofs was that, when mounting an +entry with a large number of offsets, possibly with nesting, we needed +to mount and umount all of the offsets as a single unit. Not really a +problem, except for people with a large number of offsets in map entries. +This mechanism is used for the well known "hosts" map and we have seen +cases (in 2.4) where the available number of mounts are exhausted or +where the number of privileged ports available is exhausted. + +In version 5 we mount only as we go down the tree of offsets and +similarly for expiring them which resolves the above problem. There is +somewhat more detail to the implementation but it isn't needed for the +sake of the problem explanation. The one important detail is that these +offsets are implemented using the same mechanism as the direct mounts +above and so the mount points can be covered by a mount. + +The current autofs implementation uses an ioctl file descriptor opened +on the mount point for control operations. The references held by the +descriptor are accounted for in checks made to determine if a mount is +in use and is also used to access autofs file system information held +in the mount super block. So the use of a file handle needs to be +retained. + + +The Solution +============ + +To be able to restart autofs leaving existing direct, indirect and +offset mounts in place we need to be able to obtain a file handle +for these potentially covered autofs mount points. Rather than just +implement an isolated operation it was decided to re-implement the +existing ioctl interface and add new operations to provide this +functionality. + +In addition, to be able to reconstruct a mount tree that has busy mounts, +the uid and gid of the last user that triggered the mount needs to be +available because these can be used as macro substitution variables in +autofs maps. They are recorded at mount request time and an operation +has been added to retrieve them. + +Since we're re-implementing the control interface, a couple of other +problems with the existing interface have been addressed. First, when +a mount or expire operation completes a status is returned to the +kernel by either a "send ready" or a "send fail" operation. The +"send fail" operation of the ioctl interface could only ever send +ENOENT so the re-implementation allows user space to send an actual +status. Another expensive operation in user space, for those using +very large maps, is discovering if a mount is present. Usually this +involves scanning /proc/mounts and since it needs to be done quite +often it can introduce significant overhead when there are many entries +in the mount table. An operation to lookup the mount status of a mount +point dentry (covered or not) has also been added. + +Current kernel development policy recommends avoiding the use of the +ioctl mechanism in favor of systems such as Netlink. An implementation +using this system was attempted to evaluate its suitability and it was +found to be inadequate, in this case. The Generic Netlink system was +used for this as raw Netlink would lead to a significant increase in +complexity. There's no question that the Generic Netlink system is an +elegant solution for common case ioctl functions but it's not a complete +replacement probably because its primary purpose in life is to be a +message bus implementation rather than specifically an ioctl replacement. +While it would be possible to work around this there is one concern +that lead to the decision to not use it. This is that the autofs +expire in the daemon has become far to complex because umount +candidates are enumerated, almost for no other reason than to "count" +the number of times to call the expire ioctl. This involves scanning +the mount table which has proved to be a big overhead for users with +large maps. The best way to improve this is try and get back to the +way the expire was done long ago. That is, when an expire request is +issued for a mount (file handle) we should continually call back to +the daemon until we can't umount any more mounts, then return the +appropriate status to the daemon. At the moment we just expire one +mount at a time. A Generic Netlink implementation would exclude this +possibility for future development due to the requirements of the +message bus architecture. + + +autofs4 Miscellaneous Device mount control interface +==================================================== + +The control interface is opening a device node, typically /dev/autofs. + +All the ioctls use a common structure to pass the needed parameter +information and return operation results: + +struct autofs_dev_ioctl { + __u32 ver_major; + __u32 ver_minor; + __u32 size; /* total size of data passed in + * including this struct */ + __s32 ioctlfd; /* automount command fd */ + + __u32 arg1; /* Command parameters */ + __u32 arg2; + + char path[0]; +}; + +The ioctlfd field is a mount point file descriptor of an autofs mount +point. It is returned by the open call and is used by all calls except +the check for whether a given path is a mount point, where it may +optionally be used to check a specific mount corresponding to a given +mount point file descriptor, and when requesting the uid and gid of the +last successful mount on a directory within the autofs file system. + +The fields arg1 and arg2 are used to communicate parameters and results of +calls made as described below. + +The path field is used to pass a path where it is needed and the size field +is used account for the increased structure length when translating the +structure sent from user space. + +This structure can be initialized before setting specific fields by using +the void function call init_autofs_dev_ioctl(struct autofs_dev_ioctl *). + +All of the ioctls perform a copy of this structure from user space to +kernel space and return -EINVAL if the size parameter is smaller than +the structure size itself, -ENOMEM if the kernel memory allocation fails +or -EFAULT if the copy itself fails. Other checks include a version check +of the compiled in user space version against the module version and a +mismatch results in a -EINVAL return. If the size field is greater than +the structure size then a path is assumed to be present and is checked to +ensure it begins with a "/" and is NULL terminated, otherwise -EINVAL is +returned. Following these checks, for all ioctl commands except +AUTOFS_DEV_IOCTL_VERSION_CMD, AUTOFS_DEV_IOCTL_OPENMOUNT_CMD and +AUTOFS_DEV_IOCTL_CLOSEMOUNT_CMD the ioctlfd is validated and if it is +not a valid descriptor or doesn't correspond to an autofs mount point +an error of -EBADF, -ENOTTY or -EINVAL (not an autofs descriptor) is +returned. + + +The ioctls +========== + +An example of an implementation which uses this interface can be seen +in autofs version 5.0.4 and later in file lib/dev-ioctl-lib.c of the +distribution tar available for download from kernel.org in directory +/pub/linux/daemons/autofs/v5. + +The device node ioctl operations implemented by this interface are: + + +AUTOFS_DEV_IOCTL_VERSION +------------------------ + +Get the major and minor version of the autofs4 device ioctl kernel module +implementation. It requires an initialized struct autofs_dev_ioctl as an +input parameter and sets the version information in the passed in structure. +It returns 0 on success or the error -EINVAL if a version mismatch is +detected. + + +AUTOFS_DEV_IOCTL_PROTOVER_CMD and AUTOFS_DEV_IOCTL_PROTOSUBVER_CMD +------------------------------------------------------------------ + +Get the major and minor version of the autofs4 protocol version understood +by loaded module. This call requires an initialized struct autofs_dev_ioctl +with the ioctlfd field set to a valid autofs mount point descriptor +and sets the requested version number in structure field arg1. These +commands return 0 on success or one of the negative error codes if +validation fails. + + +AUTOFS_DEV_IOCTL_OPENMOUNT and AUTOFS_DEV_IOCTL_CLOSEMOUNT +---------------------------------------------------------- + +Obtain and release a file descriptor for an autofs managed mount point +path. The open call requires an initialized struct autofs_dev_ioctl with +the the path field set and the size field adjusted appropriately as well +as the arg1 field set to the device number of the autofs mount. The +device number can be obtained from the mount options shown in +/proc/mounts. The close call requires an initialized struct +autofs_dev_ioct with the ioctlfd field set to the descriptor obtained +from the open call. The release of the file descriptor can also be done +with close(2) so any open descriptors will also be closed at process exit. +The close call is included in the implemented operations largely for +completeness and to provide for a consistent user space implementation. + + +AUTOFS_DEV_IOCTL_READY_CMD and AUTOFS_DEV_IOCTL_FAIL_CMD +-------------------------------------------------------- + +Return mount and expire result status from user space to the kernel. +Both of these calls require an initialized struct autofs_dev_ioctl +with the ioctlfd field set to the descriptor obtained from the open +call and the arg1 field set to the wait queue token number, received +by user space in the foregoing mount or expire request. The arg2 field +is set to the status to be returned. For the ready call this is always +0 and for the fail call it is set to the errno of the operation. + + +AUTOFS_DEV_IOCTL_SETPIPEFD_CMD +------------------------------ + +Set the pipe file descriptor used for kernel communication to the daemon. +Normally this is set at mount time using an option but when reconnecting +to a existing mount we need to use this to tell the autofs mount about +the new kernel pipe descriptor. In order to protect mounts against +incorrectly setting the pipe descriptor we also require that the autofs +mount be catatonic (see next call). + +The call requires an initialized struct autofs_dev_ioctl with the +ioctlfd field set to the descriptor obtained from the open call and +the arg1 field set to descriptor of the pipe. On success the call +also sets the process group id used to identify the controlling process +(eg. the owning automount(8) daemon) to the process group of the caller. + + +AUTOFS_DEV_IOCTL_CATATONIC_CMD +------------------------------ + +Make the autofs mount point catatonic. The autofs mount will no longer +issue mount requests, the kernel communication pipe descriptor is released +and any remaining waits in the queue released. + +The call requires an initialized struct autofs_dev_ioctl with the +ioctlfd field set to the descriptor obtained from the open call. + + +AUTOFS_DEV_IOCTL_TIMEOUT_CMD +---------------------------- + +Set the expire timeout for mounts within an autofs mount point. + +The call requires an initialized struct autofs_dev_ioctl with the +ioctlfd field set to the descriptor obtained from the open call. + + +AUTOFS_DEV_IOCTL_REQUESTER_CMD +------------------------------ + +Return the uid and gid of the last process to successfully trigger a the +mount on the given path dentry. + +The call requires an initialized struct autofs_dev_ioctl with the path +field set to the mount point in question and the size field adjusted +appropriately as well as the arg1 field set to the device number of the +containing autofs mount. Upon return the struct field arg1 contains the +uid and arg2 the gid. + +When reconstructing an autofs mount tree with active mounts we need to +re-connect to mounts that may have used the original process uid and +gid (or string variations of them) for mount lookups within the map entry. +This call provides the ability to obtain this uid and gid so they may be +used by user space for the mount map lookups. + + +AUTOFS_DEV_IOCTL_EXPIRE_CMD +--------------------------- + +Issue an expire request to the kernel for an autofs mount. Typically +this ioctl is called until no further expire candidates are found. + +The call requires an initialized struct autofs_dev_ioctl with the +ioctlfd field set to the descriptor obtained from the open call. In +addition an immediate expire, independent of the mount timeout, can be +requested by setting the arg1 field to 1. If no expire candidates can +be found the ioctl returns -1 with errno set to EAGAIN. + +This call causes the kernel module to check the mount corresponding +to the given ioctlfd for mounts that can be expired, issues an expire +request back to the daemon and waits for completion. + +AUTOFS_DEV_IOCTL_ASKUMOUNT_CMD +------------------------------ + +Checks if an autofs mount point is in use. + +The call requires an initialized struct autofs_dev_ioctl with the +ioctlfd field set to the descriptor obtained from the open call and +it returns the result in the arg1 field, 1 for busy and 0 otherwise. + + +AUTOFS_DEV_IOCTL_ISMOUNTPOINT_CMD +--------------------------------- + +Check if the given path is a mountpoint. + +The call requires an initialized struct autofs_dev_ioctl. There are two +possible variations. Both use the path field set to the path of the mount +point to check and the size field adjusted appropriately. One uses the +ioctlfd field to identify a specific mount point to check while the other +variation uses the path and optionally arg1 set to an autofs mount type. +The call returns 1 if this is a mount point and sets arg1 to the device +number of the mount and field arg2 to the relevant super block magic +number (described below) or 0 if it isn't a mountpoint. In both cases +the the device number (as returned by new_encode_dev()) is returned +in field arg1. + +If supplied with a file descriptor we're looking for a specific mount, +not necessarily at the top of the mounted stack. In this case the path +the descriptor corresponds to is considered a mountpoint if it is itself +a mountpoint or contains a mount, such as a multi-mount without a root +mount. In this case we return 1 if the descriptor corresponds to a mount +point and and also returns the super magic of the covering mount if there +is one or 0 if it isn't a mountpoint. + +If a path is supplied (and the ioctlfd field is set to -1) then the path +is looked up and is checked to see if it is the root of a mount. If a +type is also given we are looking for a particular autofs mount and if +a match isn't found a fail is returned. If the the located path is the +root of a mount 1 is returned along with the super magic of the mount +or 0 otherwise. + diff --git a/Documentation/filesystems/automount-support.txt b/Documentation/filesystems/automount-support.txt new file mode 100644 index 00000000..7cac200e --- /dev/null +++ b/Documentation/filesystems/automount-support.txt @@ -0,0 +1,118 @@ +Support is available for filesystems that wish to do automounting support (such +as kAFS which can be found in fs/afs/). This facility includes allowing +in-kernel mounts to be performed and mountpoint degradation to be +requested. The latter can also be requested by userspace. + + +====================== +IN-KERNEL AUTOMOUNTING +====================== + +A filesystem can now mount another filesystem on one of its directories by the +following procedure: + + (1) Give the directory a follow_link() operation. + + When the directory is accessed, the follow_link op will be called, and + it will be provided with the location of the mountpoint in the nameidata + structure (vfsmount and dentry). + + (2) Have the follow_link() op do the following steps: + + (a) Call vfs_kern_mount() to call the appropriate filesystem to set up a + superblock and gain a vfsmount structure representing it. + + (b) Copy the nameidata provided as an argument and substitute the dentry + argument into it the copy. + + (c) Call do_add_mount() to install the new vfsmount into the namespace's + mountpoint tree, thus making it accessible to userspace. Use the + nameidata set up in (b) as the destination. + + If the mountpoint will be automatically expired, then do_add_mount() + should also be given the location of an expiration list (see further + down). + + (d) Release the path in the nameidata argument and substitute in the new + vfsmount and its root dentry. The ref counts on these will need + incrementing. + +Then from userspace, you can just do something like: + + [root@andromeda root]# mount -t afs \#root.afs. /afs + [root@andromeda root]# ls /afs + asd cambridge cambridge.redhat.com grand.central.org + [root@andromeda root]# ls /afs/cambridge + afsdoc + [root@andromeda root]# ls /afs/cambridge/afsdoc/ + ChangeLog html LICENSE pdf RELNOTES-1.2.2 + +And then if you look in the mountpoint catalogue, you'll see something like: + + [root@andromeda root]# cat /proc/mounts + ... + #root.afs. /afs afs rw 0 0 + #root.cell. /afs/cambridge.redhat.com afs rw 0 0 + #afsdoc. /afs/cambridge.redhat.com/afsdoc afs rw 0 0 + + +=========================== +AUTOMATIC MOUNTPOINT EXPIRY +=========================== + +Automatic expiration of mountpoints is easy, provided you've mounted the +mountpoint to be expired in the automounting procedure outlined above. + +To do expiration, you need to follow these steps: + + (3) Create at least one list off which the vfsmounts to be expired can be + hung. Access to this list will be governed by the vfsmount_lock. + + (4) In step (2c) above, the call to do_add_mount() should be provided with a + pointer to this list. It will hang the vfsmount off of it if it succeeds. + + (5) When you want mountpoints to be expired, call mark_mounts_for_expiry() + with a pointer to this list. This will process the list, marking every + vfsmount thereon for potential expiry on the next call. + + If a vfsmount was already flagged for expiry, and if its usage count is 1 + (it's only referenced by its parent vfsmount), then it will be deleted + from the namespace and thrown away (effectively unmounted). + + It may prove simplest to simply call this at regular intervals, using + some sort of timed event to drive it. + +The expiration flag is cleared by calls to mntput. This means that expiration +will only happen on the second expiration request after the last time the +mountpoint was accessed. + +If a mountpoint is moved, it gets removed from the expiration list. If a bind +mount is made on an expirable mount, the new vfsmount will not be on the +expiration list and will not expire. + +If a namespace is copied, all mountpoints contained therein will be copied, +and the copies of those that are on an expiration list will be added to the +same expiration list. + + +======================= +USERSPACE DRIVEN EXPIRY +======================= + +As an alternative, it is possible for userspace to request expiry of any +mountpoint (though some will be rejected - the current process's idea of the +rootfs for example). It does this by passing the MNT_EXPIRE flag to +umount(). This flag is considered incompatible with MNT_FORCE and MNT_DETACH. + +If the mountpoint in question is in referenced by something other than +umount() or its parent mountpoint, an EBUSY error will be returned and the +mountpoint will not be marked for expiration or unmounted. + +If the mountpoint was not already marked for expiry at that time, an EAGAIN +error will be given and it won't be unmounted. + +Otherwise if it was already marked and it wasn't referenced, unmounting will +take place as usual. + +Again, the expiration flag is cleared every time anything other than umount() +looks at a mountpoint. diff --git a/Documentation/filesystems/befs.txt b/Documentation/filesystems/befs.txt new file mode 100644 index 00000000..6e49c363 --- /dev/null +++ b/Documentation/filesystems/befs.txt @@ -0,0 +1,117 @@ +BeOS filesystem for Linux + +Document last updated: Dec 6, 2001 + +WARNING +======= +Make sure you understand that this is alpha software. This means that the +implementation is neither complete nor well-tested. + +I DISCLAIM ALL RESPONSIBILITY FOR ANY POSSIBLE BAD EFFECTS OF THIS CODE! + +LICENSE +===== +This software is covered by the GNU General Public License. +See the file COPYING for the complete text of the license. +Or the GNU website: <http://www.gnu.org/licenses/licenses.html> + +AUTHOR +===== +The largest part of the code written by Will Dyson <will_dyson@pobox.com> +He has been working on the code since Aug 13, 2001. See the changelog for +details. + +Original Author: Makoto Kato <m_kato@ga2.so-net.ne.jp> +His original code can still be found at: +<http://hp.vector.co.jp/authors/VA008030/bfs/> +Does anyone know of a more current email address for Makoto? He doesn't +respond to the address given above... + +Current maintainer: Sergey S. Kostyliov <rathamahata@php4.ru> + +WHAT IS THIS DRIVER? +================== +This module implements the native filesystem of BeOS http://www.beincorporated.com/ +for the linux 2.4.1 and later kernels. Currently it is a read-only +implementation. + +Which is it, BFS or BEFS? +================ +Be, Inc said, "BeOS Filesystem is officially called BFS, not BeFS". +But Unixware Boot Filesystem is called bfs, too. And they are already in +the kernel. Because of this naming conflict, on Linux the BeOS +filesystem is called befs. + +HOW TO INSTALL +============== +step 1. Install the BeFS patch into the source code tree of linux. + +Apply the patchfile to your kernel source tree. +Assuming that your kernel source is in /foo/bar/linux and the patchfile +is called patch-befs-xxx, you would do the following: + + cd /foo/bar/linux + patch -p1 < /path/to/patch-befs-xxx + +if the patching step fails (i.e. there are rejected hunks), you can try to +figure it out yourself (it shouldn't be hard), or mail the maintainer +(Will Dyson <will_dyson@pobox.com>) for help. + +step 2. Configuration & make kernel + +The linux kernel has many compile-time options. Most of them are beyond the +scope of this document. I suggest the Kernel-HOWTO document as a good general +reference on this topic. http://www.linuxdocs.org/HOWTOs/Kernel-HOWTO-4.html + +However, to use the BeFS module, you must enable it at configure time. + + cd /foo/bar/linux + make menuconfig (or xconfig) + +The BeFS module is not a standard part of the linux kernel, so you must first +enable support for experimental code under the "Code maturity level" menu. + +Then, under the "Filesystems" menu will be an option called "BeFS +filesystem (experimental)", or something like that. Enable that option +(it is fine to make it a module). + +Save your kernel configuration and then build your kernel. + +step 3. Install + +See the kernel howto <http://www.linux.com/howto/Kernel-HOWTO.html> for +instructions on this critical step. + +USING BFS +========= +To use the BeOS filesystem, use filesystem type 'befs'. + +ex) + mount -t befs /dev/fd0 /beos + +MOUNT OPTIONS +============= +uid=nnn All files in the partition will be owned by user id nnn. +gid=nnn All files in the partition will be in group nnn. +iocharset=xxx Use xxx as the name of the NLS translation table. +debug The driver will output debugging information to the syslog. + +HOW TO GET LASTEST VERSION +========================== + +The latest version is currently available at: +<http://befs-driver.sourceforge.net/> + +ANY KNOWN BUGS? +=========== +As of Jan 20, 2002: + + None + +SPECIAL THANKS +============== +Dominic Giampalo ... Writing "Practical file system design with Be filesystem" +Hiroyuki Yamada ... Testing LinuxPPC. + + + diff --git a/Documentation/filesystems/bfs.txt b/Documentation/filesystems/bfs.txt new file mode 100644 index 00000000..78043d5a --- /dev/null +++ b/Documentation/filesystems/bfs.txt @@ -0,0 +1,57 @@ +BFS FILESYSTEM FOR LINUX +======================== + +The BFS filesystem is used by SCO UnixWare OS for the /stand slice, which +usually contains the kernel image and a few other files required for the +boot process. + +In order to access /stand partition under Linux you obviously need to +know the partition number and the kernel must support UnixWare disk slices +(CONFIG_UNIXWARE_DISKLABEL config option). However BFS support does not +depend on having UnixWare disklabel support because one can also mount +BFS filesystem via loopback: + +# losetup /dev/loop0 stand.img +# mount -t bfs /dev/loop0 /mnt/stand + +where stand.img is a file containing the image of BFS filesystem. +When you have finished using it and umounted you need to also deallocate +/dev/loop0 device by: + +# losetup -d /dev/loop0 + +You can simplify mounting by just typing: + +# mount -t bfs -o loop stand.img /mnt/stand + +this will allocate the first available loopback device (and load loop.o +kernel module if necessary) automatically. If the loopback driver is not +loaded automatically, make sure that you have compiled the module and +that modprobe is functioning. Beware that umount will not deallocate +/dev/loopN device if /etc/mtab file on your system is a symbolic link to +/proc/mounts. You will need to do it manually using "-d" switch of +losetup(8). Read losetup(8) manpage for more info. + +To create the BFS image under UnixWare you need to find out first which +slice contains it. The command prtvtoc(1M) is your friend: + +# prtvtoc /dev/rdsk/c0b0t0d0s0 + +(assuming your root disk is on target=0, lun=0, bus=0, controller=0). Then you +look for the slice with tag "STAND", which is usually slice 10. With this +information you can use dd(1) to create the BFS image: + +# umount /stand +# dd if=/dev/rdsk/c0b0t0d0sa of=stand.img bs=512 + +Just in case, you can verify that you have done the right thing by checking +the magic number: + +# od -Ad -tx4 stand.img | more + +The first 4 bytes should be 0x1badface. + +If you have any patches, questions or suggestions regarding this BFS +implementation please contact the author: + +Tigran Aivazian <tigran@aivazian.fsnet.co.uk> diff --git a/Documentation/filesystems/btrfs.txt b/Documentation/filesystems/btrfs.txt new file mode 100644 index 00000000..64087c34 --- /dev/null +++ b/Documentation/filesystems/btrfs.txt @@ -0,0 +1,91 @@ + + BTRFS + ===== + +Btrfs is a new copy on write filesystem for Linux aimed at +implementing advanced features while focusing on fault tolerance, +repair and easy administration. Initially developed by Oracle, Btrfs +is licensed under the GPL and open for contribution from anyone. + +Linux has a wealth of filesystems to choose from, but we are facing a +number of challenges with scaling to the large storage subsystems that +are becoming common in today's data centers. Filesystems need to scale +in their ability to address and manage large storage, and also in +their ability to detect, repair and tolerate errors in the data stored +on disk. Btrfs is under heavy development, and is not suitable for +any uses other than benchmarking and review. The Btrfs disk format is +not yet finalized. + +The main Btrfs features include: + + * Extent based file storage (2^64 max file size) + * Space efficient packing of small files + * Space efficient indexed directories + * Dynamic inode allocation + * Writable snapshots + * Subvolumes (separate internal filesystem roots) + * Object level mirroring and striping + * Checksums on data and metadata (multiple algorithms available) + * Compression + * Integrated multiple device support, with several raid algorithms + * Online filesystem check (not yet implemented) + * Very fast offline filesystem check + * Efficient incremental backup and FS mirroring (not yet implemented) + * Online filesystem defragmentation + + + + MAILING LIST + ============ + +There is a Btrfs mailing list hosted on vger.kernel.org. You can +find details on how to subscribe here: + +http://vger.kernel.org/vger-lists.html#linux-btrfs + +Mailing list archives are available from gmane: + +http://dir.gmane.org/gmane.comp.file-systems.btrfs + + + + IRC + === + +Discussion of Btrfs also occurs on the #btrfs channel of the Freenode +IRC network. + + + + UTILITIES + ========= + +Userspace tools for creating and manipulating Btrfs file systems are +available from the git repository at the following location: + + http://git.kernel.org/?p=linux/kernel/git/mason/btrfs-progs-unstable.git + git://git.kernel.org/pub/scm/linux/kernel/git/mason/btrfs-progs-unstable.git + +These include the following tools: + +mkfs.btrfs: create a filesystem + +btrfsctl: control program to create snapshots and subvolumes: + + mount /dev/sda2 /mnt + btrfsctl -s new_subvol_name /mnt + btrfsctl -s snapshot_of_default /mnt/default + btrfsctl -s snapshot_of_new_subvol /mnt/new_subvol_name + btrfsctl -s snapshot_of_a_snapshot /mnt/snapshot_of_new_subvol + ls /mnt + default snapshot_of_a_snapshot snapshot_of_new_subvol + new_subvol_name snapshot_of_default + + Snapshots and subvolumes cannot be deleted right now, but you can + rm -rf all the files and directories inside them. + +btrfsck: do a limited check of the FS extent trees. + +btrfs-debug-tree: print all of the FS metadata in text form. Example: + + btrfs-debug-tree /dev/sda2 >& big_output_file diff --git a/Documentation/filesystems/caching/backend-api.txt b/Documentation/filesystems/caching/backend-api.txt new file mode 100644 index 00000000..382d52cd --- /dev/null +++ b/Documentation/filesystems/caching/backend-api.txt @@ -0,0 +1,658 @@ + ========================== + FS-CACHE CACHE BACKEND API + ========================== + +The FS-Cache system provides an API by which actual caches can be supplied to +FS-Cache for it to then serve out to network filesystems and other interested +parties. + +This API is declared in <linux/fscache-cache.h>. + + +==================================== +INITIALISING AND REGISTERING A CACHE +==================================== + +To start off, a cache definition must be initialised and registered for each +cache the backend wants to make available. For instance, CacheFS does this in +the fill_super() operation on mounting. + +The cache definition (struct fscache_cache) should be initialised by calling: + + void fscache_init_cache(struct fscache_cache *cache, + struct fscache_cache_ops *ops, + const char *idfmt, + ...); + +Where: + + (*) "cache" is a pointer to the cache definition; + + (*) "ops" is a pointer to the table of operations that the backend supports on + this cache; and + + (*) "idfmt" is a format and printf-style arguments for constructing a label + for the cache. + + +The cache should then be registered with FS-Cache by passing a pointer to the +previously initialised cache definition to: + + int fscache_add_cache(struct fscache_cache *cache, + struct fscache_object *fsdef, + const char *tagname); + +Two extra arguments should also be supplied: + + (*) "fsdef" which should point to the object representation for the FS-Cache + master index in this cache. Netfs primary index entries will be created + here. FS-Cache keeps the caller's reference to the index object if + successful and will release it upon withdrawal of the cache. + + (*) "tagname" which, if given, should be a text string naming this cache. If + this is NULL, the identifier will be used instead. For CacheFS, the + identifier is set to name the underlying block device and the tag can be + supplied by mount. + +This function may return -ENOMEM if it ran out of memory or -EEXIST if the tag +is already in use. 0 will be returned on success. + + +===================== +UNREGISTERING A CACHE +===================== + +A cache can be withdrawn from the system by calling this function with a +pointer to the cache definition: + + void fscache_withdraw_cache(struct fscache_cache *cache); + +In CacheFS's case, this is called by put_super(). + + +======== +SECURITY +======== + +The cache methods are executed one of two contexts: + + (1) that of the userspace process that issued the netfs operation that caused + the cache method to be invoked, or + + (2) that of one of the processes in the FS-Cache thread pool. + +In either case, this may not be an appropriate context in which to access the +cache. + +The calling process's fsuid, fsgid and SELinux security identities may need to +be masqueraded for the duration of the cache driver's access to the cache. +This is left to the cache to handle; FS-Cache makes no effort in this regard. + + +=================================== +CONTROL AND STATISTICS PRESENTATION +=================================== + +The cache may present data to the outside world through FS-Cache's interfaces +in sysfs and procfs - the former for control and the latter for statistics. + +A sysfs directory called /sys/fs/fscache/<cachetag>/ is created if CONFIG_SYSFS +is enabled. This is accessible through the kobject struct fscache_cache::kobj +and is for use by the cache as it sees fit. + + +======================== +RELEVANT DATA STRUCTURES +======================== + + (*) Index/Data file FS-Cache representation cookie: + + struct fscache_cookie { + struct fscache_object_def *def; + struct fscache_netfs *netfs; + void *netfs_data; + ... + }; + + The fields that might be of use to the backend describe the object + definition, the netfs definition and the netfs's data for this cookie. + The object definition contain functions supplied by the netfs for loading + and matching index entries; these are required to provide some of the + cache operations. + + + (*) In-cache object representation: + + struct fscache_object { + int debug_id; + enum { + FSCACHE_OBJECT_RECYCLING, + ... + } state; + spinlock_t lock + struct fscache_cache *cache; + struct fscache_cookie *cookie; + ... + }; + + Structures of this type should be allocated by the cache backend and + passed to FS-Cache when requested by the appropriate cache operation. In + the case of CacheFS, they're embedded in CacheFS's internal object + structures. + + The debug_id is a simple integer that can be used in debugging messages + that refer to a particular object. In such a case it should be printed + using "OBJ%x" to be consistent with FS-Cache. + + Each object contains a pointer to the cookie that represents the object it + is backing. An object should retired when put_object() is called if it is + in state FSCACHE_OBJECT_RECYCLING. The fscache_object struct should be + initialised by calling fscache_object_init(object). + + + (*) FS-Cache operation record: + + struct fscache_operation { + atomic_t usage; + struct fscache_object *object; + unsigned long flags; + #define FSCACHE_OP_EXCLUSIVE + void (*processor)(struct fscache_operation *op); + void (*release)(struct fscache_operation *op); + ... + }; + + FS-Cache has a pool of threads that it uses to give CPU time to the + various asynchronous operations that need to be done as part of driving + the cache. These are represented by the above structure. The processor + method is called to give the op CPU time, and the release method to get + rid of it when its usage count reaches 0. + + An operation can be made exclusive upon an object by setting the + appropriate flag before enqueuing it with fscache_enqueue_operation(). If + an operation needs more processing time, it should be enqueued again. + + + (*) FS-Cache retrieval operation record: + + struct fscache_retrieval { + struct fscache_operation op; + struct address_space *mapping; + struct list_head *to_do; + ... + }; + + A structure of this type is allocated by FS-Cache to record retrieval and + allocation requests made by the netfs. This struct is then passed to the + backend to do the operation. The backend may get extra refs to it by + calling fscache_get_retrieval() and refs may be discarded by calling + fscache_put_retrieval(). + + A retrieval operation can be used by the backend to do retrieval work. To + do this, the retrieval->op.processor method pointer should be set + appropriately by the backend and fscache_enqueue_retrieval() called to + submit it to the thread pool. CacheFiles, for example, uses this to queue + page examination when it detects PG_lock being cleared. + + The to_do field is an empty list available for the cache backend to use as + it sees fit. + + + (*) FS-Cache storage operation record: + + struct fscache_storage { + struct fscache_operation op; + pgoff_t store_limit; + ... + }; + + A structure of this type is allocated by FS-Cache to record outstanding + writes to be made. FS-Cache itself enqueues this operation and invokes + the write_page() method on the object at appropriate times to effect + storage. + + +================ +CACHE OPERATIONS +================ + +The cache backend provides FS-Cache with a table of operations that can be +performed on the denizens of the cache. These are held in a structure of type: + + struct fscache_cache_ops + + (*) Name of cache provider [mandatory]: + + const char *name + + This isn't strictly an operation, but should be pointed at a string naming + the backend. + + + (*) Allocate a new object [mandatory]: + + struct fscache_object *(*alloc_object)(struct fscache_cache *cache, + struct fscache_cookie *cookie) + + This method is used to allocate a cache object representation to back a + cookie in a particular cache. fscache_object_init() should be called on + the object to initialise it prior to returning. + + This function may also be used to parse the index key to be used for + multiple lookup calls to turn it into a more convenient form. FS-Cache + will call the lookup_complete() method to allow the cache to release the + form once lookup is complete or aborted. + + + (*) Look up and create object [mandatory]: + + void (*lookup_object)(struct fscache_object *object) + + This method is used to look up an object, given that the object is already + allocated and attached to the cookie. This should instantiate that object + in the cache if it can. + + The method should call fscache_object_lookup_negative() as soon as + possible if it determines the object doesn't exist in the cache. If the + object is found to exist and the netfs indicates that it is valid then + fscache_obtained_object() should be called once the object is in a + position to have data stored in it. Similarly, fscache_obtained_object() + should also be called once a non-present object has been created. + + If a lookup error occurs, fscache_object_lookup_error() should be called + to abort the lookup of that object. + + + (*) Release lookup data [mandatory]: + + void (*lookup_complete)(struct fscache_object *object) + + This method is called to ask the cache to release any resources it was + using to perform a lookup. + + + (*) Increment object refcount [mandatory]: + + struct fscache_object *(*grab_object)(struct fscache_object *object) + + This method is called to increment the reference count on an object. It + may fail (for instance if the cache is being withdrawn) by returning NULL. + It should return the object pointer if successful. + + + (*) Lock/Unlock object [mandatory]: + + void (*lock_object)(struct fscache_object *object) + void (*unlock_object)(struct fscache_object *object) + + These methods are used to exclusively lock an object. It must be possible + to schedule with the lock held, so a spinlock isn't sufficient. + + + (*) Pin/Unpin object [optional]: + + int (*pin_object)(struct fscache_object *object) + void (*unpin_object)(struct fscache_object *object) + + These methods are used to pin an object into the cache. Once pinned an + object cannot be reclaimed to make space. Return -ENOSPC if there's not + enough space in the cache to permit this. + + + (*) Update object [mandatory]: + + int (*update_object)(struct fscache_object *object) + + This is called to update the index entry for the specified object. The + new information should be in object->cookie->netfs_data. This can be + obtained by calling object->cookie->def->get_aux()/get_attr(). + + + (*) Discard object [mandatory]: + + void (*drop_object)(struct fscache_object *object) + + This method is called to indicate that an object has been unbound from its + cookie, and that the cache should release the object's resources and + retire it if it's in state FSCACHE_OBJECT_RECYCLING. + + This method should not attempt to release any references held by the + caller. The caller will invoke the put_object() method as appropriate. + + + (*) Release object reference [mandatory]: + + void (*put_object)(struct fscache_object *object) + + This method is used to discard a reference to an object. The object may + be freed when all the references to it are released. + + + (*) Synchronise a cache [mandatory]: + + void (*sync)(struct fscache_cache *cache) + + This is called to ask the backend to synchronise a cache with its backing + device. + + + (*) Dissociate a cache [mandatory]: + + void (*dissociate_pages)(struct fscache_cache *cache) + + This is called to ask a cache to perform any page dissociations as part of + cache withdrawal. + + + (*) Notification that the attributes on a netfs file changed [mandatory]: + + int (*attr_changed)(struct fscache_object *object); + + This is called to indicate to the cache that certain attributes on a netfs + file have changed (for example the maximum size a file may reach). The + cache can read these from the netfs by calling the cookie's get_attr() + method. + + The cache may use the file size information to reserve space on the cache. + It should also call fscache_set_store_limit() to indicate to FS-Cache the + highest byte it's willing to store for an object. + + This method may return -ve if an error occurred or the cache object cannot + be expanded. In such a case, the object will be withdrawn from service. + + This operation is run asynchronously from FS-Cache's thread pool, and + storage and retrieval operations from the netfs are excluded during the + execution of this operation. + + + (*) Reserve cache space for an object's data [optional]: + + int (*reserve_space)(struct fscache_object *object, loff_t size); + + This is called to request that cache space be reserved to hold the data + for an object and the metadata used to track it. Zero size should be + taken as request to cancel a reservation. + + This should return 0 if successful, -ENOSPC if there isn't enough space + available, or -ENOMEM or -EIO on other errors. + + The reservation may exceed the current size of the object, thus permitting + future expansion. If the amount of space consumed by an object would + exceed the reservation, it's permitted to refuse requests to allocate + pages, but not required. An object may be pruned down to its reservation + size if larger than that already. + + + (*) Request page be read from cache [mandatory]: + + int (*read_or_alloc_page)(struct fscache_retrieval *op, + struct page *page, + gfp_t gfp) + + This is called to attempt to read a netfs page from the cache, or to + reserve a backing block if not. FS-Cache will have done as much checking + as it can before calling, but most of the work belongs to the backend. + + If there's no page in the cache, then -ENODATA should be returned if the + backend managed to reserve a backing block; -ENOBUFS or -ENOMEM if it + didn't. + + If there is suitable data in the cache, then a read operation should be + queued and 0 returned. When the read finishes, fscache_end_io() should be + called. + + The fscache_mark_pages_cached() should be called for the page if any cache + metadata is retained. This will indicate to the netfs that the page needs + explicit uncaching. This operation takes a pagevec, thus allowing several + pages to be marked at once. + + The retrieval record pointed to by op should be retained for each page + queued and released when I/O on the page has been formally ended. + fscache_get/put_retrieval() are available for this purpose. + + The retrieval record may be used to get CPU time via the FS-Cache thread + pool. If this is desired, the op->op.processor should be set to point to + the appropriate processing routine, and fscache_enqueue_retrieval() should + be called at an appropriate point to request CPU time. For instance, the + retrieval routine could be enqueued upon the completion of a disk read. + The to_do field in the retrieval record is provided to aid in this. + + If an I/O error occurs, fscache_io_error() should be called and -ENOBUFS + returned if possible or fscache_end_io() called with a suitable error + code.. + + + (*) Request pages be read from cache [mandatory]: + + int (*read_or_alloc_pages)(struct fscache_retrieval *op, + struct list_head *pages, + unsigned *nr_pages, + gfp_t gfp) + + This is like the read_or_alloc_page() method, except it is handed a list + of pages instead of one page. Any pages on which a read operation is + started must be added to the page cache for the specified mapping and also + to the LRU. Such pages must also be removed from the pages list and + *nr_pages decremented per page. + + If there was an error such as -ENOMEM, then that should be returned; else + if one or more pages couldn't be read or allocated, then -ENOBUFS should + be returned; else if one or more pages couldn't be read, then -ENODATA + should be returned. If all the pages are dispatched then 0 should be + returned. + + + (*) Request page be allocated in the cache [mandatory]: + + int (*allocate_page)(struct fscache_retrieval *op, + struct page *page, + gfp_t gfp) + + This is like the read_or_alloc_page() method, except that it shouldn't + read from the cache, even if there's data there that could be retrieved. + It should, however, set up any internal metadata required such that + the write_page() method can write to the cache. + + If there's no backing block available, then -ENOBUFS should be returned + (or -ENOMEM if there were other problems). If a block is successfully + allocated, then the netfs page should be marked and 0 returned. + + + (*) Request pages be allocated in the cache [mandatory]: + + int (*allocate_pages)(struct fscache_retrieval *op, + struct list_head *pages, + unsigned *nr_pages, + gfp_t gfp) + + This is an multiple page version of the allocate_page() method. pages and + nr_pages should be treated as for the read_or_alloc_pages() method. + + + (*) Request page be written to cache [mandatory]: + + int (*write_page)(struct fscache_storage *op, + struct page *page); + + This is called to write from a page on which there was a previously + successful read_or_alloc_page() call or similar. FS-Cache filters out + pages that don't have mappings. + + This method is called asynchronously from the FS-Cache thread pool. It is + not required to actually store anything, provided -ENODATA is then + returned to the next read of this page. + + If an error occurred, then a negative error code should be returned, + otherwise zero should be returned. FS-Cache will take appropriate action + in response to an error, such as withdrawing this object. + + If this method returns success then FS-Cache will inform the netfs + appropriately. + + + (*) Discard retained per-page metadata [mandatory]: + + void (*uncache_page)(struct fscache_object *object, struct page *page) + + This is called when a netfs page is being evicted from the pagecache. The + cache backend should tear down any internal representation or tracking it + maintains for this page. + + +================== +FS-CACHE UTILITIES +================== + +FS-Cache provides some utilities that a cache backend may make use of: + + (*) Note occurrence of an I/O error in a cache: + + void fscache_io_error(struct fscache_cache *cache) + + This tells FS-Cache that an I/O error occurred in the cache. After this + has been called, only resource dissociation operations (object and page + release) will be passed from the netfs to the cache backend for the + specified cache. + + This does not actually withdraw the cache. That must be done separately. + + + (*) Invoke the retrieval I/O completion function: + + void fscache_end_io(struct fscache_retrieval *op, struct page *page, + int error); + + This is called to note the end of an attempt to retrieve a page. The + error value should be 0 if successful and an error otherwise. + + + (*) Set highest store limit: + + void fscache_set_store_limit(struct fscache_object *object, + loff_t i_size); + + This sets the limit FS-Cache imposes on the highest byte it's willing to + try and store for a netfs. Any page over this limit is automatically + rejected by fscache_read_alloc_page() and co with -ENOBUFS. + + + (*) Mark pages as being cached: + + void fscache_mark_pages_cached(struct fscache_retrieval *op, + struct pagevec *pagevec); + + This marks a set of pages as being cached. After this has been called, + the netfs must call fscache_uncache_page() to unmark the pages. + + + (*) Perform coherency check on an object: + + enum fscache_checkaux fscache_check_aux(struct fscache_object *object, + const void *data, + uint16_t datalen); + + This asks the netfs to perform a coherency check on an object that has + just been looked up. The cookie attached to the object will determine the + netfs to use. data and datalen should specify where the auxiliary data + retrieved from the cache can be found. + + One of three values will be returned: + + (*) FSCACHE_CHECKAUX_OKAY + + The coherency data indicates the object is valid as is. + + (*) FSCACHE_CHECKAUX_NEEDS_UPDATE + + The coherency data needs updating, but otherwise the object is + valid. + + (*) FSCACHE_CHECKAUX_OBSOLETE + + The coherency data indicates that the object is obsolete and should + be discarded. + + + (*) Initialise a freshly allocated object: + + void fscache_object_init(struct fscache_object *object); + + This initialises all the fields in an object representation. + + + (*) Indicate the destruction of an object: + + void fscache_object_destroyed(struct fscache_cache *cache); + + This must be called to inform FS-Cache that an object that belonged to a + cache has been destroyed and deallocated. This will allow continuation + of the cache withdrawal process when it is stopped pending destruction of + all the objects. + + + (*) Indicate negative lookup on an object: + + void fscache_object_lookup_negative(struct fscache_object *object); + + This is called to indicate to FS-Cache that a lookup process for an object + found a negative result. + + This changes the state of an object to permit reads pending on lookup + completion to go off and start fetching data from the netfs server as it's + known at this point that there can't be any data in the cache. + + This may be called multiple times on an object. Only the first call is + significant - all subsequent calls are ignored. + + + (*) Indicate an object has been obtained: + + void fscache_obtained_object(struct fscache_object *object); + + This is called to indicate to FS-Cache that a lookup process for an object + produced a positive result, or that an object was created. This should + only be called once for any particular object. + + This changes the state of an object to indicate: + + (1) if no call to fscache_object_lookup_negative() has been made on + this object, that there may be data available, and that reads can + now go and look for it; and + + (2) that writes may now proceed against this object. + + + (*) Indicate that object lookup failed: + + void fscache_object_lookup_error(struct fscache_object *object); + + This marks an object as having encountered a fatal error (usually EIO) + and causes it to move into a state whereby it will be withdrawn as soon + as possible. + + + (*) Get and release references on a retrieval record: + + void fscache_get_retrieval(struct fscache_retrieval *op); + void fscache_put_retrieval(struct fscache_retrieval *op); + + These two functions are used to retain a retrieval record whilst doing + asynchronous data retrieval and block allocation. + + + (*) Enqueue a retrieval record for processing. + + void fscache_enqueue_retrieval(struct fscache_retrieval *op); + + This enqueues a retrieval record for processing by the FS-Cache thread + pool. One of the threads in the pool will invoke the retrieval record's + op->op.processor callback function. This function may be called from + within the callback function. + + + (*) List of object state names: + + const char *fscache_object_states[]; + + For debugging purposes, this may be used to turn the state that an object + is in into a text string for display purposes. diff --git a/Documentation/filesystems/caching/cachefiles.txt b/Documentation/filesystems/caching/cachefiles.txt new file mode 100644 index 00000000..748a1ae4 --- /dev/null +++ b/Documentation/filesystems/caching/cachefiles.txt @@ -0,0 +1,501 @@ + =============================================== + CacheFiles: CACHE ON ALREADY MOUNTED FILESYSTEM + =============================================== + +Contents: + + (*) Overview. + + (*) Requirements. + + (*) Configuration. + + (*) Starting the cache. + + (*) Things to avoid. + + (*) Cache culling. + + (*) Cache structure. + + (*) Security model and SELinux. + + (*) A note on security. + + (*) Statistical information. + + (*) Debugging. + + +======== +OVERVIEW +======== + +CacheFiles is a caching backend that's meant to use as a cache a directory on +an already mounted filesystem of a local type (such as Ext3). + +CacheFiles uses a userspace daemon to do some of the cache management - such as +reaping stale nodes and culling. This is called cachefilesd and lives in +/sbin. + +The filesystem and data integrity of the cache are only as good as those of the +filesystem providing the backing services. Note that CacheFiles does not +attempt to journal anything since the journalling interfaces of the various +filesystems are very specific in nature. + +CacheFiles creates a misc character device - "/dev/cachefiles" - that is used +to communication with the daemon. Only one thing may have this open at once, +and whilst it is open, a cache is at least partially in existence. The daemon +opens this and sends commands down it to control the cache. + +CacheFiles is currently limited to a single cache. + +CacheFiles attempts to maintain at least a certain percentage of free space on +the filesystem, shrinking the cache by culling the objects it contains to make +space if necessary - see the "Cache Culling" section. This means it can be +placed on the same medium as a live set of data, and will expand to make use of +spare space and automatically contract when the set of data requires more +space. + + +============ +REQUIREMENTS +============ + +The use of CacheFiles and its daemon requires the following features to be +available in the system and in the cache filesystem: + + - dnotify. + + - extended attributes (xattrs). + + - openat() and friends. + + - bmap() support on files in the filesystem (FIBMAP ioctl). + + - The use of bmap() to detect a partial page at the end of the file. + +It is strongly recommended that the "dir_index" option is enabled on Ext3 +filesystems being used as a cache. + + +============= +CONFIGURATION +============= + +The cache is configured by a script in /etc/cachefilesd.conf. These commands +set up cache ready for use. The following script commands are available: + + (*) brun <N>% + (*) bcull <N>% + (*) bstop <N>% + (*) frun <N>% + (*) fcull <N>% + (*) fstop <N>% + + Configure the culling limits. Optional. See the section on culling + The defaults are 7% (run), 5% (cull) and 1% (stop) respectively. + + The commands beginning with a 'b' are file space (block) limits, those + beginning with an 'f' are file count limits. + + (*) dir <path> + + Specify the directory containing the root of the cache. Mandatory. + + (*) tag <name> + + Specify a tag to FS-Cache to use in distinguishing multiple caches. + Optional. The default is "CacheFiles". + + (*) debug <mask> + + Specify a numeric bitmask to control debugging in the kernel module. + Optional. The default is zero (all off). The following values can be + OR'd into the mask to collect various information: + + 1 Turn on trace of function entry (_enter() macros) + 2 Turn on trace of function exit (_leave() macros) + 4 Turn on trace of internal debug points (_debug()) + + This mask can also be set through sysfs, eg: + + echo 5 >/sys/modules/cachefiles/parameters/debug + + +================== +STARTING THE CACHE +================== + +The cache is started by running the daemon. The daemon opens the cache device, +configures the cache and tells it to begin caching. At that point the cache +binds to fscache and the cache becomes live. + +The daemon is run as follows: + + /sbin/cachefilesd [-d]* [-s] [-n] [-f <configfile>] + +The flags are: + + (*) -d + + Increase the debugging level. This can be specified multiple times and + is cumulative with itself. + + (*) -s + + Send messages to stderr instead of syslog. + + (*) -n + + Don't daemonise and go into background. + + (*) -f <configfile> + + Use an alternative configuration file rather than the default one. + + +=============== +THINGS TO AVOID +=============== + +Do not mount other things within the cache as this will cause problems. The +kernel module contains its own very cut-down path walking facility that ignores +mountpoints, but the daemon can't avoid them. + +Do not create, rename or unlink files and directories in the cache whilst the +cache is active, as this may cause the state to become uncertain. + +Renaming files in the cache might make objects appear to be other objects (the +filename is part of the lookup key). + +Do not change or remove the extended attributes attached to cache files by the +cache as this will cause the cache state management to get confused. + +Do not create files or directories in the cache, lest the cache get confused or +serve incorrect data. + +Do not chmod files in the cache. The module creates things with minimal +permissions to prevent random users being able to access them directly. + + +============= +CACHE CULLING +============= + +The cache may need culling occasionally to make space. This involves +discarding objects from the cache that have been used less recently than +anything else. Culling is based on the access time of data objects. Empty +directories are culled if not in use. + +Cache culling is done on the basis of the percentage of blocks and the +percentage of files available in the underlying filesystem. There are six +"limits": + + (*) brun + (*) frun + + If the amount of free space and the number of available files in the cache + rises above both these limits, then culling is turned off. + + (*) bcull + (*) fcull + + If the amount of available space or the number of available files in the + cache falls below either of these limits, then culling is started. + + (*) bstop + (*) fstop + + If the amount of available space or the number of available files in the + cache falls below either of these limits, then no further allocation of + disk space or files is permitted until culling has raised things above + these limits again. + +These must be configured thusly: + + 0 <= bstop < bcull < brun < 100 + 0 <= fstop < fcull < frun < 100 + +Note that these are percentages of available space and available files, and do +_not_ appear as 100 minus the percentage displayed by the "df" program. + +The userspace daemon scans the cache to build up a table of cullable objects. +These are then culled in least recently used order. A new scan of the cache is +started as soon as space is made in the table. Objects will be skipped if +their atimes have changed or if the kernel module says it is still using them. + + +=============== +CACHE STRUCTURE +=============== + +The CacheFiles module will create two directories in the directory it was +given: + + (*) cache/ + + (*) graveyard/ + +The active cache objects all reside in the first directory. The CacheFiles +kernel module moves any retired or culled objects that it can't simply unlink +to the graveyard from which the daemon will actually delete them. + +The daemon uses dnotify to monitor the graveyard directory, and will delete +anything that appears therein. + + +The module represents index objects as directories with the filename "I..." or +"J...". Note that the "cache/" directory is itself a special index. + +Data objects are represented as files if they have no children, or directories +if they do. Their filenames all begin "D..." or "E...". If represented as a +directory, data objects will have a file in the directory called "data" that +actually holds the data. + +Special objects are similar to data objects, except their filenames begin +"S..." or "T...". + + +If an object has children, then it will be represented as a directory. +Immediately in the representative directory are a collection of directories +named for hash values of the child object keys with an '@' prepended. Into +this directory, if possible, will be placed the representations of the child +objects: + + INDEX INDEX INDEX DATA FILES + ========= ========== ================================= ================ + cache/@4a/I03nfs/@30/Ji000000000000000--fHg8hi8400 + cache/@4a/I03nfs/@30/Ji000000000000000--fHg8hi8400/@75/Es0g000w...DB1ry + cache/@4a/I03nfs/@30/Ji000000000000000--fHg8hi8400/@75/Es0g000w...N22ry + cache/@4a/I03nfs/@30/Ji000000000000000--fHg8hi8400/@75/Es0g000w...FP1ry + + +If the key is so long that it exceeds NAME_MAX with the decorations added on to +it, then it will be cut into pieces, the first few of which will be used to +make a nest of directories, and the last one of which will be the objects +inside the last directory. The names of the intermediate directories will have +'+' prepended: + + J1223/@23/+xy...z/+kl...m/Epqr + + +Note that keys are raw data, and not only may they exceed NAME_MAX in size, +they may also contain things like '/' and NUL characters, and so they may not +be suitable for turning directly into a filename. + +To handle this, CacheFiles will use a suitably printable filename directly and +"base-64" encode ones that aren't directly suitable. The two versions of +object filenames indicate the encoding: + + OBJECT TYPE PRINTABLE ENCODED + =============== =============== =============== + Index "I..." "J..." + Data "D..." "E..." + Special "S..." "T..." + +Intermediate directories are always "@" or "+" as appropriate. + + +Each object in the cache has an extended attribute label that holds the object +type ID (required to distinguish special objects) and the auxiliary data from +the netfs. The latter is used to detect stale objects in the cache and update +or retire them. + + +Note that CacheFiles will erase from the cache any file it doesn't recognise or +any file of an incorrect type (such as a FIFO file or a device file). + + +========================== +SECURITY MODEL AND SELINUX +========================== + +CacheFiles is implemented to deal properly with the LSM security features of +the Linux kernel and the SELinux facility. + +One of the problems that CacheFiles faces is that it is generally acting on +behalf of a process, and running in that process's context, and that includes a +security context that is not appropriate for accessing the cache - either +because the files in the cache are inaccessible to that process, or because if +the process creates a file in the cache, that file may be inaccessible to other +processes. + +The way CacheFiles works is to temporarily change the security context (fsuid, +fsgid and actor security label) that the process acts as - without changing the +security context of the process when it the target of an operation performed by +some other process (so signalling and suchlike still work correctly). + + +When the CacheFiles module is asked to bind to its cache, it: + + (1) Finds the security label attached to the root cache directory and uses + that as the security label with which it will create files. By default, + this is: + + cachefiles_var_t + + (2) Finds the security label of the process which issued the bind request + (presumed to be the cachefilesd daemon), which by default will be: + + cachefilesd_t + + and asks LSM to supply a security ID as which it should act given the + daemon's label. By default, this will be: + + cachefiles_kernel_t + + SELinux transitions the daemon's security ID to the module's security ID + based on a rule of this form in the policy. + + type_transition <daemon's-ID> kernel_t : process <module's-ID>; + + For instance: + + type_transition cachefilesd_t kernel_t : process cachefiles_kernel_t; + + +The module's security ID gives it permission to create, move and remove files +and directories in the cache, to find and access directories and files in the +cache, to set and access extended attributes on cache objects, and to read and +write files in the cache. + +The daemon's security ID gives it only a very restricted set of permissions: it +may scan directories, stat files and erase files and directories. It may +not read or write files in the cache, and so it is precluded from accessing the +data cached therein; nor is it permitted to create new files in the cache. + + +There are policy source files available in: + + http://people.redhat.com/~dhowells/fscache/cachefilesd-0.8.tar.bz2 + +and later versions. In that tarball, see the files: + + cachefilesd.te + cachefilesd.fc + cachefilesd.if + +They are built and installed directly by the RPM. + +If a non-RPM based system is being used, then copy the above files to their own +directory and run: + + make -f /usr/share/selinux/devel/Makefile + semodule -i cachefilesd.pp + +You will need checkpolicy and selinux-policy-devel installed prior to the +build. + + +By default, the cache is located in /var/fscache, but if it is desirable that +it should be elsewhere, than either the above policy files must be altered, or +an auxiliary policy must be installed to label the alternate location of the +cache. + +For instructions on how to add an auxiliary policy to enable the cache to be +located elsewhere when SELinux is in enforcing mode, please see: + + /usr/share/doc/cachefilesd-*/move-cache.txt + +When the cachefilesd rpm is installed; alternatively, the document can be found +in the sources. + + +================== +A NOTE ON SECURITY +================== + +CacheFiles makes use of the split security in the task_struct. It allocates +its own task_security structure, and redirects current->cred to point to it +when it acts on behalf of another process, in that process's context. + +The reason it does this is that it calls vfs_mkdir() and suchlike rather than +bypassing security and calling inode ops directly. Therefore the VFS and LSM +may deny the CacheFiles access to the cache data because under some +circumstances the caching code is running in the security context of whatever +process issued the original syscall on the netfs. + +Furthermore, should CacheFiles create a file or directory, the security +parameters with that object is created (UID, GID, security label) would be +derived from that process that issued the system call, thus potentially +preventing other processes from accessing the cache - including CacheFiles's +cache management daemon (cachefilesd). + +What is required is to temporarily override the security of the process that +issued the system call. We can't, however, just do an in-place change of the +security data as that affects the process as an object, not just as a subject. +This means it may lose signals or ptrace events for example, and affects what +the process looks like in /proc. + +So CacheFiles makes use of a logical split in the security between the +objective security (task->real_cred) and the subjective security (task->cred). +The objective security holds the intrinsic security properties of a process and +is never overridden. This is what appears in /proc, and is what is used when a +process is the target of an operation by some other process (SIGKILL for +example). + +The subjective security holds the active security properties of a process, and +may be overridden. This is not seen externally, and is used whan a process +acts upon another object, for example SIGKILLing another process or opening a +file. + +LSM hooks exist that allow SELinux (or Smack or whatever) to reject a request +for CacheFiles to run in a context of a specific security label, or to create +files and directories with another security label. + + +======================= +STATISTICAL INFORMATION +======================= + +If FS-Cache is compiled with the following option enabled: + + CONFIG_CACHEFILES_HISTOGRAM=y + +then it will gather certain statistics and display them through a proc file. + + (*) /proc/fs/cachefiles/histogram + + cat /proc/fs/cachefiles/histogram + JIFS SECS LOOKUPS MKDIRS CREATES + ===== ===== ========= ========= ========= + + This shows the breakdown of the number of times each amount of time + between 0 jiffies and HZ-1 jiffies a variety of tasks took to run. The + columns are as follows: + + COLUMN TIME MEASUREMENT + ======= ======================================================= + LOOKUPS Length of time to perform a lookup on the backing fs + MKDIRS Length of time to perform a mkdir on the backing fs + CREATES Length of time to perform a create on the backing fs + + Each row shows the number of events that took a particular range of times. + Each step is 1 jiffy in size. The JIFS column indicates the particular + jiffy range covered, and the SECS field the equivalent number of seconds. + + +========= +DEBUGGING +========= + +If CONFIG_CACHEFILES_DEBUG is enabled, the CacheFiles facility can have runtime +debugging enabled by adjusting the value in: + + /sys/module/cachefiles/parameters/debug + +This is a bitmask of debugging streams to enable: + + BIT VALUE STREAM POINT + ======= ======= =============================== ======================= + 0 1 General Function entry trace + 1 2 Function exit trace + 2 4 General + +The appropriate set of values should be OR'd together and the result written to +the control file. For example: + + echo $((1|4|8)) >/sys/module/cachefiles/parameters/debug + +will turn on all function entry debugging. diff --git a/Documentation/filesystems/caching/fscache.txt b/Documentation/filesystems/caching/fscache.txt new file mode 100644 index 00000000..770267af --- /dev/null +++ b/Documentation/filesystems/caching/fscache.txt @@ -0,0 +1,443 @@ + ========================== + General Filesystem Caching + ========================== + +======== +OVERVIEW +======== + +This facility is a general purpose cache for network filesystems, though it +could be used for caching other things such as ISO9660 filesystems too. + +FS-Cache mediates between cache backends (such as CacheFS) and network +filesystems: + + +---------+ + | | +--------------+ + | NFS |--+ | | + | | | +-->| CacheFS | + +---------+ | +----------+ | | /dev/hda5 | + | | | | +--------------+ + +---------+ +-->| | | + | | | |--+ + | AFS |----->| FS-Cache | + | | | |--+ + +---------+ +-->| | | + | | | | +--------------+ + +---------+ | +----------+ | | | + | | | +-->| CacheFiles | + | ISOFS |--+ | /var/cache | + | | +--------------+ + +---------+ + +Or to look at it another way, FS-Cache is a module that provides a caching +facility to a network filesystem such that the cache is transparent to the +user: + + +---------+ + | | + | Server | + | | + +---------+ + | NETWORK + ~~~~~|~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + | + | +----------+ + V | | + +---------+ | | + | | | | + | NFS |----->| FS-Cache | + | | | |--+ + +---------+ | | | +--------------+ +--------------+ + | | | | | | | | + V +----------+ +-->| CacheFiles |-->| Ext3 | + +---------+ | /var/cache | | /dev/sda6 | + | | +--------------+ +--------------+ + | VFS | ^ ^ + | | | | + +---------+ +--------------+ | + | KERNEL SPACE | | + ~~~~~|~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~|~~~~~~|~~~~ + | USER SPACE | | + V | | + +---------+ +--------------+ + | | | | + | Process | | cachefilesd | + | | | | + +---------+ +--------------+ + + +FS-Cache does not follow the idea of completely loading every netfs file +opened in its entirety into a cache before permitting it to be accessed and +then serving the pages out of that cache rather than the netfs inode because: + + (1) It must be practical to operate without a cache. + + (2) The size of any accessible file must not be limited to the size of the + cache. + + (3) The combined size of all opened files (this includes mapped libraries) + must not be limited to the size of the cache. + + (4) The user should not be forced to download an entire file just to do a + one-off access of a small portion of it (such as might be done with the + "file" program). + +It instead serves the cache out in PAGE_SIZE chunks as and when requested by +the netfs('s) using it. + + +FS-Cache provides the following facilities: + + (1) More than one cache can be used at once. Caches can be selected + explicitly by use of tags. + + (2) Caches can be added / removed at any time. + + (3) The netfs is provided with an interface that allows either party to + withdraw caching facilities from a file (required for (2)). + + (4) The interface to the netfs returns as few errors as possible, preferring + rather to let the netfs remain oblivious. + + (5) Cookies are used to represent indices, files and other objects to the + netfs. The simplest cookie is just a NULL pointer - indicating nothing + cached there. + + (6) The netfs is allowed to propose - dynamically - any index hierarchy it + desires, though it must be aware that the index search function is + recursive, stack space is limited, and indices can only be children of + indices. + + (7) Data I/O is done direct to and from the netfs's pages. The netfs + indicates that page A is at index B of the data-file represented by cookie + C, and that it should be read or written. The cache backend may or may + not start I/O on that page, but if it does, a netfs callback will be + invoked to indicate completion. The I/O may be either synchronous or + asynchronous. + + (8) Cookies can be "retired" upon release. At this point FS-Cache will mark + them as obsolete and the index hierarchy rooted at that point will get + recycled. + + (9) The netfs provides a "match" function for index searches. In addition to + saying whether a match was made or not, this can also specify that an + entry should be updated or deleted. + +(10) As much as possible is done asynchronously. + + +FS-Cache maintains a virtual indexing tree in which all indices, files, objects +and pages are kept. Bits of this tree may actually reside in one or more +caches. + + FSDEF + | + +------------------------------------+ + | | + NFS AFS + | | + +--------------------------+ +-----------+ + | | | | + homedir mirror afs.org redhat.com + | | | + +------------+ +---------------+ +----------+ + | | | | | | + 00001 00002 00007 00125 vol00001 vol00002 + | | | | | + +---+---+ +-----+ +---+ +------+------+ +-----+----+ + | | | | | | | | | | | | | +PG0 PG1 PG2 PG0 XATTR PG0 PG1 DIRENT DIRENT DIRENT R/W R/O Bak + | | + PG0 +-------+ + | | + 00001 00003 + | + +---+---+ + | | | + PG0 PG1 PG2 + +In the example above, you can see two netfs's being backed: NFS and AFS. These +have different index hierarchies: + + (*) The NFS primary index contains per-server indices. Each server index is + indexed by NFS file handles to get data file objects. Each data file + objects can have an array of pages, but may also have further child + objects, such as extended attributes and directory entries. Extended + attribute objects themselves have page-array contents. + + (*) The AFS primary index contains per-cell indices. Each cell index contains + per-logical-volume indices. Each of volume index contains up to three + indices for the read-write, read-only and backup mirrors of those volumes. + Each of these contains vnode data file objects, each of which contains an + array of pages. + +The very top index is the FS-Cache master index in which individual netfs's +have entries. + +Any index object may reside in more than one cache, provided it only has index +children. Any index with non-index object children will be assumed to only +reside in one cache. + + +The netfs API to FS-Cache can be found in: + + Documentation/filesystems/caching/netfs-api.txt + +The cache backend API to FS-Cache can be found in: + + Documentation/filesystems/caching/backend-api.txt + +A description of the internal representations and object state machine can be +found in: + + Documentation/filesystems/caching/object.txt + + +======================= +STATISTICAL INFORMATION +======================= + +If FS-Cache is compiled with the following options enabled: + + CONFIG_FSCACHE_STATS=y + CONFIG_FSCACHE_HISTOGRAM=y + +then it will gather certain statistics and display them through a number of +proc files. + + (*) /proc/fs/fscache/stats + + This shows counts of a number of events that can happen in FS-Cache: + + CLASS EVENT MEANING + ======= ======= ======================================================= + Cookies idx=N Number of index cookies allocated + dat=N Number of data storage cookies allocated + spc=N Number of special cookies allocated + Objects alc=N Number of objects allocated + nal=N Number of object allocation failures + avl=N Number of objects that reached the available state + ded=N Number of objects that reached the dead state + ChkAux non=N Number of objects that didn't have a coherency check + ok=N Number of objects that passed a coherency check + upd=N Number of objects that needed a coherency data update + obs=N Number of objects that were declared obsolete + Pages mrk=N Number of pages marked as being cached + unc=N Number of uncache page requests seen + Acquire n=N Number of acquire cookie requests seen + nul=N Number of acq reqs given a NULL parent + noc=N Number of acq reqs rejected due to no cache available + ok=N Number of acq reqs succeeded + nbf=N Number of acq reqs rejected due to error + oom=N Number of acq reqs failed on ENOMEM + Lookups n=N Number of lookup calls made on cache backends + neg=N Number of negative lookups made + pos=N Number of positive lookups made + crt=N Number of objects created by lookup + tmo=N Number of lookups timed out and requeued + Updates n=N Number of update cookie requests seen + nul=N Number of upd reqs given a NULL parent + run=N Number of upd reqs granted CPU time + Relinqs n=N Number of relinquish cookie requests seen + nul=N Number of rlq reqs given a NULL parent + wcr=N Number of rlq reqs waited on completion of creation + AttrChg n=N Number of attribute changed requests seen + ok=N Number of attr changed requests queued + nbf=N Number of attr changed rejected -ENOBUFS + oom=N Number of attr changed failed -ENOMEM + run=N Number of attr changed ops given CPU time + Allocs n=N Number of allocation requests seen + ok=N Number of successful alloc reqs + wt=N Number of alloc reqs that waited on lookup completion + nbf=N Number of alloc reqs rejected -ENOBUFS + int=N Number of alloc reqs aborted -ERESTARTSYS + ops=N Number of alloc reqs submitted + owt=N Number of alloc reqs waited for CPU time + abt=N Number of alloc reqs aborted due to object death + Retrvls n=N Number of retrieval (read) requests seen + ok=N Number of successful retr reqs + wt=N Number of retr reqs that waited on lookup completion + nod=N Number of retr reqs returned -ENODATA + nbf=N Number of retr reqs rejected -ENOBUFS + int=N Number of retr reqs aborted -ERESTARTSYS + oom=N Number of retr reqs failed -ENOMEM + ops=N Number of retr reqs submitted + owt=N Number of retr reqs waited for CPU time + abt=N Number of retr reqs aborted due to object death + Stores n=N Number of storage (write) requests seen + ok=N Number of successful store reqs + agn=N Number of store reqs on a page already pending storage + nbf=N Number of store reqs rejected -ENOBUFS + oom=N Number of store reqs failed -ENOMEM + ops=N Number of store reqs submitted + run=N Number of store reqs granted CPU time + pgs=N Number of pages given store req processing time + rxd=N Number of store reqs deleted from tracking tree + olm=N Number of store reqs over store limit + VmScan nos=N Number of release reqs against pages with no pending store + gon=N Number of release reqs against pages stored by time lock granted + bsy=N Number of release reqs ignored due to in-progress store + can=N Number of page stores cancelled due to release req + Ops pend=N Number of times async ops added to pending queues + run=N Number of times async ops given CPU time + enq=N Number of times async ops queued for processing + can=N Number of async ops cancelled + rej=N Number of async ops rejected due to object lookup/create failure + dfr=N Number of async ops queued for deferred release + rel=N Number of async ops released + gc=N Number of deferred-release async ops garbage collected + CacheOp alo=N Number of in-progress alloc_object() cache ops + luo=N Number of in-progress lookup_object() cache ops + luc=N Number of in-progress lookup_complete() cache ops + gro=N Number of in-progress grab_object() cache ops + upo=N Number of in-progress update_object() cache ops + dro=N Number of in-progress drop_object() cache ops + pto=N Number of in-progress put_object() cache ops + syn=N Number of in-progress sync_cache() cache ops + atc=N Number of in-progress attr_changed() cache ops + rap=N Number of in-progress read_or_alloc_page() cache ops + ras=N Number of in-progress read_or_alloc_pages() cache ops + alp=N Number of in-progress allocate_page() cache ops + als=N Number of in-progress allocate_pages() cache ops + wrp=N Number of in-progress write_page() cache ops + ucp=N Number of in-progress uncache_page() cache ops + dsp=N Number of in-progress dissociate_pages() cache ops + + + (*) /proc/fs/fscache/histogram + + cat /proc/fs/fscache/histogram + JIFS SECS OBJ INST OP RUNS OBJ RUNS RETRV DLY RETRIEVLS + ===== ===== ========= ========= ========= ========= ========= + + This shows the breakdown of the number of times each amount of time + between 0 jiffies and HZ-1 jiffies a variety of tasks took to run. The + columns are as follows: + + COLUMN TIME MEASUREMENT + ======= ======================================================= + OBJ INST Length of time to instantiate an object + OP RUNS Length of time a call to process an operation took + OBJ RUNS Length of time a call to process an object event took + RETRV DLY Time between an requesting a read and lookup completing + RETRIEVLS Time between beginning and end of a retrieval + + Each row shows the number of events that took a particular range of times. + Each step is 1 jiffy in size. The JIFS column indicates the particular + jiffy range covered, and the SECS field the equivalent number of seconds. + + +=========== +OBJECT LIST +=========== + +If CONFIG_FSCACHE_OBJECT_LIST is enabled, the FS-Cache facility will maintain a +list of all the objects currently allocated and allow them to be viewed +through: + + /proc/fs/fscache/objects + +This will look something like: + + [root@andromeda ~]# head /proc/fs/fscache/objects + OBJECT PARENT STAT CHLDN OPS OOP IPR EX READS EM EV F S | NETFS_COOKIE_DEF TY FL NETFS_DATA OBJECT_KEY, AUX_DATA + ======== ======== ==== ===== === === === == ===== == == = = | ================ == == ================ ================ + 17e4b 2 ACTV 0 0 0 0 0 0 7b 4 0 0 | NFS.fh DT 0 ffff88001dd82820 010006017edcf8bbc93b43298fdfbe71e50b57b13a172c0117f38472, e567634700000000000000000000000063f2404a000000000000000000000000c9030000000000000000000063f2404a + 1693a 2 ACTV 0 0 0 0 0 0 7b 4 0 0 | NFS.fh DT 0 ffff88002db23380 010006017edcf8bbc93b43298fdfbe71e50b57b1e0162c01a2df0ea6, 420ebc4a000000000000000000000000420ebc4a0000000000000000000000000e1801000000000000000000420ebc4a + +where the first set of columns before the '|' describe the object: + + COLUMN DESCRIPTION + ======= =============================================================== + OBJECT Object debugging ID (appears as OBJ%x in some debug messages) + PARENT Debugging ID of parent object + STAT Object state + CHLDN Number of child objects of this object + OPS Number of outstanding operations on this object + OOP Number of outstanding child object management operations + IPR + EX Number of outstanding exclusive operations + READS Number of outstanding read operations + EM Object's event mask + EV Events raised on this object + F Object flags + S Object work item busy state mask (1:pending 2:running) + +and the second set of columns describe the object's cookie, if present: + + COLUMN DESCRIPTION + =============== ======================================================= + NETFS_COOKIE_DEF Name of netfs cookie definition + TY Cookie type (IX - index, DT - data, hex - special) + FL Cookie flags + NETFS_DATA Netfs private data stored in the cookie + OBJECT_KEY Object key } 1 column, with separating comma + AUX_DATA Object aux data } presence may be configured + +The data shown may be filtered by attaching the a key to an appropriate keyring +before viewing the file. Something like: + + keyctl add user fscache:objlist <restrictions> @s + +where <restrictions> are a selection of the following letters: + + K Show hexdump of object key (don't show if not given) + A Show hexdump of object aux data (don't show if not given) + +and the following paired letters: + + C Show objects that have a cookie + c Show objects that don't have a cookie + B Show objects that are busy + b Show objects that aren't busy + W Show objects that have pending writes + w Show objects that don't have pending writes + R Show objects that have outstanding reads + r Show objects that don't have outstanding reads + S Show objects that have work queued + s Show objects that don't have work queued + +If neither side of a letter pair is given, then both are implied. For example: + + keyctl add user fscache:objlist KB @s + +shows objects that are busy, and lists their object keys, but does not dump +their auxiliary data. It also implies "CcWwRrSs", but as 'B' is given, 'b' is +not implied. + +By default all objects and all fields will be shown. + + +========= +DEBUGGING +========= + +If CONFIG_FSCACHE_DEBUG is enabled, the FS-Cache facility can have runtime +debugging enabled by adjusting the value in: + + /sys/module/fscache/parameters/debug + +This is a bitmask of debugging streams to enable: + + BIT VALUE STREAM POINT + ======= ======= =============================== ======================= + 0 1 Cache management Function entry trace + 1 2 Function exit trace + 2 4 General + 3 8 Cookie management Function entry trace + 4 16 Function exit trace + 5 32 General + 6 64 Page handling Function entry trace + 7 128 Function exit trace + 8 256 General + 9 512 Operation management Function entry trace + 10 1024 Function exit trace + 11 2048 General + +The appropriate set of values should be OR'd together and the result written to +the control file. For example: + + echo $((1|8|64)) >/sys/module/fscache/parameters/debug + +will turn on all function entry debugging. diff --git a/Documentation/filesystems/caching/netfs-api.txt b/Documentation/filesystems/caching/netfs-api.txt new file mode 100644 index 00000000..7cc6bf28 --- /dev/null +++ b/Documentation/filesystems/caching/netfs-api.txt @@ -0,0 +1,813 @@ + =============================== + FS-CACHE NETWORK FILESYSTEM API + =============================== + +There's an API by which a network filesystem can make use of the FS-Cache +facilities. This is based around a number of principles: + + (1) Caches can store a number of different object types. There are two main + object types: indices and files. The first is a special type used by + FS-Cache to make finding objects faster and to make retiring of groups of + objects easier. + + (2) Every index, file or other object is represented by a cookie. This cookie + may or may not have anything associated with it, but the netfs doesn't + need to care. + + (3) Barring the top-level index (one entry per cached netfs), the index + hierarchy for each netfs is structured according the whim of the netfs. + +This API is declared in <linux/fscache.h>. + +This document contains the following sections: + + (1) Network filesystem definition + (2) Index definition + (3) Object definition + (4) Network filesystem (un)registration + (5) Cache tag lookup + (6) Index registration + (7) Data file registration + (8) Miscellaneous object registration + (9) Setting the data file size + (10) Page alloc/read/write + (11) Page uncaching + (12) Index and data file update + (13) Miscellaneous cookie operations + (14) Cookie unregistration + (15) Index and data file invalidation + (16) FS-Cache specific page flags. + + +============================= +NETWORK FILESYSTEM DEFINITION +============================= + +FS-Cache needs a description of the network filesystem. This is specified +using a record of the following structure: + + struct fscache_netfs { + uint32_t version; + const char *name; + struct fscache_cookie *primary_index; + ... + }; + +This first two fields should be filled in before registration, and the third +will be filled in by the registration function; any other fields should just be +ignored and are for internal use only. + +The fields are: + + (1) The name of the netfs (used as the key in the toplevel index). + + (2) The version of the netfs (if the name matches but the version doesn't, the + entire in-cache hierarchy for this netfs will be scrapped and begun + afresh). + + (3) The cookie representing the primary index will be allocated according to + another parameter passed into the registration function. + +For example, kAFS (linux/fs/afs/) uses the following definitions to describe +itself: + + struct fscache_netfs afs_cache_netfs = { + .version = 0, + .name = "afs", + }; + + +================ +INDEX DEFINITION +================ + +Indices are used for two purposes: + + (1) To aid the finding of a file based on a series of keys (such as AFS's + "cell", "volume ID", "vnode ID"). + + (2) To make it easier to discard a subset of all the files cached based around + a particular key - for instance to mirror the removal of an AFS volume. + +However, since it's unlikely that any two netfs's are going to want to define +their index hierarchies in quite the same way, FS-Cache tries to impose as few +restraints as possible on how an index is structured and where it is placed in +the tree. The netfs can even mix indices and data files at the same level, but +it's not recommended. + +Each index entry consists of a key of indeterminate length plus some auxiliary +data, also of indeterminate length. + +There are some limits on indices: + + (1) Any index containing non-index objects should be restricted to a single + cache. Any such objects created within an index will be created in the + first cache only. The cache in which an index is created can be + controlled by cache tags (see below). + + (2) The entry data must be atomically journallable, so it is limited to about + 400 bytes at present. At least 400 bytes will be available. + + (3) The depth of the index tree should be judged with care as the search + function is recursive. Too many layers will run the kernel out of stack. + + +================= +OBJECT DEFINITION +================= + +To define an object, a structure of the following type should be filled out: + + struct fscache_cookie_def + { + uint8_t name[16]; + uint8_t type; + + struct fscache_cache_tag *(*select_cache)( + const void *parent_netfs_data, + const void *cookie_netfs_data); + + uint16_t (*get_key)(const void *cookie_netfs_data, + void *buffer, + uint16_t bufmax); + + void (*get_attr)(const void *cookie_netfs_data, + uint64_t *size); + + uint16_t (*get_aux)(const void *cookie_netfs_data, + void *buffer, + uint16_t bufmax); + + enum fscache_checkaux (*check_aux)(void *cookie_netfs_data, + const void *data, + uint16_t datalen); + + void (*get_context)(void *cookie_netfs_data, void *context); + + void (*put_context)(void *cookie_netfs_data, void *context); + + void (*mark_pages_cached)(void *cookie_netfs_data, + struct address_space *mapping, + struct pagevec *cached_pvec); + + void (*now_uncached)(void *cookie_netfs_data); + }; + +This has the following fields: + + (1) The type of the object [mandatory]. + + This is one of the following values: + + (*) FSCACHE_COOKIE_TYPE_INDEX + + This defines an index, which is a special FS-Cache type. + + (*) FSCACHE_COOKIE_TYPE_DATAFILE + + This defines an ordinary data file. + + (*) Any other value between 2 and 255 + + This defines an extraordinary object such as an XATTR. + + (2) The name of the object type (NUL terminated unless all 16 chars are used) + [optional]. + + (3) A function to select the cache in which to store an index [optional]. + + This function is invoked when an index needs to be instantiated in a cache + during the instantiation of a non-index object. Only the immediate index + parent for the non-index object will be queried. Any indices above that + in the hierarchy may be stored in multiple caches. This function does not + need to be supplied for any non-index object or any index that will only + have index children. + + If this function is not supplied or if it returns NULL then the first + cache in the parent's list will be chosen, or failing that, the first + cache in the master list. + + (4) A function to retrieve an object's key from the netfs [mandatory]. + + This function will be called with the netfs data that was passed to the + cookie acquisition function and the maximum length of key data that it may + provide. It should write the required key data into the given buffer and + return the quantity it wrote. + + (5) A function to retrieve attribute data from the netfs [optional]. + + This function will be called with the netfs data that was passed to the + cookie acquisition function. It should return the size of the file if + this is a data file. The size may be used to govern how much cache must + be reserved for this file in the cache. + + If the function is absent, a file size of 0 is assumed. + + (6) A function to retrieve auxiliary data from the netfs [optional]. + + This function will be called with the netfs data that was passed to the + cookie acquisition function and the maximum length of auxiliary data that + it may provide. It should write the auxiliary data into the given buffer + and return the quantity it wrote. + + If this function is absent, the auxiliary data length will be set to 0. + + The length of the auxiliary data buffer may be dependent on the key + length. A netfs mustn't rely on being able to provide more than 400 bytes + for both. + + (7) A function to check the auxiliary data [optional]. + + This function will be called to check that a match found in the cache for + this object is valid. For instance with AFS it could check the auxiliary + data against the data version number returned by the server to determine + whether the index entry in a cache is still valid. + + If this function is absent, it will be assumed that matching objects in a + cache are always valid. + + If present, the function should return one of the following values: + + (*) FSCACHE_CHECKAUX_OKAY - the entry is okay as is + (*) FSCACHE_CHECKAUX_NEEDS_UPDATE - the entry requires update + (*) FSCACHE_CHECKAUX_OBSOLETE - the entry should be deleted + + This function can also be used to extract data from the auxiliary data in + the cache and copy it into the netfs's structures. + + (8) A pair of functions to manage contexts for the completion callback + [optional]. + + The cache read/write functions are passed a context which is then passed + to the I/O completion callback function. To ensure this context remains + valid until after the I/O completion is called, two functions may be + provided: one to get an extra reference on the context, and one to drop a + reference to it. + + If the context is not used or is a type of object that won't go out of + scope, then these functions are not required. These functions are not + required for indices as indices may not contain data. These functions may + be called in interrupt context and so may not sleep. + + (9) A function to mark a page as retaining cache metadata [optional]. + + This is called by the cache to indicate that it is retaining in-memory + information for this page and that the netfs should uncache the page when + it has finished. This does not indicate whether there's data on the disk + or not. Note that several pages at once may be presented for marking. + + The PG_fscache bit is set on the pages before this function would be + called, so the function need not be provided if this is sufficient. + + This function is not required for indices as they're not permitted data. + +(10) A function to unmark all the pages retaining cache metadata [mandatory]. + + This is called by FS-Cache to indicate that a backing store is being + unbound from a cookie and that all the marks on the pages should be + cleared to prevent confusion. Note that the cache will have torn down all + its tracking information so that the pages don't need to be explicitly + uncached. + + This function is not required for indices as they're not permitted data. + + +=================================== +NETWORK FILESYSTEM (UN)REGISTRATION +=================================== + +The first step is to declare the network filesystem to the cache. This also +involves specifying the layout of the primary index (for AFS, this would be the +"cell" level). + +The registration function is: + + int fscache_register_netfs(struct fscache_netfs *netfs); + +It just takes a pointer to the netfs definition. It returns 0 or an error as +appropriate. + +For kAFS, registration is done as follows: + + ret = fscache_register_netfs(&afs_cache_netfs); + +The last step is, of course, unregistration: + + void fscache_unregister_netfs(struct fscache_netfs *netfs); + + +================ +CACHE TAG LOOKUP +================ + +FS-Cache permits the use of more than one cache. To permit particular index +subtrees to be bound to particular caches, the second step is to look up cache +representation tags. This step is optional; it can be left entirely up to +FS-Cache as to which cache should be used. The problem with doing that is that +FS-Cache will always pick the first cache that was registered. + +To get the representation for a named tag: + + struct fscache_cache_tag *fscache_lookup_cache_tag(const char *name); + +This takes a text string as the name and returns a representation of a tag. It +will never return an error. It may return a dummy tag, however, if it runs out +of memory; this will inhibit caching with this tag. + +Any representation so obtained must be released by passing it to this function: + + void fscache_release_cache_tag(struct fscache_cache_tag *tag); + +The tag will be retrieved by FS-Cache when it calls the object definition +operation select_cache(). + + +================== +INDEX REGISTRATION +================== + +The third step is to inform FS-Cache about part of an index hierarchy that can +be used to locate files. This is done by requesting a cookie for each index in +the path to the file: + + struct fscache_cookie * + fscache_acquire_cookie(struct fscache_cookie *parent, + const struct fscache_object_def *def, + void *netfs_data); + +This function creates an index entry in the index represented by parent, +filling in the index entry by calling the operations pointed to by def. + +Note that this function never returns an error - all errors are handled +internally. It may, however, return NULL to indicate no cookie. It is quite +acceptable to pass this token back to this function as the parent to another +acquisition (or even to the relinquish cookie, read page and write page +functions - see below). + +Note also that no indices are actually created in a cache until a non-index +object needs to be created somewhere down the hierarchy. Furthermore, an index +may be created in several different caches independently at different times. +This is all handled transparently, and the netfs doesn't see any of it. + +For example, with AFS, a cell would be added to the primary index. This index +entry would have a dependent inode containing a volume location index for the +volume mappings within this cell: + + cell->cache = + fscache_acquire_cookie(afs_cache_netfs.primary_index, + &afs_cell_cache_index_def, + cell); + +Then when a volume location was accessed, it would be entered into the cell's +index and an inode would be allocated that acts as a volume type and hash chain +combination: + + vlocation->cache = + fscache_acquire_cookie(cell->cache, + &afs_vlocation_cache_index_def, + vlocation); + +And then a particular flavour of volume (R/O for example) could be added to +that index, creating another index for vnodes (AFS inode equivalents): + + volume->cache = + fscache_acquire_cookie(vlocation->cache, + &afs_volume_cache_index_def, + volume); + + +====================== +DATA FILE REGISTRATION +====================== + +The fourth step is to request a data file be created in the cache. This is +identical to index cookie acquisition. The only difference is that the type in +the object definition should be something other than index type. + + vnode->cache = + fscache_acquire_cookie(volume->cache, + &afs_vnode_cache_object_def, + vnode); + + +================================= +MISCELLANEOUS OBJECT REGISTRATION +================================= + +An optional step is to request an object of miscellaneous type be created in +the cache. This is almost identical to index cookie acquisition. The only +difference is that the type in the object definition should be something other +than index type. Whilst the parent object could be an index, it's more likely +it would be some other type of object such as a data file. + + xattr->cache = + fscache_acquire_cookie(vnode->cache, + &afs_xattr_cache_object_def, + xattr); + +Miscellaneous objects might be used to store extended attributes or directory +entries for example. + + +========================== +SETTING THE DATA FILE SIZE +========================== + +The fifth step is to set the physical attributes of the file, such as its size. +This doesn't automatically reserve any space in the cache, but permits the +cache to adjust its metadata for data tracking appropriately: + + int fscache_attr_changed(struct fscache_cookie *cookie); + +The cache will return -ENOBUFS if there is no backing cache or if there is no +space to allocate any extra metadata required in the cache. The attributes +will be accessed with the get_attr() cookie definition operation. + +Note that attempts to read or write data pages in the cache over this size may +be rebuffed with -ENOBUFS. + +This operation schedules an attribute adjustment to happen asynchronously at +some point in the future, and as such, it may happen after the function returns +to the caller. The attribute adjustment excludes read and write operations. + + +===================== +PAGE READ/ALLOC/WRITE +===================== + +And the sixth step is to store and retrieve pages in the cache. There are +three functions that are used to do this. + +Note: + + (1) A page should not be re-read or re-allocated without uncaching it first. + + (2) A read or allocated page must be uncached when the netfs page is released + from the pagecache. + + (3) A page should only be written to the cache if previous read or allocated. + +This permits the cache to maintain its page tracking in proper order. + + +PAGE READ +--------- + +Firstly, the netfs should ask FS-Cache to examine the caches and read the +contents cached for a particular page of a particular file if present, or else +allocate space to store the contents if not: + + typedef + void (*fscache_rw_complete_t)(struct page *page, + void *context, + int error); + + int fscache_read_or_alloc_page(struct fscache_cookie *cookie, + struct page *page, + fscache_rw_complete_t end_io_func, + void *context, + gfp_t gfp); + +The cookie argument must specify a cookie for an object that isn't an index, +the page specified will have the data loaded into it (and is also used to +specify the page number), and the gfp argument is used to control how any +memory allocations made are satisfied. + +If the cookie indicates the inode is not cached: + + (1) The function will return -ENOBUFS. + +Else if there's a copy of the page resident in the cache: + + (1) The mark_pages_cached() cookie operation will be called on that page. + + (2) The function will submit a request to read the data from the cache's + backing device directly into the page specified. + + (3) The function will return 0. + + (4) When the read is complete, end_io_func() will be invoked with: + + (*) The netfs data supplied when the cookie was created. + + (*) The page descriptor. + + (*) The context argument passed to the above function. This will be + maintained with the get_context/put_context functions mentioned above. + + (*) An argument that's 0 on success or negative for an error code. + + If an error occurs, it should be assumed that the page contains no usable + data. + + end_io_func() will be called in process context if the read is results in + an error, but it might be called in interrupt context if the read is + successful. + +Otherwise, if there's not a copy available in cache, but the cache may be able +to store the page: + + (1) The mark_pages_cached() cookie operation will be called on that page. + + (2) A block may be reserved in the cache and attached to the object at the + appropriate place. + + (3) The function will return -ENODATA. + +This function may also return -ENOMEM or -EINTR, in which case it won't have +read any data from the cache. + + +PAGE ALLOCATE +------------- + +Alternatively, if there's not expected to be any data in the cache for a page +because the file has been extended, a block can simply be allocated instead: + + int fscache_alloc_page(struct fscache_cookie *cookie, + struct page *page, + gfp_t gfp); + +This is similar to the fscache_read_or_alloc_page() function, except that it +never reads from the cache. It will return 0 if a block has been allocated, +rather than -ENODATA as the other would. One or the other must be performed +before writing to the cache. + +The mark_pages_cached() cookie operation will be called on the page if +successful. + + +PAGE WRITE +---------- + +Secondly, if the netfs changes the contents of the page (either due to an +initial download or if a user performs a write), then the page should be +written back to the cache: + + int fscache_write_page(struct fscache_cookie *cookie, + struct page *page, + gfp_t gfp); + +The cookie argument must specify a data file cookie, the page specified should +contain the data to be written (and is also used to specify the page number), +and the gfp argument is used to control how any memory allocations made are +satisfied. + +The page must have first been read or allocated successfully and must not have +been uncached before writing is performed. + +If the cookie indicates the inode is not cached then: + + (1) The function will return -ENOBUFS. + +Else if space can be allocated in the cache to hold this page: + + (1) PG_fscache_write will be set on the page. + + (2) The function will submit a request to write the data to cache's backing + device directly from the page specified. + + (3) The function will return 0. + + (4) When the write is complete PG_fscache_write is cleared on the page and + anyone waiting for that bit will be woken up. + +Else if there's no space available in the cache, -ENOBUFS will be returned. It +is also possible for the PG_fscache_write bit to be cleared when no write took +place if unforeseen circumstances arose (such as a disk error). + +Writing takes place asynchronously. + + +MULTIPLE PAGE READ +------------------ + +A facility is provided to read several pages at once, as requested by the +readpages() address space operation: + + int fscache_read_or_alloc_pages(struct fscache_cookie *cookie, + struct address_space *mapping, + struct list_head *pages, + int *nr_pages, + fscache_rw_complete_t end_io_func, + void *context, + gfp_t gfp); + +This works in a similar way to fscache_read_or_alloc_page(), except: + + (1) Any page it can retrieve data for is removed from pages and nr_pages and + dispatched for reading to the disk. Reads of adjacent pages on disk may + be merged for greater efficiency. + + (2) The mark_pages_cached() cookie operation will be called on several pages + at once if they're being read or allocated. + + (3) If there was an general error, then that error will be returned. + + Else if some pages couldn't be allocated or read, then -ENOBUFS will be + returned. + + Else if some pages couldn't be read but were allocated, then -ENODATA will + be returned. + + Otherwise, if all pages had reads dispatched, then 0 will be returned, the + list will be empty and *nr_pages will be 0. + + (4) end_io_func will be called once for each page being read as the reads + complete. It will be called in process context if error != 0, but it may + be called in interrupt context if there is no error. + +Note that a return of -ENODATA, -ENOBUFS or any other error does not preclude +some of the pages being read and some being allocated. Those pages will have +been marked appropriately and will need uncaching. + + +============== +PAGE UNCACHING +============== + +To uncache a page, this function should be called: + + void fscache_uncache_page(struct fscache_cookie *cookie, + struct page *page); + +This function permits the cache to release any in-memory representation it +might be holding for this netfs page. This function must be called once for +each page on which the read or write page functions above have been called to +make sure the cache's in-memory tracking information gets torn down. + +Note that pages can't be explicitly deleted from the a data file. The whole +data file must be retired (see the relinquish cookie function below). + +Furthermore, note that this does not cancel the asynchronous read or write +operation started by the read/alloc and write functions, so the page +invalidation functions must use: + + bool fscache_check_page_write(struct fscache_cookie *cookie, + struct page *page); + +to see if a page is being written to the cache, and: + + void fscache_wait_on_page_write(struct fscache_cookie *cookie, + struct page *page); + +to wait for it to finish if it is. + + +When releasepage() is being implemented, a special FS-Cache function exists to +manage the heuristics of coping with vmscan trying to eject pages, which may +conflict with the cache trying to write pages to the cache (which may itself +need to allocate memory): + + bool fscache_maybe_release_page(struct fscache_cookie *cookie, + struct page *page, + gfp_t gfp); + +This takes the netfs cookie, and the page and gfp arguments as supplied to +releasepage(). It will return false if the page cannot be released yet for +some reason and if it returns true, the page has been uncached and can now be +released. + +To make a page available for release, this function may wait for an outstanding +storage request to complete, or it may attempt to cancel the storage request - +in which case the page will not be stored in the cache this time. + + +BULK INODE PAGE UNCACHE +----------------------- + +A convenience routine is provided to perform an uncache on all the pages +attached to an inode. This assumes that the pages on the inode correspond on a +1:1 basis with the pages in the cache. + + void fscache_uncache_all_inode_pages(struct fscache_cookie *cookie, + struct inode *inode); + +This takes the netfs cookie that the pages were cached with and the inode that +the pages are attached to. This function will wait for pages to finish being +written to the cache and for the cache to finish with the page generally. No +error is returned. + + +========================== +INDEX AND DATA FILE UPDATE +========================== + +To request an update of the index data for an index or other object, the +following function should be called: + + void fscache_update_cookie(struct fscache_cookie *cookie); + +This function will refer back to the netfs_data pointer stored in the cookie by +the acquisition function to obtain the data to write into each revised index +entry. The update method in the parent index definition will be called to +transfer the data. + +Note that partial updates may happen automatically at other times, such as when +data blocks are added to a data file object. + + +=============================== +MISCELLANEOUS COOKIE OPERATIONS +=============================== + +There are a number of operations that can be used to control cookies: + + (*) Cookie pinning: + + int fscache_pin_cookie(struct fscache_cookie *cookie); + void fscache_unpin_cookie(struct fscache_cookie *cookie); + + These operations permit data cookies to be pinned into the cache and to + have the pinning removed. They are not permitted on index cookies. + + The pinning function will return 0 if successful, -ENOBUFS in the cookie + isn't backed by a cache, -EOPNOTSUPP if the cache doesn't support pinning, + -ENOSPC if there isn't enough space to honour the operation, -ENOMEM or + -EIO if there's any other problem. + + (*) Data space reservation: + + int fscache_reserve_space(struct fscache_cookie *cookie, loff_t size); + + This permits a netfs to request cache space be reserved to store up to the + given amount of a file. It is permitted to ask for more than the current + size of the file to allow for future file expansion. + + If size is given as zero then the reservation will be cancelled. + + The function will return 0 if successful, -ENOBUFS in the cookie isn't + backed by a cache, -EOPNOTSUPP if the cache doesn't support reservations, + -ENOSPC if there isn't enough space to honour the operation, -ENOMEM or + -EIO if there's any other problem. + + Note that this doesn't pin an object in a cache; it can still be culled to + make space if it's not in use. + + +===================== +COOKIE UNREGISTRATION +===================== + +To get rid of a cookie, this function should be called. + + void fscache_relinquish_cookie(struct fscache_cookie *cookie, + int retire); + +If retire is non-zero, then the object will be marked for recycling, and all +copies of it will be removed from all active caches in which it is present. +Not only that but all child objects will also be retired. + +If retire is zero, then the object may be available again when next the +acquisition function is called. Retirement here will overrule the pinning on a +cookie. + +One very important note - relinquish must NOT be called for a cookie unless all +the cookies for "child" indices, objects and pages have been relinquished +first. + + +================================ +INDEX AND DATA FILE INVALIDATION +================================ + +There is no direct way to invalidate an index subtree or a data file. To do +this, the caller should relinquish and retire the cookie they have, and then +acquire a new one. + + +=========================== +FS-CACHE SPECIFIC PAGE FLAG +=========================== + +FS-Cache makes use of a page flag, PG_private_2, for its own purpose. This is +given the alternative name PG_fscache. + +PG_fscache is used to indicate that the page is known by the cache, and that +the cache must be informed if the page is going to go away. It's an indication +to the netfs that the cache has an interest in this page, where an interest may +be a pointer to it, resources allocated or reserved for it, or I/O in progress +upon it. + +The netfs can use this information in methods such as releasepage() to +determine whether it needs to uncache a page or update it. + +Furthermore, if this bit is set, releasepage() and invalidatepage() operations +will be called on a page to get rid of it, even if PG_private is not set. This +allows caching to attempted on a page before read_cache_pages() to be called +after fscache_read_or_alloc_pages() as the former will try and release pages it +was given under certain circumstances. + +This bit does not overlap with such as PG_private. This means that FS-Cache +can be used with a filesystem that uses the block buffering code. + +There are a number of operations defined on this flag: + + int PageFsCache(struct page *page); + void SetPageFsCache(struct page *page) + void ClearPageFsCache(struct page *page) + int TestSetPageFsCache(struct page *page) + int TestClearPageFsCache(struct page *page) + +These functions are bit test, bit set, bit clear, bit test and set and bit +test and clear operations on PG_fscache. diff --git a/Documentation/filesystems/caching/object.txt b/Documentation/filesystems/caching/object.txt new file mode 100644 index 00000000..e8b0a35d --- /dev/null +++ b/Documentation/filesystems/caching/object.txt @@ -0,0 +1,313 @@ + ==================================================== + IN-KERNEL CACHE OBJECT REPRESENTATION AND MANAGEMENT + ==================================================== + +By: David Howells <dhowells@redhat.com> + +Contents: + + (*) Representation + + (*) Object management state machine. + + - Provision of cpu time. + - Locking simplification. + + (*) The set of states. + + (*) The set of events. + + +============== +REPRESENTATION +============== + +FS-Cache maintains an in-kernel representation of each object that a netfs is +currently interested in. Such objects are represented by the fscache_cookie +struct and are referred to as cookies. + +FS-Cache also maintains a separate in-kernel representation of the objects that +a cache backend is currently actively caching. Such objects are represented by +the fscache_object struct. The cache backends allocate these upon request, and +are expected to embed them in their own representations. These are referred to +as objects. + +There is a 1:N relationship between cookies and objects. A cookie may be +represented by multiple objects - an index may exist in more than one cache - +or even by no objects (it may not be cached). + +Furthermore, both cookies and objects are hierarchical. The two hierarchies +correspond, but the cookies tree is a superset of the union of the object trees +of multiple caches: + + NETFS INDEX TREE : CACHE 1 : CACHE 2 + : : + : +-----------+ : + +----------->| IObject | : + +-----------+ | : +-----------+ : + | ICookie |-------+ : | : + +-----------+ | : | : +-----------+ + | +------------------------------>| IObject | + | : | : +-----------+ + | : V : | + | : +-----------+ : | + V +----------->| IObject | : | + +-----------+ | : +-----------+ : | + | ICookie |-------+ : | : V + +-----------+ | : | : +-----------+ + | +------------------------------>| IObject | + +-----+-----+ : | : +-----------+ + | | : | : | + V | : V : | + +-----------+ | : +-----------+ : | + | ICookie |------------------------->| IObject | : | + +-----------+ | : +-----------+ : | + | V : | : V + | +-----------+ : | : +-----------+ + | | ICookie |-------------------------------->| IObject | + | +-----------+ : | : +-----------+ + V | : V : | + +-----------+ | : +-----------+ : | + | DCookie |------------------------->| DObject | : | + +-----------+ | : +-----------+ : | + | : : | + +-------+-------+ : : | + | | : : | + V V : : V + +-----------+ +-----------+ : : +-----------+ + | DCookie | | DCookie |------------------------>| DObject | + +-----------+ +-----------+ : : +-----------+ + : : + +In the above illustration, ICookie and IObject represent indices and DCookie +and DObject represent data storage objects. Indices may have representation in +multiple caches, but currently, non-index objects may not. Objects of any type +may also be entirely unrepresented. + +As far as the netfs API goes, the netfs is only actually permitted to see +pointers to the cookies. The cookies themselves and any objects attached to +those cookies are hidden from it. + + +=============================== +OBJECT MANAGEMENT STATE MACHINE +=============================== + +Within FS-Cache, each active object is managed by its own individual state +machine. The state for an object is kept in the fscache_object struct, in +object->state. A cookie may point to a set of objects that are in different +states. + +Each state has an action associated with it that is invoked when the machine +wakes up in that state. There are four logical sets of states: + + (1) Preparation: states that wait for the parent objects to become ready. The + representations are hierarchical, and it is expected that an object must + be created or accessed with respect to its parent object. + + (2) Initialisation: states that perform lookups in the cache and validate + what's found and that create on disk any missing metadata. + + (3) Normal running: states that allow netfs operations on objects to proceed + and that update the state of objects. + + (4) Termination: states that detach objects from their netfs cookies, that + delete objects from disk, that handle disk and system errors and that free + up in-memory resources. + + +In most cases, transitioning between states is in response to signalled events. +When a state has finished processing, it will usually set the mask of events in +which it is interested (object->event_mask) and relinquish the worker thread. +Then when an event is raised (by calling fscache_raise_event()), if the event +is not masked, the object will be queued for processing (by calling +fscache_enqueue_object()). + + +PROVISION OF CPU TIME +--------------------- + +The work to be done by the various states is given CPU time by the threads of +the slow work facility (see Documentation/slow-work.txt). This is used in +preference to the workqueue facility because: + + (1) Threads may be completely occupied for very long periods of time by a + particular work item. These state actions may be doing sequences of + synchronous, journalled disk accesses (lookup, mkdir, create, setxattr, + getxattr, truncate, unlink, rmdir, rename). + + (2) Threads may do little actual work, but may rather spend a lot of time + sleeping on I/O. This means that single-threaded and 1-per-CPU-threaded + workqueues don't necessarily have the right numbers of threads. + + +LOCKING SIMPLIFICATION +---------------------- + +Because only one worker thread may be operating on any particular object's +state machine at once, this simplifies the locking, particularly with respect +to disconnecting the netfs's representation of a cache object (fscache_cookie) +from the cache backend's representation (fscache_object) - which may be +requested from either end. + + +================= +THE SET OF STATES +================= + +The object state machine has a set of states that it can be in. There are +preparation states in which the object sets itself up and waits for its parent +object to transit to a state that allows access to its children: + + (1) State FSCACHE_OBJECT_INIT. + + Initialise the object and wait for the parent object to become active. In + the cache, it is expected that it will not be possible to look an object + up from the parent object, until that parent object itself has been looked + up. + +There are initialisation states in which the object sets itself up and accesses +disk for the object metadata: + + (2) State FSCACHE_OBJECT_LOOKING_UP. + + Look up the object on disk, using the parent as a starting point. + FS-Cache expects the cache backend to probe the cache to see whether this + object is represented there, and if it is, to see if it's valid (coherency + management). + + The cache should call fscache_object_lookup_negative() to indicate lookup + failure for whatever reason, and should call fscache_obtained_object() to + indicate success. + + At the completion of lookup, FS-Cache will let the netfs go ahead with + read operations, no matter whether the file is yet cached. If not yet + cached, read operations will be immediately rejected with ENODATA until + the first known page is uncached - as to that point there can be no data + to be read out of the cache for that file that isn't currently also held + in the pagecache. + + (3) State FSCACHE_OBJECT_CREATING. + + Create an object on disk, using the parent as a starting point. This + happens if the lookup failed to find the object, or if the object's + coherency data indicated what's on disk is out of date. In this state, + FS-Cache expects the cache to create + + The cache should call fscache_obtained_object() if creation completes + successfully, fscache_object_lookup_negative() otherwise. + + At the completion of creation, FS-Cache will start processing write + operations the netfs has queued for an object. If creation failed, the + write ops will be transparently discarded, and nothing recorded in the + cache. + +There are some normal running states in which the object spends its time +servicing netfs requests: + + (4) State FSCACHE_OBJECT_AVAILABLE. + + A transient state in which pending operations are started, child objects + are permitted to advance from FSCACHE_OBJECT_INIT state, and temporary + lookup data is freed. + + (5) State FSCACHE_OBJECT_ACTIVE. + + The normal running state. In this state, requests the netfs makes will be + passed on to the cache. + + (6) State FSCACHE_OBJECT_UPDATING. + + The state machine comes here to update the object in the cache from the + netfs's records. This involves updating the auxiliary data that is used + to maintain coherency. + +And there are terminal states in which an object cleans itself up, deallocates +memory and potentially deletes stuff from disk: + + (7) State FSCACHE_OBJECT_LC_DYING. + + The object comes here if it is dying because of a lookup or creation + error. This would be due to a disk error or system error of some sort. + Temporary data is cleaned up, and the parent is released. + + (8) State FSCACHE_OBJECT_DYING. + + The object comes here if it is dying due to an error, because its parent + cookie has been relinquished by the netfs or because the cache is being + withdrawn. + + Any child objects waiting on this one are given CPU time so that they too + can destroy themselves. This object waits for all its children to go away + before advancing to the next state. + + (9) State FSCACHE_OBJECT_ABORT_INIT. + + The object comes to this state if it was waiting on its parent in + FSCACHE_OBJECT_INIT, but its parent died. The object will destroy itself + so that the parent may proceed from the FSCACHE_OBJECT_DYING state. + +(10) State FSCACHE_OBJECT_RELEASING. +(11) State FSCACHE_OBJECT_RECYCLING. + + The object comes to one of these two states when dying once it is rid of + all its children, if it is dying because the netfs relinquished its + cookie. In the first state, the cached data is expected to persist, and + in the second it will be deleted. + +(12) State FSCACHE_OBJECT_WITHDRAWING. + + The object transits to this state if the cache decides it wants to + withdraw the object from service, perhaps to make space, but also due to + error or just because the whole cache is being withdrawn. + +(13) State FSCACHE_OBJECT_DEAD. + + The object transits to this state when the in-memory object record is + ready to be deleted. The object processor shouldn't ever see an object in + this state. + + +THE SET OF EVENTS +----------------- + +There are a number of events that can be raised to an object state machine: + + (*) FSCACHE_OBJECT_EV_UPDATE + + The netfs requested that an object be updated. The state machine will ask + the cache backend to update the object, and the cache backend will ask the + netfs for details of the change through its cookie definition ops. + + (*) FSCACHE_OBJECT_EV_CLEARED + + This is signalled in two circumstances: + + (a) when an object's last child object is dropped and + + (b) when the last operation outstanding on an object is completed. + + This is used to proceed from the dying state. + + (*) FSCACHE_OBJECT_EV_ERROR + + This is signalled when an I/O error occurs during the processing of some + object. + + (*) FSCACHE_OBJECT_EV_RELEASE + (*) FSCACHE_OBJECT_EV_RETIRE + + These are signalled when the netfs relinquishes a cookie it was using. + The event selected depends on whether the netfs asks for the backing + object to be retired (deleted) or retained. + + (*) FSCACHE_OBJECT_EV_WITHDRAW + + This is signalled when the cache backend wants to withdraw an object. + This means that the object will have to be detached from the netfs's + cookie. + +Because the withdrawing releasing/retiring events are all handled by the object +state machine, it doesn't matter if there's a collision with both ends trying +to sever the connection at the same time. The state machine can just pick +which one it wants to honour, and that effects the other. diff --git a/Documentation/filesystems/caching/operations.txt b/Documentation/filesystems/caching/operations.txt new file mode 100644 index 00000000..b6b070c5 --- /dev/null +++ b/Documentation/filesystems/caching/operations.txt @@ -0,0 +1,213 @@ + ================================ + ASYNCHRONOUS OPERATIONS HANDLING + ================================ + +By: David Howells <dhowells@redhat.com> + +Contents: + + (*) Overview. + + (*) Operation record initialisation. + + (*) Parameters. + + (*) Procedure. + + (*) Asynchronous callback. + + +======== +OVERVIEW +======== + +FS-Cache has an asynchronous operations handling facility that it uses for its +data storage and retrieval routines. Its operations are represented by +fscache_operation structs, though these are usually embedded into some other +structure. + +This facility is available to and expected to be be used by the cache backends, +and FS-Cache will create operations and pass them off to the appropriate cache +backend for completion. + +To make use of this facility, <linux/fscache-cache.h> should be #included. + + +=============================== +OPERATION RECORD INITIALISATION +=============================== + +An operation is recorded in an fscache_operation struct: + + struct fscache_operation { + union { + struct work_struct fast_work; + struct slow_work slow_work; + }; + unsigned long flags; + fscache_operation_processor_t processor; + ... + }; + +Someone wanting to issue an operation should allocate something with this +struct embedded in it. They should initialise it by calling: + + void fscache_operation_init(struct fscache_operation *op, + fscache_operation_release_t release); + +with the operation to be initialised and the release function to use. + +The op->flags parameter should be set to indicate the CPU time provision and +the exclusivity (see the Parameters section). + +The op->fast_work, op->slow_work and op->processor flags should be set as +appropriate for the CPU time provision (see the Parameters section). + +FSCACHE_OP_WAITING may be set in op->flags prior to each submission of the +operation and waited for afterwards. + + +========== +PARAMETERS +========== + +There are a number of parameters that can be set in the operation record's flag +parameter. There are three options for the provision of CPU time in these +operations: + + (1) The operation may be done synchronously (FSCACHE_OP_MYTHREAD). A thread + may decide it wants to handle an operation itself without deferring it to + another thread. + + This is, for example, used in read operations for calling readpages() on + the backing filesystem in CacheFiles. Although readpages() does an + asynchronous data fetch, the determination of whether pages exist is done + synchronously - and the netfs does not proceed until this has been + determined. + + If this option is to be used, FSCACHE_OP_WAITING must be set in op->flags + before submitting the operation, and the operating thread must wait for it + to be cleared before proceeding: + + wait_on_bit(&op->flags, FSCACHE_OP_WAITING, + fscache_wait_bit, TASK_UNINTERRUPTIBLE); + + + (2) The operation may be fast asynchronous (FSCACHE_OP_FAST), in which case it + will be given to keventd to process. Such an operation is not permitted + to sleep on I/O. + + This is, for example, used by CacheFiles to copy data from a backing fs + page to a netfs page after the backing fs has read the page in. + + If this option is used, op->fast_work and op->processor must be + initialised before submitting the operation: + + INIT_WORK(&op->fast_work, do_some_work); + + + (3) The operation may be slow asynchronous (FSCACHE_OP_SLOW), in which case it + will be given to the slow work facility to process. Such an operation is + permitted to sleep on I/O. + + This is, for example, used by FS-Cache to handle background writes of + pages that have just been fetched from a remote server. + + If this option is used, op->slow_work and op->processor must be + initialised before submitting the operation: + + fscache_operation_init_slow(op, processor) + + +Furthermore, operations may be one of two types: + + (1) Exclusive (FSCACHE_OP_EXCLUSIVE). Operations of this type may not run in + conjunction with any other operation on the object being operated upon. + + An example of this is the attribute change operation, in which the file + being written to may need truncation. + + (2) Shareable. Operations of this type may be running simultaneously. It's + up to the operation implementation to prevent interference between other + operations running at the same time. + + +========= +PROCEDURE +========= + +Operations are used through the following procedure: + + (1) The submitting thread must allocate the operation and initialise it + itself. Normally this would be part of a more specific structure with the + generic op embedded within. + + (2) The submitting thread must then submit the operation for processing using + one of the following two functions: + + int fscache_submit_op(struct fscache_object *object, + struct fscache_operation *op); + + int fscache_submit_exclusive_op(struct fscache_object *object, + struct fscache_operation *op); + + The first function should be used to submit non-exclusive ops and the + second to submit exclusive ones. The caller must still set the + FSCACHE_OP_EXCLUSIVE flag. + + If successful, both functions will assign the operation to the specified + object and return 0. -ENOBUFS will be returned if the object specified is + permanently unavailable. + + The operation manager will defer operations on an object that is still + undergoing lookup or creation. The operation will also be deferred if an + operation of conflicting exclusivity is in progress on the object. + + If the operation is asynchronous, the manager will retain a reference to + it, so the caller should put their reference to it by passing it to: + + void fscache_put_operation(struct fscache_operation *op); + + (3) If the submitting thread wants to do the work itself, and has marked the + operation with FSCACHE_OP_MYTHREAD, then it should monitor + FSCACHE_OP_WAITING as described above and check the state of the object if + necessary (the object might have died whilst the thread was waiting). + + When it has finished doing its processing, it should call + fscache_put_operation() on it. + + (4) The operation holds an effective lock upon the object, preventing other + exclusive ops conflicting until it is released. The operation can be + enqueued for further immediate asynchronous processing by adjusting the + CPU time provisioning option if necessary, eg: + + op->flags &= ~FSCACHE_OP_TYPE; + op->flags |= ~FSCACHE_OP_FAST; + + and calling: + + void fscache_enqueue_operation(struct fscache_operation *op) + + This can be used to allow other things to have use of the worker thread + pools. + + +===================== +ASYNCHRONOUS CALLBACK +===================== + +When used in asynchronous mode, the worker thread pool will invoke the +processor method with a pointer to the operation. This should then get at the +container struct by using container_of(): + + static void fscache_write_op(struct fscache_operation *_op) + { + struct fscache_storage *op = + container_of(_op, struct fscache_storage, op); + ... + } + +The caller holds a reference on the operation, and will invoke +fscache_put_operation() when the processor function returns. The processor +function is at liberty to call fscache_enqueue_operation() or to take extra +references. diff --git a/Documentation/filesystems/ceph.txt b/Documentation/filesystems/ceph.txt new file mode 100644 index 00000000..763d8ebb --- /dev/null +++ b/Documentation/filesystems/ceph.txt @@ -0,0 +1,140 @@ +Ceph Distributed File System +============================ + +Ceph is a distributed network file system designed to provide good +performance, reliability, and scalability. + +Basic features include: + + * POSIX semantics + * Seamless scaling from 1 to many thousands of nodes + * High availability and reliability. No single point of failure. + * N-way replication of data across storage nodes + * Fast recovery from node failures + * Automatic rebalancing of data on node addition/removal + * Easy deployment: most FS components are userspace daemons + +Also, + * Flexible snapshots (on any directory) + * Recursive accounting (nested files, directories, bytes) + +In contrast to cluster filesystems like GFS, OCFS2, and GPFS that rely +on symmetric access by all clients to shared block devices, Ceph +separates data and metadata management into independent server +clusters, similar to Lustre. Unlike Lustre, however, metadata and +storage nodes run entirely as user space daemons. Storage nodes +utilize btrfs to store data objects, leveraging its advanced features +(checksumming, metadata replication, etc.). File data is striped +across storage nodes in large chunks to distribute workload and +facilitate high throughputs. When storage nodes fail, data is +re-replicated in a distributed fashion by the storage nodes themselves +(with some minimal coordination from a cluster monitor), making the +system extremely efficient and scalable. + +Metadata servers effectively form a large, consistent, distributed +in-memory cache above the file namespace that is extremely scalable, +dynamically redistributes metadata in response to workload changes, +and can tolerate arbitrary (well, non-Byzantine) node failures. The +metadata server takes a somewhat unconventional approach to metadata +storage to significantly improve performance for common workloads. In +particular, inodes with only a single link are embedded in +directories, allowing entire directories of dentries and inodes to be +loaded into its cache with a single I/O operation. The contents of +extremely large directories can be fragmented and managed by +independent metadata servers, allowing scalable concurrent access. + +The system offers automatic data rebalancing/migration when scaling +from a small cluster of just a few nodes to many hundreds, without +requiring an administrator carve the data set into static volumes or +go through the tedious process of migrating data between servers. +When the file system approaches full, new nodes can be easily added +and things will "just work." + +Ceph includes flexible snapshot mechanism that allows a user to create +a snapshot on any subdirectory (and its nested contents) in the +system. Snapshot creation and deletion are as simple as 'mkdir +.snap/foo' and 'rmdir .snap/foo'. + +Ceph also provides some recursive accounting on directories for nested +files and bytes. That is, a 'getfattr -d foo' on any directory in the +system will reveal the total number of nested regular files and +subdirectories, and a summation of all nested file sizes. This makes +the identification of large disk space consumers relatively quick, as +no 'du' or similar recursive scan of the file system is required. + + +Mount Syntax +============ + +The basic mount syntax is: + + # mount -t ceph monip[:port][,monip2[:port]...]:/[subdir] mnt + +You only need to specify a single monitor, as the client will get the +full list when it connects. (However, if the monitor you specify +happens to be down, the mount won't succeed.) The port can be left +off if the monitor is using the default. So if the monitor is at +1.2.3.4, + + # mount -t ceph 1.2.3.4:/ /mnt/ceph + +is sufficient. If /sbin/mount.ceph is installed, a hostname can be +used instead of an IP address. + + + +Mount Options +============= + + ip=A.B.C.D[:N] + Specify the IP and/or port the client should bind to locally. + There is normally not much reason to do this. If the IP is not + specified, the client's IP address is determined by looking at the + address its connection to the monitor originates from. + + wsize=X + Specify the maximum write size in bytes. By default there is no + maximum. Ceph will normally size writes based on the file stripe + size. + + rsize=X + Specify the maximum readahead. + + mount_timeout=X + Specify the timeout value for mount (in seconds), in the case + of a non-responsive Ceph file system. The default is 30 + seconds. + + rbytes + When stat() is called on a directory, set st_size to 'rbytes', + the summation of file sizes over all files nested beneath that + directory. This is the default. + + norbytes + When stat() is called on a directory, set st_size to the + number of entries in that directory. + + nocrc + Disable CRC32C calculation for data writes. If set, the storage node + must rely on TCP's error correction to detect data corruption + in the data payload. + + noasyncreaddir + Disable client's use its local cache to satisfy readdir + requests. (This does not change correctness; the client uses + cached metadata only when a lease or capability ensures it is + valid.) + + +More Information +================ + +For more information on Ceph, see the home page at + http://ceph.newdream.net/ + +The Linux kernel client source tree is available at + git://ceph.newdream.net/git/ceph-client.git + git://git.kernel.org/pub/scm/linux/kernel/git/sage/ceph-client.git + +and the source for the full system is at + git://ceph.newdream.net/git/ceph.git diff --git a/Documentation/filesystems/cifs.txt b/Documentation/filesystems/cifs.txt new file mode 100644 index 00000000..49cc923a --- /dev/null +++ b/Documentation/filesystems/cifs.txt @@ -0,0 +1,51 @@ + This is the client VFS module for the Common Internet File System + (CIFS) protocol which is the successor to the Server Message Block + (SMB) protocol, the native file sharing mechanism for most early + PC operating systems. CIFS is fully supported by current network + file servers such as Windows 2000, Windows 2003 (including + Windows XP) as well by Samba (which provides excellent CIFS + server support for Linux and many other operating systems), so + this network filesystem client can mount to a wide variety of + servers. The smbfs module should be used instead of this cifs module + for mounting to older SMB servers such as OS/2. The smbfs and cifs + modules can coexist and do not conflict. The CIFS VFS filesystem + module is designed to work well with servers that implement the + newer versions (dialects) of the SMB/CIFS protocol such as Samba, + the program written by Andrew Tridgell that turns any Unix host + into a SMB/CIFS file server. + + The intent of this module is to provide the most advanced network + file system function for CIFS compliant servers, including better + POSIX compliance, secure per-user session establishment, high + performance safe distributed caching (oplock), optional packet + signing, large files, Unicode support and other internationalization + improvements. Since both Samba server and this filesystem client support + the CIFS Unix extensions, the combination can provide a reasonable + alternative to NFSv4 for fileserving in some Linux to Linux environments, + not just in Linux to Windows environments. + + This filesystem has an optional mount utility (mount.cifs) that can + be obtained from the project page and installed in the path in the same + directory with the other mount helpers (such as mount.smbfs). + Mounting using the cifs filesystem without installing the mount helper + requires specifying the server's ip address. + + For Linux 2.4: + mount //anything/here /mnt_target -o + user=username,pass=password,unc=//ip_address_of_server/sharename + + For Linux 2.5: + mount //ip_address_of_server/sharename /mnt_target -o user=username, pass=password + + + For more information on the module see the project page at + + http://us1.samba.org/samba/Linux_CIFS_client.html + + For more information on CIFS see: + + http://www.snia.org/tech_activities/CIFS + + or the Samba site: + + http://www.samba.org diff --git a/Documentation/filesystems/coda.txt b/Documentation/filesystems/coda.txt new file mode 100644 index 00000000..61311356 --- /dev/null +++ b/Documentation/filesystems/coda.txt @@ -0,0 +1,1673 @@ +NOTE: +This is one of the technical documents describing a component of +Coda -- this document describes the client kernel-Venus interface. + +For more information: + http://www.coda.cs.cmu.edu +For user level software needed to run Coda: + ftp://ftp.coda.cs.cmu.edu + +To run Coda you need to get a user level cache manager for the client, +named Venus, as well as tools to manipulate ACLs, to log in, etc. The +client needs to have the Coda filesystem selected in the kernel +configuration. + +The server needs a user level server and at present does not depend on +kernel support. + + + + + + + + The Venus kernel interface + Peter J. Braam + v1.0, Nov 9, 1997 + + This document describes the communication between Venus and kernel + level filesystem code needed for the operation of the Coda file sys- + tem. This document version is meant to describe the current interface + (version 1.0) as well as improvements we envisage. + ______________________________________________________________________ + + Table of Contents + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + + 1. Introduction + + 2. Servicing Coda filesystem calls + + 3. The message layer + + 3.1 Implementation details + + 4. The interface at the call level + + 4.1 Data structures shared by the kernel and Venus + 4.2 The pioctl interface + 4.3 root + 4.4 lookup + 4.5 getattr + 4.6 setattr + 4.7 access + 4.8 create + 4.9 mkdir + 4.10 link + 4.11 symlink + 4.12 remove + 4.13 rmdir + 4.14 readlink + 4.15 open + 4.16 close + 4.17 ioctl + 4.18 rename + 4.19 readdir + 4.20 vget + 4.21 fsync + 4.22 inactive + 4.23 rdwr + 4.24 odymount + 4.25 ody_lookup + 4.26 ody_expand + 4.27 prefetch + 4.28 signal + + 5. The minicache and downcalls + + 5.1 INVALIDATE + 5.2 FLUSH + 5.3 PURGEUSER + 5.4 ZAPFILE + 5.5 ZAPDIR + 5.6 ZAPVNODE + 5.7 PURGEFID + 5.8 REPLACE + + 6. Initialization and cleanup + + 6.1 Requirements + + + ______________________________________________________________________ + 0wpage + + 11.. IInnttrroodduuccttiioonn + + + + A key component in the Coda Distributed File System is the cache + manager, _V_e_n_u_s. + + + When processes on a Coda enabled system access files in the Coda + filesystem, requests are directed at the filesystem layer in the + operating system. The operating system will communicate with Venus to + service the request for the process. Venus manages a persistent + client cache and makes remote procedure calls to Coda file servers and + related servers (such as authentication servers) to service these + requests it receives from the operating system. When Venus has + serviced a request it replies to the operating system with appropriate + return codes, and other data related to the request. Optionally the + kernel support for Coda may maintain a minicache of recently processed + requests to limit the number of interactions with Venus. Venus + possesses the facility to inform the kernel when elements from its + minicache are no longer valid. + + This document describes precisely this communication between the + kernel and Venus. The definitions of so called upcalls and downcalls + will be given with the format of the data they handle. We shall also + describe the semantic invariants resulting from the calls. + + Historically Coda was implemented in a BSD file system in Mach 2.6. + The interface between the kernel and Venus is very similar to the BSD + VFS interface. Similar functionality is provided, and the format of + the parameters and returned data is very similar to the BSD VFS. This + leads to an almost natural environment for implementing a kernel-level + filesystem driver for Coda in a BSD system. However, other operating + systems such as Linux and Windows 95 and NT have virtual filesystem + with different interfaces. + + To implement Coda on these systems some reverse engineering of the + Venus/Kernel protocol is necessary. Also it came to light that other + systems could profit significantly from certain small optimizations + and modifications to the protocol. To facilitate this work as well as + to make future ports easier, communication between Venus and the + kernel should be documented in great detail. This is the aim of this + document. + + 0wpage + + 22.. SSeerrvviicciinngg CCooddaa ffiilleessyysstteemm ccaallllss + + The service of a request for a Coda file system service originates in + a process PP which accessing a Coda file. It makes a system call which + traps to the OS kernel. Examples of such calls trapping to the kernel + are _r_e_a_d_, _w_r_i_t_e_, _o_p_e_n_, _c_l_o_s_e_, _c_r_e_a_t_e_, _m_k_d_i_r_, _r_m_d_i_r_, _c_h_m_o_d in a Unix + context. Similar calls exist in the Win32 environment, and are named + _C_r_e_a_t_e_F_i_l_e_, . + + Generally the operating system handles the request in a virtual + filesystem (VFS) layer, which is named I/O Manager in NT and IFS + manager in Windows 95. The VFS is responsible for partial processing + of the request and for locating the specific filesystem(s) which will + service parts of the request. Usually the information in the path + assists in locating the correct FS drivers. Sometimes after extensive + pre-processing, the VFS starts invoking exported routines in the FS + driver. This is the point where the FS specific processing of the + request starts, and here the Coda specific kernel code comes into + play. + + The FS layer for Coda must expose and implement several interfaces. + First and foremost the VFS must be able to make all necessary calls to + the Coda FS layer, so the Coda FS driver must expose the VFS interface + as applicable in the operating system. These differ very significantly + among operating systems, but share features such as facilities to + read/write and create and remove objects. The Coda FS layer services + such VFS requests by invoking one or more well defined services + offered by the cache manager Venus. When the replies from Venus have + come back to the FS driver, servicing of the VFS call continues and + finishes with a reply to the kernel's VFS. Finally the VFS layer + returns to the process. + + As a result of this design a basic interface exposed by the FS driver + must allow Venus to manage message traffic. In particular Venus must + be able to retrieve and place messages and to be notified of the + arrival of a new message. The notification must be through a mechanism + which does not block Venus since Venus must attend to other tasks even + when no messages are waiting or being processed. + + + + + + + Interfaces of the Coda FS Driver + + Furthermore the FS layer provides for a special path of communication + between a user process and Venus, called the pioctl interface. The + pioctl interface is used for Coda specific services, such as + requesting detailed information about the persistent cache managed by + Venus. Here the involvement of the kernel is minimal. It identifies + the calling process and passes the information on to Venus. When + Venus replies the response is passed back to the caller in unmodified + form. + + Finally Venus allows the kernel FS driver to cache the results from + certain services. This is done to avoid excessive context switches + and results in an efficient system. However, Venus may acquire + information, for example from the network which implies that cached + information must be flushed or replaced. Venus then makes a downcall + to the Coda FS layer to request flushes or updates in the cache. The + kernel FS driver handles such requests synchronously. + + Among these interfaces the VFS interface and the facility to place, + receive and be notified of messages are platform specific. We will + not go into the calls exported to the VFS layer but we will state the + requirements of the message exchange mechanism. + + 0wpage + + 33.. TThhee mmeessssaaggee llaayyeerr + + + + At the lowest level the communication between Venus and the FS driver + proceeds through messages. The synchronization between processes + requesting Coda file service and Venus relies on blocking and waking + up processes. The Coda FS driver processes VFS- and pioctl-requests + on behalf of a process P, creates messages for Venus, awaits replies + and finally returns to the caller. The implementation of the exchange + of messages is platform specific, but the semantics have (so far) + appeared to be generally applicable. Data buffers are created by the + FS Driver in kernel memory on behalf of P and copied to user memory in + Venus. + + The FS Driver while servicing P makes upcalls to Venus. Such an + upcall is dispatched to Venus by creating a message structure. The + structure contains the identification of P, the message sequence + number, the size of the request and a pointer to the data in kernel + memory for the request. Since the data buffer is re-used to hold the + reply from Venus, there is a field for the size of the reply. A flags + field is used in the message to precisely record the status of the + message. Additional platform dependent structures involve pointers to + determine the position of the message on queues and pointers to + synchronization objects. In the upcall routine the message structure + is filled in, flags are set to 0, and it is placed on the _p_e_n_d_i_n_g + queue. The routine calling upcall is responsible for allocating the + data buffer; its structure will be described in the next section. + + A facility must exist to notify Venus that the message has been + created, and implemented using available synchronization objects in + the OS. This notification is done in the upcall context of the process + P. When the message is on the pending queue, process P cannot proceed + in upcall. The (kernel mode) processing of P in the filesystem + request routine must be suspended until Venus has replied. Therefore + the calling thread in P is blocked in upcall. A pointer in the + message structure will locate the synchronization object on which P is + sleeping. + + Venus detects the notification that a message has arrived, and the FS + driver allow Venus to retrieve the message with a getmsg_from_kernel + call. This action finishes in the kernel by putting the message on the + queue of processing messages and setting flags to READ. Venus is + passed the contents of the data buffer. The getmsg_from_kernel call + now returns and Venus processes the request. + + At some later point the FS driver receives a message from Venus, + namely when Venus calls sendmsg_to_kernel. At this moment the Coda FS + driver looks at the contents of the message and decides if: + + + +o the message is a reply for a suspended thread P. If so it removes + the message from the processing queue and marks the message as + WRITTEN. Finally, the FS driver unblocks P (still in the kernel + mode context of Venus) and the sendmsg_to_kernel call returns to + Venus. The process P will be scheduled at some point and continues + processing its upcall with the data buffer replaced with the reply + from Venus. + + +o The message is a _d_o_w_n_c_a_l_l. A downcall is a request from Venus to + the FS Driver. The FS driver processes the request immediately + (usually a cache eviction or replacement) and when it finishes + sendmsg_to_kernel returns. + + Now P awakes and continues processing upcall. There are some + subtleties to take account of. First P will determine if it was woken + up in upcall by a signal from some other source (for example an + attempt to terminate P) or as is normally the case by Venus in its + sendmsg_to_kernel call. In the normal case, the upcall routine will + deallocate the message structure and return. The FS routine can proceed + with its processing. + + + + + + + + Sleeping and IPC arrangements + + In case P is woken up by a signal and not by Venus, it will first look + at the flags field. If the message is not yet READ, the process P can + handle its signal without notifying Venus. If Venus has READ, and + the request should not be processed, P can send Venus a signal message + to indicate that it should disregard the previous message. Such + signals are put in the queue at the head, and read first by Venus. If + the message is already marked as WRITTEN it is too late to stop the + processing. The VFS routine will now continue. (-- If a VFS request + involves more than one upcall, this can lead to complicated state, an + extra field "handle_signals" could be added in the message structure + to indicate points of no return have been passed.--) + + + + 33..11.. IImmpplleemmeennttaattiioonn ddeettaaiillss + + The Unix implementation of this mechanism has been through the + implementation of a character device associated with Coda. Venus + retrieves messages by doing a read on the device, replies are sent + with a write and notification is through the select system call on the + file descriptor for the device. The process P is kept waiting on an + interruptible wait queue object. + + In Windows NT and the DPMI Windows 95 implementation a DeviceIoControl + call is used. The DeviceIoControl call is designed to copy buffers + from user memory to kernel memory with OPCODES. The sendmsg_to_kernel + is issued as a synchronous call, while the getmsg_from_kernel call is + asynchronous. Windows EventObjects are used for notification of + message arrival. The process P is kept waiting on a KernelEvent + object in NT and a semaphore in Windows 95. + + 0wpage + + 44.. TThhee iinntteerrffaaccee aatt tthhee ccaallll lleevveell + + + This section describes the upcalls a Coda FS driver can make to Venus. + Each of these upcalls make use of two structures: inputArgs and + outputArgs. In pseudo BNF form the structures take the following + form: + + + struct inputArgs { + u_long opcode; + u_long unique; /* Keep multiple outstanding msgs distinct */ + u_short pid; /* Common to all */ + u_short pgid; /* Common to all */ + struct CodaCred cred; /* Common to all */ + + <union "in" of call dependent parts of inputArgs> + }; + + struct outputArgs { + u_long opcode; + u_long unique; /* Keep multiple outstanding msgs distinct */ + u_long result; + + <union "out" of call dependent parts of inputArgs> + }; + + + + Before going on let us elucidate the role of the various fields. The + inputArgs start with the opcode which defines the type of service + requested from Venus. There are approximately 30 upcalls at present + which we will discuss. The unique field labels the inputArg with a + unique number which will identify the message uniquely. A process and + process group id are passed. Finally the credentials of the caller + are included. + + Before delving into the specific calls we need to discuss a variety of + data structures shared by the kernel and Venus. + + + + + 44..11.. DDaattaa ssttrruuccttuurreess sshhaarreedd bbyy tthhee kkeerrnneell aanndd VVeennuuss + + + The CodaCred structure defines a variety of user and group ids as + they are set for the calling process. The vuid_t and guid_t are 32 bit + unsigned integers. It also defines group membership in an array. On + Unix the CodaCred has proven sufficient to implement good security + semantics for Coda but the structure may have to undergo modification + for the Windows environment when these mature. + + struct CodaCred { + vuid_t cr_uid, cr_euid, cr_suid, cr_fsuid; /* Real, effective, set, fs uid*/ + vgid_t cr_gid, cr_egid, cr_sgid, cr_fsgid; /* same for groups */ + vgid_t cr_groups[NGROUPS]; /* Group membership for caller */ + }; + + + + NNOOTTEE It is questionable if we need CodaCreds in Venus. Finally Venus + doesn't know about groups, although it does create files with the + default uid/gid. Perhaps the list of group membership is superfluous. + + + The next item is the fundamental identifier used to identify Coda + files, the ViceFid. A fid of a file uniquely defines a file or + directory in the Coda filesystem within a _c_e_l_l. (-- A _c_e_l_l is a + group of Coda servers acting under the aegis of a single system + control machine or SCM. See the Coda Administration manual for a + detailed description of the role of the SCM.--) + + + typedef struct ViceFid { + VolumeId Volume; + VnodeId Vnode; + Unique_t Unique; + } ViceFid; + + + + Each of the constituent fields: VolumeId, VnodeId and Unique_t are + unsigned 32 bit integers. We envisage that a further field will need + to be prefixed to identify the Coda cell; this will probably take the + form of a Ipv6 size IP address naming the Coda cell through DNS. + + The next important structure shared between Venus and the kernel is + the attributes of the file. The following structure is used to + exchange information. It has room for future extensions such as + support for device files (currently not present in Coda). + + + + + + + + + + + + + + + + + + + struct coda_vattr { + enum coda_vtype va_type; /* vnode type (for create) */ + u_short va_mode; /* files access mode and type */ + short va_nlink; /* number of references to file */ + vuid_t va_uid; /* owner user id */ + vgid_t va_gid; /* owner group id */ + long va_fsid; /* file system id (dev for now) */ + long va_fileid; /* file id */ + u_quad_t va_size; /* file size in bytes */ + long va_blocksize; /* blocksize preferred for i/o */ + struct timespec va_atime; /* time of last access */ + struct timespec va_mtime; /* time of last modification */ + struct timespec va_ctime; /* time file changed */ + u_long va_gen; /* generation number of file */ + u_long va_flags; /* flags defined for file */ + dev_t va_rdev; /* device special file represents */ + u_quad_t va_bytes; /* bytes of disk space held by file */ + u_quad_t va_filerev; /* file modification number */ + u_int va_vaflags; /* operations flags, see below */ + long va_spare; /* remain quad aligned */ + }; + + + + + 44..22.. TThhee ppiiooccttll iinntteerrffaaccee + + + Coda specific requests can be made by application through the pioctl + interface. The pioctl is implemented as an ordinary ioctl on a + fictitious file /coda/.CONTROL. The pioctl call opens this file, gets + a file handle and makes the ioctl call. Finally it closes the file. + + The kernel involvement in this is limited to providing the facility to + open and close and pass the ioctl message _a_n_d to verify that a path in + the pioctl data buffers is a file in a Coda filesystem. + + The kernel is handed a data packet of the form: + + struct { + const char *path; + struct ViceIoctl vidata; + int follow; + } data; + + + + where + + + struct ViceIoctl { + caddr_t in, out; /* Data to be transferred in, or out */ + short in_size; /* Size of input buffer <= 2K */ + short out_size; /* Maximum size of output buffer, <= 2K */ + }; + + + + The path must be a Coda file, otherwise the ioctl upcall will not be + made. + + NNOOTTEE The data structures and code are a mess. We need to clean this + up. + + We now proceed to document the individual calls: + + 0wpage + + 44..33.. rroooott + + + AArrgguummeennttss + + iinn empty + + oouutt + + struct cfs_root_out { + ViceFid VFid; + } cfs_root; + + + + DDeessccrriippttiioonn This call is made to Venus during the initialization of + the Coda filesystem. If the result is zero, the cfs_root structure + contains the ViceFid of the root of the Coda filesystem. If a non-zero + result is generated, its value is a platform dependent error code + indicating the difficulty Venus encountered in locating the root of + the Coda filesystem. + + 0wpage + + 44..44.. llooookkuupp + + + SSuummmmaarryy Find the ViceFid and type of an object in a directory if it + exists. + + AArrgguummeennttss + + iinn + + struct cfs_lookup_in { + ViceFid VFid; + char *name; /* Place holder for data. */ + } cfs_lookup; + + + + oouutt + + struct cfs_lookup_out { + ViceFid VFid; + int vtype; + } cfs_lookup; + + + + DDeessccrriippttiioonn This call is made to determine the ViceFid and filetype of + a directory entry. The directory entry requested carries name name + and Venus will search the directory identified by cfs_lookup_in.VFid. + The result may indicate that the name does not exist, or that + difficulty was encountered in finding it (e.g. due to disconnection). + If the result is zero, the field cfs_lookup_out.VFid contains the + targets ViceFid and cfs_lookup_out.vtype the coda_vtype giving the + type of object the name designates. + + The name of the object is an 8 bit character string of maximum length + CFS_MAXNAMLEN, currently set to 256 (including a 0 terminator.) + + It is extremely important to realize that Venus bitwise ors the field + cfs_lookup.vtype with CFS_NOCACHE to indicate that the object should + not be put in the kernel name cache. + + NNOOTTEE The type of the vtype is currently wrong. It should be + coda_vtype. Linux does not take note of CFS_NOCACHE. It should. + + 0wpage + + 44..55.. ggeettaattttrr + + + SSuummmmaarryy Get the attributes of a file. + + AArrgguummeennttss + + iinn + + struct cfs_getattr_in { + ViceFid VFid; + struct coda_vattr attr; /* XXXXX */ + } cfs_getattr; + + + + oouutt + + struct cfs_getattr_out { + struct coda_vattr attr; + } cfs_getattr; + + + + DDeessccrriippttiioonn This call returns the attributes of the file identified by + fid. + + EErrrroorrss Errors can occur if the object with fid does not exist, is + unaccessible or if the caller does not have permission to fetch + attributes. + + NNoottee Many kernel FS drivers (Linux, NT and Windows 95) need to acquire + the attributes as well as the Fid for the instantiation of an internal + "inode" or "FileHandle". A significant improvement in performance on + such systems could be made by combining the _l_o_o_k_u_p and _g_e_t_a_t_t_r calls + both at the Venus/kernel interaction level and at the RPC level. + + The vattr structure included in the input arguments is superfluous and + should be removed. + + 0wpage + + 44..66.. sseettaattttrr + + + SSuummmmaarryy Set the attributes of a file. + + AArrgguummeennttss + + iinn + + struct cfs_setattr_in { + ViceFid VFid; + struct coda_vattr attr; + } cfs_setattr; + + + + + oouutt + empty + + DDeessccrriippttiioonn The structure attr is filled with attributes to be changed + in BSD style. Attributes not to be changed are set to -1, apart from + vtype which is set to VNON. Other are set to the value to be assigned. + The only attributes which the FS driver may request to change are the + mode, owner, groupid, atime, mtime and ctime. The return value + indicates success or failure. + + EErrrroorrss A variety of errors can occur. The object may not exist, may + be inaccessible, or permission may not be granted by Venus. + + 0wpage + + 44..77.. aacccceessss + + + SSuummmmaarryy + + AArrgguummeennttss + + iinn + + struct cfs_access_in { + ViceFid VFid; + int flags; + } cfs_access; + + + + oouutt + empty + + DDeessccrriippttiioonn Verify if access to the object identified by VFid for + operations described by flags is permitted. The result indicates if + access will be granted. It is important to remember that Coda uses + ACLs to enforce protection and that ultimately the servers, not the + clients enforce the security of the system. The result of this call + will depend on whether a _t_o_k_e_n is held by the user. + + EErrrroorrss The object may not exist, or the ACL describing the protection + may not be accessible. + + 0wpage + + 44..88.. ccrreeaattee + + + SSuummmmaarryy Invoked to create a file + + AArrgguummeennttss + + iinn + + struct cfs_create_in { + ViceFid VFid; + struct coda_vattr attr; + int excl; + int mode; + char *name; /* Place holder for data. */ + } cfs_create; + + + + + oouutt + + struct cfs_create_out { + ViceFid VFid; + struct coda_vattr attr; + } cfs_create; + + + + DDeessccrriippttiioonn This upcall is invoked to request creation of a file. + The file will be created in the directory identified by VFid, its name + will be name, and the mode will be mode. If excl is set an error will + be returned if the file already exists. If the size field in attr is + set to zero the file will be truncated. The uid and gid of the file + are set by converting the CodaCred to a uid using a macro CRTOUID + (this macro is platform dependent). Upon success the VFid and + attributes of the file are returned. The Coda FS Driver will normally + instantiate a vnode, inode or file handle at kernel level for the new + object. + + + EErrrroorrss A variety of errors can occur. Permissions may be insufficient. + If the object exists and is not a file the error EISDIR is returned + under Unix. + + NNOOTTEE The packing of parameters is very inefficient and appears to + indicate confusion between the system call creat and the VFS operation + create. The VFS operation create is only called to create new objects. + This create call differs from the Unix one in that it is not invoked + to return a file descriptor. The truncate and exclusive options, + together with the mode, could simply be part of the mode as it is + under Unix. There should be no flags argument; this is used in open + (2) to return a file descriptor for READ or WRITE mode. + + The attributes of the directory should be returned too, since the size + and mtime changed. + + 0wpage + + 44..99.. mmkkddiirr + + + SSuummmmaarryy Create a new directory. + + AArrgguummeennttss + + iinn + + struct cfs_mkdir_in { + ViceFid VFid; + struct coda_vattr attr; + char *name; /* Place holder for data. */ + } cfs_mkdir; + + + + oouutt + + struct cfs_mkdir_out { + ViceFid VFid; + struct coda_vattr attr; + } cfs_mkdir; + + + + + DDeessccrriippttiioonn This call is similar to create but creates a directory. + Only the mode field in the input parameters is used for creation. + Upon successful creation, the attr returned contains the attributes of + the new directory. + + EErrrroorrss As for create. + + NNOOTTEE The input parameter should be changed to mode instead of + attributes. + + The attributes of the parent should be returned since the size and + mtime changes. + + 0wpage + + 44..1100.. lliinnkk + + + SSuummmmaarryy Create a link to an existing file. + + AArrgguummeennttss + + iinn + + struct cfs_link_in { + ViceFid sourceFid; /* cnode to link *to* */ + ViceFid destFid; /* Directory in which to place link */ + char *tname; /* Place holder for data. */ + } cfs_link; + + + + oouutt + empty + + DDeessccrriippttiioonn This call creates a link to the sourceFid in the directory + identified by destFid with name tname. The source must reside in the + target's parent, i.e. the source must be have parent destFid, i.e. Coda + does not support cross directory hard links. Only the return value is + relevant. It indicates success or the type of failure. + + EErrrroorrss The usual errors can occur.0wpage + + 44..1111.. ssyymmlliinnkk + + + SSuummmmaarryy create a symbolic link + + AArrgguummeennttss + + iinn + + struct cfs_symlink_in { + ViceFid VFid; /* Directory to put symlink in */ + char *srcname; + struct coda_vattr attr; + char *tname; + } cfs_symlink; + + + + oouutt + none + + DDeessccrriippttiioonn Create a symbolic link. The link is to be placed in the + directory identified by VFid and named tname. It should point to the + pathname srcname. The attributes of the newly created object are to + be set to attr. + + EErrrroorrss + + NNOOTTEE The attributes of the target directory should be returned since + its size changed. + + 0wpage + + 44..1122.. rreemmoovvee + + + SSuummmmaarryy Remove a file + + AArrgguummeennttss + + iinn + + struct cfs_remove_in { + ViceFid VFid; + char *name; /* Place holder for data. */ + } cfs_remove; + + + + oouutt + none + + DDeessccrriippttiioonn Remove file named cfs_remove_in.name in directory + identified by VFid. + + EErrrroorrss + + NNOOTTEE The attributes of the directory should be returned since its + mtime and size may change. + + 0wpage + + 44..1133.. rrmmddiirr + + + SSuummmmaarryy Remove a directory + + AArrgguummeennttss + + iinn + + struct cfs_rmdir_in { + ViceFid VFid; + char *name; /* Place holder for data. */ + } cfs_rmdir; + + + + oouutt + none + + DDeessccrriippttiioonn Remove the directory with name name from the directory + identified by VFid. + + EErrrroorrss + + NNOOTTEE The attributes of the parent directory should be returned since + its mtime and size may change. + + 0wpage + + 44..1144.. rreeaaddlliinnkk + + + SSuummmmaarryy Read the value of a symbolic link. + + AArrgguummeennttss + + iinn + + struct cfs_readlink_in { + ViceFid VFid; + } cfs_readlink; + + + + oouutt + + struct cfs_readlink_out { + int count; + caddr_t data; /* Place holder for data. */ + } cfs_readlink; + + + + DDeessccrriippttiioonn This routine reads the contents of symbolic link + identified by VFid into the buffer data. The buffer data must be able + to hold any name up to CFS_MAXNAMLEN (PATH or NAM??). + + EErrrroorrss No unusual errors. + + 0wpage + + 44..1155.. ooppeenn + + + SSuummmmaarryy Open a file. + + AArrgguummeennttss + + iinn + + struct cfs_open_in { + ViceFid VFid; + int flags; + } cfs_open; + + + + oouutt + + struct cfs_open_out { + dev_t dev; + ino_t inode; + } cfs_open; + + + + DDeessccrriippttiioonn This request asks Venus to place the file identified by + VFid in its cache and to note that the calling process wishes to open + it with flags as in open(2). The return value to the kernel differs + for Unix and Windows systems. For Unix systems the Coda FS Driver is + informed of the device and inode number of the container file in the + fields dev and inode. For Windows the path of the container file is + returned to the kernel. + EErrrroorrss + + NNOOTTEE Currently the cfs_open_out structure is not properly adapted to + deal with the Windows case. It might be best to implement two + upcalls, one to open aiming at a container file name, the other at a + container file inode. + + 0wpage + + 44..1166.. cclloossee + + + SSuummmmaarryy Close a file, update it on the servers. + + AArrgguummeennttss + + iinn + + struct cfs_close_in { + ViceFid VFid; + int flags; + } cfs_close; + + + + oouutt + none + + DDeessccrriippttiioonn Close the file identified by VFid. + + EErrrroorrss + + NNOOTTEE The flags argument is bogus and not used. However, Venus' code + has room to deal with an execp input field, probably this field should + be used to inform Venus that the file was closed but is still memory + mapped for execution. There are comments about fetching versus not + fetching the data in Venus vproc_vfscalls. This seems silly. If a + file is being closed, the data in the container file is to be the new + data. Here again the execp flag might be in play to create confusion: + currently Venus might think a file can be flushed from the cache when + it is still memory mapped. This needs to be understood. + + 0wpage + + 44..1177.. iiooccttll + + + SSuummmmaarryy Do an ioctl on a file. This includes the pioctl interface. + + AArrgguummeennttss + + iinn + + struct cfs_ioctl_in { + ViceFid VFid; + int cmd; + int len; + int rwflag; + char *data; /* Place holder for data. */ + } cfs_ioctl; + + + + oouutt + + + struct cfs_ioctl_out { + int len; + caddr_t data; /* Place holder for data. */ + } cfs_ioctl; + + + + DDeessccrriippttiioonn Do an ioctl operation on a file. The command, len and + data arguments are filled as usual. flags is not used by Venus. + + EErrrroorrss + + NNOOTTEE Another bogus parameter. flags is not used. What is the + business about PREFETCHING in the Venus code? + + + 0wpage + + 44..1188.. rreennaammee + + + SSuummmmaarryy Rename a fid. + + AArrgguummeennttss + + iinn + + struct cfs_rename_in { + ViceFid sourceFid; + char *srcname; + ViceFid destFid; + char *destname; + } cfs_rename; + + + + oouutt + none + + DDeessccrriippttiioonn Rename the object with name srcname in directory + sourceFid to destname in destFid. It is important that the names + srcname and destname are 0 terminated strings. Strings in Unix + kernels are not always null terminated. + + EErrrroorrss + + 0wpage + + 44..1199.. rreeaaddddiirr + + + SSuummmmaarryy Read directory entries. + + AArrgguummeennttss + + iinn + + struct cfs_readdir_in { + ViceFid VFid; + int count; + int offset; + } cfs_readdir; + + + + + oouutt + + struct cfs_readdir_out { + int size; + caddr_t data; /* Place holder for data. */ + } cfs_readdir; + + + + DDeessccrriippttiioonn Read directory entries from VFid starting at offset and + read at most count bytes. Returns the data in data and returns + the size in size. + + EErrrroorrss + + NNOOTTEE This call is not used. Readdir operations exploit container + files. We will re-evaluate this during the directory revamp which is + about to take place. + + 0wpage + + 44..2200.. vvggeett + + + SSuummmmaarryy instructs Venus to do an FSDB->Get. + + AArrgguummeennttss + + iinn + + struct cfs_vget_in { + ViceFid VFid; + } cfs_vget; + + + + oouutt + + struct cfs_vget_out { + ViceFid VFid; + int vtype; + } cfs_vget; + + + + DDeessccrriippttiioonn This upcall asks Venus to do a get operation on an fsobj + labelled by VFid. + + EErrrroorrss + + NNOOTTEE This operation is not used. However, it is extremely useful + since it can be used to deal with read/write memory mapped files. + These can be "pinned" in the Venus cache using vget and released with + inactive. + + 0wpage + + 44..2211.. ffssyynncc + + + SSuummmmaarryy Tell Venus to update the RVM attributes of a file. + + AArrgguummeennttss + + iinn + + struct cfs_fsync_in { + ViceFid VFid; + } cfs_fsync; + + + + oouutt + none + + DDeessccrriippttiioonn Ask Venus to update RVM attributes of object VFid. This + should be called as part of kernel level fsync type calls. The + result indicates if the syncing was successful. + + EErrrroorrss + + NNOOTTEE Linux does not implement this call. It should. + + 0wpage + + 44..2222.. iinnaaccttiivvee + + + SSuummmmaarryy Tell Venus a vnode is no longer in use. + + AArrgguummeennttss + + iinn + + struct cfs_inactive_in { + ViceFid VFid; + } cfs_inactive; + + + + oouutt + none + + DDeessccrriippttiioonn This operation returns EOPNOTSUPP. + + EErrrroorrss + + NNOOTTEE This should perhaps be removed. + + 0wpage + + 44..2233.. rrddwwrr + + + SSuummmmaarryy Read or write from a file + + AArrgguummeennttss + + iinn + + struct cfs_rdwr_in { + ViceFid VFid; + int rwflag; + int count; + int offset; + int ioflag; + caddr_t data; /* Place holder for data. */ + } cfs_rdwr; + + + + + oouutt + + struct cfs_rdwr_out { + int rwflag; + int count; + caddr_t data; /* Place holder for data. */ + } cfs_rdwr; + + + + DDeessccrriippttiioonn This upcall asks Venus to read or write from a file. + + EErrrroorrss + + NNOOTTEE It should be removed since it is against the Coda philosophy that + read/write operations never reach Venus. I have been told the + operation does not work. It is not currently used. + + + 0wpage + + 44..2244.. ooddyymmoouunntt + + + SSuummmmaarryy Allows mounting multiple Coda "filesystems" on one Unix mount + point. + + AArrgguummeennttss + + iinn + + struct ody_mount_in { + char *name; /* Place holder for data. */ + } ody_mount; + + + + oouutt + + struct ody_mount_out { + ViceFid VFid; + } ody_mount; + + + + DDeessccrriippttiioonn Asks Venus to return the rootfid of a Coda system named + name. The fid is returned in VFid. + + EErrrroorrss + + NNOOTTEE This call was used by David for dynamic sets. It should be + removed since it causes a jungle of pointers in the VFS mounting area. + It is not used by Coda proper. Call is not implemented by Venus. + + 0wpage + + 44..2255.. ooddyy__llooookkuupp + + + SSuummmmaarryy Looks up something. + + AArrgguummeennttss + + iinn irrelevant + + + oouutt + irrelevant + + DDeessccrriippttiioonn + + EErrrroorrss + + NNOOTTEE Gut it. Call is not implemented by Venus. + + 0wpage + + 44..2266.. ooddyy__eexxppaanndd + + + SSuummmmaarryy expands something in a dynamic set. + + AArrgguummeennttss + + iinn irrelevant + + oouutt + irrelevant + + DDeessccrriippttiioonn + + EErrrroorrss + + NNOOTTEE Gut it. Call is not implemented by Venus. + + 0wpage + + 44..2277.. pprreeffeettcchh + + + SSuummmmaarryy Prefetch a dynamic set. + + AArrgguummeennttss + + iinn Not documented. + + oouutt + Not documented. + + DDeessccrriippttiioonn Venus worker.cc has support for this call, although it is + noted that it doesn't work. Not surprising, since the kernel does not + have support for it. (ODY_PREFETCH is not a defined operation). + + EErrrroorrss + + NNOOTTEE Gut it. It isn't working and isn't used by Coda. + + + 0wpage + + 44..2288.. ssiiggnnaall + + + SSuummmmaarryy Send Venus a signal about an upcall. + + AArrgguummeennttss + + iinn none + + oouutt + not applicable. + + DDeessccrriippttiioonn This is an out-of-band upcall to Venus to inform Venus + that the calling process received a signal after Venus read the + message from the input queue. Venus is supposed to clean up the + operation. + + EErrrroorrss No reply is given. + + NNOOTTEE We need to better understand what Venus needs to clean up and if + it is doing this correctly. Also we need to handle multiple upcall + per system call situations correctly. It would be important to know + what state changes in Venus take place after an upcall for which the + kernel is responsible for notifying Venus to clean up (e.g. open + definitely is such a state change, but many others are maybe not). + + 0wpage + + 55.. TThhee mmiinniiccaacchhee aanndd ddoowwnnccaallllss + + + The Coda FS Driver can cache results of lookup and access upcalls, to + limit the frequency of upcalls. Upcalls carry a price since a process + context switch needs to take place. The counterpart of caching the + information is that Venus will notify the FS Driver that cached + entries must be flushed or renamed. + + The kernel code generally has to maintain a structure which links the + internal file handles (called vnodes in BSD, inodes in Linux and + FileHandles in Windows) with the ViceFid's which Venus maintains. The + reason is that frequent translations back and forth are needed in + order to make upcalls and use the results of upcalls. Such linking + objects are called ccnnooddeess. + + The current minicache implementations have cache entries which record + the following: + + 1. the name of the file + + 2. the cnode of the directory containing the object + + 3. a list of CodaCred's for which the lookup is permitted. + + 4. the cnode of the object + + The lookup call in the Coda FS Driver may request the cnode of the + desired object from the cache, by passing its name, directory and the + CodaCred's of the caller. The cache will return the cnode or indicate + that it cannot be found. The Coda FS Driver must be careful to + invalidate cache entries when it modifies or removes objects. + + When Venus obtains information that indicates that cache entries are + no longer valid, it will make a downcall to the kernel. Downcalls are + intercepted by the Coda FS Driver and lead to cache invalidations of + the kind described below. The Coda FS Driver does not return an error + unless the downcall data could not be read into kernel memory. + + + 55..11.. IINNVVAALLIIDDAATTEE + + + No information is available on this call. + + + 55..22.. FFLLUUSSHH + + + + AArrgguummeennttss None + + SSuummmmaarryy Flush the name cache entirely. + + DDeessccrriippttiioonn Venus issues this call upon startup and when it dies. This + is to prevent stale cache information being held. Some operating + systems allow the kernel name cache to be switched off dynamically. + When this is done, this downcall is made. + + + 55..33.. PPUURRGGEEUUSSEERR + + + AArrgguummeennttss + + struct cfs_purgeuser_out {/* CFS_PURGEUSER is a venus->kernel call */ + struct CodaCred cred; + } cfs_purgeuser; + + + + DDeessccrriippttiioonn Remove all entries in the cache carrying the Cred. This + call is issued when tokens for a user expire or are flushed. + + + 55..44.. ZZAAPPFFIILLEE + + + AArrgguummeennttss + + struct cfs_zapfile_out { /* CFS_ZAPFILE is a venus->kernel call */ + ViceFid CodaFid; + } cfs_zapfile; + + + + DDeessccrriippttiioonn Remove all entries which have the (dir vnode, name) pair. + This is issued as a result of an invalidation of cached attributes of + a vnode. + + NNOOTTEE Call is not named correctly in NetBSD and Mach. The minicache + zapfile routine takes different arguments. Linux does not implement + the invalidation of attributes correctly. + + + + 55..55.. ZZAAPPDDIIRR + + + AArrgguummeennttss + + struct cfs_zapdir_out { /* CFS_ZAPDIR is a venus->kernel call */ + ViceFid CodaFid; + } cfs_zapdir; + + + + DDeessccrriippttiioonn Remove all entries in the cache lying in a directory + CodaFid, and all children of this directory. This call is issued when + Venus receives a callback on the directory. + + + 55..66.. ZZAAPPVVNNOODDEE + + + + AArrgguummeennttss + + struct cfs_zapvnode_out { /* CFS_ZAPVNODE is a venus->kernel call */ + struct CodaCred cred; + ViceFid VFid; + } cfs_zapvnode; + + + + DDeessccrriippttiioonn Remove all entries in the cache carrying the cred and VFid + as in the arguments. This downcall is probably never issued. + + + 55..77.. PPUURRGGEEFFIIDD + + + SSuummmmaarryy + + AArrgguummeennttss + + struct cfs_purgefid_out { /* CFS_PURGEFID is a venus->kernel call */ + ViceFid CodaFid; + } cfs_purgefid; + + + + DDeessccrriippttiioonn Flush the attribute for the file. If it is a dir (odd + vnode), purge its children from the namecache and remove the file from the + namecache. + + + + 55..88.. RREEPPLLAACCEE + + + SSuummmmaarryy Replace the Fid's for a collection of names. + + AArrgguummeennttss + + struct cfs_replace_out { /* cfs_replace is a venus->kernel call */ + ViceFid NewFid; + ViceFid OldFid; + } cfs_replace; + + + + DDeessccrriippttiioonn This routine replaces a ViceFid in the name cache with + another. It is added to allow Venus during reintegration to replace + locally allocated temp fids while disconnected with global fids even + when the reference counts on those fids are not zero. + + 0wpage + + 66.. IInniittiiaalliizzaattiioonn aanndd cclleeaannuupp + + + This section gives brief hints as to desirable features for the Coda + FS Driver at startup and upon shutdown or Venus failures. Before + entering the discussion it is useful to repeat that the Coda FS Driver + maintains the following data: + + + 1. message queues + + 2. cnodes + + 3. name cache entries + + The name cache entries are entirely private to the driver, so they + can easily be manipulated. The message queues will generally have + clear points of initialization and destruction. The cnodes are + much more delicate. User processes hold reference counts in Coda + filesystems and it can be difficult to clean up the cnodes. + + It can expect requests through: + + 1. the message subsystem + + 2. the VFS layer + + 3. pioctl interface + + Currently the _p_i_o_c_t_l passes through the VFS for Coda so we can + treat these similarly. + + + 66..11.. RReeqquuiirreemmeennttss + + + The following requirements should be accommodated: + + 1. The message queues should have open and close routines. On Unix + the opening of the character devices are such routines. + + +o Before opening, no messages can be placed. + + +o Opening will remove any old messages still pending. + + +o Close will notify any sleeping processes that their upcall cannot + be completed. + + +o Close will free all memory allocated by the message queues. + + + 2. At open the namecache shall be initialized to empty state. + + 3. Before the message queues are open, all VFS operations will fail. + Fortunately this can be achieved by making sure than mounting the + Coda filesystem cannot succeed before opening. + + 4. After closing of the queues, no VFS operations can succeed. Here + one needs to be careful, since a few operations (lookup, + read/write, readdir) can proceed without upcalls. These must be + explicitly blocked. + + 5. Upon closing the namecache shall be flushed and disabled. + + 6. All memory held by cnodes can be freed without relying on upcalls. + + 7. Unmounting the file system can be done without relying on upcalls. + + 8. Mounting the Coda filesystem should fail gracefully if Venus cannot + get the rootfid or the attributes of the rootfid. The latter is + best implemented by Venus fetching these objects before attempting + to mount. + + NNOOTTEE NetBSD in particular but also Linux have not implemented the + above requirements fully. For smooth operation this needs to be + corrected. + + + diff --git a/Documentation/filesystems/configfs/Makefile b/Documentation/filesystems/configfs/Makefile new file mode 100644 index 00000000..be7ec5e6 --- /dev/null +++ b/Documentation/filesystems/configfs/Makefile @@ -0,0 +1,3 @@ +ifneq ($(CONFIG_CONFIGFS_FS),) +obj-m += configfs_example_explicit.o configfs_example_macros.o +endif diff --git a/Documentation/filesystems/configfs/configfs.txt b/Documentation/filesystems/configfs/configfs.txt new file mode 100644 index 00000000..dd57bb6b --- /dev/null +++ b/Documentation/filesystems/configfs/configfs.txt @@ -0,0 +1,484 @@ + +configfs - Userspace-driven kernel object configuration. + +Joel Becker <joel.becker@oracle.com> + +Updated: 31 March 2005 + +Copyright (c) 2005 Oracle Corporation, + Joel Becker <joel.becker@oracle.com> + + +[What is configfs?] + +configfs is a ram-based filesystem that provides the converse of +sysfs's functionality. Where sysfs is a filesystem-based view of +kernel objects, configfs is a filesystem-based manager of kernel +objects, or config_items. + +With sysfs, an object is created in kernel (for example, when a device +is discovered) and it is registered with sysfs. Its attributes then +appear in sysfs, allowing userspace to read the attributes via +readdir(3)/read(2). It may allow some attributes to be modified via +write(2). The important point is that the object is created and +destroyed in kernel, the kernel controls the lifecycle of the sysfs +representation, and sysfs is merely a window on all this. + +A configfs config_item is created via an explicit userspace operation: +mkdir(2). It is destroyed via rmdir(2). The attributes appear at +mkdir(2) time, and can be read or modified via read(2) and write(2). +As with sysfs, readdir(3) queries the list of items and/or attributes. +symlink(2) can be used to group items together. Unlike sysfs, the +lifetime of the representation is completely driven by userspace. The +kernel modules backing the items must respond to this. + +Both sysfs and configfs can and should exist together on the same +system. One is not a replacement for the other. + +[Using configfs] + +configfs can be compiled as a module or into the kernel. You can access +it by doing + + mount -t configfs none /config + +The configfs tree will be empty unless client modules are also loaded. +These are modules that register their item types with configfs as +subsystems. Once a client subsystem is loaded, it will appear as a +subdirectory (or more than one) under /config. Like sysfs, the +configfs tree is always there, whether mounted on /config or not. + +An item is created via mkdir(2). The item's attributes will also +appear at this time. readdir(3) can determine what the attributes are, +read(2) can query their default values, and write(2) can store new +values. Like sysfs, attributes should be ASCII text files, preferably +with only one value per file. The same efficiency caveats from sysfs +apply. Don't mix more than one attribute in one attribute file. + +Like sysfs, configfs expects write(2) to store the entire buffer at +once. When writing to configfs attributes, userspace processes should +first read the entire file, modify the portions they wish to change, and +then write the entire buffer back. Attribute files have a maximum size +of one page (PAGE_SIZE, 4096 on i386). + +When an item needs to be destroyed, remove it with rmdir(2). An +item cannot be destroyed if any other item has a link to it (via +symlink(2)). Links can be removed via unlink(2). + +[Configuring FakeNBD: an Example] + +Imagine there's a Network Block Device (NBD) driver that allows you to +access remote block devices. Call it FakeNBD. FakeNBD uses configfs +for its configuration. Obviously, there will be a nice program that +sysadmins use to configure FakeNBD, but somehow that program has to tell +the driver about it. Here's where configfs comes in. + +When the FakeNBD driver is loaded, it registers itself with configfs. +readdir(3) sees this just fine: + + # ls /config + fakenbd + +A fakenbd connection can be created with mkdir(2). The name is +arbitrary, but likely the tool will make some use of the name. Perhaps +it is a uuid or a disk name: + + # mkdir /config/fakenbd/disk1 + # ls /config/fakenbd/disk1 + target device rw + +The target attribute contains the IP address of the server FakeNBD will +connect to. The device attribute is the device on the server. +Predictably, the rw attribute determines whether the connection is +read-only or read-write. + + # echo 10.0.0.1 > /config/fakenbd/disk1/target + # echo /dev/sda1 > /config/fakenbd/disk1/device + # echo 1 > /config/fakenbd/disk1/rw + +That's it. That's all there is. Now the device is configured, via the +shell no less. + +[Coding With configfs] + +Every object in configfs is a config_item. A config_item reflects an +object in the subsystem. It has attributes that match values on that +object. configfs handles the filesystem representation of that object +and its attributes, allowing the subsystem to ignore all but the +basic show/store interaction. + +Items are created and destroyed inside a config_group. A group is a +collection of items that share the same attributes and operations. +Items are created by mkdir(2) and removed by rmdir(2), but configfs +handles that. The group has a set of operations to perform these tasks + +A subsystem is the top level of a client module. During initialization, +the client module registers the subsystem with configfs, the subsystem +appears as a directory at the top of the configfs filesystem. A +subsystem is also a config_group, and can do everything a config_group +can. + +[struct config_item] + + struct config_item { + char *ci_name; + char ci_namebuf[UOBJ_NAME_LEN]; + struct kref ci_kref; + struct list_head ci_entry; + struct config_item *ci_parent; + struct config_group *ci_group; + struct config_item_type *ci_type; + struct dentry *ci_dentry; + }; + + void config_item_init(struct config_item *); + void config_item_init_type_name(struct config_item *, + const char *name, + struct config_item_type *type); + struct config_item *config_item_get(struct config_item *); + void config_item_put(struct config_item *); + +Generally, struct config_item is embedded in a container structure, a +structure that actually represents what the subsystem is doing. The +config_item portion of that structure is how the object interacts with +configfs. + +Whether statically defined in a source file or created by a parent +config_group, a config_item must have one of the _init() functions +called on it. This initializes the reference count and sets up the +appropriate fields. + +All users of a config_item should have a reference on it via +config_item_get(), and drop the reference when they are done via +config_item_put(). + +By itself, a config_item cannot do much more than appear in configfs. +Usually a subsystem wants the item to display and/or store attributes, +among other things. For that, it needs a type. + +[struct config_item_type] + + struct configfs_item_operations { + void (*release)(struct config_item *); + ssize_t (*show_attribute)(struct config_item *, + struct configfs_attribute *, + char *); + ssize_t (*store_attribute)(struct config_item *, + struct configfs_attribute *, + const char *, size_t); + int (*allow_link)(struct config_item *src, + struct config_item *target); + int (*drop_link)(struct config_item *src, + struct config_item *target); + }; + + struct config_item_type { + struct module *ct_owner; + struct configfs_item_operations *ct_item_ops; + struct configfs_group_operations *ct_group_ops; + struct configfs_attribute **ct_attrs; + }; + +The most basic function of a config_item_type is to define what +operations can be performed on a config_item. All items that have been +allocated dynamically will need to provide the ct_item_ops->release() +method. This method is called when the config_item's reference count +reaches zero. Items that wish to display an attribute need to provide +the ct_item_ops->show_attribute() method. Similarly, storing a new +attribute value uses the store_attribute() method. + +[struct configfs_attribute] + + struct configfs_attribute { + char *ca_name; + struct module *ca_owner; + mode_t ca_mode; + }; + +When a config_item wants an attribute to appear as a file in the item's +configfs directory, it must define a configfs_attribute describing it. +It then adds the attribute to the NULL-terminated array +config_item_type->ct_attrs. When the item appears in configfs, the +attribute file will appear with the configfs_attribute->ca_name +filename. configfs_attribute->ca_mode specifies the file permissions. + +If an attribute is readable and the config_item provides a +ct_item_ops->show_attribute() method, that method will be called +whenever userspace asks for a read(2) on the attribute. The converse +will happen for write(2). + +[struct config_group] + +A config_item cannot live in a vacuum. The only way one can be created +is via mkdir(2) on a config_group. This will trigger creation of a +child item. + + struct config_group { + struct config_item cg_item; + struct list_head cg_children; + struct configfs_subsystem *cg_subsys; + struct config_group **default_groups; + }; + + void config_group_init(struct config_group *group); + void config_group_init_type_name(struct config_group *group, + const char *name, + struct config_item_type *type); + + +The config_group structure contains a config_item. Properly configuring +that item means that a group can behave as an item in its own right. +However, it can do more: it can create child items or groups. This is +accomplished via the group operations specified on the group's +config_item_type. + + struct configfs_group_operations { + struct config_item *(*make_item)(struct config_group *group, + const char *name); + struct config_group *(*make_group)(struct config_group *group, + const char *name); + int (*commit_item)(struct config_item *item); + void (*disconnect_notify)(struct config_group *group, + struct config_item *item); + void (*drop_item)(struct config_group *group, + struct config_item *item); + }; + +A group creates child items by providing the +ct_group_ops->make_item() method. If provided, this method is called from mkdir(2) in the group's directory. The subsystem allocates a new +config_item (or more likely, its container structure), initializes it, +and returns it to configfs. Configfs will then populate the filesystem +tree to reflect the new item. + +If the subsystem wants the child to be a group itself, the subsystem +provides ct_group_ops->make_group(). Everything else behaves the same, +using the group _init() functions on the group. + +Finally, when userspace calls rmdir(2) on the item or group, +ct_group_ops->drop_item() is called. As a config_group is also a +config_item, it is not necessary for a separate drop_group() method. +The subsystem must config_item_put() the reference that was initialized +upon item allocation. If a subsystem has no work to do, it may omit +the ct_group_ops->drop_item() method, and configfs will call +config_item_put() on the item on behalf of the subsystem. + +IMPORTANT: drop_item() is void, and as such cannot fail. When rmdir(2) +is called, configfs WILL remove the item from the filesystem tree +(assuming that it has no children to keep it busy). The subsystem is +responsible for responding to this. If the subsystem has references to +the item in other threads, the memory is safe. It may take some time +for the item to actually disappear from the subsystem's usage. But it +is gone from configfs. + +When drop_item() is called, the item's linkage has already been torn +down. It no longer has a reference on its parent and has no place in +the item hierarchy. If a client needs to do some cleanup before this +teardown happens, the subsystem can implement the +ct_group_ops->disconnect_notify() method. The method is called after +configfs has removed the item from the filesystem view but before the +item is removed from its parent group. Like drop_item(), +disconnect_notify() is void and cannot fail. Client subsystems should +not drop any references here, as they still must do it in drop_item(). + +A config_group cannot be removed while it still has child items. This +is implemented in the configfs rmdir(2) code. ->drop_item() will not be +called, as the item has not been dropped. rmdir(2) will fail, as the +directory is not empty. + +[struct configfs_subsystem] + +A subsystem must register itself, usually at module_init time. This +tells configfs to make the subsystem appear in the file tree. + + struct configfs_subsystem { + struct config_group su_group; + struct mutex su_mutex; + }; + + int configfs_register_subsystem(struct configfs_subsystem *subsys); + void configfs_unregister_subsystem(struct configfs_subsystem *subsys); + + A subsystem consists of a toplevel config_group and a mutex. +The group is where child config_items are created. For a subsystem, +this group is usually defined statically. Before calling +configfs_register_subsystem(), the subsystem must have initialized the +group via the usual group _init() functions, and it must also have +initialized the mutex. + When the register call returns, the subsystem is live, and it +will be visible via configfs. At that point, mkdir(2) can be called and +the subsystem must be ready for it. + +[An Example] + +The best example of these basic concepts is the simple_children +subsystem/group and the simple_child item in configfs_example_explicit.c +and configfs_example_macros.c. It shows a trivial object displaying and +storing an attribute, and a simple group creating and destroying these +children. + +The only difference between configfs_example_explicit.c and +configfs_example_macros.c is how the attributes of the childless item +are defined. The childless item has extended attributes, each with +their own show()/store() operation. This follows a convention commonly +used in sysfs. configfs_example_explicit.c creates these attributes +by explicitly defining the structures involved. Conversely +configfs_example_macros.c uses some convenience macros from configfs.h +to define the attributes. These macros are similar to their sysfs +counterparts. + +[Hierarchy Navigation and the Subsystem Mutex] + +There is an extra bonus that configfs provides. The config_groups and +config_items are arranged in a hierarchy due to the fact that they +appear in a filesystem. A subsystem is NEVER to touch the filesystem +parts, but the subsystem might be interested in this hierarchy. For +this reason, the hierarchy is mirrored via the config_group->cg_children +and config_item->ci_parent structure members. + +A subsystem can navigate the cg_children list and the ci_parent pointer +to see the tree created by the subsystem. This can race with configfs' +management of the hierarchy, so configfs uses the subsystem mutex to +protect modifications. Whenever a subsystem wants to navigate the +hierarchy, it must do so under the protection of the subsystem +mutex. + +A subsystem will be prevented from acquiring the mutex while a newly +allocated item has not been linked into this hierarchy. Similarly, it +will not be able to acquire the mutex while a dropping item has not +yet been unlinked. This means that an item's ci_parent pointer will +never be NULL while the item is in configfs, and that an item will only +be in its parent's cg_children list for the same duration. This allows +a subsystem to trust ci_parent and cg_children while they hold the +mutex. + +[Item Aggregation Via symlink(2)] + +configfs provides a simple group via the group->item parent/child +relationship. Often, however, a larger environment requires aggregation +outside of the parent/child connection. This is implemented via +symlink(2). + +A config_item may provide the ct_item_ops->allow_link() and +ct_item_ops->drop_link() methods. If the ->allow_link() method exists, +symlink(2) may be called with the config_item as the source of the link. +These links are only allowed between configfs config_items. Any +symlink(2) attempt outside the configfs filesystem will be denied. + +When symlink(2) is called, the source config_item's ->allow_link() +method is called with itself and a target item. If the source item +allows linking to target item, it returns 0. A source item may wish to +reject a link if it only wants links to a certain type of object (say, +in its own subsystem). + +When unlink(2) is called on the symbolic link, the source item is +notified via the ->drop_link() method. Like the ->drop_item() method, +this is a void function and cannot return failure. The subsystem is +responsible for responding to the change. + +A config_item cannot be removed while it links to any other item, nor +can it be removed while an item links to it. Dangling symlinks are not +allowed in configfs. + +[Automatically Created Subgroups] + +A new config_group may want to have two types of child config_items. +While this could be codified by magic names in ->make_item(), it is much +more explicit to have a method whereby userspace sees this divergence. + +Rather than have a group where some items behave differently than +others, configfs provides a method whereby one or many subgroups are +automatically created inside the parent at its creation. Thus, +mkdir("parent") results in "parent", "parent/subgroup1", up through +"parent/subgroupN". Items of type 1 can now be created in +"parent/subgroup1", and items of type N can be created in +"parent/subgroupN". + +These automatic subgroups, or default groups, do not preclude other +children of the parent group. If ct_group_ops->make_group() exists, +other child groups can be created on the parent group directly. + +A configfs subsystem specifies default groups by filling in the +NULL-terminated array default_groups on the config_group structure. +Each group in that array is populated in the configfs tree at the same +time as the parent group. Similarly, they are removed at the same time +as the parent. No extra notification is provided. When a ->drop_item() +method call notifies the subsystem the parent group is going away, it +also means every default group child associated with that parent group. + +As a consequence of this, default_groups cannot be removed directly via +rmdir(2). They also are not considered when rmdir(2) on the parent +group is checking for children. + +[Dependent Subsystems] + +Sometimes other drivers depend on particular configfs items. For +example, ocfs2 mounts depend on a heartbeat region item. If that +region item is removed with rmdir(2), the ocfs2 mount must BUG or go +readonly. Not happy. + +configfs provides two additional API calls: configfs_depend_item() and +configfs_undepend_item(). A client driver can call +configfs_depend_item() on an existing item to tell configfs that it is +depended on. configfs will then return -EBUSY from rmdir(2) for that +item. When the item is no longer depended on, the client driver calls +configfs_undepend_item() on it. + +These API cannot be called underneath any configfs callbacks, as +they will conflict. They can block and allocate. A client driver +probably shouldn't calling them of its own gumption. Rather it should +be providing an API that external subsystems call. + +How does this work? Imagine the ocfs2 mount process. When it mounts, +it asks for a heartbeat region item. This is done via a call into the +heartbeat code. Inside the heartbeat code, the region item is looked +up. Here, the heartbeat code calls configfs_depend_item(). If it +succeeds, then heartbeat knows the region is safe to give to ocfs2. +If it fails, it was being torn down anyway, and heartbeat can gracefully +pass up an error. + +[Committable Items] + +NOTE: Committable items are currently unimplemented. + +Some config_items cannot have a valid initial state. That is, no +default values can be specified for the item's attributes such that the +item can do its work. Userspace must configure one or more attributes, +after which the subsystem can start whatever entity this item +represents. + +Consider the FakeNBD device from above. Without a target address *and* +a target device, the subsystem has no idea what block device to import. +The simple example assumes that the subsystem merely waits until all the +appropriate attributes are configured, and then connects. This will, +indeed, work, but now every attribute store must check if the attributes +are initialized. Every attribute store must fire off the connection if +that condition is met. + +Far better would be an explicit action notifying the subsystem that the +config_item is ready to go. More importantly, an explicit action allows +the subsystem to provide feedback as to whether the attributes are +initialized in a way that makes sense. configfs provides this as +committable items. + +configfs still uses only normal filesystem operations. An item is +committed via rename(2). The item is moved from a directory where it +can be modified to a directory where it cannot. + +Any group that provides the ct_group_ops->commit_item() method has +committable items. When this group appears in configfs, mkdir(2) will +not work directly in the group. Instead, the group will have two +subdirectories: "live" and "pending". The "live" directory does not +support mkdir(2) or rmdir(2) either. It only allows rename(2). The +"pending" directory does allow mkdir(2) and rmdir(2). An item is +created in the "pending" directory. Its attributes can be modified at +will. Userspace commits the item by renaming it into the "live" +directory. At this point, the subsystem receives the ->commit_item() +callback. If all required attributes are filled to satisfaction, the +method returns zero and the item is moved to the "live" directory. + +As rmdir(2) does not work in the "live" directory, an item must be +shutdown, or "uncommitted". Again, this is done via rename(2), this +time from the "live" directory back to the "pending" one. The subsystem +is notified by the ct_group_ops->uncommit_object() method. + + diff --git a/Documentation/filesystems/configfs/configfs_example_explicit.c b/Documentation/filesystems/configfs/configfs_example_explicit.c new file mode 100644 index 00000000..1420233d --- /dev/null +++ b/Documentation/filesystems/configfs/configfs_example_explicit.c @@ -0,0 +1,483 @@ +/* + * vim: noexpandtab ts=8 sts=0 sw=8: + * + * configfs_example_explicit.c - This file is a demonstration module + * containing a number of configfs subsystems. It explicitly defines + * each structure without using the helper macros defined in + * configfs.h. + * + * This program is free software; you can redistribute it and/or + * modify it under the terms of the GNU General Public + * License as published by the Free Software Foundation; either + * version 2 of the License, or (at your option) any later version. + * + * This program is distributed in the hope that it will be useful, + * but WITHOUT ANY WARRANTY; without even the implied warranty of + * MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the GNU + * General Public License for more details. + * + * You should have received a copy of the GNU General Public + * License along with this program; if not, write to the + * Free Software Foundation, Inc., 59 Temple Place - Suite 330, + * Boston, MA 021110-1307, USA. + * + * Based on sysfs: + * sysfs is Copyright (C) 2001, 2002, 2003 Patrick Mochel + * + * configfs Copyright (C) 2005 Oracle. All rights reserved. + */ + +#include <linux/init.h> +#include <linux/module.h> +#include <linux/slab.h> + +#include <linux/configfs.h> + + + +/* + * 01-childless + * + * This first example is a childless subsystem. It cannot create + * any config_items. It just has attributes. + * + * Note that we are enclosing the configfs_subsystem inside a container. + * This is not necessary if a subsystem has no attributes directly + * on the subsystem. See the next example, 02-simple-children, for + * such a subsystem. + */ + +struct childless { + struct configfs_subsystem subsys; + int showme; + int storeme; +}; + +struct childless_attribute { + struct configfs_attribute attr; + ssize_t (*show)(struct childless *, char *); + ssize_t (*store)(struct childless *, const char *, size_t); +}; + +static inline struct childless *to_childless(struct config_item *item) +{ + return item ? container_of(to_configfs_subsystem(to_config_group(item)), struct childless, subsys) : NULL; +} + +static ssize_t childless_showme_read(struct childless *childless, + char *page) +{ + ssize_t pos; + + pos = sprintf(page, "%d\n", childless->showme); + childless->showme++; + + return pos; +} + +static ssize_t childless_storeme_read(struct childless *childless, + char *page) +{ + return sprintf(page, "%d\n", childless->storeme); +} + +static ssize_t childless_storeme_write(struct childless *childless, + const char *page, + size_t count) +{ + unsigned long tmp; + char *p = (char *) page; + + tmp = simple_strtoul(p, &p, 10); + if ((*p != '\0') && (*p != '\n')) + return -EINVAL; + + if (tmp > INT_MAX) + return -ERANGE; + + childless->storeme = tmp; + + return count; +} + +static ssize_t childless_description_read(struct childless *childless, + char *page) +{ + return sprintf(page, +"[01-childless]\n" +"\n" +"The childless subsystem is the simplest possible subsystem in\n" +"configfs. It does not support the creation of child config_items.\n" +"It only has a few attributes. In fact, it isn't much different\n" +"than a directory in /proc.\n"); +} + +static struct childless_attribute childless_attr_showme = { + .attr = { .ca_owner = THIS_MODULE, .ca_name = "showme", .ca_mode = S_IRUGO }, + .show = childless_showme_read, +}; +static struct childless_attribute childless_attr_storeme = { + .attr = { .ca_owner = THIS_MODULE, .ca_name = "storeme", .ca_mode = S_IRUGO | S_IWUSR }, + .show = childless_storeme_read, + .store = childless_storeme_write, +}; +static struct childless_attribute childless_attr_description = { + .attr = { .ca_owner = THIS_MODULE, .ca_name = "description", .ca_mode = S_IRUGO }, + .show = childless_description_read, +}; + +static struct configfs_attribute *childless_attrs[] = { + &childless_attr_showme.attr, + &childless_attr_storeme.attr, + &childless_attr_description.attr, + NULL, +}; + +static ssize_t childless_attr_show(struct config_item *item, + struct configfs_attribute *attr, + char *page) +{ + struct childless *childless = to_childless(item); + struct childless_attribute *childless_attr = + container_of(attr, struct childless_attribute, attr); + ssize_t ret = 0; + + if (childless_attr->show) + ret = childless_attr->show(childless, page); + return ret; +} + +static ssize_t childless_attr_store(struct config_item *item, + struct configfs_attribute *attr, + const char *page, size_t count) +{ + struct childless *childless = to_childless(item); + struct childless_attribute *childless_attr = + container_of(attr, struct childless_attribute, attr); + ssize_t ret = -EINVAL; + + if (childless_attr->store) + ret = childless_attr->store(childless, page, count); + return ret; +} + +static struct configfs_item_operations childless_item_ops = { + .show_attribute = childless_attr_show, + .store_attribute = childless_attr_store, +}; + +static struct config_item_type childless_type = { + .ct_item_ops = &childless_item_ops, + .ct_attrs = childless_attrs, + .ct_owner = THIS_MODULE, +}; + +static struct childless childless_subsys = { + .subsys = { + .su_group = { + .cg_item = { + .ci_namebuf = "01-childless", + .ci_type = &childless_type, + }, + }, + }, +}; + + +/* ----------------------------------------------------------------- */ + +/* + * 02-simple-children + * + * This example merely has a simple one-attribute child. Note that + * there is no extra attribute structure, as the child's attribute is + * known from the get-go. Also, there is no container for the + * subsystem, as it has no attributes of its own. + */ + +struct simple_child { + struct config_item item; + int storeme; +}; + +static inline struct simple_child *to_simple_child(struct config_item *item) +{ + return item ? container_of(item, struct simple_child, item) : NULL; +} + +static struct configfs_attribute simple_child_attr_storeme = { + .ca_owner = THIS_MODULE, + .ca_name = "storeme", + .ca_mode = S_IRUGO | S_IWUSR, +}; + +static struct configfs_attribute *simple_child_attrs[] = { + &simple_child_attr_storeme, + NULL, +}; + +static ssize_t simple_child_attr_show(struct config_item *item, + struct configfs_attribute *attr, + char *page) +{ + ssize_t count; + struct simple_child *simple_child = to_simple_child(item); + + count = sprintf(page, "%d\n", simple_child->storeme); + + return count; +} + +static ssize_t simple_child_attr_store(struct config_item *item, + struct configfs_attribute *attr, + const char *page, size_t count) +{ + struct simple_child *simple_child = to_simple_child(item); + unsigned long tmp; + char *p = (char *) page; + + tmp = simple_strtoul(p, &p, 10); + if (!p || (*p && (*p != '\n'))) + return -EINVAL; + + if (tmp > INT_MAX) + return -ERANGE; + + simple_child->storeme = tmp; + + return count; +} + +static void simple_child_release(struct config_item *item) +{ + kfree(to_simple_child(item)); +} + +static struct configfs_item_operations simple_child_item_ops = { + .release = simple_child_release, + .show_attribute = simple_child_attr_show, + .store_attribute = simple_child_attr_store, +}; + +static struct config_item_type simple_child_type = { + .ct_item_ops = &simple_child_item_ops, + .ct_attrs = simple_child_attrs, + .ct_owner = THIS_MODULE, +}; + + +struct simple_children { + struct config_group group; +}; + +static inline struct simple_children *to_simple_children(struct config_item *item) +{ + return item ? container_of(to_config_group(item), struct simple_children, group) : NULL; +} + +static struct config_item *simple_children_make_item(struct config_group *group, const char *name) +{ + struct simple_child *simple_child; + + simple_child = kzalloc(sizeof(struct simple_child), GFP_KERNEL); + if (!simple_child) + return ERR_PTR(-ENOMEM); + + config_item_init_type_name(&simple_child->item, name, + &simple_child_type); + + simple_child->storeme = 0; + + return &simple_child->item; +} + +static struct configfs_attribute simple_children_attr_description = { + .ca_owner = THIS_MODULE, + .ca_name = "description", + .ca_mode = S_IRUGO, +}; + +static struct configfs_attribute *simple_children_attrs[] = { + &simple_children_attr_description, + NULL, +}; + +static ssize_t simple_children_attr_show(struct config_item *item, + struct configfs_attribute *attr, + char *page) +{ + return sprintf(page, +"[02-simple-children]\n" +"\n" +"This subsystem allows the creation of child config_items. These\n" +"items have only one attribute that is readable and writeable.\n"); +} + +static void simple_children_release(struct config_item *item) +{ + kfree(to_simple_children(item)); +} + +static struct configfs_item_operations simple_children_item_ops = { + .release = simple_children_release, + .show_attribute = simple_children_attr_show, +}; + +/* + * Note that, since no extra work is required on ->drop_item(), + * no ->drop_item() is provided. + */ +static struct configfs_group_operations simple_children_group_ops = { + .make_item = simple_children_make_item, +}; + +static struct config_item_type simple_children_type = { + .ct_item_ops = &simple_children_item_ops, + .ct_group_ops = &simple_children_group_ops, + .ct_attrs = simple_children_attrs, + .ct_owner = THIS_MODULE, +}; + +static struct configfs_subsystem simple_children_subsys = { + .su_group = { + .cg_item = { + .ci_namebuf = "02-simple-children", + .ci_type = &simple_children_type, + }, + }, +}; + + +/* ----------------------------------------------------------------- */ + +/* + * 03-group-children + * + * This example reuses the simple_children group from above. However, + * the simple_children group is not the subsystem itself, it is a + * child of the subsystem. Creation of a group in the subsystem creates + * a new simple_children group. That group can then have simple_child + * children of its own. + */ + +static struct config_group *group_children_make_group(struct config_group *group, const char *name) +{ + struct simple_children *simple_children; + + simple_children = kzalloc(sizeof(struct simple_children), + GFP_KERNEL); + if (!simple_children) + return ERR_PTR(-ENOMEM); + + config_group_init_type_name(&simple_children->group, name, + &simple_children_type); + + return &simple_children->group; +} + +static struct configfs_attribute group_children_attr_description = { + .ca_owner = THIS_MODULE, + .ca_name = "description", + .ca_mode = S_IRUGO, +}; + +static struct configfs_attribute *group_children_attrs[] = { + &group_children_attr_description, + NULL, +}; + +static ssize_t group_children_attr_show(struct config_item *item, + struct configfs_attribute *attr, + char *page) +{ + return sprintf(page, +"[03-group-children]\n" +"\n" +"This subsystem allows the creation of child config_groups. These\n" +"groups are like the subsystem simple-children.\n"); +} + +static struct configfs_item_operations group_children_item_ops = { + .show_attribute = group_children_attr_show, +}; + +/* + * Note that, since no extra work is required on ->drop_item(), + * no ->drop_item() is provided. + */ +static struct configfs_group_operations group_children_group_ops = { + .make_group = group_children_make_group, +}; + +static struct config_item_type group_children_type = { + .ct_item_ops = &group_children_item_ops, + .ct_group_ops = &group_children_group_ops, + .ct_attrs = group_children_attrs, + .ct_owner = THIS_MODULE, +}; + +static struct configfs_subsystem group_children_subsys = { + .su_group = { + .cg_item = { + .ci_namebuf = "03-group-children", + .ci_type = &group_children_type, + }, + }, +}; + +/* ----------------------------------------------------------------- */ + +/* + * We're now done with our subsystem definitions. + * For convenience in this module, here's a list of them all. It + * allows the init function to easily register them. Most modules + * will only have one subsystem, and will only call register_subsystem + * on it directly. + */ +static struct configfs_subsystem *example_subsys[] = { + &childless_subsys.subsys, + &simple_children_subsys, + &group_children_subsys, + NULL, +}; + +static int __init configfs_example_init(void) +{ + int ret; + int i; + struct configfs_subsystem *subsys; + + for (i = 0; example_subsys[i]; i++) { + subsys = example_subsys[i]; + + config_group_init(&subsys->su_group); + mutex_init(&subsys->su_mutex); + ret = configfs_register_subsystem(subsys); + if (ret) { + printk(KERN_ERR "Error %d while registering subsystem %s\n", + ret, + subsys->su_group.cg_item.ci_namebuf); + goto out_unregister; + } + } + + return 0; + +out_unregister: + for (i--; i >= 0; i--) + configfs_unregister_subsystem(example_subsys[i]); + + return ret; +} + +static void __exit configfs_example_exit(void) +{ + int i; + + for (i = 0; example_subsys[i]; i++) + configfs_unregister_subsystem(example_subsys[i]); +} + +module_init(configfs_example_init); +module_exit(configfs_example_exit); +MODULE_LICENSE("GPL"); diff --git a/Documentation/filesystems/configfs/configfs_example_macros.c b/Documentation/filesystems/configfs/configfs_example_macros.c new file mode 100644 index 00000000..327dfbc6 --- /dev/null +++ b/Documentation/filesystems/configfs/configfs_example_macros.c @@ -0,0 +1,446 @@ +/* + * vim: noexpandtab ts=8 sts=0 sw=8: + * + * configfs_example_macros.c - This file is a demonstration module + * containing a number of configfs subsystems. It uses the helper + * macros defined by configfs.h + * + * This program is free software; you can redistribute it and/or + * modify it under the terms of the GNU General Public + * License as published by the Free Software Foundation; either + * version 2 of the License, or (at your option) any later version. + * + * This program is distributed in the hope that it will be useful, + * but WITHOUT ANY WARRANTY; without even the implied warranty of + * MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the GNU + * General Public License for more details. + * + * You should have received a copy of the GNU General Public + * License along with this program; if not, write to the + * Free Software Foundation, Inc., 59 Temple Place - Suite 330, + * Boston, MA 021110-1307, USA. + * + * Based on sysfs: + * sysfs is Copyright (C) 2001, 2002, 2003 Patrick Mochel + * + * configfs Copyright (C) 2005 Oracle. All rights reserved. + */ + +#include <linux/init.h> +#include <linux/module.h> +#include <linux/slab.h> + +#include <linux/configfs.h> + + + +/* + * 01-childless + * + * This first example is a childless subsystem. It cannot create + * any config_items. It just has attributes. + * + * Note that we are enclosing the configfs_subsystem inside a container. + * This is not necessary if a subsystem has no attributes directly + * on the subsystem. See the next example, 02-simple-children, for + * such a subsystem. + */ + +struct childless { + struct configfs_subsystem subsys; + int showme; + int storeme; +}; + +static inline struct childless *to_childless(struct config_item *item) +{ + return item ? container_of(to_configfs_subsystem(to_config_group(item)), struct childless, subsys) : NULL; +} + +CONFIGFS_ATTR_STRUCT(childless); +#define CHILDLESS_ATTR(_name, _mode, _show, _store) \ +struct childless_attribute childless_attr_##_name = __CONFIGFS_ATTR(_name, _mode, _show, _store) +#define CHILDLESS_ATTR_RO(_name, _show) \ +struct childless_attribute childless_attr_##_name = __CONFIGFS_ATTR_RO(_name, _show); + +static ssize_t childless_showme_read(struct childless *childless, + char *page) +{ + ssize_t pos; + + pos = sprintf(page, "%d\n", childless->showme); + childless->showme++; + + return pos; +} + +static ssize_t childless_storeme_read(struct childless *childless, + char *page) +{ + return sprintf(page, "%d\n", childless->storeme); +} + +static ssize_t childless_storeme_write(struct childless *childless, + const char *page, + size_t count) +{ + unsigned long tmp; + char *p = (char *) page; + + tmp = simple_strtoul(p, &p, 10); + if (!p || (*p && (*p != '\n'))) + return -EINVAL; + + if (tmp > INT_MAX) + return -ERANGE; + + childless->storeme = tmp; + + return count; +} + +static ssize_t childless_description_read(struct childless *childless, + char *page) +{ + return sprintf(page, +"[01-childless]\n" +"\n" +"The childless subsystem is the simplest possible subsystem in\n" +"configfs. It does not support the creation of child config_items.\n" +"It only has a few attributes. In fact, it isn't much different\n" +"than a directory in /proc.\n"); +} + +CHILDLESS_ATTR_RO(showme, childless_showme_read); +CHILDLESS_ATTR(storeme, S_IRUGO | S_IWUSR, childless_storeme_read, + childless_storeme_write); +CHILDLESS_ATTR_RO(description, childless_description_read); + +static struct configfs_attribute *childless_attrs[] = { + &childless_attr_showme.attr, + &childless_attr_storeme.attr, + &childless_attr_description.attr, + NULL, +}; + +CONFIGFS_ATTR_OPS(childless); +static struct configfs_item_operations childless_item_ops = { + .show_attribute = childless_attr_show, + .store_attribute = childless_attr_store, +}; + +static struct config_item_type childless_type = { + .ct_item_ops = &childless_item_ops, + .ct_attrs = childless_attrs, + .ct_owner = THIS_MODULE, +}; + +static struct childless childless_subsys = { + .subsys = { + .su_group = { + .cg_item = { + .ci_namebuf = "01-childless", + .ci_type = &childless_type, + }, + }, + }, +}; + + +/* ----------------------------------------------------------------- */ + +/* + * 02-simple-children + * + * This example merely has a simple one-attribute child. Note that + * there is no extra attribute structure, as the child's attribute is + * known from the get-go. Also, there is no container for the + * subsystem, as it has no attributes of its own. + */ + +struct simple_child { + struct config_item item; + int storeme; +}; + +static inline struct simple_child *to_simple_child(struct config_item *item) +{ + return item ? container_of(item, struct simple_child, item) : NULL; +} + +static struct configfs_attribute simple_child_attr_storeme = { + .ca_owner = THIS_MODULE, + .ca_name = "storeme", + .ca_mode = S_IRUGO | S_IWUSR, +}; + +static struct configfs_attribute *simple_child_attrs[] = { + &simple_child_attr_storeme, + NULL, +}; + +static ssize_t simple_child_attr_show(struct config_item *item, + struct configfs_attribute *attr, + char *page) +{ + ssize_t count; + struct simple_child *simple_child = to_simple_child(item); + + count = sprintf(page, "%d\n", simple_child->storeme); + + return count; +} + +static ssize_t simple_child_attr_store(struct config_item *item, + struct configfs_attribute *attr, + const char *page, size_t count) +{ + struct simple_child *simple_child = to_simple_child(item); + unsigned long tmp; + char *p = (char *) page; + + tmp = simple_strtoul(p, &p, 10); + if (!p || (*p && (*p != '\n'))) + return -EINVAL; + + if (tmp > INT_MAX) + return -ERANGE; + + simple_child->storeme = tmp; + + return count; +} + +static void simple_child_release(struct config_item *item) +{ + kfree(to_simple_child(item)); +} + +static struct configfs_item_operations simple_child_item_ops = { + .release = simple_child_release, + .show_attribute = simple_child_attr_show, + .store_attribute = simple_child_attr_store, +}; + +static struct config_item_type simple_child_type = { + .ct_item_ops = &simple_child_item_ops, + .ct_attrs = simple_child_attrs, + .ct_owner = THIS_MODULE, +}; + + +struct simple_children { + struct config_group group; +}; + +static inline struct simple_children *to_simple_children(struct config_item *item) +{ + return item ? container_of(to_config_group(item), struct simple_children, group) : NULL; +} + +static struct config_item *simple_children_make_item(struct config_group *group, const char *name) +{ + struct simple_child *simple_child; + + simple_child = kzalloc(sizeof(struct simple_child), GFP_KERNEL); + if (!simple_child) + return ERR_PTR(-ENOMEM); + + config_item_init_type_name(&simple_child->item, name, + &simple_child_type); + + simple_child->storeme = 0; + + return &simple_child->item; +} + +static struct configfs_attribute simple_children_attr_description = { + .ca_owner = THIS_MODULE, + .ca_name = "description", + .ca_mode = S_IRUGO, +}; + +static struct configfs_attribute *simple_children_attrs[] = { + &simple_children_attr_description, + NULL, +}; + +static ssize_t simple_children_attr_show(struct config_item *item, + struct configfs_attribute *attr, + char *page) +{ + return sprintf(page, +"[02-simple-children]\n" +"\n" +"This subsystem allows the creation of child config_items. These\n" +"items have only one attribute that is readable and writeable.\n"); +} + +static void simple_children_release(struct config_item *item) +{ + kfree(to_simple_children(item)); +} + +static struct configfs_item_operations simple_children_item_ops = { + .release = simple_children_release, + .show_attribute = simple_children_attr_show, +}; + +/* + * Note that, since no extra work is required on ->drop_item(), + * no ->drop_item() is provided. + */ +static struct configfs_group_operations simple_children_group_ops = { + .make_item = simple_children_make_item, +}; + +static struct config_item_type simple_children_type = { + .ct_item_ops = &simple_children_item_ops, + .ct_group_ops = &simple_children_group_ops, + .ct_attrs = simple_children_attrs, + .ct_owner = THIS_MODULE, +}; + +static struct configfs_subsystem simple_children_subsys = { + .su_group = { + .cg_item = { + .ci_namebuf = "02-simple-children", + .ci_type = &simple_children_type, + }, + }, +}; + + +/* ----------------------------------------------------------------- */ + +/* + * 03-group-children + * + * This example reuses the simple_children group from above. However, + * the simple_children group is not the subsystem itself, it is a + * child of the subsystem. Creation of a group in the subsystem creates + * a new simple_children group. That group can then have simple_child + * children of its own. + */ + +static struct config_group *group_children_make_group(struct config_group *group, const char *name) +{ + struct simple_children *simple_children; + + simple_children = kzalloc(sizeof(struct simple_children), + GFP_KERNEL); + if (!simple_children) + return ERR_PTR(-ENOMEM); + + config_group_init_type_name(&simple_children->group, name, + &simple_children_type); + + return &simple_children->group; +} + +static struct configfs_attribute group_children_attr_description = { + .ca_owner = THIS_MODULE, + .ca_name = "description", + .ca_mode = S_IRUGO, +}; + +static struct configfs_attribute *group_children_attrs[] = { + &group_children_attr_description, + NULL, +}; + +static ssize_t group_children_attr_show(struct config_item *item, + struct configfs_attribute *attr, + char *page) +{ + return sprintf(page, +"[03-group-children]\n" +"\n" +"This subsystem allows the creation of child config_groups. These\n" +"groups are like the subsystem simple-children.\n"); +} + +static struct configfs_item_operations group_children_item_ops = { + .show_attribute = group_children_attr_show, +}; + +/* + * Note that, since no extra work is required on ->drop_item(), + * no ->drop_item() is provided. + */ +static struct configfs_group_operations group_children_group_ops = { + .make_group = group_children_make_group, +}; + +static struct config_item_type group_children_type = { + .ct_item_ops = &group_children_item_ops, + .ct_group_ops = &group_children_group_ops, + .ct_attrs = group_children_attrs, + .ct_owner = THIS_MODULE, +}; + +static struct configfs_subsystem group_children_subsys = { + .su_group = { + .cg_item = { + .ci_namebuf = "03-group-children", + .ci_type = &group_children_type, + }, + }, +}; + +/* ----------------------------------------------------------------- */ + +/* + * We're now done with our subsystem definitions. + * For convenience in this module, here's a list of them all. It + * allows the init function to easily register them. Most modules + * will only have one subsystem, and will only call register_subsystem + * on it directly. + */ +static struct configfs_subsystem *example_subsys[] = { + &childless_subsys.subsys, + &simple_children_subsys, + &group_children_subsys, + NULL, +}; + +static int __init configfs_example_init(void) +{ + int ret; + int i; + struct configfs_subsystem *subsys; + + for (i = 0; example_subsys[i]; i++) { + subsys = example_subsys[i]; + + config_group_init(&subsys->su_group); + mutex_init(&subsys->su_mutex); + ret = configfs_register_subsystem(subsys); + if (ret) { + printk(KERN_ERR "Error %d while registering subsystem %s\n", + ret, + subsys->su_group.cg_item.ci_namebuf); + goto out_unregister; + } + } + + return 0; + +out_unregister: + for (i--; i >= 0; i--) + configfs_unregister_subsystem(example_subsys[i]); + + return ret; +} + +static void __exit configfs_example_exit(void) +{ + int i; + + for (i = 0; example_subsys[i]; i++) + configfs_unregister_subsystem(example_subsys[i]); +} + +module_init(configfs_example_init); +module_exit(configfs_example_exit); +MODULE_LICENSE("GPL"); diff --git a/Documentation/filesystems/cramfs.txt b/Documentation/filesystems/cramfs.txt new file mode 100644 index 00000000..31f53f0a --- /dev/null +++ b/Documentation/filesystems/cramfs.txt @@ -0,0 +1,76 @@ + + Cramfs - cram a filesystem onto a small ROM + +cramfs is designed to be simple and small, and to compress things well. + +It uses the zlib routines to compress a file one page at a time, and +allows random page access. The meta-data is not compressed, but is +expressed in a very terse representation to make it use much less +diskspace than traditional filesystems. + +You can't write to a cramfs filesystem (making it compressible and +compact also makes it _very_ hard to update on-the-fly), so you have to +create the disk image with the "mkcramfs" utility. + + +Usage Notes +----------- + +File sizes are limited to less than 16MB. + +Maximum filesystem size is a little over 256MB. (The last file on the +filesystem is allowed to extend past 256MB.) + +Only the low 8 bits of gid are stored. The current version of +mkcramfs simply truncates to 8 bits, which is a potential security +issue. + +Hard links are supported, but hard linked files +will still have a link count of 1 in the cramfs image. + +Cramfs directories have no `.' or `..' entries. Directories (like +every other file on cramfs) always have a link count of 1. (There's +no need to use -noleaf in `find', btw.) + +No timestamps are stored in a cramfs, so these default to the epoch +(1970 GMT). Recently-accessed files may have updated timestamps, but +the update lasts only as long as the inode is cached in memory, after +which the timestamp reverts to 1970, i.e. moves backwards in time. + +Currently, cramfs must be written and read with architectures of the +same endianness, and can be read only by kernels with PAGE_CACHE_SIZE +== 4096. At least the latter of these is a bug, but it hasn't been +decided what the best fix is. For the moment if you have larger pages +you can just change the #define in mkcramfs.c, so long as you don't +mind the filesystem becoming unreadable to future kernels. + + +For /usr/share/magic +-------------------- + +0 ulelong 0x28cd3d45 Linux cramfs offset 0 +>4 ulelong x size %d +>8 ulelong x flags 0x%x +>12 ulelong x future 0x%x +>16 string >\0 signature "%.16s" +>32 ulelong x fsid.crc 0x%x +>36 ulelong x fsid.edition %d +>40 ulelong x fsid.blocks %d +>44 ulelong x fsid.files %d +>48 string >\0 name "%.16s" +512 ulelong 0x28cd3d45 Linux cramfs offset 512 +>516 ulelong x size %d +>520 ulelong x flags 0x%x +>524 ulelong x future 0x%x +>528 string >\0 signature "%.16s" +>544 ulelong x fsid.crc 0x%x +>548 ulelong x fsid.edition %d +>552 ulelong x fsid.blocks %d +>556 ulelong x fsid.files %d +>560 string >\0 name "%.16s" + + +Hacker Notes +------------ + +See fs/cramfs/README for filesystem layout and implementation notes. diff --git a/Documentation/filesystems/debugfs.txt b/Documentation/filesystems/debugfs.txt new file mode 100644 index 00000000..ed52af60 --- /dev/null +++ b/Documentation/filesystems/debugfs.txt @@ -0,0 +1,158 @@ +Copyright 2009 Jonathan Corbet <corbet@lwn.net> + +Debugfs exists as a simple way for kernel developers to make information +available to user space. Unlike /proc, which is only meant for information +about a process, or sysfs, which has strict one-value-per-file rules, +debugfs has no rules at all. Developers can put any information they want +there. The debugfs filesystem is also intended to not serve as a stable +ABI to user space; in theory, there are no stability constraints placed on +files exported there. The real world is not always so simple, though [1]; +even debugfs interfaces are best designed with the idea that they will need +to be maintained forever. + +Debugfs is typically mounted with a command like: + + mount -t debugfs none /sys/kernel/debug + +(Or an equivalent /etc/fstab line). + +Note that the debugfs API is exported GPL-only to modules. + +Code using debugfs should include <linux/debugfs.h>. Then, the first order +of business will be to create at least one directory to hold a set of +debugfs files: + + struct dentry *debugfs_create_dir(const char *name, struct dentry *parent); + +This call, if successful, will make a directory called name underneath the +indicated parent directory. If parent is NULL, the directory will be +created in the debugfs root. On success, the return value is a struct +dentry pointer which can be used to create files in the directory (and to +clean it up at the end). A NULL return value indicates that something went +wrong. If ERR_PTR(-ENODEV) is returned, that is an indication that the +kernel has been built without debugfs support and none of the functions +described below will work. + +The most general way to create a file within a debugfs directory is with: + + struct dentry *debugfs_create_file(const char *name, mode_t mode, + struct dentry *parent, void *data, + const struct file_operations *fops); + +Here, name is the name of the file to create, mode describes the access +permissions the file should have, parent indicates the directory which +should hold the file, data will be stored in the i_private field of the +resulting inode structure, and fops is a set of file operations which +implement the file's behavior. At a minimum, the read() and/or write() +operations should be provided; others can be included as needed. Again, +the return value will be a dentry pointer to the created file, NULL for +error, or ERR_PTR(-ENODEV) if debugfs support is missing. + +In a number of cases, the creation of a set of file operations is not +actually necessary; the debugfs code provides a number of helper functions +for simple situations. Files containing a single integer value can be +created with any of: + + struct dentry *debugfs_create_u8(const char *name, mode_t mode, + struct dentry *parent, u8 *value); + struct dentry *debugfs_create_u16(const char *name, mode_t mode, + struct dentry *parent, u16 *value); + struct dentry *debugfs_create_u32(const char *name, mode_t mode, + struct dentry *parent, u32 *value); + struct dentry *debugfs_create_u64(const char *name, mode_t mode, + struct dentry *parent, u64 *value); + +These files support both reading and writing the given value; if a specific +file should not be written to, simply set the mode bits accordingly. The +values in these files are in decimal; if hexadecimal is more appropriate, +the following functions can be used instead: + + struct dentry *debugfs_create_x8(const char *name, mode_t mode, + struct dentry *parent, u8 *value); + struct dentry *debugfs_create_x16(const char *name, mode_t mode, + struct dentry *parent, u16 *value); + struct dentry *debugfs_create_x32(const char *name, mode_t mode, + struct dentry *parent, u32 *value); + +Note that there is no debugfs_create_x64(). + +These functions are useful as long as the developer knows the size of the +value to be exported. Some types can have different widths on different +architectures, though, complicating the situation somewhat. There is a +function meant to help out in one special case: + + struct dentry *debugfs_create_size_t(const char *name, mode_t mode, + struct dentry *parent, + size_t *value); + +As might be expected, this function will create a debugfs file to represent +a variable of type size_t. + +Boolean values can be placed in debugfs with: + + struct dentry *debugfs_create_bool(const char *name, mode_t mode, + struct dentry *parent, u32 *value); + +A read on the resulting file will yield either Y (for non-zero values) or +N, followed by a newline. If written to, it will accept either upper- or +lower-case values, or 1 or 0. Any other input will be silently ignored. + +Finally, a block of arbitrary binary data can be exported with: + + struct debugfs_blob_wrapper { + void *data; + unsigned long size; + }; + + struct dentry *debugfs_create_blob(const char *name, mode_t mode, + struct dentry *parent, + struct debugfs_blob_wrapper *blob); + +A read of this file will return the data pointed to by the +debugfs_blob_wrapper structure. Some drivers use "blobs" as a simple way +to return several lines of (static) formatted text output. This function +can be used to export binary information, but there does not appear to be +any code which does so in the mainline. Note that all files created with +debugfs_create_blob() are read-only. + +There are a couple of other directory-oriented helper functions: + + struct dentry *debugfs_rename(struct dentry *old_dir, + struct dentry *old_dentry, + struct dentry *new_dir, + const char *new_name); + + struct dentry *debugfs_create_symlink(const char *name, + struct dentry *parent, + const char *target); + +A call to debugfs_rename() will give a new name to an existing debugfs +file, possibly in a different directory. The new_name must not exist prior +to the call; the return value is old_dentry with updated information. +Symbolic links can be created with debugfs_create_symlink(). + +There is one important thing that all debugfs users must take into account: +there is no automatic cleanup of any directories created in debugfs. If a +module is unloaded without explicitly removing debugfs entries, the result +will be a lot of stale pointers and no end of highly antisocial behavior. +So all debugfs users - at least those which can be built as modules - must +be prepared to remove all files and directories they create there. A file +can be removed with: + + void debugfs_remove(struct dentry *dentry); + +The dentry value can be NULL, in which case nothing will be removed. + +Once upon a time, debugfs users were required to remember the dentry +pointer for every debugfs file they created so that all files could be +cleaned up. We live in more civilized times now, though, and debugfs users +can call: + + void debugfs_remove_recursive(struct dentry *dentry); + +If this function is passed a pointer for the dentry corresponding to the +top-level directory, the entire hierarchy below that directory will be +removed. + +Notes: + [1] http://lwn.net/Articles/309298/ diff --git a/Documentation/filesystems/devpts.txt b/Documentation/filesystems/devpts.txt new file mode 100644 index 00000000..68dffd87 --- /dev/null +++ b/Documentation/filesystems/devpts.txt @@ -0,0 +1,132 @@ + +To support containers, we now allow multiple instances of devpts filesystem, +such that indices of ptys allocated in one instance are independent of indices +allocated in other instances of devpts. + +To preserve backward compatibility, this support for multiple instances is +enabled only if: + + - CONFIG_DEVPTS_MULTIPLE_INSTANCES=y, and + - '-o newinstance' mount option is specified while mounting devpts + +IOW, devpts now supports both single-instance and multi-instance semantics. + +If CONFIG_DEVPTS_MULTIPLE_INSTANCES=n, there is no change in behavior and +this referred to as the "legacy" mode. In this mode, the new mount options +(-o newinstance and -o ptmxmode) will be ignored with a 'bogus option' message +on console. + +If CONFIG_DEVPTS_MULTIPLE_INSTANCES=y and devpts is mounted without the +'newinstance' option (as in current start-up scripts) the new mount binds +to the initial kernel mount of devpts. This mode is referred to as the +'single-instance' mode and the current, single-instance semantics are +preserved, i.e PTYs are common across the system. + +The only difference between this single-instance mode and the legacy mode +is the presence of new, '/dev/pts/ptmx' node with permissions 0000, which +can safely be ignored. + +If CONFIG_DEVPTS_MULTIPLE_INSTANCES=y and 'newinstance' option is specified, +the mount is considered to be in the multi-instance mode and a new instance +of the devpts fs is created. Any ptys created in this instance are independent +of ptys in other instances of devpts. Like in the single-instance mode, the +/dev/pts/ptmx node is present. To effectively use the multi-instance mode, +open of /dev/ptmx must be a redirected to '/dev/pts/ptmx' using a symlink or +bind-mount. + +Eg: A container startup script could do the following: + + $ chmod 0666 /dev/pts/ptmx + $ rm /dev/ptmx + $ ln -s pts/ptmx /dev/ptmx + $ ns_exec -cm /bin/bash + + # We are now in new container + + $ umount /dev/pts + $ mount -t devpts -o newinstance lxcpts /dev/pts + $ sshd -p 1234 + +where 'ns_exec -cm /bin/bash' calls clone() with CLONE_NEWNS flag and execs +/bin/bash in the child process. A pty created by the sshd is not visible in +the original mount of /dev/pts. + +User-space changes +------------------ + +In multi-instance mode (i.e '-o newinstance' mount option is specified at least +once), following user-space issues should be noted. + +1. If -o newinstance mount option is never used, /dev/pts/ptmx can be ignored + and no change is needed to system-startup scripts. + +2. To effectively use multi-instance mode (i.e -o newinstance is specified) + administrators or startup scripts should "redirect" open of /dev/ptmx to + /dev/pts/ptmx using either a bind mount or symlink. + + $ mount -t devpts -o newinstance devpts /dev/pts + + followed by either + + $ rm /dev/ptmx + $ ln -s pts/ptmx /dev/ptmx + $ chmod 666 /dev/pts/ptmx + or + $ mount -o bind /dev/pts/ptmx /dev/ptmx + +3. The '/dev/ptmx -> pts/ptmx' symlink is the preferred method since it + enables better error-reporting and treats both single-instance and + multi-instance mounts similarly. + + But this method requires that system-startup scripts set the mode of + /dev/pts/ptmx correctly (default mode is 0000). The scripts can set the + mode by, either + + - adding ptmxmode mount option to devpts entry in /etc/fstab, or + - using 'chmod 0666 /dev/pts/ptmx' + +4. If multi-instance mode mount is needed for containers, but the system + startup scripts have not yet been updated, container-startup scripts + should bind mount /dev/ptmx to /dev/pts/ptmx to avoid breaking single- + instance mounts. + + Or, in general, container-startup scripts should use: + + mount -t devpts -o newinstance -o ptmxmode=0666 devpts /dev/pts + if [ ! -L /dev/ptmx ]; then + mount -o bind /dev/pts/ptmx /dev/ptmx + fi + + When all devpts mounts are multi-instance, /dev/ptmx can permanently be + a symlink to pts/ptmx and the bind mount can be ignored. + +5. A multi-instance mount that is not accompanied by the /dev/ptmx to + /dev/pts/ptmx redirection would result in an unusable/unreachable pty. + + mount -t devpts -o newinstance lxcpts /dev/pts + + immediately followed by: + + open("/dev/ptmx") + + would create a pty, say /dev/pts/7, in the initial kernel mount. + But /dev/pts/7 would be invisible in the new mount. + +6. The permissions for /dev/pts/ptmx node should be specified when mounting + /dev/pts, using the '-o ptmxmode=%o' mount option (default is 0000). + + mount -t devpts -o newinstance -o ptmxmode=0644 devpts /dev/pts + + The permissions can be later be changed as usual with 'chmod'. + + chmod 666 /dev/pts/ptmx + +7. A mount of devpts without the 'newinstance' option results in binding to + initial kernel mount. This behavior while preserving legacy semantics, + does not provide strict isolation in a container environment. i.e by + mounting devpts without the 'newinstance' option, a container could + get visibility into the 'host' or root container's devpts. + + To workaround this and have strict isolation, all mounts of devpts, + including the mount in the root container, should use the newinstance + option. diff --git a/Documentation/filesystems/directory-locking b/Documentation/filesystems/directory-locking new file mode 100644 index 00000000..ff7b611a --- /dev/null +++ b/Documentation/filesystems/directory-locking @@ -0,0 +1,114 @@ + Locking scheme used for directory operations is based on two +kinds of locks - per-inode (->i_mutex) and per-filesystem +(->s_vfs_rename_mutex). + + For our purposes all operations fall in 5 classes: + +1) read access. Locking rules: caller locks directory we are accessing. + +2) object creation. Locking rules: same as above. + +3) object removal. Locking rules: caller locks parent, finds victim, +locks victim and calls the method. + +4) rename() that is _not_ cross-directory. Locking rules: caller locks +the parent, finds source and target, if target already exists - locks it +and then calls the method. + +5) link creation. Locking rules: + * lock parent + * check that source is not a directory + * lock source + * call the method. + +6) cross-directory rename. The trickiest in the whole bunch. Locking +rules: + * lock the filesystem + * lock parents in "ancestors first" order. + * find source and target. + * if old parent is equal to or is a descendent of target + fail with -ENOTEMPTY + * if new parent is equal to or is a descendent of source + fail with -ELOOP + * if target exists - lock it. + * call the method. + + +The rules above obviously guarantee that all directories that are going to be +read, modified or removed by method will be locked by caller. + + +If no directory is its own ancestor, the scheme above is deadlock-free. +Proof: + + First of all, at any moment we have a partial ordering of the +objects - A < B iff A is an ancestor of B. + + That ordering can change. However, the following is true: + +(1) if object removal or non-cross-directory rename holds lock on A and + attempts to acquire lock on B, A will remain the parent of B until we + acquire the lock on B. (Proof: only cross-directory rename can change + the parent of object and it would have to lock the parent). + +(2) if cross-directory rename holds the lock on filesystem, order will not + change until rename acquires all locks. (Proof: other cross-directory + renames will be blocked on filesystem lock and we don't start changing + the order until we had acquired all locks). + +(3) any operation holds at most one lock on non-directory object and + that lock is acquired after all other locks. (Proof: see descriptions + of operations). + + Now consider the minimal deadlock. Each process is blocked on +attempt to acquire some lock and already holds at least one lock. Let's +consider the set of contended locks. First of all, filesystem lock is +not contended, since any process blocked on it is not holding any locks. +Thus all processes are blocked on ->i_mutex. + + Non-directory objects are not contended due to (3). Thus link +creation can't be a part of deadlock - it can't be blocked on source +and it means that it doesn't hold any locks. + + Any contended object is either held by cross-directory rename or +has a child that is also contended. Indeed, suppose that it is held by +operation other than cross-directory rename. Then the lock this operation +is blocked on belongs to child of that object due to (1). + + It means that one of the operations is cross-directory rename. +Otherwise the set of contended objects would be infinite - each of them +would have a contended child and we had assumed that no object is its +own descendent. Moreover, there is exactly one cross-directory rename +(see above). + + Consider the object blocking the cross-directory rename. One +of its descendents is locked by cross-directory rename (otherwise we +would again have an infinite set of contended objects). But that +means that cross-directory rename is taking locks out of order. Due +to (2) the order hadn't changed since we had acquired filesystem lock. +But locking rules for cross-directory rename guarantee that we do not +try to acquire lock on descendent before the lock on ancestor. +Contradiction. I.e. deadlock is impossible. Q.E.D. + + + These operations are guaranteed to avoid loop creation. Indeed, +the only operation that could introduce loops is cross-directory rename. +Since the only new (parent, child) pair added by rename() is (new parent, +source), such loop would have to contain these objects and the rest of it +would have to exist before rename(). I.e. at the moment of loop creation +rename() responsible for that would be holding filesystem lock and new parent +would have to be equal to or a descendent of source. But that means that +new parent had been equal to or a descendent of source since the moment when +we had acquired filesystem lock and rename() would fail with -ELOOP in that +case. + + While this locking scheme works for arbitrary DAGs, it relies on +ability to check that directory is a descendent of another object. Current +implementation assumes that directory graph is a tree. This assumption is +also preserved by all operations (cross-directory rename on a tree that would +not introduce a cycle will leave it a tree and link() fails for directories). + + Notice that "directory" in the above == "anything that might have +children", so if we are going to introduce hybrid objects we will need +either to make sure that link(2) doesn't work for them or to make changes +in is_subdir() that would make it work even in presence of such beasts. diff --git a/Documentation/filesystems/dlmfs.txt b/Documentation/filesystems/dlmfs.txt new file mode 100644 index 00000000..1b528b2a --- /dev/null +++ b/Documentation/filesystems/dlmfs.txt @@ -0,0 +1,130 @@ +dlmfs +================== +A minimal DLM userspace interface implemented via a virtual file +system. + +dlmfs is built with OCFS2 as it requires most of its infrastructure. + +Project web page: http://oss.oracle.com/projects/ocfs2 +Tools web page: http://oss.oracle.com/projects/ocfs2-tools +OCFS2 mailing lists: http://oss.oracle.com/projects/ocfs2/mailman/ + +All code copyright 2005 Oracle except when otherwise noted. + +CREDITS +======= + +Some code taken from ramfs which is Copyright (C) 2000 Linus Torvalds +and Transmeta Corp. + +Mark Fasheh <mark.fasheh@oracle.com> + +Caveats +======= +- Right now it only works with the OCFS2 DLM, though support for other + DLM implementations should not be a major issue. + +Mount options +============= +None + +Usage +===== + +If you're just interested in OCFS2, then please see ocfs2.txt. The +rest of this document will be geared towards those who want to use +dlmfs for easy to setup and easy to use clustered locking in +userspace. + +Setup +===== + +dlmfs requires that the OCFS2 cluster infrastructure be in +place. Please download ocfs2-tools from the above url and configure a +cluster. + +You'll want to start heartbeating on a volume which all the nodes in +your lockspace can access. The easiest way to do this is via +ocfs2_hb_ctl (distributed with ocfs2-tools). Right now it requires +that an OCFS2 file system be in place so that it can automatically +find its heartbeat area, though it will eventually support heartbeat +against raw disks. + +Please see the ocfs2_hb_ctl and mkfs.ocfs2 manual pages distributed +with ocfs2-tools. + +Once you're heartbeating, DLM lock 'domains' can be easily created / +destroyed and locks within them accessed. + +Locking +======= + +Users may access dlmfs via standard file system calls, or they can use +'libo2dlm' (distributed with ocfs2-tools) which abstracts the file +system calls and presents a more traditional locking api. + +dlmfs handles lock caching automatically for the user, so a lock +request for an already acquired lock will not generate another DLM +call. Userspace programs are assumed to handle their own local +locking. + +Two levels of locks are supported - Shared Read, and Exclusive. +Also supported is a Trylock operation. + +For information on the libo2dlm interface, please see o2dlm.h, +distributed with ocfs2-tools. + +Lock value blocks can be read and written to a resource via read(2) +and write(2) against the fd obtained via your open(2) call. The +maximum currently supported LVB length is 64 bytes (though that is an +OCFS2 DLM limitation). Through this mechanism, users of dlmfs can share +small amounts of data amongst their nodes. + +mkdir(2) signals dlmfs to join a domain (which will have the same name +as the resulting directory) + +rmdir(2) signals dlmfs to leave the domain + +Locks for a given domain are represented by regular inodes inside the +domain directory. Locking against them is done via the open(2) system +call. + +The open(2) call will not return until your lock has been granted or +an error has occurred, unless it has been instructed to do a trylock +operation. If the lock succeeds, you'll get an fd. + +open(2) with O_CREAT to ensure the resource inode is created - dlmfs does +not automatically create inodes for existing lock resources. + +Open Flag Lock Request Type +--------- ----------------- +O_RDONLY Shared Read +O_RDWR Exclusive + +Open Flag Resulting Locking Behavior +--------- -------------------------- +O_NONBLOCK Trylock operation + +You must provide exactly one of O_RDONLY or O_RDWR. + +If O_NONBLOCK is also provided and the trylock operation was valid but +could not lock the resource then open(2) will return ETXTBUSY. + +close(2) drops the lock associated with your fd. + +Modes passed to mkdir(2) or open(2) are adhered to locally. Chown is +supported locally as well. This means you can use them to restrict +access to the resources via dlmfs on your local node only. + +The resource LVB may be read from the fd in either Shared Read or +Exclusive modes via the read(2) system call. It can be written via +write(2) only when open in Exclusive mode. + +Once written, an LVB will be visible to other nodes who obtain Read +Only or higher level locks on the resource. + +See Also +======== +http://opendlm.sourceforge.net/cvsmirror/opendlm/docs/dlmbook_final.pdf + +For more information on the VMS distributed locking API. diff --git a/Documentation/filesystems/dnotify.txt b/Documentation/filesystems/dnotify.txt new file mode 100644 index 00000000..6baf88f4 --- /dev/null +++ b/Documentation/filesystems/dnotify.txt @@ -0,0 +1,70 @@ + Linux Directory Notification + ============================ + + Stephen Rothwell <sfr@canb.auug.org.au> + +The intention of directory notification is to allow user applications +to be notified when a directory, or any of the files in it, are changed. +The basic mechanism involves the application registering for notification +on a directory using a fcntl(2) call and the notifications themselves +being delivered using signals. + +The application decides which "events" it wants to be notified about. +The currently defined events are: + + DN_ACCESS A file in the directory was accessed (read) + DN_MODIFY A file in the directory was modified (write,truncate) + DN_CREATE A file was created in the directory + DN_DELETE A file was unlinked from directory + DN_RENAME A file in the directory was renamed + DN_ATTRIB A file in the directory had its attributes + changed (chmod,chown) + +Usually, the application must reregister after each notification, but +if DN_MULTISHOT is or'ed with the event mask, then the registration will +remain until explicitly removed (by registering for no events). + +By default, SIGIO will be delivered to the process and no other useful +information. However, if the F_SETSIG fcntl(2) call is used to let the +kernel know which signal to deliver, a siginfo structure will be passed to +the signal handler and the si_fd member of that structure will contain the +file descriptor associated with the directory in which the event occurred. + +Preferably the application will choose one of the real time signals +(SIGRTMIN + <n>) so that the notifications may be queued. This is +especially important if DN_MULTISHOT is specified. Note that SIGRTMIN +is often blocked, so it is better to use (at least) SIGRTMIN + 1. + +Implementation expectations (features and bugs :-)) +--------------------------- + +The notification should work for any local access to files even if the +actual file system is on a remote server. This implies that remote +access to files served by local user mode servers should be notified. +Also, remote accesses to files served by a local kernel NFS server should +be notified. + +In order to make the impact on the file system code as small as possible, +the problem of hard links to files has been ignored. So if a file (x) +exists in two directories (a and b) then a change to the file using the +name "a/x" should be notified to a program expecting notifications on +directory "a", but will not be notified to one expecting notifications on +directory "b". + +Also, files that are unlinked, will still cause notifications in the +last directory that they were linked to. + +Configuration +------------- + +Dnotify is controlled via the CONFIG_DNOTIFY configuration option. When +disabled, fcntl(fd, F_NOTIFY, ...) will return -EINVAL. + +Example +------- +See Documentation/filesystems/dnotify_test.c for an example. + +NOTE +---- +Beginning with Linux 2.6.13, dnotify has been replaced by inotify. +See Documentation/filesystems/inotify.txt for more information on it. diff --git a/Documentation/filesystems/dnotify_test.c b/Documentation/filesystems/dnotify_test.c new file mode 100644 index 00000000..8b37b4a1 --- /dev/null +++ b/Documentation/filesystems/dnotify_test.c @@ -0,0 +1,34 @@ +#define _GNU_SOURCE /* needed to get the defines */ +#include <fcntl.h> /* in glibc 2.2 this has the needed + values defined */ +#include <signal.h> +#include <stdio.h> +#include <unistd.h> + +static volatile int event_fd; + +static void handler(int sig, siginfo_t *si, void *data) +{ + event_fd = si->si_fd; +} + +int main(void) +{ + struct sigaction act; + int fd; + + act.sa_sigaction = handler; + sigemptyset(&act.sa_mask); + act.sa_flags = SA_SIGINFO; + sigaction(SIGRTMIN + 1, &act, NULL); + + fd = open(".", O_RDONLY); + fcntl(fd, F_SETSIG, SIGRTMIN + 1); + fcntl(fd, F_NOTIFY, DN_MODIFY|DN_CREATE|DN_MULTISHOT); + /* we will now be notified if any of the files + in "." is modified or new files are created */ + while (1) { + pause(); + printf("Got event on fd=%d\n", event_fd); + } +} diff --git a/Documentation/filesystems/ecryptfs.txt b/Documentation/filesystems/ecryptfs.txt new file mode 100644 index 00000000..01d8a083 --- /dev/null +++ b/Documentation/filesystems/ecryptfs.txt @@ -0,0 +1,77 @@ +eCryptfs: A stacked cryptographic filesystem for Linux + +eCryptfs is free software. Please see the file COPYING for details. +For documentation, please see the files in the doc/ subdirectory. For +building and installation instructions please see the INSTALL file. + +Maintainer: Phillip Hellewell +Lead developer: Michael A. Halcrow <mhalcrow@us.ibm.com> +Developers: Michael C. Thompson + Kent Yoder +Web Site: http://ecryptfs.sf.net + +This software is currently undergoing development. Make sure to +maintain a backup copy of any data you write into eCryptfs. + +eCryptfs requires the userspace tools downloadable from the +SourceForge site: + +http://sourceforge.net/projects/ecryptfs/ + +Userspace requirements include: + - David Howells' userspace keyring headers and libraries (version + 1.0 or higher), obtainable from + http://people.redhat.com/~dhowells/keyutils/ + - Libgcrypt + + +NOTES + +In the beta/experimental releases of eCryptfs, when you upgrade +eCryptfs, you should copy the files to an unencrypted location and +then copy the files back into the new eCryptfs mount to migrate the +files. + + +MOUNT-WIDE PASSPHRASE + +Create a new directory into which eCryptfs will write its encrypted +files (i.e., /root/crypt). Then, create the mount point directory +(i.e., /mnt/crypt). Now it's time to mount eCryptfs: + +mount -t ecryptfs /root/crypt /mnt/crypt + +You should be prompted for a passphrase and a salt (the salt may be +blank). + +Try writing a new file: + +echo "Hello, World" > /mnt/crypt/hello.txt + +The operation will complete. Notice that there is a new file in +/root/crypt that is at least 12288 bytes in size (depending on your +host page size). This is the encrypted underlying file for what you +just wrote. To test reading, from start to finish, you need to clear +the user session keyring: + +keyctl clear @u + +Then umount /mnt/crypt and mount again per the instructions given +above. + +cat /mnt/crypt/hello.txt + + +NOTES + +eCryptfs version 0.1 should only be mounted on (1) empty directories +or (2) directories containing files only created by eCryptfs. If you +mount a directory that has pre-existing files not created by eCryptfs, +then behavior is undefined. Do not run eCryptfs in higher verbosity +levels unless you are doing so for the sole purpose of debugging or +development, since secret values will be written out to the system log +in that case. + + +Mike Halcrow +mhalcrow@us.ibm.com diff --git a/Documentation/filesystems/exofs.txt b/Documentation/filesystems/exofs.txt new file mode 100644 index 00000000..23583a13 --- /dev/null +++ b/Documentation/filesystems/exofs.txt @@ -0,0 +1,185 @@ +=============================================================================== +WHAT IS EXOFS? +=============================================================================== + +exofs is a file system that uses an OSD and exports the API of a normal Linux +file system. Users access exofs like any other local file system, and exofs +will in turn issue commands to the local OSD initiator. + +OSD is a new T10 command set that views storage devices not as a large/flat +array of sectors but as a container of objects, each having a length, quota, +time attributes and more. Each object is addressed by a 64bit ID, and is +contained in a 64bit ID partition. Each object has associated attributes +attached to it, which are integral part of the object and provide metadata about +the object. The standard defines some common obligatory attributes, but user +attributes can be added as needed. + +=============================================================================== +ENVIRONMENT +=============================================================================== + +To use this file system, you need to have an object store to run it on. You +may download a target from: +http://open-osd.org + +See Documentation/scsi/osd.txt for how to setup a working osd environment. + +=============================================================================== +USAGE +=============================================================================== + +1. Download and compile exofs and open-osd initiator: + You need an external Kernel source tree or kernel headers from your + distribution. (anything based on 2.6.26 or later). + + a. download open-osd including exofs source using: + [parent-directory]$ git clone git://git.open-osd.org/open-osd.git + + b. Build the library module like this: + [parent-directory]$ make -C KSRC=$(KER_DIR) open-osd + + This will build both the open-osd initiator as well as the exofs kernel + module. Use whatever parameters you compiled your Kernel with and + $(KER_DIR) above pointing to the Kernel you compile against. See the file + open-osd/top-level-Makefile for an example. + +2. Get the OSD initiator and target set up properly, and login to the target. + See Documentation/scsi/osd.txt for farther instructions. Also see ./do-osd + for example script that does all these steps. + +3. Insmod the exofs.ko module: + [exofs]$ insmod exofs.ko + +4. Make sure the directory where you want to mount exists. If not, create it. + (For example, mkdir /mnt/exofs) + +5. At first run you will need to invoke the mkfs.exofs application + + As an example, this will create the file system on: + /dev/osd0 partition ID 65536 + + mkfs.exofs --pid=65536 --format /dev/osd0 + + The --format is optional. If not specified, no OSD_FORMAT will be + performed and a clean file system will be created in the specified pid, + in the available space of the target. (Use --format=size_in_meg to limit + the total LUN space available) + + If pid already exists, it will be deleted and a new one will be created in + its place. Be careful. + + An exofs lives inside a single OSD partition. You can create multiple exofs + filesystems on the same device using multiple pids. + + (run mkfs.exofs without any parameters for usage help message) + +6. Mount the file system. + + For example, to mount /dev/osd0, partition ID 0x10000 on /mnt/exofs: + + mount -t exofs -o pid=65536 /dev/osd0 /mnt/exofs/ + +7. For reference (See do-exofs example script): + do-exofs start - an example of how to perform the above steps. + do-exofs stop - an example of how to unmount the file system. + do-exofs format - an example of how to format and mkfs a new exofs. + +8. Extra compilation flags (uncomment in fs/exofs/Kbuild): + CONFIG_EXOFS_DEBUG - for debug messages and extra checks. + +=============================================================================== +exofs mount options +=============================================================================== +Similar to any mount command: + mount -t exofs -o exofs_options /dev/osdX mount_exofs_directory + +Where: + -t exofs: specifies the exofs file system + + /dev/osdX: X is a decimal number. /dev/osdX was created after a successful + login into an OSD target. + + mount_exofs_directory: The directory to mount the file system on + + exofs specific options: Options are separated by commas (,) + pid=<integer> - The partition number to mount/create as + container of the filesystem. + This option is mandatory. integer can be + Hex by pre-pending an 0x to the number. + osdname=<id> - Mount by a device's osdname. + osdname is usually a 36 character uuid of the + form "d2683732-c906-4ee1-9dbd-c10c27bb40df". + It is one of the device's uuid specified in the + mkfs.exofs format command. + If this option is specified then the /dev/osdX + above can be empty and is ignored. + to=<integer> - Timeout in ticks for a single command. + default is (60 * HZ) [for debugging only] + +=============================================================================== +DESIGN +=============================================================================== + +* The file system control block (AKA on-disk superblock) resides in an object + with a special ID (defined in common.h). + Information included in the file system control block is used to fill the + in-memory superblock structure at mount time. This object is created before + the file system is used by mkexofs.c. It contains information such as: + - The file system's magic number + - The next inode number to be allocated + +* Each file resides in its own object and contains the data (and it will be + possible to extend the file over multiple objects, though this has not been + implemented yet). + +* A directory is treated as a file, and essentially contains a list of <file + name, inode #> pairs for files that are found in that directory. The object + IDs correspond to the files' inode numbers and will be allocated according to + a bitmap (stored in a separate object). Now they are allocated using a + counter. + +* Each file's control block (AKA on-disk inode) is stored in its object's + attributes. This applies to both regular files and other types (directories, + device files, symlinks, etc.). + +* Credentials are generated per object (inode and superblock) when they are + created in memory (read from disk or created). The credential works for all + operations and is used as long as the object remains in memory. + +* Async OSD operations are used whenever possible, but the target may execute + them out of order. The operations that concern us are create, delete, + readpage, writepage, update_inode, and truncate. The following pairs of + operations should execute in the order written, and we need to prevent them + from executing in reverse order: + - The following are handled with the OBJ_CREATED and OBJ_2BCREATED + flags. OBJ_CREATED is set when we know the object exists on the OSD - + in create's callback function, and when we successfully do a + read_inode. + OBJ_2BCREATED is set in the beginning of the create function, so we + know that we should wait. + - create/delete: delete should wait until the object is created + on the OSD. + - create/readpage: readpage should be able to return a page + full of zeroes in this case. If there was a write already + en-route (i.e. create, writepage, readpage) then the page + would be locked, and so it would really be the same as + create/writepage. + - create/writepage: if writepage is called for a sync write, it + should wait until the object is created on the OSD. + Otherwise, it should just return. + - create/truncate: truncate should wait until the object is + created on the OSD. + - create/update_inode: update_inode should wait until the + object is created on the OSD. + - Handled by VFS locks: + - readpage/delete: shouldn't happen because of page lock. + - writepage/delete: shouldn't happen because of page lock. + - readpage/writepage: shouldn't happen because of page lock. + +=============================================================================== +LICENSE/COPYRIGHT +=============================================================================== +The exofs file system is based on ext2 v0.5b (distributed with the Linux kernel +version 2.6.10). All files include the original copyrights, and the license +is GPL version 2 (only version 2, as is true for the Linux kernel). The +Linux kernel can be downloaded from www.kernel.org. diff --git a/Documentation/filesystems/ext2.txt b/Documentation/filesystems/ext2.txt new file mode 100644 index 00000000..67639f90 --- /dev/null +++ b/Documentation/filesystems/ext2.txt @@ -0,0 +1,383 @@ + +The Second Extended Filesystem +============================== + +ext2 was originally released in January 1993. Written by R\'emy Card, +Theodore Ts'o and Stephen Tweedie, it was a major rewrite of the +Extended Filesystem. It is currently still (April 2001) the predominant +filesystem in use by Linux. There are also implementations available +for NetBSD, FreeBSD, the GNU HURD, Windows 95/98/NT, OS/2 and RISC OS. + +Options +======= + +Most defaults are determined by the filesystem superblock, and can be +set using tune2fs(8). Kernel-determined defaults are indicated by (*). + +bsddf (*) Makes `df' act like BSD. +minixdf Makes `df' act like Minix. + +check=none, nocheck (*) Don't do extra checking of bitmaps on mount + (check=normal and check=strict options removed) + +debug Extra debugging information is sent to the + kernel syslog. Useful for developers. + +errors=continue Keep going on a filesystem error. +errors=remount-ro Remount the filesystem read-only on an error. +errors=panic Panic and halt the machine if an error occurs. + +grpid, bsdgroups Give objects the same group ID as their parent. +nogrpid, sysvgroups New objects have the group ID of their creator. + +nouid32 Use 16-bit UIDs and GIDs. + +oldalloc Enable the old block allocator. Orlov should + have better performance, we'd like to get some + feedback if it's the contrary for you. +orlov (*) Use the Orlov block allocator. + (See http://lwn.net/Articles/14633/ and + http://lwn.net/Articles/14446/.) + +resuid=n The user ID which may use the reserved blocks. +resgid=n The group ID which may use the reserved blocks. + +sb=n Use alternate superblock at this location. + +user_xattr Enable "user." POSIX Extended Attributes + (requires CONFIG_EXT2_FS_XATTR). + See also http://acl.bestbits.at +nouser_xattr Don't support "user." extended attributes. + +acl Enable POSIX Access Control Lists support + (requires CONFIG_EXT2_FS_POSIX_ACL). + See also http://acl.bestbits.at +noacl Don't support POSIX ACLs. + +nobh Do not attach buffer_heads to file pagecache. + +xip Use execute in place (no caching) if possible + +grpquota,noquota,quota,usrquota Quota options are silently ignored by ext2. + + +Specification +============= + +ext2 shares many properties with traditional Unix filesystems. It has +the concepts of blocks, inodes and directories. It has space in the +specification for Access Control Lists (ACLs), fragments, undeletion and +compression though these are not yet implemented (some are available as +separate patches). There is also a versioning mechanism to allow new +features (such as journalling) to be added in a maximally compatible +manner. + +Blocks +------ + +The space in the device or file is split up into blocks. These are +a fixed size, of 1024, 2048 or 4096 bytes (8192 bytes on Alpha systems), +which is decided when the filesystem is created. Smaller blocks mean +less wasted space per file, but require slightly more accounting overhead, +and also impose other limits on the size of files and the filesystem. + +Block Groups +------------ + +Blocks are clustered into block groups in order to reduce fragmentation +and minimise the amount of head seeking when reading a large amount +of consecutive data. Information about each block group is kept in a +descriptor table stored in the block(s) immediately after the superblock. +Two blocks near the start of each group are reserved for the block usage +bitmap and the inode usage bitmap which show which blocks and inodes +are in use. Since each bitmap is limited to a single block, this means +that the maximum size of a block group is 8 times the size of a block. + +The block(s) following the bitmaps in each block group are designated +as the inode table for that block group and the remainder are the data +blocks. The block allocation algorithm attempts to allocate data blocks +in the same block group as the inode which contains them. + +The Superblock +-------------- + +The superblock contains all the information about the configuration of +the filing system. The primary copy of the superblock is stored at an +offset of 1024 bytes from the start of the device, and it is essential +to mounting the filesystem. Since it is so important, backup copies of +the superblock are stored in block groups throughout the filesystem. +The first version of ext2 (revision 0) stores a copy at the start of +every block group, along with backups of the group descriptor block(s). +Because this can consume a considerable amount of space for large +filesystems, later revisions can optionally reduce the number of backup +copies by only putting backups in specific groups (this is the sparse +superblock feature). The groups chosen are 0, 1 and powers of 3, 5 and 7. + +The information in the superblock contains fields such as the total +number of inodes and blocks in the filesystem and how many are free, +how many inodes and blocks are in each block group, when the filesystem +was mounted (and if it was cleanly unmounted), when it was modified, +what version of the filesystem it is (see the Revisions section below) +and which OS created it. + +If the filesystem is revision 1 or higher, then there are extra fields, +such as a volume name, a unique identification number, the inode size, +and space for optional filesystem features to store configuration info. + +All fields in the superblock (as in all other ext2 structures) are stored +on the disc in little endian format, so a filesystem is portable between +machines without having to know what machine it was created on. + +Inodes +------ + +The inode (index node) is a fundamental concept in the ext2 filesystem. +Each object in the filesystem is represented by an inode. The inode +structure contains pointers to the filesystem blocks which contain the +data held in the object and all of the metadata about an object except +its name. The metadata about an object includes the permissions, owner, +group, flags, size, number of blocks used, access time, change time, +modification time, deletion time, number of links, fragments, version +(for NFS) and extended attributes (EAs) and/or Access Control Lists (ACLs). + +There are some reserved fields which are currently unused in the inode +structure and several which are overloaded. One field is reserved for the +directory ACL if the inode is a directory and alternately for the top 32 +bits of the file size if the inode is a regular file (allowing file sizes +larger than 2GB). The translator field is unused under Linux, but is used +by the HURD to reference the inode of a program which will be used to +interpret this object. Most of the remaining reserved fields have been +used up for both Linux and the HURD for larger owner and group fields, +The HURD also has a larger mode field so it uses another of the remaining +fields to store the extra more bits. + +There are pointers to the first 12 blocks which contain the file's data +in the inode. There is a pointer to an indirect block (which contains +pointers to the next set of blocks), a pointer to a doubly-indirect +block (which contains pointers to indirect blocks) and a pointer to a +trebly-indirect block (which contains pointers to doubly-indirect blocks). + +The flags field contains some ext2-specific flags which aren't catered +for by the standard chmod flags. These flags can be listed with lsattr +and changed with the chattr command, and allow specific filesystem +behaviour on a per-file basis. There are flags for secure deletion, +undeletable, compression, synchronous updates, immutability, append-only, +dumpable, no-atime, indexed directories, and data-journaling. Not all +of these are supported yet. + +Directories +----------- + +A directory is a filesystem object and has an inode just like a file. +It is a specially formatted file containing records which associate +each name with an inode number. Later revisions of the filesystem also +encode the type of the object (file, directory, symlink, device, fifo, +socket) to avoid the need to check the inode itself for this information +(support for taking advantage of this feature does not yet exist in +Glibc 2.2). + +The inode allocation code tries to assign inodes which are in the same +block group as the directory in which they are first created. + +The current implementation of ext2 uses a singly-linked list to store +the filenames in the directory; a pending enhancement uses hashing of the +filenames to allow lookup without the need to scan the entire directory. + +The current implementation never removes empty directory blocks once they +have been allocated to hold more files. + +Special files +------------- + +Symbolic links are also filesystem objects with inodes. They deserve +special mention because the data for them is stored within the inode +itself if the symlink is less than 60 bytes long. It uses the fields +which would normally be used to store the pointers to data blocks. +This is a worthwhile optimisation as it we avoid allocating a full +block for the symlink, and most symlinks are less than 60 characters long. + +Character and block special devices never have data blocks assigned to +them. Instead, their device number is stored in the inode, again reusing +the fields which would be used to point to the data blocks. + +Reserved Space +-------------- + +In ext2, there is a mechanism for reserving a certain number of blocks +for a particular user (normally the super-user). This is intended to +allow for the system to continue functioning even if non-privileged users +fill up all the space available to them (this is independent of filesystem +quotas). It also keeps the filesystem from filling up entirely which +helps combat fragmentation. + +Filesystem check +---------------- + +At boot time, most systems run a consistency check (e2fsck) on their +filesystems. The superblock of the ext2 filesystem contains several +fields which indicate whether fsck should actually run (since checking +the filesystem at boot can take a long time if it is large). fsck will +run if the filesystem was not cleanly unmounted, if the maximum mount +count has been exceeded or if the maximum time between checks has been +exceeded. + +Feature Compatibility +--------------------- + +The compatibility feature mechanism used in ext2 is sophisticated. +It safely allows features to be added to the filesystem, without +unnecessarily sacrificing compatibility with older versions of the +filesystem code. The feature compatibility mechanism is not supported by +the original revision 0 (EXT2_GOOD_OLD_REV) of ext2, but was introduced in +revision 1. There are three 32-bit fields, one for compatible features +(COMPAT), one for read-only compatible (RO_COMPAT) features and one for +incompatible (INCOMPAT) features. + +These feature flags have specific meanings for the kernel as follows: + +A COMPAT flag indicates that a feature is present in the filesystem, +but the on-disk format is 100% compatible with older on-disk formats, so +a kernel which didn't know anything about this feature could read/write +the filesystem without any chance of corrupting the filesystem (or even +making it inconsistent). This is essentially just a flag which says +"this filesystem has a (hidden) feature" that the kernel or e2fsck may +want to be aware of (more on e2fsck and feature flags later). The ext3 +HAS_JOURNAL feature is a COMPAT flag because the ext3 journal is simply +a regular file with data blocks in it so the kernel does not need to +take any special notice of it if it doesn't understand ext3 journaling. + +An RO_COMPAT flag indicates that the on-disk format is 100% compatible +with older on-disk formats for reading (i.e. the feature does not change +the visible on-disk format). However, an old kernel writing to such a +filesystem would/could corrupt the filesystem, so this is prevented. The +most common such feature, SPARSE_SUPER, is an RO_COMPAT feature because +sparse groups allow file data blocks where superblock/group descriptor +backups used to live, and ext2_free_blocks() refuses to free these blocks, +which would leading to inconsistent bitmaps. An old kernel would also +get an error if it tried to free a series of blocks which crossed a group +boundary, but this is a legitimate layout in a SPARSE_SUPER filesystem. + +An INCOMPAT flag indicates the on-disk format has changed in some +way that makes it unreadable by older kernels, or would otherwise +cause a problem if an old kernel tried to mount it. FILETYPE is an +INCOMPAT flag because older kernels would think a filename was longer +than 256 characters, which would lead to corrupt directory listings. +The COMPRESSION flag is an obvious INCOMPAT flag - if the kernel +doesn't understand compression, you would just get garbage back from +read() instead of it automatically decompressing your data. The ext3 +RECOVER flag is needed to prevent a kernel which does not understand the +ext3 journal from mounting the filesystem without replaying the journal. + +For e2fsck, it needs to be more strict with the handling of these +flags than the kernel. If it doesn't understand ANY of the COMPAT, +RO_COMPAT, or INCOMPAT flags it will refuse to check the filesystem, +because it has no way of verifying whether a given feature is valid +or not. Allowing e2fsck to succeed on a filesystem with an unknown +feature is a false sense of security for the user. Refusing to check +a filesystem with unknown features is a good incentive for the user to +update to the latest e2fsck. This also means that anyone adding feature +flags to ext2 also needs to update e2fsck to verify these features. + +Metadata +-------- + +It is frequently claimed that the ext2 implementation of writing +asynchronous metadata is faster than the ffs synchronous metadata +scheme but less reliable. Both methods are equally resolvable by their +respective fsck programs. + +If you're exceptionally paranoid, there are 3 ways of making metadata +writes synchronous on ext2: + +per-file if you have the program source: use the O_SYNC flag to open() +per-file if you don't have the source: use "chattr +S" on the file +per-filesystem: add the "sync" option to mount (or in /etc/fstab) + +the first and last are not ext2 specific but do force the metadata to +be written synchronously. See also Journaling below. + +Limitations +----------- + +There are various limits imposed by the on-disk layout of ext2. Other +limits are imposed by the current implementation of the kernel code. +Many of the limits are determined at the time the filesystem is first +created, and depend upon the block size chosen. The ratio of inodes to +data blocks is fixed at filesystem creation time, so the only way to +increase the number of inodes is to increase the size of the filesystem. +No tools currently exist which can change the ratio of inodes to blocks. + +Most of these limits could be overcome with slight changes in the on-disk +format and using a compatibility flag to signal the format change (at +the expense of some compatibility). + +Filesystem block size: 1kB 2kB 4kB 8kB + +File size limit: 16GB 256GB 2048GB 2048GB +Filesystem size limit: 2047GB 8192GB 16384GB 32768GB + +There is a 2.4 kernel limit of 2048GB for a single block device, so no +filesystem larger than that can be created at this time. There is also +an upper limit on the block size imposed by the page size of the kernel, +so 8kB blocks are only allowed on Alpha systems (and other architectures +which support larger pages). + +There is an upper limit of 32000 subdirectories in a single directory. + +There is a "soft" upper limit of about 10-15k files in a single directory +with the current linear linked-list directory implementation. This limit +stems from performance problems when creating and deleting (and also +finding) files in such large directories. Using a hashed directory index +(under development) allows 100k-1M+ files in a single directory without +performance problems (although RAM size becomes an issue at this point). + +The (meaningless) absolute upper limit of files in a single directory +(imposed by the file size, the realistic limit is obviously much less) +is over 130 trillion files. It would be higher except there are not +enough 4-character names to make up unique directory entries, so they +have to be 8 character filenames, even then we are fairly close to +running out of unique filenames. + +Journaling +---------- + +A journaling extension to the ext2 code has been developed by Stephen +Tweedie. It avoids the risks of metadata corruption and the need to +wait for e2fsck to complete after a crash, without requiring a change +to the on-disk ext2 layout. In a nutshell, the journal is a regular +file which stores whole metadata (and optionally data) blocks that have +been modified, prior to writing them into the filesystem. This means +it is possible to add a journal to an existing ext2 filesystem without +the need for data conversion. + +When changes to the filesystem (e.g. a file is renamed) they are stored in +a transaction in the journal and can either be complete or incomplete at +the time of a crash. If a transaction is complete at the time of a crash +(or in the normal case where the system does not crash), then any blocks +in that transaction are guaranteed to represent a valid filesystem state, +and are copied into the filesystem. If a transaction is incomplete at +the time of the crash, then there is no guarantee of consistency for +the blocks in that transaction so they are discarded (which means any +filesystem changes they represent are also lost). +Check Documentation/filesystems/ext3.txt if you want to read more about +ext3 and journaling. + +References +========== + +The kernel source file:/usr/src/linux/fs/ext2/ +e2fsprogs (e2fsck) http://e2fsprogs.sourceforge.net/ +Design & Implementation http://e2fsprogs.sourceforge.net/ext2intro.html +Journaling (ext3) ftp://ftp.uk.linux.org/pub/linux/sct/fs/jfs/ +Filesystem Resizing http://ext2resize.sourceforge.net/ +Compression (*) http://e2compr.sourceforge.net/ + +Implementations for: +Windows 95/98/NT/2000 http://www.chrysocome.net/explore2fs +Windows 95 (*) http://www.yipton.net/content.html#FSDEXT2 +DOS client (*) ftp://metalab.unc.edu/pub/Linux/system/filesystems/ext2/ +OS/2 (+) ftp://metalab.unc.edu/pub/Linux/system/filesystems/ext2/ +RISC OS client http://www.esw-heim.tu-clausthal.de/~marco/smorbrod/IscaFS/ + +(*) no longer actively developed/supported (as of Apr 2001) +(+) no longer actively developed/supported (as of Mar 2009) diff --git a/Documentation/filesystems/ext3.txt b/Documentation/filesystems/ext3.txt new file mode 100644 index 00000000..272f80d5 --- /dev/null +++ b/Documentation/filesystems/ext3.txt @@ -0,0 +1,231 @@ + +Ext3 Filesystem +=============== + +Ext3 was originally released in September 1999. Written by Stephen Tweedie +for the 2.2 branch, and ported to 2.4 kernels by Peter Braam, Andreas Dilger, +Andrew Morton, Alexander Viro, Ted Ts'o and Stephen Tweedie. + +Ext3 is the ext2 filesystem enhanced with journalling capabilities. + +Options +======= + +When mounting an ext3 filesystem, the following option are accepted: +(*) == default + +ro Mount filesystem read only. Note that ext3 will replay + the journal (and thus write to the partition) even when + mounted "read only". Mount options "ro,noload" can be + used to prevent writes to the filesystem. + +journal=update Update the ext3 file system's journal to the current + format. + +journal=inum When a journal already exists, this option is ignored. + Otherwise, it specifies the number of the inode which + will represent the ext3 file system's journal file. + +journal_dev=devnum When the external journal device's major/minor numbers + have changed, this option allows the user to specify + the new journal location. The journal device is + identified through its new major/minor numbers encoded + in devnum. + +norecovery Don't load the journal on mounting. Note that this forces +noload mount of inconsistent filesystem, which can lead to + various problems. + +data=journal All data are committed into the journal prior to being + written into the main file system. + +data=ordered (*) All data are forced directly out to the main file + system prior to its metadata being committed to the + journal. + +data=writeback Data ordering is not preserved, data may be written + into the main file system after its metadata has been + committed to the journal. + +commit=nrsec (*) Ext3 can be told to sync all its data and metadata + every 'nrsec' seconds. The default value is 5 seconds. + This means that if you lose your power, you will lose + as much as the latest 5 seconds of work (your + filesystem will not be damaged though, thanks to the + journaling). This default value (or any low value) + will hurt performance, but it's good for data-safety. + Setting it to 0 will have the same effect as leaving + it at the default (5 seconds). + Setting it to very large values will improve + performance. + +barrier=<0(*)|1> This enables/disables the use of write barriers in +barrier the jbd code. barrier=0 disables, barrier=1 enables. +nobarrier (*) This also requires an IO stack which can support + barriers, and if jbd gets an error on a barrier + write, it will disable again with a warning. + Write barriers enforce proper on-disk ordering + of journal commits, making volatile disk write caches + safe to use, at some performance penalty. If + your disks are battery-backed in one way or another, + disabling barriers may safely improve performance. + The mount options "barrier" and "nobarrier" can + also be used to enable or disable barriers, for + consistency with other ext3 mount options. + +orlov (*) This enables the new Orlov block allocator. It is + enabled by default. + +oldalloc This disables the Orlov block allocator and enables + the old block allocator. Orlov should have better + performance - we'd like to get some feedback if it's + the contrary for you. + +user_xattr Enables Extended User Attributes. Additionally, you + need to have extended attribute support enabled in the + kernel configuration (CONFIG_EXT3_FS_XATTR). See the + attr(5) manual page and http://acl.bestbits.at/ to + learn more about extended attributes. + +nouser_xattr Disables Extended User Attributes. + +acl Enables POSIX Access Control Lists support. + Additionally, you need to have ACL support enabled in + the kernel configuration (CONFIG_EXT3_FS_POSIX_ACL). + See the acl(5) manual page and http://acl.bestbits.at/ + for more information. + +noacl This option disables POSIX Access Control List + support. + +reservation + +noreservation + +bsddf (*) Make 'df' act like BSD. +minixdf Make 'df' act like Minix. + +check=none Don't do extra checking of bitmaps on mount. +nocheck + +debug Extra debugging information is sent to syslog. + +errors=remount-ro Remount the filesystem read-only on an error. +errors=continue Keep going on a filesystem error. +errors=panic Panic and halt the machine if an error occurs. + (These mount options override the errors behavior + specified in the superblock, which can be + configured using tune2fs.) + +data_err=ignore(*) Just print an error message if an error occurs + in a file data buffer in ordered mode. +data_err=abort Abort the journal if an error occurs in a file + data buffer in ordered mode. + +grpid Give objects the same group ID as their creator. +bsdgroups + +nogrpid (*) New objects have the group ID of their creator. +sysvgroups + +resgid=n The group ID which may use the reserved blocks. + +resuid=n The user ID which may use the reserved blocks. + +sb=n Use alternate superblock at this location. + +quota These options are ignored by the filesystem. They +noquota are used only by quota tools to recognize volumes +grpquota where quota should be turned on. See documentation +usrquota in the quota-tools package for more details + (http://sourceforge.net/projects/linuxquota). + +jqfmt=<quota type> These options tell filesystem details about quota +usrjquota=<file> so that quota information can be properly updated +grpjquota=<file> during journal replay. They replace the above + quota options. See documentation in the quota-tools + package for more details + (http://sourceforge.net/projects/linuxquota). + +bh (*) ext3 associates buffer heads to data pages to +nobh (a) cache disk block mapping information + (b) link pages into transaction to provide + ordering guarantees. + "bh" option forces use of buffer heads. + "nobh" option tries to avoid associating buffer + heads (supported only for "writeback" mode). + + +Specification +============= +Ext3 shares all disk implementation with the ext2 filesystem, and adds +transactions capabilities to ext2. Journaling is done by the Journaling Block +Device layer. + +Journaling Block Device layer +----------------------------- +The Journaling Block Device layer (JBD) isn't ext3 specific. It was designed +to add journaling capabilities to a block device. The ext3 filesystem code +will inform the JBD of modifications it is performing (called a transaction). +The journal supports the transactions start and stop, and in case of a crash, +the journal can replay the transactions to quickly put the partition back into +a consistent state. + +Handles represent a single atomic update to a filesystem. JBD can handle an +external journal on a block device. + +Data Mode +--------- +There are 3 different data modes: + +* writeback mode +In data=writeback mode, ext3 does not journal data at all. This mode provides +a similar level of journaling as that of XFS, JFS, and ReiserFS in its default +mode - metadata journaling. A crash+recovery can cause incorrect data to +appear in files which were written shortly before the crash. This mode will +typically provide the best ext3 performance. + +* ordered mode +In data=ordered mode, ext3 only officially journals metadata, but it logically +groups metadata and data blocks into a single unit called a transaction. When +it's time to write the new metadata out to disk, the associated data blocks +are written first. In general, this mode performs slightly slower than +writeback but significantly faster than journal mode. + +* journal mode +data=journal mode provides full data and metadata journaling. All new data is +written to the journal first, and then to its final location. +In the event of a crash, the journal can be replayed, bringing both data and +metadata into a consistent state. This mode is the slowest except when data +needs to be read from and written to disk at the same time where it +outperforms all other modes. + +Compatibility +------------- + +Ext2 partitions can be easily convert to ext3, with `tune2fs -j <dev>`. +Ext3 is fully compatible with Ext2. Ext3 partitions can easily be mounted as +Ext2. + + +External Tools +============== +See manual pages to learn more. + +tune2fs: create a ext3 journal on a ext2 partition with the -j flag. +mke2fs: create a ext3 partition with the -j flag. +debugfs: ext2 and ext3 file system debugger. +ext2online: online (mounted) ext2 and ext3 filesystem resizer + + +References +========== + +kernel source: <file:fs/ext3/> + <file:fs/jbd/> + +programs: http://e2fsprogs.sourceforge.net/ + http://ext2resize.sourceforge.net + +useful links: http://www.ibm.com/developerworks/library/l-fs7.html + http://www.ibm.com/developerworks/library/l-fs8.html diff --git a/Documentation/filesystems/ext4.txt b/Documentation/filesystems/ext4.txt new file mode 100644 index 00000000..3ae9bc94 --- /dev/null +++ b/Documentation/filesystems/ext4.txt @@ -0,0 +1,615 @@ + +Ext4 Filesystem +=============== + +Ext4 is an an advanced level of the ext3 filesystem which incorporates +scalability and reliability enhancements for supporting large filesystems +(64 bit) in keeping with increasing disk capacities and state-of-the-art +feature requirements. + +Mailing list: linux-ext4@vger.kernel.org +Web site: http://ext4.wiki.kernel.org + + +1. Quick usage instructions: +=========================== + +Note: More extensive information for getting started with ext4 can be + found at the ext4 wiki site at the URL: + http://ext4.wiki.kernel.org/index.php/Ext4_Howto + + - Compile and install the latest version of e2fsprogs (as of this + writing version 1.41.3) from: + + http://sourceforge.net/project/showfiles.php?group_id=2406 + + or + + ftp://ftp.kernel.org/pub/linux/kernel/people/tytso/e2fsprogs/ + + or grab the latest git repository from: + + git://git.kernel.org/pub/scm/fs/ext2/e2fsprogs.git + + - Note that it is highly important to install the mke2fs.conf file + that comes with the e2fsprogs 1.41.x sources in /etc/mke2fs.conf. If + you have edited the /etc/mke2fs.conf file installed on your system, + you will need to merge your changes with the version from e2fsprogs + 1.41.x. + + - Create a new filesystem using the ext4 filesystem type: + + # mke2fs -t ext4 /dev/hda1 + + Or to configure an existing ext3 filesystem to support extents: + + # tune2fs -O extents /dev/hda1 + + If the filesystem was created with 128 byte inodes, it can be + converted to use 256 byte for greater efficiency via: + + # tune2fs -I 256 /dev/hda1 + + (Note: we currently do not have tools to convert an ext4 + filesystem back to ext3; so please do not do try this on production + filesystems.) + + - Mounting: + + # mount -t ext4 /dev/hda1 /wherever + + - When comparing performance with other filesystems, it's always + important to try multiple workloads; very often a subtle change in a + workload parameter can completely change the ranking of which + filesystems do well compared to others. When comparing versus ext3, + note that ext4 enables write barriers by default, while ext3 does + not enable write barriers by default. So it is useful to use + explicitly specify whether barriers are enabled or not when via the + '-o barriers=[0|1]' mount option for both ext3 and ext4 filesystems + for a fair comparison. When tuning ext3 for best benchmark numbers, + it is often worthwhile to try changing the data journaling mode; '-o + data=writeback,nobh' can be faster for some workloads. (Note + however that running mounted with data=writeback can potentially + leave stale data exposed in recently written files in case of an + unclean shutdown, which could be a security exposure in some + situations.) Configuring the filesystem with a large journal can + also be helpful for metadata-intensive workloads. + +2. Features +=========== + +2.1 Currently available + +* ability to use filesystems > 16TB (e2fsprogs support not available yet) +* extent format reduces metadata overhead (RAM, IO for access, transactions) +* extent format more robust in face of on-disk corruption due to magics, +* internal redundancy in tree +* improved file allocation (multi-block alloc) +* lift 32000 subdirectory limit imposed by i_links_count[1] +* nsec timestamps for mtime, atime, ctime, create time +* inode version field on disk (NFSv4, Lustre) +* reduced e2fsck time via uninit_bg feature +* journal checksumming for robustness, performance +* persistent file preallocation (e.g for streaming media, databases) +* ability to pack bitmaps and inode tables into larger virtual groups via the + flex_bg feature +* large file support +* Inode allocation using large virtual block groups via flex_bg +* delayed allocation +* large block (up to pagesize) support +* efficient new ordered mode in JBD2 and ext4(avoid using buffer head to force + the ordering) + +[1] Filesystems with a block size of 1k may see a limit imposed by the +directory hash tree having a maximum depth of two. + +2.2 Candidate features for future inclusion + +* Online defrag (patches available but not well tested) +* reduced mke2fs time via lazy itable initialization in conjunction with + the uninit_bg feature (capability to do this is available in e2fsprogs + but a kernel thread to do lazy zeroing of unused inode table blocks + after filesystem is first mounted is required for safety) + +There are several others under discussion, whether they all make it in is +partly a function of how much time everyone has to work on them. Features like +metadata checksumming have been discussed and planned for a bit but no patches +exist yet so I'm not sure they're in the near-term roadmap. + +The big performance win will come with mballoc, delalloc and flex_bg +grouping of bitmaps and inode tables. Some test results available here: + + - http://www.bullopensource.org/ext4/20080818-ffsb/ffsb-write-2.6.27-rc1.html + - http://www.bullopensource.org/ext4/20080818-ffsb/ffsb-readwrite-2.6.27-rc1.html + +3. Options +========== + +When mounting an ext4 filesystem, the following option are accepted: +(*) == default + +ro Mount filesystem read only. Note that ext4 will + replay the journal (and thus write to the + partition) even when mounted "read only". The + mount options "ro,noload" can be used to prevent + writes to the filesystem. + +journal_checksum Enable checksumming of the journal transactions. + This will allow the recovery code in e2fsck and the + kernel to detect corruption in the kernel. It is a + compatible change and will be ignored by older kernels. + +journal_async_commit Commit block can be written to disk without waiting + for descriptor blocks. If enabled older kernels cannot + mount the device. This will enable 'journal_checksum' + internally. + +journal=update Update the ext4 file system's journal to the current + format. + +journal_dev=devnum When the external journal device's major/minor numbers + have changed, this option allows the user to specify + the new journal location. The journal device is + identified through its new major/minor numbers encoded + in devnum. + +norecovery Don't load the journal on mounting. Note that +noload if the filesystem was not unmounted cleanly, + skipping the journal replay will lead to the + filesystem containing inconsistencies that can + lead to any number of problems. + +data=journal All data are committed into the journal prior to being + written into the main file system. + +data=ordered (*) All data are forced directly out to the main file + system prior to its metadata being committed to the + journal. + +data=writeback Data ordering is not preserved, data may be written + into the main file system after its metadata has been + committed to the journal. + +commit=nrsec (*) Ext4 can be told to sync all its data and metadata + every 'nrsec' seconds. The default value is 5 seconds. + This means that if you lose your power, you will lose + as much as the latest 5 seconds of work (your + filesystem will not be damaged though, thanks to the + journaling). This default value (or any low value) + will hurt performance, but it's good for data-safety. + Setting it to 0 will have the same effect as leaving + it at the default (5 seconds). + Setting it to very large values will improve + performance. + +barrier=<0|1(*)> This enables/disables the use of write barriers in +barrier(*) the jbd code. barrier=0 disables, barrier=1 enables. +nobarrier This also requires an IO stack which can support + barriers, and if jbd gets an error on a barrier + write, it will disable again with a warning. + Write barriers enforce proper on-disk ordering + of journal commits, making volatile disk write caches + safe to use, at some performance penalty. If + your disks are battery-backed in one way or another, + disabling barriers may safely improve performance. + The mount options "barrier" and "nobarrier" can + also be used to enable or disable barriers, for + consistency with other ext4 mount options. + +inode_readahead_blks=n This tuning parameter controls the maximum + number of inode table blocks that ext4's inode + table readahead algorithm will pre-read into + the buffer cache. The default value is 32 blocks. + +orlov (*) This enables the new Orlov block allocator. It is + enabled by default. + +oldalloc This disables the Orlov block allocator and enables + the old block allocator. Orlov should have better + performance - we'd like to get some feedback if it's + the contrary for you. + +user_xattr Enables Extended User Attributes. Additionally, you + need to have extended attribute support enabled in the + kernel configuration (CONFIG_EXT4_FS_XATTR). See the + attr(5) manual page and http://acl.bestbits.at/ to + learn more about extended attributes. + +nouser_xattr Disables Extended User Attributes. + +acl Enables POSIX Access Control Lists support. + Additionally, you need to have ACL support enabled in + the kernel configuration (CONFIG_EXT4_FS_POSIX_ACL). + See the acl(5) manual page and http://acl.bestbits.at/ + for more information. + +noacl This option disables POSIX Access Control List + support. + +bsddf (*) Make 'df' act like BSD. +minixdf Make 'df' act like Minix. + +debug Extra debugging information is sent to syslog. + +abort Simulate the effects of calling ext4_abort() for + debugging purposes. This is normally used while + remounting a filesystem which is already mounted. + +errors=remount-ro Remount the filesystem read-only on an error. +errors=continue Keep going on a filesystem error. +errors=panic Panic and halt the machine if an error occurs. + (These mount options override the errors behavior + specified in the superblock, which can be configured + using tune2fs) + +data_err=ignore(*) Just print an error message if an error occurs + in a file data buffer in ordered mode. +data_err=abort Abort the journal if an error occurs in a file + data buffer in ordered mode. + +grpid Give objects the same group ID as their creator. +bsdgroups + +nogrpid (*) New objects have the group ID of their creator. +sysvgroups + +resgid=n The group ID which may use the reserved blocks. + +resuid=n The user ID which may use the reserved blocks. + +sb=n Use alternate superblock at this location. + +quota These options are ignored by the filesystem. They +noquota are used only by quota tools to recognize volumes +grpquota where quota should be turned on. See documentation +usrquota in the quota-tools package for more details + (http://sourceforge.net/projects/linuxquota). + +jqfmt=<quota type> These options tell filesystem details about quota +usrjquota=<file> so that quota information can be properly updated +grpjquota=<file> during journal replay. They replace the above + quota options. See documentation in the quota-tools + package for more details + (http://sourceforge.net/projects/linuxquota). + +bh (*) ext4 associates buffer heads to data pages to +nobh (a) cache disk block mapping information + (b) link pages into transaction to provide + ordering guarantees. + "bh" option forces use of buffer heads. + "nobh" option tries to avoid associating buffer + heads (supported only for "writeback" mode). + +stripe=n Number of filesystem blocks that mballoc will try + to use for allocation size and alignment. For RAID5/6 + systems this should be the number of data + disks * RAID chunk size in file system blocks. + +delalloc (*) Defer block allocation until just before ext4 + writes out the block(s) in question. This + allows ext4 to better allocation decisions + more efficiently. +nodelalloc Disable delayed allocation. Blocks are allocated + when the data is copied from userspace to the + page cache, either via the write(2) system call + or when an mmap'ed page which was previously + unallocated is written for the first time. + +max_batch_time=usec Maximum amount of time ext4 should wait for + additional filesystem operations to be batch + together with a synchronous write operation. + Since a synchronous write operation is going to + force a commit and then a wait for the I/O + complete, it doesn't cost much, and can be a + huge throughput win, we wait for a small amount + of time to see if any other transactions can + piggyback on the synchronous write. The + algorithm used is designed to automatically tune + for the speed of the disk, by measuring the + amount of time (on average) that it takes to + finish committing a transaction. Call this time + the "commit time". If the time that the + transaction has been running is less than the + commit time, ext4 will try sleeping for the + commit time to see if other operations will join + the transaction. The commit time is capped by + the max_batch_time, which defaults to 15000us + (15ms). This optimization can be turned off + entirely by setting max_batch_time to 0. + +min_batch_time=usec This parameter sets the commit time (as + described above) to be at least min_batch_time. + It defaults to zero microseconds. Increasing + this parameter may improve the throughput of + multi-threaded, synchronous workloads on very + fast disks, at the cost of increasing latency. + +journal_ioprio=prio The I/O priority (from 0 to 7, where 0 is the + highest priorty) which should be used for I/O + operations submitted by kjournald2 during a + commit operation. This defaults to 3, which is + a slightly higher priority than the default I/O + priority. + +auto_da_alloc(*) Many broken applications don't use fsync() when +noauto_da_alloc replacing existing files via patterns such as + fd = open("foo.new")/write(fd,..)/close(fd)/ + rename("foo.new", "foo"), or worse yet, + fd = open("foo", O_TRUNC)/write(fd,..)/close(fd). + If auto_da_alloc is enabled, ext4 will detect + the replace-via-rename and replace-via-truncate + patterns and force that any delayed allocation + blocks are allocated such that at the next + journal commit, in the default data=ordered + mode, the data blocks of the new file are forced + to disk before the rename() operation is + committed. This provides roughly the same level + of guarantees as ext3, and avoids the + "zero-length" problem that can happen when a + system crashes before the delayed allocation + blocks are forced to disk. + +noinit_itable Do not initialize any uninitialized inode table + blocks in the background. This feature may be + used by installation CD's so that the install + process can complete as quickly as possible; the + inode table initialization process would then be + deferred until the next time the file system + is unmounted. + +init_itable=n The lazy itable init code will wait n times the + number of milliseconds it took to zero out the + previous block group's inode table. This + minimizes the impact on the systme performance + while file system's inode table is being initialized. + +discard Controls whether ext4 should issue discard/TRIM +nodiscard(*) commands to the underlying block device when + blocks are freed. This is useful for SSD devices + and sparse/thinly-provisioned LUNs, but it is off + by default until sufficient testing has been done. + +nouid32 Disables 32-bit UIDs and GIDs. This is for + interoperability with older kernels which only + store and expect 16-bit values. + +resize Allows to resize filesystem to the end of the last + existing block group, further resize has to be done + with resize2fs either online, or offline. It can be + used only with conjunction with remount. + +block_validity This options allows to enables/disables the in-kernel +noblock_validity facility for tracking filesystem metadata blocks + within internal data structures. This allows multi- + block allocator and other routines to quickly locate + extents which might overlap with filesystem metadata + blocks. This option is intended for debugging + purposes and since it negatively affects the + performance, it is off by default. + +dioread_lock Controls whether or not ext4 should use the DIO read +dioread_nolock locking. If the dioread_nolock option is specified + ext4 will allocate uninitialized extent before buffer + write and convert the extent to initialized after IO + completes. This approach allows ext4 code to avoid + using inode mutex, which improves scalability on high + speed storages. However this does not work with nobh + option and the mount will fail. Nor does it work with + data journaling and dioread_nolock option will be + ignored with kernel warning. Note that dioread_nolock + code path is only used for extent-based files. + Because of the restrictions this options comprises + it is off by default (e.g. dioread_lock). + +i_version Enable 64-bit inode version support. This option is + off by default. + +Data Mode +========= +There are 3 different data modes: + +* writeback mode +In data=writeback mode, ext4 does not journal data at all. This mode provides +a similar level of journaling as that of XFS, JFS, and ReiserFS in its default +mode - metadata journaling. A crash+recovery can cause incorrect data to +appear in files which were written shortly before the crash. This mode will +typically provide the best ext4 performance. + +* ordered mode +In data=ordered mode, ext4 only officially journals metadata, but it logically +groups metadata information related to data changes with the data blocks into a +single unit called a transaction. When it's time to write the new metadata +out to disk, the associated data blocks are written first. In general, +this mode performs slightly slower than writeback but significantly faster than journal mode. + +* journal mode +data=journal mode provides full data and metadata journaling. All new data is +written to the journal first, and then to its final location. +In the event of a crash, the journal can be replayed, bringing both data and +metadata into a consistent state. This mode is the slowest except when data +needs to be read from and written to disk at the same time where it +outperforms all others modes. Currently ext4 does not have delayed +allocation support if this data journalling mode is selected. + +/proc entries +============= + +Information about mounted ext4 file systems can be found in +/proc/fs/ext4. Each mounted filesystem will have a directory in +/proc/fs/ext4 based on its device name (i.e., /proc/fs/ext4/hdc or +/proc/fs/ext4/dm-0). The files in each per-device directory are shown +in table below. + +Files in /proc/fs/ext4/<devname> +.............................................................................. + File Content + mb_groups details of multiblock allocator buddy cache of free blocks +.............................................................................. + +/sys entries +============ + +Information about mounted ext4 file systems can be found in +/sys/fs/ext4. Each mounted filesystem will have a directory in +/sys/fs/ext4 based on its device name (i.e., /sys/fs/ext4/hdc or +/sys/fs/ext4/dm-0). The files in each per-device directory are shown +in table below. + +Files in /sys/fs/ext4/<devname> +(see also Documentation/ABI/testing/sysfs-fs-ext4) +.............................................................................. + File Content + + delayed_allocation_blocks This file is read-only and shows the number of + blocks that are dirty in the page cache, but + which do not have their location in the + filesystem allocated yet. + + inode_goal Tuning parameter which (if non-zero) controls + the goal inode used by the inode allocator in + preference to all other allocation heuristics. + This is intended for debugging use only, and + should be 0 on production systems. + + inode_readahead_blks Tuning parameter which controls the maximum + number of inode table blocks that ext4's inode + table readahead algorithm will pre-read into + the buffer cache + + lifetime_write_kbytes This file is read-only and shows the number of + kilobytes of data that have been written to this + filesystem since it was created. + + max_writeback_mb_bump The maximum number of megabytes the writeback + code will try to write out before move on to + another inode. + + mb_group_prealloc The multiblock allocator will round up allocation + requests to a multiple of this tuning parameter if + the stripe size is not set in the ext4 superblock + + mb_max_to_scan The maximum number of extents the multiblock + allocator will search to find the best extent + + mb_min_to_scan The minimum number of extents the multiblock + allocator will search to find the best extent + + mb_order2_req Tuning parameter which controls the minimum size + for requests (as a power of 2) where the buddy + cache is used + + mb_stats Controls whether the multiblock allocator should + collect statistics, which are shown during the + unmount. 1 means to collect statistics, 0 means + not to collect statistics + + mb_stream_req Files which have fewer blocks than this tunable + parameter will have their blocks allocated out + of a block group specific preallocation pool, so + that small files are packed closely together. + Each large file will have its blocks allocated + out of its own unique preallocation pool. + + session_write_kbytes This file is read-only and shows the number of + kilobytes of data that have been written to this + filesystem since it was mounted. +.............................................................................. + +Ioctls +====== + +There is some Ext4 specific functionality which can be accessed by applications +through the system call interfaces. The list of all Ext4 specific ioctls are +shown in the table below. + +Table of Ext4 specific ioctls +.............................................................................. + Ioctl Description + EXT4_IOC_GETFLAGS Get additional attributes associated with inode. + The ioctl argument is an integer bitfield, with + bit values described in ext4.h. This ioctl is an + alias for FS_IOC_GETFLAGS. + + EXT4_IOC_SETFLAGS Set additional attributes associated with inode. + The ioctl argument is an integer bitfield, with + bit values described in ext4.h. This ioctl is an + alias for FS_IOC_SETFLAGS. + + EXT4_IOC_GETVERSION + EXT4_IOC_GETVERSION_OLD + Get the inode i_generation number stored for + each inode. The i_generation number is normally + changed only when new inode is created and it is + particularly useful for network filesystems. The + '_OLD' version of this ioctl is an alias for + FS_IOC_GETVERSION. + + EXT4_IOC_SETVERSION + EXT4_IOC_SETVERSION_OLD + Set the inode i_generation number stored for + each inode. The '_OLD' version of this ioctl + is an alias for FS_IOC_SETVERSION. + + EXT4_IOC_GROUP_EXTEND This ioctl has the same purpose as the resize + mount option. It allows to resize filesystem + to the end of the last existing block group, + further resize has to be done with resize2fs, + either online, or offline. The argument points + to the unsigned logn number representing the + filesystem new block count. + + EXT4_IOC_MOVE_EXT Move the block extents from orig_fd (the one + this ioctl is pointing to) to the donor_fd (the + one specified in move_extent structure passed + as an argument to this ioctl). Then, exchange + inode metadata between orig_fd and donor_fd. + This is especially useful for online + defragmentation, because the allocator has the + opportunity to allocate moved blocks better, + ideally into one contiguous extent. + + EXT4_IOC_GROUP_ADD Add a new group descriptor to an existing or + new group descriptor block. The new group + descriptor is described by ext4_new_group_input + structure, which is passed as an argument to + this ioctl. This is especially useful in + conjunction with EXT4_IOC_GROUP_EXTEND, + which allows online resize of the filesystem + to the end of the last existing block group. + Those two ioctls combined is used in userspace + online resize tool (e.g. resize2fs). + + EXT4_IOC_MIGRATE This ioctl operates on the filesystem itself. + It converts (migrates) ext3 indirect block mapped + inode to ext4 extent mapped inode by walking + through indirect block mapping of the original + inode and converting contiguous block ranges + into ext4 extents of the temporary inode. Then, + inodes are swapped. This ioctl might help, when + migrating from ext3 to ext4 filesystem, however + suggestion is to create fresh ext4 filesystem + and copy data from the backup. Note, that + filesystem has to support extents for this ioctl + to work. + + EXT4_IOC_ALLOC_DA_BLKS Force all of the delay allocated blocks to be + allocated to preserve application-expected ext3 + behaviour. Note that this will also start + triggering a write of the data blocks, but this + behaviour may change in the future as it is + not necessary and has been done this way only + for sake of simplicity. +.............................................................................. + +References +========== + +kernel source: <file:fs/ext4/> + <file:fs/jbd2/> + +programs: http://e2fsprogs.sourceforge.net/ + +useful links: http://fedoraproject.org/wiki/ext3-devel + http://www.bullopensource.org/ext4/ + http://ext4.wiki.kernel.org/index.php/Main_Page + http://fedoraproject.org/wiki/Features/Ext4 diff --git a/Documentation/filesystems/fiemap.txt b/Documentation/filesystems/fiemap.txt new file mode 100644 index 00000000..1b805a0e --- /dev/null +++ b/Documentation/filesystems/fiemap.txt @@ -0,0 +1,228 @@ +============ +Fiemap Ioctl +============ + +The fiemap ioctl is an efficient method for userspace to get file +extent mappings. Instead of block-by-block mapping (such as bmap), fiemap +returns a list of extents. + + +Request Basics +-------------- + +A fiemap request is encoded within struct fiemap: + +struct fiemap { + __u64 fm_start; /* logical offset (inclusive) at + * which to start mapping (in) */ + __u64 fm_length; /* logical length of mapping which + * userspace cares about (in) */ + __u32 fm_flags; /* FIEMAP_FLAG_* flags for request (in/out) */ + __u32 fm_mapped_extents; /* number of extents that were + * mapped (out) */ + __u32 fm_extent_count; /* size of fm_extents array (in) */ + __u32 fm_reserved; + struct fiemap_extent fm_extents[0]; /* array of mapped extents (out) */ +}; + + +fm_start, and fm_length specify the logical range within the file +which the process would like mappings for. Extents returned mirror +those on disk - that is, the logical offset of the 1st returned extent +may start before fm_start, and the range covered by the last returned +extent may end after fm_length. All offsets and lengths are in bytes. + +Certain flags to modify the way in which mappings are looked up can be +set in fm_flags. If the kernel doesn't understand some particular +flags, it will return EBADR and the contents of fm_flags will contain +the set of flags which caused the error. If the kernel is compatible +with all flags passed, the contents of fm_flags will be unmodified. +It is up to userspace to determine whether rejection of a particular +flag is fatal to its operation. This scheme is intended to allow the +fiemap interface to grow in the future but without losing +compatibility with old software. + +fm_extent_count specifies the number of elements in the fm_extents[] array +that can be used to return extents. If fm_extent_count is zero, then the +fm_extents[] array is ignored (no extents will be returned), and the +fm_mapped_extents count will hold the number of extents needed in +fm_extents[] to hold the file's current mapping. Note that there is +nothing to prevent the file from changing between calls to FIEMAP. + +The following flags can be set in fm_flags: + +* FIEMAP_FLAG_SYNC +If this flag is set, the kernel will sync the file before mapping extents. + +* FIEMAP_FLAG_XATTR +If this flag is set, the extents returned will describe the inodes +extended attribute lookup tree, instead of its data tree. + + +Extent Mapping +-------------- + +Extent information is returned within the embedded fm_extents array +which userspace must allocate along with the fiemap structure. The +number of elements in the fiemap_extents[] array should be passed via +fm_extent_count. The number of extents mapped by kernel will be +returned via fm_mapped_extents. If the number of fiemap_extents +allocated is less than would be required to map the requested range, +the maximum number of extents that can be mapped in the fm_extent[] +array will be returned and fm_mapped_extents will be equal to +fm_extent_count. In that case, the last extent in the array will not +complete the requested range and will not have the FIEMAP_EXTENT_LAST +flag set (see the next section on extent flags). + +Each extent is described by a single fiemap_extent structure as +returned in fm_extents. + +struct fiemap_extent { + __u64 fe_logical; /* logical offset in bytes for the start of + * the extent */ + __u64 fe_physical; /* physical offset in bytes for the start + * of the extent */ + __u64 fe_length; /* length in bytes for the extent */ + __u64 fe_reserved64[2]; + __u32 fe_flags; /* FIEMAP_EXTENT_* flags for this extent */ + __u32 fe_reserved[3]; +}; + +All offsets and lengths are in bytes and mirror those on disk. It is valid +for an extents logical offset to start before the request or its logical +length to extend past the request. Unless FIEMAP_EXTENT_NOT_ALIGNED is +returned, fe_logical, fe_physical, and fe_length will be aligned to the +block size of the file system. With the exception of extents flagged as +FIEMAP_EXTENT_MERGED, adjacent extents will not be merged. + +The fe_flags field contains flags which describe the extent returned. +A special flag, FIEMAP_EXTENT_LAST is always set on the last extent in +the file so that the process making fiemap calls can determine when no +more extents are available, without having to call the ioctl again. + +Some flags are intentionally vague and will always be set in the +presence of other more specific flags. This way a program looking for +a general property does not have to know all existing and future flags +which imply that property. + +For example, if FIEMAP_EXTENT_DATA_INLINE or FIEMAP_EXTENT_DATA_TAIL +are set, FIEMAP_EXTENT_NOT_ALIGNED will also be set. A program looking +for inline or tail-packed data can key on the specific flag. Software +which simply cares not to try operating on non-aligned extents +however, can just key on FIEMAP_EXTENT_NOT_ALIGNED, and not have to +worry about all present and future flags which might imply unaligned +data. Note that the opposite is not true - it would be valid for +FIEMAP_EXTENT_NOT_ALIGNED to appear alone. + +* FIEMAP_EXTENT_LAST +This is the last extent in the file. A mapping attempt past this +extent will return nothing. + +* FIEMAP_EXTENT_UNKNOWN +The location of this extent is currently unknown. This may indicate +the data is stored on an inaccessible volume or that no storage has +been allocated for the file yet. + +* FIEMAP_EXTENT_DELALLOC + - This will also set FIEMAP_EXTENT_UNKNOWN. +Delayed allocation - while there is data for this extent, its +physical location has not been allocated yet. + +* FIEMAP_EXTENT_ENCODED +This extent does not consist of plain filesystem blocks but is +encoded (e.g. encrypted or compressed). Reading the data in this +extent via I/O to the block device will have undefined results. + +Note that it is *always* undefined to try to update the data +in-place by writing to the indicated location without the +assistance of the filesystem, or to access the data using the +information returned by the FIEMAP interface while the filesystem +is mounted. In other words, user applications may only read the +extent data via I/O to the block device while the filesystem is +unmounted, and then only if the FIEMAP_EXTENT_ENCODED flag is +clear; user applications must not try reading or writing to the +filesystem via the block device under any other circumstances. + +* FIEMAP_EXTENT_DATA_ENCRYPTED + - This will also set FIEMAP_EXTENT_ENCODED +The data in this extent has been encrypted by the file system. + +* FIEMAP_EXTENT_NOT_ALIGNED +Extent offsets and length are not guaranteed to be block aligned. + +* FIEMAP_EXTENT_DATA_INLINE + This will also set FIEMAP_EXTENT_NOT_ALIGNED +Data is located within a meta data block. + +* FIEMAP_EXTENT_DATA_TAIL + This will also set FIEMAP_EXTENT_NOT_ALIGNED +Data is packed into a block with data from other files. + +* FIEMAP_EXTENT_UNWRITTEN +Unwritten extent - the extent is allocated but its data has not been +initialized. This indicates the extent's data will be all zero if read +through the filesystem but the contents are undefined if read directly from +the device. + +* FIEMAP_EXTENT_MERGED +This will be set when a file does not support extents, i.e., it uses a block +based addressing scheme. Since returning an extent for each block back to +userspace would be highly inefficient, the kernel will try to merge most +adjacent blocks into 'extents'. + + +VFS -> File System Implementation +--------------------------------- + +File systems wishing to support fiemap must implement a ->fiemap callback on +their inode_operations structure. The fs ->fiemap call is responsible for +defining its set of supported fiemap flags, and calling a helper function on +each discovered extent: + +struct inode_operations { + ... + + int (*fiemap)(struct inode *, struct fiemap_extent_info *, u64 start, + u64 len); + +->fiemap is passed struct fiemap_extent_info which describes the +fiemap request: + +struct fiemap_extent_info { + unsigned int fi_flags; /* Flags as passed from user */ + unsigned int fi_extents_mapped; /* Number of mapped extents */ + unsigned int fi_extents_max; /* Size of fiemap_extent array */ + struct fiemap_extent *fi_extents_start; /* Start of fiemap_extent array */ +}; + +It is intended that the file system should not need to access any of this +structure directly. + + +Flag checking should be done at the beginning of the ->fiemap callback via the +fiemap_check_flags() helper: + +int fiemap_check_flags(struct fiemap_extent_info *fieinfo, u32 fs_flags); + +The struct fieinfo should be passed in as received from ioctl_fiemap(). The +set of fiemap flags which the fs understands should be passed via fs_flags. If +fiemap_check_flags finds invalid user flags, it will place the bad values in +fieinfo->fi_flags and return -EBADR. If the file system gets -EBADR, from +fiemap_check_flags(), it should immediately exit, returning that error back to +ioctl_fiemap(). + + +For each extent in the request range, the file system should call +the helper function, fiemap_fill_next_extent(): + +int fiemap_fill_next_extent(struct fiemap_extent_info *info, u64 logical, + u64 phys, u64 len, u32 flags, u32 dev); + +fiemap_fill_next_extent() will use the passed values to populate the +next free extent in the fm_extents array. 'General' extent flags will +automatically be set from specific flags on behalf of the calling file +system so that the userspace API is not broken. + +fiemap_fill_next_extent() returns 0 on success, and 1 when the +user-supplied fm_extents array is full. If an error is encountered +while copying the extent to user memory, -EFAULT will be returned. diff --git a/Documentation/filesystems/files.txt b/Documentation/filesystems/files.txt new file mode 100644 index 00000000..ac2facc5 --- /dev/null +++ b/Documentation/filesystems/files.txt @@ -0,0 +1,123 @@ +File management in the Linux kernel +----------------------------------- + +This document describes how locking for files (struct file) +and file descriptor table (struct files) works. + +Up until 2.6.12, the file descriptor table has been protected +with a lock (files->file_lock) and reference count (files->count). +->file_lock protected accesses to all the file related fields +of the table. ->count was used for sharing the file descriptor +table between tasks cloned with CLONE_FILES flag. Typically +this would be the case for posix threads. As with the common +refcounting model in the kernel, the last task doing +a put_files_struct() frees the file descriptor (fd) table. +The files (struct file) themselves are protected using +reference count (->f_count). + +In the new lock-free model of file descriptor management, +the reference counting is similar, but the locking is +based on RCU. The file descriptor table contains multiple +elements - the fd sets (open_fds and close_on_exec, the +array of file pointers, the sizes of the sets and the array +etc.). In order for the updates to appear atomic to +a lock-free reader, all the elements of the file descriptor +table are in a separate structure - struct fdtable. +files_struct contains a pointer to struct fdtable through +which the actual fd table is accessed. Initially the +fdtable is embedded in files_struct itself. On a subsequent +expansion of fdtable, a new fdtable structure is allocated +and files->fdtab points to the new structure. The fdtable +structure is freed with RCU and lock-free readers either +see the old fdtable or the new fdtable making the update +appear atomic. Here are the locking rules for +the fdtable structure - + +1. All references to the fdtable must be done through + the files_fdtable() macro : + + struct fdtable *fdt; + + rcu_read_lock(); + + fdt = files_fdtable(files); + .... + if (n <= fdt->max_fds) + .... + ... + rcu_read_unlock(); + + files_fdtable() uses rcu_dereference() macro which takes care of + the memory barrier requirements for lock-free dereference. + The fdtable pointer must be read within the read-side + critical section. + +2. Reading of the fdtable as described above must be protected + by rcu_read_lock()/rcu_read_unlock(). + +3. For any update to the fd table, files->file_lock must + be held. + +4. To look up the file structure given an fd, a reader + must use either fcheck() or fcheck_files() APIs. These + take care of barrier requirements due to lock-free lookup. + An example : + + struct file *file; + + rcu_read_lock(); + file = fcheck(fd); + if (file) { + ... + } + .... + rcu_read_unlock(); + +5. Handling of the file structures is special. Since the look-up + of the fd (fget()/fget_light()) are lock-free, it is possible + that look-up may race with the last put() operation on the + file structure. This is avoided using atomic_long_inc_not_zero() + on ->f_count : + + rcu_read_lock(); + file = fcheck_files(files, fd); + if (file) { + if (atomic_long_inc_not_zero(&file->f_count)) + *fput_needed = 1; + else + /* Didn't get the reference, someone's freed */ + file = NULL; + } + rcu_read_unlock(); + .... + return file; + + atomic_long_inc_not_zero() detects if refcounts is already zero or + goes to zero during increment. If it does, we fail + fget()/fget_light(). + +6. Since both fdtable and file structures can be looked up + lock-free, they must be installed using rcu_assign_pointer() + API. If they are looked up lock-free, rcu_dereference() + must be used. However it is advisable to use files_fdtable() + and fcheck()/fcheck_files() which take care of these issues. + +7. While updating, the fdtable pointer must be looked up while + holding files->file_lock. If ->file_lock is dropped, then + another thread expand the files thereby creating a new + fdtable and making the earlier fdtable pointer stale. + For example : + + spin_lock(&files->file_lock); + fd = locate_fd(files, file, start); + if (fd >= 0) { + /* locate_fd() may have expanded fdtable, load the ptr */ + fdt = files_fdtable(files); + FD_SET(fd, fdt->open_fds); + FD_CLR(fd, fdt->close_on_exec); + spin_unlock(&files->file_lock); + ..... + + Since locate_fd() can drop ->file_lock (and reacquire ->file_lock), + the fdtable pointer (fdt) must be loaded after locate_fd(). + diff --git a/Documentation/filesystems/fuse.txt b/Documentation/filesystems/fuse.txt new file mode 100644 index 00000000..13af4a49 --- /dev/null +++ b/Documentation/filesystems/fuse.txt @@ -0,0 +1,423 @@ +Definitions +~~~~~~~~~~~ + +Userspace filesystem: + + A filesystem in which data and metadata are provided by an ordinary + userspace process. The filesystem can be accessed normally through + the kernel interface. + +Filesystem daemon: + + The process(es) providing the data and metadata of the filesystem. + +Non-privileged mount (or user mount): + + A userspace filesystem mounted by a non-privileged (non-root) user. + The filesystem daemon is running with the privileges of the mounting + user. NOTE: this is not the same as mounts allowed with the "user" + option in /etc/fstab, which is not discussed here. + +Filesystem connection: + + A connection between the filesystem daemon and the kernel. The + connection exists until either the daemon dies, or the filesystem is + umounted. Note that detaching (or lazy umounting) the filesystem + does _not_ break the connection, in this case it will exist until + the last reference to the filesystem is released. + +Mount owner: + + The user who does the mounting. + +User: + + The user who is performing filesystem operations. + +What is FUSE? +~~~~~~~~~~~~~ + +FUSE is a userspace filesystem framework. It consists of a kernel +module (fuse.ko), a userspace library (libfuse.*) and a mount utility +(fusermount). + +One of the most important features of FUSE is allowing secure, +non-privileged mounts. This opens up new possibilities for the use of +filesystems. A good example is sshfs: a secure network filesystem +using the sftp protocol. + +The userspace library and utilities are available from the FUSE +homepage: + + http://fuse.sourceforge.net/ + +Filesystem type +~~~~~~~~~~~~~~~ + +The filesystem type given to mount(2) can be one of the following: + +'fuse' + + This is the usual way to mount a FUSE filesystem. The first + argument of the mount system call may contain an arbitrary string, + which is not interpreted by the kernel. + +'fuseblk' + + The filesystem is block device based. The first argument of the + mount system call is interpreted as the name of the device. + +Mount options +~~~~~~~~~~~~~ + +'fd=N' + + The file descriptor to use for communication between the userspace + filesystem and the kernel. The file descriptor must have been + obtained by opening the FUSE device ('/dev/fuse'). + +'rootmode=M' + + The file mode of the filesystem's root in octal representation. + +'user_id=N' + + The numeric user id of the mount owner. + +'group_id=N' + + The numeric group id of the mount owner. + +'default_permissions' + + By default FUSE doesn't check file access permissions, the + filesystem is free to implement its access policy or leave it to + the underlying file access mechanism (e.g. in case of network + filesystems). This option enables permission checking, restricting + access based on file mode. It is usually useful together with the + 'allow_other' mount option. + +'allow_other' + + This option overrides the security measure restricting file access + to the user mounting the filesystem. This option is by default only + allowed to root, but this restriction can be removed with a + (userspace) configuration option. + +'max_read=N' + + With this option the maximum size of read operations can be set. + The default is infinite. Note that the size of read requests is + limited anyway to 32 pages (which is 128kbyte on i386). + +'blksize=N' + + Set the block size for the filesystem. The default is 512. This + option is only valid for 'fuseblk' type mounts. + +Control filesystem +~~~~~~~~~~~~~~~~~~ + +There's a control filesystem for FUSE, which can be mounted by: + + mount -t fusectl none /sys/fs/fuse/connections + +Mounting it under the '/sys/fs/fuse/connections' directory makes it +backwards compatible with earlier versions. + +Under the fuse control filesystem each connection has a directory +named by a unique number. + +For each connection the following files exist within this directory: + + 'waiting' + + The number of requests which are waiting to be transferred to + userspace or being processed by the filesystem daemon. If there is + no filesystem activity and 'waiting' is non-zero, then the + filesystem is hung or deadlocked. + + 'abort' + + Writing anything into this file will abort the filesystem + connection. This means that all waiting requests will be aborted an + error returned for all aborted and new requests. + +Only the owner of the mount may read or write these files. + +Interrupting filesystem operations +~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + +If a process issuing a FUSE filesystem request is interrupted, the +following will happen: + + 1) If the request is not yet sent to userspace AND the signal is + fatal (SIGKILL or unhandled fatal signal), then the request is + dequeued and returns immediately. + + 2) If the request is not yet sent to userspace AND the signal is not + fatal, then an 'interrupted' flag is set for the request. When + the request has been successfully transferred to userspace and + this flag is set, an INTERRUPT request is queued. + + 3) If the request is already sent to userspace, then an INTERRUPT + request is queued. + +INTERRUPT requests take precedence over other requests, so the +userspace filesystem will receive queued INTERRUPTs before any others. + +The userspace filesystem may ignore the INTERRUPT requests entirely, +or may honor them by sending a reply to the _original_ request, with +the error set to EINTR. + +It is also possible that there's a race between processing the +original request and its INTERRUPT request. There are two possibilities: + + 1) The INTERRUPT request is processed before the original request is + processed + + 2) The INTERRUPT request is processed after the original request has + been answered + +If the filesystem cannot find the original request, it should wait for +some timeout and/or a number of new requests to arrive, after which it +should reply to the INTERRUPT request with an EAGAIN error. In case +1) the INTERRUPT request will be requeued. In case 2) the INTERRUPT +reply will be ignored. + +Aborting a filesystem connection +~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + +It is possible to get into certain situations where the filesystem is +not responding. Reasons for this may be: + + a) Broken userspace filesystem implementation + + b) Network connection down + + c) Accidental deadlock + + d) Malicious deadlock + +(For more on c) and d) see later sections) + +In either of these cases it may be useful to abort the connection to +the filesystem. There are several ways to do this: + + - Kill the filesystem daemon. Works in case of a) and b) + + - Kill the filesystem daemon and all users of the filesystem. Works + in all cases except some malicious deadlocks + + - Use forced umount (umount -f). Works in all cases but only if + filesystem is still attached (it hasn't been lazy unmounted) + + - Abort filesystem through the FUSE control filesystem. Most + powerful method, always works. + +How do non-privileged mounts work? +~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + +Since the mount() system call is a privileged operation, a helper +program (fusermount) is needed, which is installed setuid root. + +The implication of providing non-privileged mounts is that the mount +owner must not be able to use this capability to compromise the +system. Obvious requirements arising from this are: + + A) mount owner should not be able to get elevated privileges with the + help of the mounted filesystem + + B) mount owner should not get illegitimate access to information from + other users' and the super user's processes + + C) mount owner should not be able to induce undesired behavior in + other users' or the super user's processes + +How are requirements fulfilled? +~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + + A) The mount owner could gain elevated privileges by either: + + 1) creating a filesystem containing a device file, then opening + this device + + 2) creating a filesystem containing a suid or sgid application, + then executing this application + + The solution is not to allow opening device files and ignore + setuid and setgid bits when executing programs. To ensure this + fusermount always adds "nosuid" and "nodev" to the mount options + for non-privileged mounts. + + B) If another user is accessing files or directories in the + filesystem, the filesystem daemon serving requests can record the + exact sequence and timing of operations performed. This + information is otherwise inaccessible to the mount owner, so this + counts as an information leak. + + The solution to this problem will be presented in point 2) of C). + + C) There are several ways in which the mount owner can induce + undesired behavior in other users' processes, such as: + + 1) mounting a filesystem over a file or directory which the mount + owner could otherwise not be able to modify (or could only + make limited modifications). + + This is solved in fusermount, by checking the access + permissions on the mountpoint and only allowing the mount if + the mount owner can do unlimited modification (has write + access to the mountpoint, and mountpoint is not a "sticky" + directory) + + 2) Even if 1) is solved the mount owner can change the behavior + of other users' processes. + + i) It can slow down or indefinitely delay the execution of a + filesystem operation creating a DoS against the user or the + whole system. For example a suid application locking a + system file, and then accessing a file on the mount owner's + filesystem could be stopped, and thus causing the system + file to be locked forever. + + ii) It can present files or directories of unlimited length, or + directory structures of unlimited depth, possibly causing a + system process to eat up diskspace, memory or other + resources, again causing DoS. + + The solution to this as well as B) is not to allow processes + to access the filesystem, which could otherwise not be + monitored or manipulated by the mount owner. Since if the + mount owner can ptrace a process, it can do all of the above + without using a FUSE mount, the same criteria as used in + ptrace can be used to check if a process is allowed to access + the filesystem or not. + + Note that the ptrace check is not strictly necessary to + prevent B/2/i, it is enough to check if mount owner has enough + privilege to send signal to the process accessing the + filesystem, since SIGSTOP can be used to get a similar effect. + +I think these limitations are unacceptable? +~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + +If a sysadmin trusts the users enough, or can ensure through other +measures, that system processes will never enter non-privileged +mounts, it can relax the last limitation with a "user_allow_other" +config option. If this config option is set, the mounting user can +add the "allow_other" mount option which disables the check for other +users' processes. + +Kernel - userspace interface +~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + +The following diagram shows how a filesystem operation (in this +example unlink) is performed in FUSE. + +NOTE: everything in this description is greatly simplified + + | "rm /mnt/fuse/file" | FUSE filesystem daemon + | | + | | >sys_read() + | | >fuse_dev_read() + | | >request_wait() + | | [sleep on fc->waitq] + | | + | >sys_unlink() | + | >fuse_unlink() | + | [get request from | + | fc->unused_list] | + | >request_send() | + | [queue req on fc->pending] | + | [wake up fc->waitq] | [woken up] + | >request_wait_answer() | + | [sleep on req->waitq] | + | | <request_wait() + | | [remove req from fc->pending] + | | [copy req to read buffer] + | | [add req to fc->processing] + | | <fuse_dev_read() + | | <sys_read() + | | + | | [perform unlink] + | | + | | >sys_write() + | | >fuse_dev_write() + | | [look up req in fc->processing] + | | [remove from fc->processing] + | | [copy write buffer to req] + | [woken up] | [wake up req->waitq] + | | <fuse_dev_write() + | | <sys_write() + | <request_wait_answer() | + | <request_send() | + | [add request to | + | fc->unused_list] | + | <fuse_unlink() | + | <sys_unlink() | + +There are a couple of ways in which to deadlock a FUSE filesystem. +Since we are talking about unprivileged userspace programs, +something must be done about these. + +Scenario 1 - Simple deadlock +----------------------------- + + | "rm /mnt/fuse/file" | FUSE filesystem daemon + | | + | >sys_unlink("/mnt/fuse/file") | + | [acquire inode semaphore | + | for "file"] | + | >fuse_unlink() | + | [sleep on req->waitq] | + | | <sys_read() + | | >sys_unlink("/mnt/fuse/file") + | | [acquire inode semaphore + | | for "file"] + | | *DEADLOCK* + +The solution for this is to allow the filesystem to be aborted. + +Scenario 2 - Tricky deadlock +---------------------------- + +This one needs a carefully crafted filesystem. It's a variation on +the above, only the call back to the filesystem is not explicit, +but is caused by a pagefault. + + | Kamikaze filesystem thread 1 | Kamikaze filesystem thread 2 + | | + | [fd = open("/mnt/fuse/file")] | [request served normally] + | [mmap fd to 'addr'] | + | [close fd] | [FLUSH triggers 'magic' flag] + | [read a byte from addr] | + | >do_page_fault() | + | [find or create page] | + | [lock page] | + | >fuse_readpage() | + | [queue READ request] | + | [sleep on req->waitq] | + | | [read request to buffer] + | | [create reply header before addr] + | | >sys_write(addr - headerlength) + | | >fuse_dev_write() + | | [look up req in fc->processing] + | | [remove from fc->processing] + | | [copy write buffer to req] + | | >do_page_fault() + | | [find or create page] + | | [lock page] + | | * DEADLOCK * + +Solution is basically the same as above. + +An additional problem is that while the write buffer is being copied +to the request, the request must not be interrupted/aborted. This is +because the destination address of the copy may not be valid after the +request has returned. + +This is solved with doing the copy atomically, and allowing abort +while the page(s) belonging to the write buffer are faulted with +get_user_pages(). The 'req->locked' flag indicates when the copy is +taking place, and abort is delayed until this flag is unset. diff --git a/Documentation/filesystems/gfs2-glocks.txt b/Documentation/filesystems/gfs2-glocks.txt new file mode 100644 index 00000000..0494f78d --- /dev/null +++ b/Documentation/filesystems/gfs2-glocks.txt @@ -0,0 +1,114 @@ + Glock internal locking rules + ------------------------------ + +This documents the basic principles of the glock state machine +internals. Each glock (struct gfs2_glock in fs/gfs2/incore.h) +has two main (internal) locks: + + 1. A spinlock (gl_spin) which protects the internal state such + as gl_state, gl_target and the list of holders (gl_holders) + 2. A non-blocking bit lock, GLF_LOCK, which is used to prevent other + threads from making calls to the DLM, etc. at the same time. If a + thread takes this lock, it must then call run_queue (usually via the + workqueue) when it releases it in order to ensure any pending tasks + are completed. + +The gl_holders list contains all the queued lock requests (not +just the holders) associated with the glock. If there are any +held locks, then they will be contiguous entries at the head +of the list. Locks are granted in strictly the order that they +are queued, except for those marked LM_FLAG_PRIORITY which are +used only during recovery, and even then only for journal locks. + +There are three lock states that users of the glock layer can request, +namely shared (SH), deferred (DF) and exclusive (EX). Those translate +to the following DLM lock modes: + +Glock mode | DLM lock mode +------------------------------ + UN | IV/NL Unlocked (no DLM lock associated with glock) or NL + SH | PR (Protected read) + DF | CW (Concurrent write) + EX | EX (Exclusive) + +Thus DF is basically a shared mode which is incompatible with the "normal" +shared lock mode, SH. In GFS2 the DF mode is used exclusively for direct I/O +operations. The glocks are basically a lock plus some routines which deal +with cache management. The following rules apply for the cache: + +Glock mode | Cache data | Cache Metadata | Dirty Data | Dirty Metadata +-------------------------------------------------------------------------- + UN | No | No | No | No + SH | Yes | Yes | No | No + DF | No | Yes | No | No + EX | Yes | Yes | Yes | Yes + +These rules are implemented using the various glock operations which +are defined for each type of glock. Not all types of glocks use +all the modes. Only inode glocks use the DF mode for example. + +Table of glock operations and per type constants: + +Field | Purpose +---------------------------------------------------------------------------- +go_xmote_th | Called before remote state change (e.g. to sync dirty data) +go_xmote_bh | Called after remote state change (e.g. to refill cache) +go_inval | Called if remote state change requires invalidating the cache +go_demote_ok | Returns boolean value of whether its ok to demote a glock + | (e.g. checks timeout, and that there is no cached data) +go_lock | Called for the first local holder of a lock +go_unlock | Called on the final local unlock of a lock +go_dump | Called to print content of object for debugfs file, or on + | error to dump glock to the log. +go_type | The type of the glock, LM_TYPE_..... +go_min_hold_time | The minimum hold time + +The minimum hold time for each lock is the time after a remote lock +grant for which we ignore remote demote requests. This is in order to +prevent a situation where locks are being bounced around the cluster +from node to node with none of the nodes making any progress. This +tends to show up most with shared mmaped files which are being written +to by multiple nodes. By delaying the demotion in response to a +remote callback, that gives the userspace program time to make +some progress before the pages are unmapped. + +There is a plan to try and remove the go_lock and go_unlock callbacks +if possible, in order to try and speed up the fast path though the locking. +Also, eventually we hope to make the glock "EX" mode locally shared +such that any local locking will be done with the i_mutex as required +rather than via the glock. + +Locking rules for glock operations: + +Operation | GLF_LOCK bit lock held | gl_spin spinlock held +----------------------------------------------------------------- +go_xmote_th | Yes | No +go_xmote_bh | Yes | No +go_inval | Yes | No +go_demote_ok | Sometimes | Yes +go_lock | Yes | No +go_unlock | Yes | No +go_dump | Sometimes | Yes + +N.B. Operations must not drop either the bit lock or the spinlock +if its held on entry. go_dump and do_demote_ok must never block. +Note that go_dump will only be called if the glock's state +indicates that it is caching uptodate data. + +Glock locking order within GFS2: + + 1. i_mutex (if required) + 2. Rename glock (for rename only) + 3. Inode glock(s) + (Parents before children, inodes at "same level" with same parent in + lock number order) + 4. Rgrp glock(s) (for (de)allocation operations) + 5. Transaction glock (via gfs2_trans_begin) for non-read operations + 6. Page lock (always last, very important!) + +There are two glocks per inode. One deals with access to the inode +itself (locking order as above), and the other, known as the iopen +glock is used in conjunction with the i_nlink field in the inode to +determine the lifetime of the inode in question. Locking of inodes +is on a per-inode basis. Locking of rgrps is on a per rgrp basis. + diff --git a/Documentation/filesystems/gfs2-uevents.txt b/Documentation/filesystems/gfs2-uevents.txt new file mode 100644 index 00000000..d8188966 --- /dev/null +++ b/Documentation/filesystems/gfs2-uevents.txt @@ -0,0 +1,100 @@ + uevents and GFS2 + ================== + +During the lifetime of a GFS2 mount, a number of uevents are generated. +This document explains what the events are and what they are used +for (by gfs_controld in gfs2-utils). + +A list of GFS2 uevents +----------------------- + +1. ADD + +The ADD event occurs at mount time. It will always be the first +uevent generated by the newly created filesystem. If the mount +is successful, an ONLINE uevent will follow. If it is not successful +then a REMOVE uevent will follow. + +The ADD uevent has two environment variables: SPECTATOR=[0|1] +and RDONLY=[0|1] that specify the spectator status (a read-only mount +with no journal assigned), and read-only (with journal assigned) status +of the filesystem respectively. + +2. ONLINE + +The ONLINE uevent is generated after a successful mount or remount. It +has the same environment variables as the ADD uevent. The ONLINE +uevent, along with the two environment variables for spectator and +RDONLY are a relatively recent addition (2.6.32-rc+) and will not +be generated by older kernels. + +3. CHANGE + +The CHANGE uevent is used in two places. One is when reporting the +successful mount of the filesystem by the first node (FIRSTMOUNT=Done). +This is used as a signal by gfs_controld that it is then ok for other +nodes in the cluster to mount the filesystem. + +The other CHANGE uevent is used to inform of the completion +of journal recovery for one of the filesystems journals. It has +two environment variables, JID= which specifies the journal id which +has just been recovered, and RECOVERY=[Done|Failed] to indicate the +success (or otherwise) of the operation. These uevents are generated +for every journal recovered, whether it is during the initial mount +process or as the result of gfs_controld requesting a specific journal +recovery via the /sys/fs/gfs2/<fsname>/lock_module/recovery file. + +Because the CHANGE uevent was used (in early versions of gfs_controld) +without checking the environment variables to discover the state, we +cannot add any more functions to it without running the risk of +someone using an older version of the user tools and breaking their +cluster. For this reason the ONLINE uevent was used when adding a new +uevent for a successful mount or remount. + +4. OFFLINE + +The OFFLINE uevent is only generated due to filesystem errors and is used +as part of the "withdraw" mechanism. Currently this doesn't give any +information about what the error is, which is something that needs to +be fixed. + +5. REMOVE + +The REMOVE uevent is generated at the end of an unsuccessful mount +or at the end of a umount of the filesystem. All REMOVE uevents will +have been preceded by at least an ADD uevent for the same fileystem, +and unlike the other uevents is generated automatically by the kernel's +kobject subsystem. + + +Information common to all GFS2 uevents (uevent environment variables) +---------------------------------------------------------------------- + +1. LOCKTABLE= + +The LOCKTABLE is a string, as supplied on the mount command +line (locktable=) or via fstab. It is used as a filesystem label +as well as providing the information for a lock_dlm mount to be +able to join the cluster. + +2. LOCKPROTO= + +The LOCKPROTO is a string, and its value depends on what is set +on the mount command line, or via fstab. It will be either +lock_nolock or lock_dlm. In the future other lock managers +may be supported. + +3. JOURNALID= + +If a journal is in use by the filesystem (journals are not +assigned for spectator mounts) then this will give the +numeric journal id in all GFS2 uevents. + +4. UUID= + +With recent versions of gfs2-utils, mkfs.gfs2 writes a UUID +into the filesystem superblock. If it exists, this will +be included in every uevent relating to the filesystem. + + + diff --git a/Documentation/filesystems/gfs2.txt b/Documentation/filesystems/gfs2.txt new file mode 100644 index 00000000..4cda9266 --- /dev/null +++ b/Documentation/filesystems/gfs2.txt @@ -0,0 +1,46 @@ +Global File System +------------------ + +http://sources.redhat.com/cluster/wiki/ + +GFS is a cluster file system. It allows a cluster of computers to +simultaneously use a block device that is shared between them (with FC, +iSCSI, NBD, etc). GFS reads and writes to the block device like a local +file system, but also uses a lock module to allow the computers coordinate +their I/O so file system consistency is maintained. One of the nifty +features of GFS is perfect consistency -- changes made to the file system +on one machine show up immediately on all other machines in the cluster. + +GFS uses interchangeable inter-node locking mechanisms, the currently +supported mechanisms are: + + lock_nolock -- allows gfs to be used as a local file system + + lock_dlm -- uses a distributed lock manager (dlm) for inter-node locking + The dlm is found at linux/fs/dlm/ + +Lock_dlm depends on user space cluster management systems found +at the URL above. + +To use gfs as a local file system, no external clustering systems are +needed, simply: + + $ mkfs -t gfs2 -p lock_nolock -j 1 /dev/block_device + $ mount -t gfs2 /dev/block_device /dir + +If you are using Fedora, you need to install the gfs2-utils package +and, for lock_dlm, you will also need to install the cman package +and write a cluster.conf as per the documentation. + +GFS2 is not on-disk compatible with previous versions of GFS, but it +is pretty close. + +The following man pages can be found at the URL above: + fsck.gfs2 to repair a filesystem + gfs2_grow to expand a filesystem online + gfs2_jadd to add journals to a filesystem online + gfs2_tool to manipulate, examine and tune a filesystem + gfs2_quota to examine and change quota values in a filesystem + gfs2_convert to convert a gfs filesystem to gfs2 in-place + mount.gfs2 to help mount(8) mount a filesystem + mkfs.gfs2 to make a filesystem diff --git a/Documentation/filesystems/hfs.txt b/Documentation/filesystems/hfs.txt new file mode 100644 index 00000000..bd0fa770 --- /dev/null +++ b/Documentation/filesystems/hfs.txt @@ -0,0 +1,83 @@ + +Macintosh HFS Filesystem for Linux +================================== + +HFS stands for ``Hierarchical File System'' and is the filesystem used +by the Mac Plus and all later Macintosh models. Earlier Macintosh +models used MFS (``Macintosh File System''), which is not supported, +MacOS 8.1 and newer support a filesystem called HFS+ that's similar to +HFS but is extended in various areas. Use the hfsplus filesystem driver +to access such filesystems from Linux. + + +Mount options +============= + +When mounting an HFS filesystem, the following options are accepted: + + creator=cccc, type=cccc + Specifies the creator/type values as shown by the MacOS finder + used for creating new files. Default values: '????'. + + uid=n, gid=n + Specifies the user/group that owns all files on the filesystems. + Default: user/group id of the mounting process. + + dir_umask=n, file_umask=n, umask=n + Specifies the umask used for all files , all directories or all + files and directories. Defaults to the umask of the mounting process. + + session=n + Select the CDROM session to mount as HFS filesystem. Defaults to + leaving that decision to the CDROM driver. This option will fail + with anything but a CDROM as underlying devices. + + part=n + Select partition number n from the devices. Does only makes + sense for CDROMS because they can't be partitioned under Linux. + For disk devices the generic partition parsing code does this + for us. Defaults to not parsing the partition table at all. + + quiet + Ignore invalid mount options instead of complaining. + + +Writing to HFS Filesystems +========================== + +HFS is not a UNIX filesystem, thus it does not have the usual features you'd +expect: + + o You can't modify the set-uid, set-gid, sticky or executable bits or the uid + and gid of files. + o You can't create hard- or symlinks, device files, sockets or FIFOs. + +HFS does on the other have the concepts of multiple forks per file. These +non-standard forks are represented as hidden additional files in the normal +filesystems namespace which is kind of a cludge and makes the semantics for +the a little strange: + + o You can't create, delete or rename resource forks of files or the + Finder's metadata. + o They are however created (with default values), deleted and renamed + along with the corresponding data fork or directory. + o Copying files to a different filesystem will loose those attributes + that are essential for MacOS to work. + + +Creating HFS filesystems +=================================== + +The hfsutils package from Robert Leslie contains a program called +hformat that can be used to create HFS filesystem. See +<http://www.mars.org/home/rob/proj/hfs/> for details. + + +Credits +======= + +The HFS drivers was written by Paul H. Hargrovea (hargrove@sccm.Stanford.EDU) +and is now maintained by Roman Zippel (roman@ardistech.com) at Ardis +Technologies. +Roman rewrote large parts of the code and brought in btree routines derived +from Brad Boyer's hfsplus driver (also maintained by Roman now). diff --git a/Documentation/filesystems/hfsplus.txt b/Documentation/filesystems/hfsplus.txt new file mode 100644 index 00000000..af1628a1 --- /dev/null +++ b/Documentation/filesystems/hfsplus.txt @@ -0,0 +1,59 @@ + +Macintosh HFSPlus Filesystem for Linux +====================================== + +HFSPlus is a filesystem first introduced in MacOS 8.1. +HFSPlus has several extensions to HFS, including 32-bit allocation +blocks, 255-character unicode filenames, and file sizes of 2^63 bytes. + + +Mount options +============= + +When mounting an HFSPlus filesystem, the following options are accepted: + + creator=cccc, type=cccc + Specifies the creator/type values as shown by the MacOS finder + used for creating new files. Default values: '????'. + + uid=n, gid=n + Specifies the user/group that owns all files on the filesystem + that have uninitialized permissions structures. + Default: user/group id of the mounting process. + + umask=n + Specifies the umask (in octal) used for files and directories + that have uninitialized permissions structures. + Default: umask of the mounting process. + + session=n + Select the CDROM session to mount as HFSPlus filesystem. Defaults to + leaving that decision to the CDROM driver. This option will fail + with anything but a CDROM as underlying devices. + + part=n + Select partition number n from the devices. This option only makes + sense for CDROMs because they can't be partitioned under Linux. + For disk devices the generic partition parsing code does this + for us. Defaults to not parsing the partition table at all. + + decompose + Decompose file name characters. + + nodecompose + Do not decompose file name characters. + + force + Used to force write access to volumes that are marked as journalled + or locked. Use at your own risk. + + nls=cccc + Encoding to use when presenting file names. + + +References +========== + +kernel source: <file:fs/hfsplus> + +Apple Technote 1150 http://developer.apple.com/technotes/tn/tn1150.html diff --git a/Documentation/filesystems/hpfs.txt b/Documentation/filesystems/hpfs.txt new file mode 100644 index 00000000..74630bd5 --- /dev/null +++ b/Documentation/filesystems/hpfs.txt @@ -0,0 +1,296 @@ +Read/Write HPFS 2.09 +1998-2004, Mikulas Patocka + +email: mikulas@artax.karlin.mff.cuni.cz +homepage: http://artax.karlin.mff.cuni.cz/~mikulas/vyplody/hpfs/index-e.cgi + +CREDITS: +Chris Smith, 1993, original read-only HPFS, some code and hpfs structures file + is taken from it +Jacques Gelinas, MSDos mmap, Inspired by fs/nfs/mmap.c (Jon Tombs 15 Aug 1993) +Werner Almesberger, 1992, 1993, MSDos option parser & CR/LF conversion + +Mount options + +uid=xxx,gid=xxx,umask=xxx (default uid=gid=0 umask=default_system_umask) + Set owner/group/mode for files that do not have it specified in extended + attributes. Mode is inverted umask - for example umask 027 gives owner + all permission, group read permission and anybody else no access. Note + that for files mode is anded with 0666. If you want files to have 'x' + rights, you must use extended attributes. +case=lower,asis (default asis) + File name lowercasing in readdir. +conv=binary,text,auto (default binary) + CR/LF -> LF conversion, if auto, decision is made according to extension + - there is a list of text extensions (I thing it's better to not convert + text file than to damage binary file). If you want to change that list, + change it in the source. Original readonly HPFS contained some strange + heuristic algorithm that I removed. I thing it's danger to let the + computer decide whether file is text or binary. For example, DJGPP + binaries contain small text message at the beginning and they could be + misidentified and damaged under some circumstances. +check=none,normal,strict (default normal) + Check level. Selecting none will cause only little speedup and big + danger. I tried to write it so that it won't crash if check=normal on + corrupted filesystems. check=strict means many superfluous checks - + used for debugging (for example it checks if file is allocated in + bitmaps when accessing it). +errors=continue,remount-ro,panic (default remount-ro) + Behaviour when filesystem errors found. +chkdsk=no,errors,always (default errors) + When to mark filesystem dirty so that OS/2 checks it. +eas=no,ro,rw (default rw) + What to do with extended attributes. 'no' - ignore them and use always + values specified in uid/gid/mode options. 'ro' - read extended + attributes but do not create them. 'rw' - create extended attributes + when you use chmod/chown/chgrp/mknod/ln -s on the filesystem. +timeshift=(-)nnn (default 0) + Shifts the time by nnn seconds. For example, if you see under linux + one hour more, than under os/2, use timeshift=-3600. + + +File names + +As in OS/2, filenames are case insensitive. However, shell thinks that names +are case sensitive, so for example when you create a file FOO, you can use +'cat FOO', 'cat Foo', 'cat foo' or 'cat F*' but not 'cat f*'. Note, that you +also won't be able to compile linux kernel (and maybe other things) on HPFS +because kernel creates different files with names like bootsect.S and +bootsect.s. When searching for file thats name has characters >= 128, codepages +are used - see below. +OS/2 ignores dots and spaces at the end of file name, so this driver does as +well. If you create 'a. ...', the file 'a' will be created, but you can still +access it under names 'a.', 'a..', 'a . . . ' etc. + + +Extended attributes + +On HPFS partitions, OS/2 can associate to each file a special information called +extended attributes. Extended attributes are pairs of (key,value) where key is +an ascii string identifying that attribute and value is any string of bytes of +variable length. OS/2 stores window and icon positions and file types there. So +why not use it for unix-specific info like file owner or access rights? This +driver can do it. If you chown/chgrp/chmod on a hpfs partition, extended +attributes with keys "UID", "GID" or "MODE" and 2-byte values are created. Only +that extended attributes those value differs from defaults specified in mount +options are created. Once created, the extended attributes are never deleted, +they're just changed. It means that when your default uid=0 and you type +something like 'chown luser file; chown root file' the file will contain +extended attribute UID=0. And when you umount the fs and mount it again with +uid=luser_uid, the file will be still owned by root! If you chmod file to 444, +extended attribute "MODE" will not be set, this special case is done by setting +read-only flag. When you mknod a block or char device, besides "MODE", the +special 4-byte extended attribute "DEV" will be created containing the device +number. Currently this driver cannot resize extended attributes - it means +that if somebody (I don't know who?) has set "UID", "GID", "MODE" or "DEV" +attributes with different sizes, they won't be rewritten and changing these +values doesn't work. + + +Symlinks + +You can do symlinks on HPFS partition, symlinks are achieved by setting extended +attribute named "SYMLINK" with symlink value. Like on ext2, you can chown and +chgrp symlinks but I don't know what is it good for. chmoding symlink results +in chmoding file where symlink points. These symlinks are just for Linux use and +incompatible with OS/2. OS/2 PmShell symlinks are not supported because they are +stored in very crazy way. They tried to do it so that link changes when file is +moved ... sometimes it works. But the link is partly stored in directory +extended attributes and partly in OS2SYS.INI. I don't want (and don't know how) +to analyze or change OS2SYS.INI. + + +Codepages + +HPFS can contain several uppercasing tables for several codepages and each +file has a pointer to codepage its name is in. However OS/2 was created in +America where people don't care much about codepages and so multiple codepages +support is quite buggy. I have Czech OS/2 working in codepage 852 on my disk. +Once I booted English OS/2 working in cp 850 and I created a file on my 852 +partition. It marked file name codepage as 850 - good. But when I again booted +Czech OS/2, the file was completely inaccessible under any name. It seems that +OS/2 uppercases the search pattern with its system code page (852) and file +name it's comparing to with its code page (850). These could never match. Is it +really what IBM developers wanted? But problems continued. When I created in +Czech OS/2 another file in that directory, that file was inaccessible too. OS/2 +probably uses different uppercasing method when searching where to place a file +(note, that files in HPFS directory must be sorted) and when searching for +a file. Finally when I opened this directory in PmShell, PmShell crashed (the +funny thing was that, when rebooted, PmShell tried to reopen this directory +again :-). chkdsk happily ignores these errors and only low-level disk +modification saved me. Never mix different language versions of OS/2 on one +system although HPFS was designed to allow that. +OK, I could implement complex codepage support to this driver but I think it +would cause more problems than benefit with such buggy implementation in OS/2. +So this driver simply uses first codepage it finds for uppercasing and +lowercasing no matter what's file codepage index. Usually all file names are in +this codepage - if you don't try to do what I described above :-) + + +Known bugs + +HPFS386 on OS/2 server is not supported. HPFS386 installed on normal OS/2 client +should work. If you have OS/2 server, use only read-only mode. I don't know how +to handle some HPFS386 structures like access control list or extended perm +list, I don't know how to delete them when file is deleted and how to not +overwrite them with extended attributes. Send me some info on these structures +and I'll make it. However, this driver should detect presence of HPFS386 +structures, remount read-only and not destroy them (I hope). + +When there's not enough space for extended attributes, they will be truncated +and no error is returned. + +OS/2 can't access files if the path is longer than about 256 chars but this +driver allows you to do it. chkdsk ignores such errors. + +Sometimes you won't be able to delete some files on a very full filesystem +(returning error ENOSPC). That's because file in non-leaf node in directory tree +(one directory, if it's large, has dirents in tree on HPFS) must be replaced +with another node when deleted. And that new file might have larger name than +the old one so the new name doesn't fit in directory node (dnode). And that +would result in directory tree splitting, that takes disk space. Workaround is +to delete other files that are leaf (probability that the file is non-leaf is +about 1/50) or to truncate file first to make some space. +You encounter this problem only if you have many directories so that +preallocated directory band is full i.e. + number_of_directories / size_of_filesystem_in_mb > 4. + +You can't delete open directories. + +You can't rename over directories (what is it good for?). + +Renaming files so that only case changes doesn't work. This driver supports it +but vfs doesn't. Something like 'mv file FILE' won't work. + +All atimes and directory mtimes are not updated. That's because of performance +reasons. If you extremely wish to update them, let me know, I'll write it (but +it will be slow). + +When the system is out of memory and swap, it may slightly corrupt filesystem +(lost files, unbalanced directories). (I guess all filesystem may do it). + +When compiled, you get warning: function declaration isn't a prototype. Does +anybody know what does it mean? + + +What does "unbalanced tree" message mean? + +Old versions of this driver created sometimes unbalanced dnode trees. OS/2 +chkdsk doesn't scream if the tree is unbalanced (and sometimes creates +unbalanced trees too :-) but both HPFS and HPFS386 contain bug that it rarely +crashes when the tree is not balanced. This driver handles unbalanced trees +correctly and writes warning if it finds them. If you see this message, this is +probably because of directories created with old version of this driver. +Workaround is to move all files from that directory to another and then back +again. Do it in Linux, not OS/2! If you see this message in directory that is +whole created by this driver, it is BUG - let me know about it. + + +Bugs in OS/2 + +When you have two (or more) lost directories pointing each to other, chkdsk +locks up when repairing filesystem. + +Sometimes (I think it's random) when you create a file with one-char name under +OS/2, OS/2 marks it as 'long'. chkdsk then removes this flag saying "Minor fs +error corrected". + +File names like "a .b" are marked as 'long' by OS/2 but chkdsk "corrects" it and +marks them as short (and writes "minor fs error corrected"). This bug is not in +HPFS386. + +Codepage bugs described above. + +If you don't install fixpacks, there are many, many more... + + +History + +0.90 First public release +0.91 Fixed bug that caused shooting to memory when write_inode was called on + open inode (rarely happened) +0.92 Fixed a little memory leak in freeing directory inodes +0.93 Fixed bug that locked up the machine when there were too many filenames + with first 15 characters same + Fixed write_file to zero file when writing behind file end +0.94 Fixed a little memory leak when trying to delete busy file or directory +0.95 Fixed a bug that i_hpfs_parent_dir was not updated when moving files +1.90 First version for 2.1.1xx kernels +1.91 Fixed a bug that chk_sectors failed when sectors were at the end of disk + Fixed a race-condition when write_inode is called while deleting file + Fixed a bug that could possibly happen (with very low probability) when + using 0xff in filenames + Rewritten locking to avoid race-conditions + Mount option 'eas' now works + Fsync no longer returns error + Files beginning with '.' are marked hidden + Remount support added + Alloc is not so slow when filesystem becomes full + Atimes are no more updated because it slows down operation + Code cleanup (removed all commented debug prints) +1.92 Corrected a bug when sync was called just before closing file +1.93 Modified, so that it works with kernels >= 2.1.131, I don't know if it + works with previous versions + Fixed a possible problem with disks > 64G (but I don't have one, so I can't + test it) + Fixed a file overflow at 2G + Added new option 'timeshift' + Changed behaviour on HPFS386: It is now possible to operate on HPFS386 in + read-only mode + Fixed a bug that slowed down alloc and prevented allocating 100% space + (this bug was not destructive) +1.94 Added workaround for one bug in Linux + Fixed one buffer leak + Fixed some incompatibilities with large extended attributes (but it's still + not 100% ok, I have no info on it and OS/2 doesn't want to create them) + Rewritten allocation + Fixed a bug with i_blocks (du sometimes didn't display correct values) + Directories have no longer archive attribute set (some programs don't like + it) + Fixed a bug that it set badly one flag in large anode tree (it was not + destructive) +1.95 Fixed one buffer leak, that could happen on corrupted filesystem + Fixed one bug in allocation in 1.94 +1.96 Added workaround for one bug in OS/2 (HPFS locked up, HPFS386 reported + error sometimes when opening directories in PMSHELL) + Fixed a possible bitmap race + Fixed possible problem on large disks + You can now delete open files + Fixed a nondestructive race in rename +1.97 Support for HPFS v3 (on large partitions) + Fixed a bug that it didn't allow creation of files > 128M (it should be 2G) +1.97.1 Changed names of global symbols + Fixed a bug when chmoding or chowning root directory +1.98 Fixed a deadlock when using old_readdir + Better directory handling; workaround for "unbalanced tree" bug in OS/2 +1.99 Corrected a possible problem when there's not enough space while deleting + file + Now it tries to truncate the file if there's not enough space when deleting + Removed a lot of redundant code +2.00 Fixed a bug in rename (it was there since 1.96) + Better anti-fragmentation strategy +2.01 Fixed problem with directory listing over NFS + Directory lseek now checks for proper parameters + Fixed race-condition in buffer code - it is in all filesystems in Linux; + when reading device (cat /dev/hda) while creating files on it, files + could be damaged +2.02 Workaround for bug in breada in Linux. breada could cause accesses beyond + end of partition +2.03 Char, block devices and pipes are correctly created + Fixed non-crashing race in unlink (Alexander Viro) + Now it works with Japanese version of OS/2 +2.04 Fixed error when ftruncate used to extend file +2.05 Fixed crash when got mount parameters without = + Fixed crash when allocation of anode failed due to full disk + Fixed some crashes when block io or inode allocation failed +2.06 Fixed some crash on corrupted disk structures + Better allocation strategy + Reschedule points added so that it doesn't lock CPU long time + It should work in read-only mode on Warp Server +2.07 More fixes for Warp Server. Now it really works +2.08 Creating new files is not so slow on large disks + An attempt to sync deleted file does not generate filesystem error +2.09 Fixed error on extremely fragmented files + + + vim: set textwidth=80: diff --git a/Documentation/filesystems/inotify.txt b/Documentation/filesystems/inotify.txt new file mode 100644 index 00000000..59a919f1 --- /dev/null +++ b/Documentation/filesystems/inotify.txt @@ -0,0 +1,269 @@ + inotify + a powerful yet simple file change notification system + + + +Document started 15 Mar 2005 by Robert Love <rml@novell.com> + + +(i) User Interface + +Inotify is controlled by a set of three system calls and normal file I/O on a +returned file descriptor. + +First step in using inotify is to initialise an inotify instance: + + int fd = inotify_init (); + +Each instance is associated with a unique, ordered queue. + +Change events are managed by "watches". A watch is an (object,mask) pair where +the object is a file or directory and the mask is a bit mask of one or more +inotify events that the application wishes to receive. See <linux/inotify.h> +for valid events. A watch is referenced by a watch descriptor, or wd. + +Watches are added via a path to the file. + +Watches on a directory will return events on any files inside of the directory. + +Adding a watch is simple: + + int wd = inotify_add_watch (fd, path, mask); + +Where "fd" is the return value from inotify_init(), path is the path to the +object to watch, and mask is the watch mask (see <linux/inotify.h>). + +You can update an existing watch in the same manner, by passing in a new mask. + +An existing watch is removed via + + int ret = inotify_rm_watch (fd, wd); + +Events are provided in the form of an inotify_event structure that is read(2) +from a given inotify instance. The filename is of dynamic length and follows +the struct. It is of size len. The filename is padded with null bytes to +ensure proper alignment. This padding is reflected in len. + +You can slurp multiple events by passing a large buffer, for example + + size_t len = read (fd, buf, BUF_LEN); + +Where "buf" is a pointer to an array of "inotify_event" structures at least +BUF_LEN bytes in size. The above example will return as many events as are +available and fit in BUF_LEN. + +Each inotify instance fd is also select()- and poll()-able. + +You can find the size of the current event queue via the standard FIONREAD +ioctl on the fd returned by inotify_init(). + +All watches are destroyed and cleaned up on close. + + +(ii) + +Prototypes: + + int inotify_init (void); + int inotify_add_watch (int fd, const char *path, __u32 mask); + int inotify_rm_watch (int fd, __u32 mask); + + +(iii) Kernel Interface + +Inotify's kernel API consists a set of functions for managing watches and an +event callback. + +To use the kernel API, you must first initialize an inotify instance with a set +of inotify_operations. You are given an opaque inotify_handle, which you use +for any further calls to inotify. + + struct inotify_handle *ih = inotify_init(my_event_handler); + +You must provide a function for processing events and a function for destroying +the inotify watch. + + void handle_event(struct inotify_watch *watch, u32 wd, u32 mask, + u32 cookie, const char *name, struct inode *inode) + + watch - the pointer to the inotify_watch that triggered this call + wd - the watch descriptor + mask - describes the event that occurred + cookie - an identifier for synchronizing events + name - the dentry name for affected files in a directory-based event + inode - the affected inode in a directory-based event + + void destroy_watch(struct inotify_watch *watch) + +You may add watches by providing a pre-allocated and initialized inotify_watch +structure and specifying the inode to watch along with an inotify event mask. +You must pin the inode during the call. You will likely wish to embed the +inotify_watch structure in a structure of your own which contains other +information about the watch. Once you add an inotify watch, it is immediately +subject to removal depending on filesystem events. You must grab a reference if +you depend on the watch hanging around after the call. + + inotify_init_watch(&my_watch->iwatch); + inotify_get_watch(&my_watch->iwatch); // optional + s32 wd = inotify_add_watch(ih, &my_watch->iwatch, inode, mask); + inotify_put_watch(&my_watch->iwatch); // optional + +You may use the watch descriptor (wd) or the address of the inotify_watch for +other inotify operations. You must not directly read or manipulate data in the +inotify_watch. Additionally, you must not call inotify_add_watch() more than +once for a given inotify_watch structure, unless you have first called either +inotify_rm_watch() or inotify_rm_wd(). + +To determine if you have already registered a watch for a given inode, you may +call inotify_find_watch(), which gives you both the wd and the watch pointer for +the inotify_watch, or an error if the watch does not exist. + + wd = inotify_find_watch(ih, inode, &watchp); + +You may use container_of() on the watch pointer to access your own data +associated with a given watch. When an existing watch is found, +inotify_find_watch() bumps the refcount before releasing its locks. You must +put that reference with: + + put_inotify_watch(watchp); + +Call inotify_find_update_watch() to update the event mask for an existing watch. +inotify_find_update_watch() returns the wd of the updated watch, or an error if +the watch does not exist. + + wd = inotify_find_update_watch(ih, inode, mask); + +An existing watch may be removed by calling either inotify_rm_watch() or +inotify_rm_wd(). + + int ret = inotify_rm_watch(ih, &my_watch->iwatch); + int ret = inotify_rm_wd(ih, wd); + +A watch may be removed while executing your event handler with the following: + + inotify_remove_watch_locked(ih, iwatch); + +Call inotify_destroy() to remove all watches from your inotify instance and +release it. If there are no outstanding references, inotify_destroy() will call +your destroy_watch op for each watch. + + inotify_destroy(ih); + +When inotify removes a watch, it sends an IN_IGNORED event to your callback. +You may use this event as an indication to free the watch memory. Note that +inotify may remove a watch due to filesystem events, as well as by your request. +If you use IN_ONESHOT, inotify will remove the watch after the first event, at +which point you may call the final inotify_put_watch. + +(iv) Kernel Interface Prototypes + + struct inotify_handle *inotify_init(struct inotify_operations *ops); + + inotify_init_watch(struct inotify_watch *watch); + + s32 inotify_add_watch(struct inotify_handle *ih, + struct inotify_watch *watch, + struct inode *inode, u32 mask); + + s32 inotify_find_watch(struct inotify_handle *ih, struct inode *inode, + struct inotify_watch **watchp); + + s32 inotify_find_update_watch(struct inotify_handle *ih, + struct inode *inode, u32 mask); + + int inotify_rm_wd(struct inotify_handle *ih, u32 wd); + + int inotify_rm_watch(struct inotify_handle *ih, + struct inotify_watch *watch); + + void inotify_remove_watch_locked(struct inotify_handle *ih, + struct inotify_watch *watch); + + void inotify_destroy(struct inotify_handle *ih); + + void get_inotify_watch(struct inotify_watch *watch); + void put_inotify_watch(struct inotify_watch *watch); + + +(v) Internal Kernel Implementation + +Each inotify instance is represented by an inotify_handle structure. +Inotify's userspace consumers also have an inotify_device which is +associated with the inotify_handle, and on which events are queued. + +Each watch is associated with an inotify_watch structure. Watches are chained +off of each associated inotify_handle and each associated inode. + +See fs/inotify.c and fs/inotify_user.c for the locking and lifetime rules. + + +(vi) Rationale + +Q: What is the design decision behind not tying the watch to the open fd of + the watched object? + +A: Watches are associated with an open inotify device, not an open file. + This solves the primary problem with dnotify: keeping the file open pins + the file and thus, worse, pins the mount. Dnotify is therefore infeasible + for use on a desktop system with removable media as the media cannot be + unmounted. Watching a file should not require that it be open. + +Q: What is the design decision behind using an-fd-per-instance as opposed to + an fd-per-watch? + +A: An fd-per-watch quickly consumes more file descriptors than are allowed, + more fd's than are feasible to manage, and more fd's than are optimally + select()-able. Yes, root can bump the per-process fd limit and yes, users + can use epoll, but requiring both is a silly and extraneous requirement. + A watch consumes less memory than an open file, separating the number + spaces is thus sensible. The current design is what user-space developers + want: Users initialize inotify, once, and add n watches, requiring but one + fd and no twiddling with fd limits. Initializing an inotify instance two + thousand times is silly. If we can implement user-space's preferences + cleanly--and we can, the idr layer makes stuff like this trivial--then we + should. + + There are other good arguments. With a single fd, there is a single + item to block on, which is mapped to a single queue of events. The single + fd returns all watch events and also any potential out-of-band data. If + every fd was a separate watch, + + - There would be no way to get event ordering. Events on file foo and + file bar would pop poll() on both fd's, but there would be no way to tell + which happened first. A single queue trivially gives you ordering. Such + ordering is crucial to existing applications such as Beagle. Imagine + "mv a b ; mv b a" events without ordering. + + - We'd have to maintain n fd's and n internal queues with state, + versus just one. It is a lot messier in the kernel. A single, linear + queue is the data structure that makes sense. + + - User-space developers prefer the current API. The Beagle guys, for + example, love it. Trust me, I asked. It is not a surprise: Who'd want + to manage and block on 1000 fd's via select? + + - No way to get out of band data. + + - 1024 is still too low. ;-) + + When you talk about designing a file change notification system that + scales to 1000s of directories, juggling 1000s of fd's just does not seem + the right interface. It is too heavy. + + Additionally, it _is_ possible to more than one instance and + juggle more than one queue and thus more than one associated fd. There + need not be a one-fd-per-process mapping; it is one-fd-per-queue and a + process can easily want more than one queue. + +Q: Why the system call approach? + +A: The poor user-space interface is the second biggest problem with dnotify. + Signals are a terrible, terrible interface for file notification. Or for + anything, for that matter. The ideal solution, from all perspectives, is a + file descriptor-based one that allows basic file I/O and poll/select. + Obtaining the fd and managing the watches could have been done either via a + device file or a family of new system calls. We decided to implement a + family of system calls because that is the preferred approach for new kernel + interfaces. The only real difference was whether we wanted to use open(2) + and ioctl(2) or a couple of new system calls. System calls beat ioctls. + diff --git a/Documentation/filesystems/isofs.txt b/Documentation/filesystems/isofs.txt new file mode 100644 index 00000000..ba0a9338 --- /dev/null +++ b/Documentation/filesystems/isofs.txt @@ -0,0 +1,48 @@ +Mount options that are the same as for msdos and vfat partitions. + + gid=nnn All files in the partition will be in group nnn. + uid=nnn All files in the partition will be owned by user id nnn. + umask=nnn The permission mask (see umask(1)) for the partition. + +Mount options that are the same as vfat partitions. These are only useful +when using discs encoded using Microsoft's Joliet extensions. + iocharset=name Character set to use for converting from Unicode to + ASCII. Joliet filenames are stored in Unicode format, but + Unix for the most part doesn't know how to deal with Unicode. + There is also an option of doing UTF-8 translations with the + utf8 option. + utf8 Encode Unicode names in UTF-8 format. Default is no. + +Mount options unique to the isofs filesystem. + block=512 Set the block size for the disk to 512 bytes + block=1024 Set the block size for the disk to 1024 bytes + block=2048 Set the block size for the disk to 2048 bytes + check=relaxed Matches filenames with different cases + check=strict Matches only filenames with the exact same case + cruft Try to handle badly formatted CDs. + map=off Do not map non-Rock Ridge filenames to lower case + map=normal Map non-Rock Ridge filenames to lower case + map=acorn As map=normal but also apply Acorn extensions if present + mode=xxx Sets the permissions on files to xxx unless Rock Ridge + extensions set the permissions otherwise + dmode=xxx Sets the permissions on directories to xxx unless Rock Ridge + extensions set the permissions otherwise + overriderockperm Set permissions on files and directories according to + 'mode' and 'dmode' even though Rock Ridge extensions are + present. + nojoliet Ignore Joliet extensions if they are present. + norock Ignore Rock Ridge extensions if they are present. + hide Completely strip hidden files from the file system. + showassoc Show files marked with the 'associated' bit + unhide Deprecated; showing hidden files is now default; + If given, it is a synonym for 'showassoc' which will + recreate previous unhide behavior + session=x Select number of session on multisession CD + sbsector=xxx Session begins from sector xxx + +Recommended documents about ISO 9660 standard are located at: +http://www.y-adagio.com/ +ftp://ftp.ecma.ch/ecma-st/Ecma-119.pdf +Quoting from the PDF "This 2nd Edition of Standard ECMA-119 is technically +identical with ISO 9660.", so it is a valid and gratis substitute of the +official ISO specification. diff --git a/Documentation/filesystems/jfs.txt b/Documentation/filesystems/jfs.txt new file mode 100644 index 00000000..26ebde77 --- /dev/null +++ b/Documentation/filesystems/jfs.txt @@ -0,0 +1,41 @@ +IBM's Journaled File System (JFS) for Linux + +JFS Homepage: http://jfs.sourceforge.net/ + +The following mount options are supported: + +iocharset=name Character set to use for converting from Unicode to + ASCII. The default is to do no conversion. Use + iocharset=utf8 for UTF-8 translations. This requires + CONFIG_NLS_UTF8 to be set in the kernel .config file. + iocharset=none specifies the default behavior explicitly. + +resize=value Resize the volume to <value> blocks. JFS only supports + growing a volume, not shrinking it. This option is only + valid during a remount, when the volume is mounted + read-write. The resize keyword with no value will grow + the volume to the full size of the partition. + +nointegrity Do not write to the journal. The primary use of this option + is to allow for higher performance when restoring a volume + from backup media. The integrity of the volume is not + guaranteed if the system abnormally abends. + +integrity Default. Commit metadata changes to the journal. Use this + option to remount a volume where the nointegrity option was + previously specified in order to restore normal behavior. + +errors=continue Keep going on a filesystem error. +errors=remount-ro Default. Remount the filesystem read-only on an error. +errors=panic Panic and halt the machine if an error occurs. + +uid=value Override on-disk uid with specified value +gid=value Override on-disk gid with specified value +umask=value Override on-disk umask with specified octal value. For + directories, the execute bit will be set if the corresponding + read bit is set. + +Please send bugs, comments, cards and letters to shaggy@linux.vnet.ibm.com. + +The JFS mailing list can be subscribed to by using the link labeled +"Mail list Subscribe" at our web page http://jfs.sourceforge.net/ diff --git a/Documentation/filesystems/locks.txt b/Documentation/filesystems/locks.txt new file mode 100644 index 00000000..fab857ac --- /dev/null +++ b/Documentation/filesystems/locks.txt @@ -0,0 +1,67 @@ + File Locking Release Notes + + Andy Walker <andy@lysaker.kvaerner.no> + + 12 May 1997 + + +1. What's New? +-------------- + +1.1 Broken Flock Emulation +-------------------------- + +The old flock(2) emulation in the kernel was swapped for proper BSD +compatible flock(2) support in the 1.3.x series of kernels. With the +release of the 2.1.x kernel series, support for the old emulation has +been totally removed, so that we don't need to carry this baggage +forever. + +This should not cause problems for anybody, since everybody using a +2.1.x kernel should have updated their C library to a suitable version +anyway (see the file "Documentation/Changes".) + +1.2 Allow Mixed Locks Again +--------------------------- + +1.2.1 Typical Problems - Sendmail +--------------------------------- +Because sendmail was unable to use the old flock() emulation, many sendmail +installations use fcntl() instead of flock(). This is true of Slackware 3.0 +for example. This gave rise to some other subtle problems if sendmail was +configured to rebuild the alias file. Sendmail tried to lock the aliases.dir +file with fcntl() at the same time as the GDBM routines tried to lock this +file with flock(). With pre 1.3.96 kernels this could result in deadlocks that, +over time, or under a very heavy mail load, would eventually cause the kernel +to lock solid with deadlocked processes. + + +1.2.2 The Solution +------------------ +The solution I have chosen, after much experimentation and discussion, +is to make flock() and fcntl() locks oblivious to each other. Both can +exists, and neither will have any effect on the other. + +I wanted the two lock styles to be cooperative, but there were so many +race and deadlock conditions that the current solution was the only +practical one. It puts us in the same position as, for example, SunOS +4.1.x and several other commercial Unices. The only OS's that support +cooperative flock()/fcntl() are those that emulate flock() using +fcntl(), with all the problems that implies. + + +1.3 Mandatory Locking As A Mount Option +--------------------------------------- + +Mandatory locking, as described in 'Documentation/filesystems/mandatory.txt' +was prior to this release a general configuration option that was valid for +all mounted filesystems. This had a number of inherent dangers, not the +least of which was the ability to freeze an NFS server by asking it to read +a file for which a mandatory lock existed. + +From this release of the kernel, mandatory locking can be turned on and off +on a per-filesystem basis, using the mount options 'mand' and 'nomand'. +The default is to disallow mandatory locking. The intention is that +mandatory locking only be enabled on a local filesystem as the specific need +arises. + diff --git a/Documentation/filesystems/logfs.txt b/Documentation/filesystems/logfs.txt new file mode 100644 index 00000000..bca42c22 --- /dev/null +++ b/Documentation/filesystems/logfs.txt @@ -0,0 +1,241 @@ + +The LogFS Flash Filesystem +========================== + +Specification +============= + +Superblocks +----------- + +Two superblocks exist at the beginning and end of the filesystem. +Each superblock is 256 Bytes large, with another 3840 Bytes reserved +for future purposes, making a total of 4096 Bytes. + +Superblock locations may differ for MTD and block devices. On MTD the +first non-bad block contains a superblock in the first 4096 Bytes and +the last non-bad block contains a superblock in the last 4096 Bytes. +On block devices, the first 4096 Bytes of the device contain the first +superblock and the last aligned 4096 Byte-block contains the second +superblock. + +For the most part, the superblocks can be considered read-only. They +are written only to correct errors detected within the superblocks, +move the journal and change the filesystem parameters through tunefs. +As a result, the superblock does not contain any fields that require +constant updates, like the amount of free space, etc. + +Segments +-------- + +The space in the device is split up into equal-sized segments. +Segments are the primary write unit of LogFS. Within each segments, +writes happen from front (low addresses) to back (high addresses. If +only a partial segment has been written, the segment number, the +current position within and optionally a write buffer are stored in +the journal. + +Segments are erased as a whole. Therefore Garbage Collection may be +required to completely free a segment before doing so. + +Journal +-------- + +The journal contains all global information about the filesystem that +is subject to frequent change. At mount time, it has to be scanned +for the most recent commit entry, which contains a list of pointers to +all currently valid entries. + +Object Store +------------ + +All space except for the superblocks and journal is part of the object +store. Each segment contains a segment header and a number of +objects, each consisting of the object header and the payload. +Objects are either inodes, directory entries (dentries), file data +blocks or indirect blocks. + +Levels +------ + +Garbage collection (GC) may fail if all data is written +indiscriminately. One requirement of GC is that data is separated +roughly according to the distance between the tree root and the data. +Effectively that means all file data is on level 0, indirect blocks +are on levels 1, 2, 3 4 or 5 for 1x, 2x, 3x, 4x or 5x indirect blocks, +respectively. Inode file data is on level 6 for the inodes and 7-11 +for indirect blocks. + +Each segment contains objects of a single level only. As a result, +each level requires its own separate segment to be open for writing. + +Inode File +---------- + +All inodes are stored in a special file, the inode file. Single +exception is the inode file's inode (master inode) which for obvious +reasons is stored in the journal instead. Instead of data blocks, the +leaf nodes of the inode files are inodes. + +Aliases +------- + +Writes in LogFS are done by means of a wandering tree. A naïve +implementation would require that for each write or a block, all +parent blocks are written as well, since the block pointers have +changed. Such an implementation would not be very efficient. + +In LogFS, the block pointer changes are cached in the journal by means +of alias entries. Each alias consists of its logical address - inode +number, block index, level and child number (index into block) - and +the changed data. Any 8-byte word can be changes in this manner. + +Currently aliases are used for block pointers, file size, file used +bytes and the height of an inodes indirect tree. + +Segment Aliases +--------------- + +Related to regular aliases, these are used to handle bad blocks. +Initially, bad blocks are handled by moving the affected segment +content to a spare segment and noting this move in the journal with a +segment alias, a simple (to, from) tupel. GC will later empty this +segment and the alias can be removed again. This is used on MTD only. + +Vim +--- + +By cleverly predicting the life time of data, it is possible to +separate long-living data from short-living data and thereby reduce +the GC overhead later. Each type of distinc life expectency (vim) can +have a separate segment open for writing. Each (level, vim) tupel can +be open just once. If an open segment with unknown vim is encountered +at mount time, it is closed and ignored henceforth. + +Indirect Tree +------------- + +Inodes in LogFS are similar to FFS-style filesystems with direct and +indirect block pointers. One difference is that LogFS uses a single +indirect pointer that can be either a 1x, 2x, etc. indirect pointer. +A height field in the inode defines the height of the indirect tree +and thereby the indirection of the pointer. + +Another difference is the addressing of indirect blocks. In LogFS, +the first 16 pointers in the first indirect block are left empty, +corresponding to the 16 direct pointers in the inode. In ext2 (maybe +others as well) the first pointer in the first indirect block +corresponds to logical block 12, skipping the 12 direct pointers. +So where ext2 is using arithmetic to better utilize space, LogFS keeps +arithmetic simple and uses compression to save space. + +Compression +----------- + +Both file data and metadata can be compressed. Compression for file +data can be enabled with chattr +c and disabled with chattr -c. Doing +so has no effect on existing data, but new data will be stored +accordingly. New inodes will inherit the compression flag of the +parent directory. + +Metadata is always compressed. However, the space accounting ignores +this and charges for the uncompressed size. Failing to do so could +result in GC failures when, after moving some data, indirect blocks +compress worse than previously. Even on a 100% full medium, GC may +not consume any extra space, so the compression gains are lost space +to the user. + +However, they are not lost space to the filesystem internals. By +cheating the user for those bytes, the filesystem gained some slack +space and GC will run less often and faster. + +Garbage Collection and Wear Leveling +------------------------------------ + +Garbage collection is invoked whenever the number of free segments +falls below a threshold. The best (known) candidate is picked based +on the least amount of valid data contained in the segment. All +remaining valid data is copied elsewhere, thereby invalidating it. + +The GC code also checks for aliases and writes then back if their +number gets too large. + +Wear leveling is done by occasionally picking a suboptimal segment for +garbage collection. If a stale segments erase count is significantly +lower than the active segments' erase counts, it will be picked. Wear +leveling is rate limited, so it will never monopolize the device for +more than one segment worth at a time. + +Values for "occasionally", "significantly lower" are compile time +constants. + +Hashed directories +------------------ + +To satisfy efficient lookup(), directory entries are hashed and +located based on the hash. In order to both support large directories +and not be overly inefficient for small directories, several hash +tables of increasing size are used. For each table, the hash value +modulo the table size gives the table index. + +Tables sizes are chosen to limit the number of indirect blocks with a +fully populated table to 0, 1, 2 or 3 respectively. So the first +table contains 16 entries, the second 512-16, etc. + +The last table is special in several ways. First its size depends on +the effective 32bit limit on telldir/seekdir cookies. Since logfs +uses the upper half of the address space for indirect blocks, the size +is limited to 2^31. Secondly the table contains hash buckets with 16 +entries each. + +Using single-entry buckets would result in birthday "attacks". At +just 2^16 used entries, hash collisions would be likely (P >= 0.5). +My math skills are insufficient to do the combinatorics for the 17x +collisions necessary to overflow a bucket, but testing showed that in +10,000 runs the lowest directory fill before a bucket overflow was +188,057,130 entries with an average of 315,149,915 entries. So for +directory sizes of up to a million, bucket overflows should be +virtually impossible under normal circumstances. + +With carefully chosen filenames, it is obviously possible to cause an +overflow with just 21 entries (4 higher tables + 16 entries + 1). So +there may be a security concern if a malicious user has write access +to a directory. + +Open For Discussion +=================== + +Device Address Space +-------------------- + +A device address space is used for caching. Both block devices and +MTD provide functions to either read a single page or write a segment. +Partial segments may be written for data integrity, but where possible +complete segments are written for performance on simple block device +flash media. + +Meta Inodes +----------- + +Inodes are stored in the inode file, which is just a regular file for +most purposes. At umount time, however, the inode file needs to +remain open until all dirty inodes are written. So +generic_shutdown_super() may not close this inode, but shouldn't +complain about remaining inodes due to the inode file either. Same +goes for mapping inode of the device address space. + +Currently logfs uses a hack that essentially copies part of fs/inode.c +code over. A general solution would be preferred. + +Indirect block mapping +---------------------- + +With compression, the block device (or mapping inode) cannot be used +to cache indirect blocks. Some other place is required. Currently +logfs uses the top half of each inode's address space. The low 8TB +(on 32bit) are filled with file data, the high 8TB are used for +indirect blocks. + +One problem is that 16TB files created on 64bit systems actually have +data in the top 8TB. But files >16TB would cause problems anyway, so +only the limit has changed. diff --git a/Documentation/filesystems/mandatory-locking.txt b/Documentation/filesystems/mandatory-locking.txt new file mode 100644 index 00000000..0979d1d2 --- /dev/null +++ b/Documentation/filesystems/mandatory-locking.txt @@ -0,0 +1,171 @@ + Mandatory File Locking For The Linux Operating System + + Andy Walker <andy@lysaker.kvaerner.no> + + 15 April 1996 + (Updated September 2007) + +0. Why you should avoid mandatory locking +----------------------------------------- + +The Linux implementation is prey to a number of difficult-to-fix race +conditions which in practice make it not dependable: + + - The write system call checks for a mandatory lock only once + at its start. It is therefore possible for a lock request to + be granted after this check but before the data is modified. + A process may then see file data change even while a mandatory + lock was held. + - Similarly, an exclusive lock may be granted on a file after + the kernel has decided to proceed with a read, but before the + read has actually completed, and the reading process may see + the file data in a state which should not have been visible + to it. + - Similar races make the claimed mutual exclusion between lock + and mmap similarly unreliable. + +1. What is mandatory locking? +------------------------------ + +Mandatory locking is kernel enforced file locking, as opposed to the more usual +cooperative file locking used to guarantee sequential access to files among +processes. File locks are applied using the flock() and fcntl() system calls +(and the lockf() library routine which is a wrapper around fcntl().) It is +normally a process' responsibility to check for locks on a file it wishes to +update, before applying its own lock, updating the file and unlocking it again. +The most commonly used example of this (and in the case of sendmail, the most +troublesome) is access to a user's mailbox. The mail user agent and the mail +transfer agent must guard against updating the mailbox at the same time, and +prevent reading the mailbox while it is being updated. + +In a perfect world all processes would use and honour a cooperative, or +"advisory" locking scheme. However, the world isn't perfect, and there's +a lot of poorly written code out there. + +In trying to address this problem, the designers of System V UNIX came up +with a "mandatory" locking scheme, whereby the operating system kernel would +block attempts by a process to write to a file that another process holds a +"read" -or- "shared" lock on, and block attempts to both read and write to a +file that a process holds a "write " -or- "exclusive" lock on. + +The System V mandatory locking scheme was intended to have as little impact as +possible on existing user code. The scheme is based on marking individual files +as candidates for mandatory locking, and using the existing fcntl()/lockf() +interface for applying locks just as if they were normal, advisory locks. + +Note 1: In saying "file" in the paragraphs above I am actually not telling +the whole truth. System V locking is based on fcntl(). The granularity of +fcntl() is such that it allows the locking of byte ranges in files, in addition +to entire files, so the mandatory locking rules also have byte level +granularity. + +Note 2: POSIX.1 does not specify any scheme for mandatory locking, despite +borrowing the fcntl() locking scheme from System V. The mandatory locking +scheme is defined by the System V Interface Definition (SVID) Version 3. + +2. Marking a file for mandatory locking +--------------------------------------- + +A file is marked as a candidate for mandatory locking by setting the group-id +bit in its file mode but removing the group-execute bit. This is an otherwise +meaningless combination, and was chosen by the System V implementors so as not +to break existing user programs. + +Note that the group-id bit is usually automatically cleared by the kernel when +a setgid file is written to. This is a security measure. The kernel has been +modified to recognize the special case of a mandatory lock candidate and to +refrain from clearing this bit. Similarly the kernel has been modified not +to run mandatory lock candidates with setgid privileges. + +3. Available implementations +---------------------------- + +I have considered the implementations of mandatory locking available with +SunOS 4.1.x, Solaris 2.x and HP-UX 9.x. + +Generally I have tried to make the most sense out of the behaviour exhibited +by these three reference systems. There are many anomalies. + +All the reference systems reject all calls to open() for a file on which +another process has outstanding mandatory locks. This is in direct +contravention of SVID 3, which states that only calls to open() with the +O_TRUNC flag set should be rejected. The Linux implementation follows the SVID +definition, which is the "Right Thing", since only calls with O_TRUNC can +modify the contents of the file. + +HP-UX even disallows open() with O_TRUNC for a file with advisory locks, not +just mandatory locks. That would appear to contravene POSIX.1. + +mmap() is another interesting case. All the operating systems mentioned +prevent mandatory locks from being applied to an mmap()'ed file, but HP-UX +also disallows advisory locks for such a file. SVID actually specifies the +paranoid HP-UX behaviour. + +In my opinion only MAP_SHARED mappings should be immune from locking, and then +only from mandatory locks - that is what is currently implemented. + +SunOS is so hopeless that it doesn't even honour the O_NONBLOCK flag for +mandatory locks, so reads and writes to locked files always block when they +should return EAGAIN. + +I'm afraid that this is such an esoteric area that the semantics described +below are just as valid as any others, so long as the main points seem to +agree. + +4. Semantics +------------ + +1. Mandatory locks can only be applied via the fcntl()/lockf() locking + interface - in other words the System V/POSIX interface. BSD style + locks using flock() never result in a mandatory lock. + +2. If a process has locked a region of a file with a mandatory read lock, then + other processes are permitted to read from that region. If any of these + processes attempts to write to the region it will block until the lock is + released, unless the process has opened the file with the O_NONBLOCK + flag in which case the system call will return immediately with the error + status EAGAIN. + +3. If a process has locked a region of a file with a mandatory write lock, all + attempts to read or write to that region block until the lock is released, + unless a process has opened the file with the O_NONBLOCK flag in which case + the system call will return immediately with the error status EAGAIN. + +4. Calls to open() with O_TRUNC, or to creat(), on a existing file that has + any mandatory locks owned by other processes will be rejected with the + error status EAGAIN. + +5. Attempts to apply a mandatory lock to a file that is memory mapped and + shared (via mmap() with MAP_SHARED) will be rejected with the error status + EAGAIN. + +6. Attempts to create a shared memory map of a file (via mmap() with MAP_SHARED) + that has any mandatory locks in effect will be rejected with the error status + EAGAIN. + +5. Which system calls are affected? +----------------------------------- + +Those which modify a file's contents, not just the inode. That gives read(), +write(), readv(), writev(), open(), creat(), mmap(), truncate() and +ftruncate(). truncate() and ftruncate() are considered to be "write" actions +for the purposes of mandatory locking. + +The affected region is usually defined as stretching from the current position +for the total number of bytes read or written. For the truncate calls it is +defined as the bytes of a file removed or added (we must also consider bytes +added, as a lock can specify just "the whole file", rather than a specific +range of bytes.) + +Note 3: I may have overlooked some system calls that need mandatory lock +checking in my eagerness to get this code out the door. Please let me know, or +better still fix the system calls yourself and submit a patch to me or Linus. + +6. Warning! +----------- + +Not even root can override a mandatory lock, so runaway processes can wreak +havoc if they lock crucial files. The way around it is to change the file +permissions (remove the setgid bit) before trying to read or write to it. +Of course, that might be a bit tricky if the system is hung :-( + diff --git a/Documentation/filesystems/ncpfs.txt b/Documentation/filesystems/ncpfs.txt new file mode 100644 index 00000000..5af164f4 --- /dev/null +++ b/Documentation/filesystems/ncpfs.txt @@ -0,0 +1,12 @@ +The ncpfs filesystem understands the NCP protocol, designed by the +Novell Corporation for their NetWare(tm) product. NCP is functionally +similar to the NFS used in the TCP/IP community. +To mount a NetWare filesystem, you need a special mount program, which +can be found in the ncpfs package. The home site for ncpfs is +ftp.gwdg.de/pub/linux/misc/ncpfs, but sunsite and its many mirrors +will have it as well. + +Related products are linware and mars_nwe, which will give Linux partial +NetWare server functionality. + +mars_nwe can be found on ftp.gwdg.de/pub/linux/misc/ncpfs. diff --git a/Documentation/filesystems/nfs/00-INDEX b/Documentation/filesystems/nfs/00-INDEX new file mode 100644 index 00000000..a57e1241 --- /dev/null +++ b/Documentation/filesystems/nfs/00-INDEX @@ -0,0 +1,20 @@ +00-INDEX + - this file (nfs-related documentation). +Exporting + - explanation of how to make filesystems exportable. +knfsd-stats.txt + - statistics which the NFS server makes available to user space. +nfs.txt + - nfs client, and DNS resolution for fs_locations. +nfs41-server.txt + - info on the Linux server implementation of NFSv4 minor version 1. +nfs-rdma.txt + - how to install and setup the Linux NFS/RDMA client and server software +nfsroot.txt + - short guide on setting up a diskless box with NFS root filesystem. +pnfs.txt + - short explanation of some of the internals of the pnfs client code +rpc-cache.txt + - introduction to the caching mechanisms in the sunrpc layer. +idmapper.txt + - information for configuring request-keys to be used by idmapper diff --git a/Documentation/filesystems/nfs/Exporting b/Documentation/filesystems/nfs/Exporting new file mode 100644 index 00000000..87019d2b --- /dev/null +++ b/Documentation/filesystems/nfs/Exporting @@ -0,0 +1,147 @@ + +Making Filesystems Exportable +============================= + +Overview +-------- + +All filesystem operations require a dentry (or two) as a starting +point. Local applications have a reference-counted hold on suitable +dentries via open file descriptors or cwd/root. However remote +applications that access a filesystem via a remote filesystem protocol +such as NFS may not be able to hold such a reference, and so need a +different way to refer to a particular dentry. As the alternative +form of reference needs to be stable across renames, truncates, and +server-reboot (among other things, though these tend to be the most +problematic), there is no simple answer like 'filename'. + +The mechanism discussed here allows each filesystem implementation to +specify how to generate an opaque (outside of the filesystem) byte +string for any dentry, and how to find an appropriate dentry for any +given opaque byte string. +This byte string will be called a "filehandle fragment" as it +corresponds to part of an NFS filehandle. + +A filesystem which supports the mapping between filehandle fragments +and dentries will be termed "exportable". + + + +Dcache Issues +------------- + +The dcache normally contains a proper prefix of any given filesystem +tree. This means that if any filesystem object is in the dcache, then +all of the ancestors of that filesystem object are also in the dcache. +As normal access is by filename this prefix is created naturally and +maintained easily (by each object maintaining a reference count on +its parent). + +However when objects are included into the dcache by interpreting a +filehandle fragment, there is no automatic creation of a path prefix +for the object. This leads to two related but distinct features of +the dcache that are not needed for normal filesystem access. + +1/ The dcache must sometimes contain objects that are not part of the + proper prefix. i.e that are not connected to the root. +2/ The dcache must be prepared for a newly found (via ->lookup) directory + to already have a (non-connected) dentry, and must be able to move + that dentry into place (based on the parent and name in the + ->lookup). This is particularly needed for directories as + it is a dcache invariant that directories only have one dentry. + +To implement these features, the dcache has: + +a/ A dentry flag DCACHE_DISCONNECTED which is set on + any dentry that might not be part of the proper prefix. + This is set when anonymous dentries are created, and cleared when a + dentry is noticed to be a child of a dentry which is in the proper + prefix. + +b/ A per-superblock list "s_anon" of dentries which are the roots of + subtrees that are not in the proper prefix. These dentries, as + well as the proper prefix, need to be released at unmount time. As + these dentries will not be hashed, they are linked together on the + d_hash list_head. + +c/ Helper routines to allocate anonymous dentries, and to help attach + loose directory dentries at lookup time. They are: + d_alloc_anon(inode) will return a dentry for the given inode. + If the inode already has a dentry, one of those is returned. + If it doesn't, a new anonymous (IS_ROOT and + DCACHE_DISCONNECTED) dentry is allocated and attached. + In the case of a directory, care is taken that only one dentry + can ever be attached. + d_splice_alias(inode, dentry) will make sure that there is a + dentry with the same name and parent as the given dentry, and + which refers to the given inode. + If the inode is a directory and already has a dentry, then that + dentry is d_moved over the given dentry. + If the passed dentry gets attached, care is taken that this is + mutually exclusive to a d_alloc_anon operation. + If the passed dentry is used, NULL is returned, else the used + dentry is returned. This corresponds to the calling pattern of + ->lookup. + + +Filesystem Issues +----------------- + +For a filesystem to be exportable it must: + + 1/ provide the filehandle fragment routines described below. + 2/ make sure that d_splice_alias is used rather than d_add + when ->lookup finds an inode for a given parent and name. + Typically the ->lookup routine will end with a: + + return d_splice_alias(inode, dentry); + } + + + + A file system implementation declares that instances of the filesystem +are exportable by setting the s_export_op field in the struct +super_block. This field must point to a "struct export_operations" +struct which has the following members: + + encode_fh (optional) + Takes a dentry and creates a filehandle fragment which can later be used + to find or create a dentry for the same object. The default + implementation creates a filehandle fragment that encodes a 32bit inode + and generation number for the inode encoded, and if necessary the + same information for the parent. + + fh_to_dentry (mandatory) + Given a filehandle fragment, this should find the implied object and + create a dentry for it (possibly with d_alloc_anon). + + fh_to_parent (optional but strongly recommended) + Given a filehandle fragment, this should find the parent of the + implied object and create a dentry for it (possibly with d_alloc_anon). + May fail if the filehandle fragment is too small. + + get_parent (optional but strongly recommended) + When given a dentry for a directory, this should return a dentry for + the parent. Quite possibly the parent dentry will have been allocated + by d_alloc_anon. The default get_parent function just returns an error + so any filehandle lookup that requires finding a parent will fail. + ->lookup("..") is *not* used as a default as it can leave ".." entries + in the dcache which are too messy to work with. + + get_name (optional) + When given a parent dentry and a child dentry, this should find a name + in the directory identified by the parent dentry, which leads to the + object identified by the child dentry. If no get_name function is + supplied, a default implementation is provided which uses vfs_readdir + to find potential names, and matches inode numbers to find the correct + match. + + +A filehandle fragment consists of an array of 1 or more 4byte words, +together with a one byte "type". +The decode_fh routine should not depend on the stated size that is +passed to it. This size may be larger than the original filehandle +generated by encode_fh, in which case it will have been padded with +nuls. Rather, the encode_fh routine should choose a "type" which +indicates the decode_fh how much of the filehandle is valid, and how +it should be interpreted. diff --git a/Documentation/filesystems/nfs/idmapper.txt b/Documentation/filesystems/nfs/idmapper.txt new file mode 100644 index 00000000..9c8fd614 --- /dev/null +++ b/Documentation/filesystems/nfs/idmapper.txt @@ -0,0 +1,67 @@ + +========= +ID Mapper +========= +Id mapper is used by NFS to translate user and group ids into names, and to +translate user and group names into ids. Part of this translation involves +performing an upcall to userspace to request the information. Id mapper will +user request-key to perform this upcall and cache the result. The program +/usr/sbin/nfs.idmap should be called by request-key, and will perform the +translation and initialize a key with the resulting information. + + NFS_USE_NEW_IDMAPPER must be selected when configuring the kernel to use this + feature. + +=========== +Configuring +=========== +The file /etc/request-key.conf will need to be modified so /sbin/request-key can +direct the upcall. The following line should be added: + +#OP TYPE DESCRIPTION CALLOUT INFO PROGRAM ARG1 ARG2 ARG3 ... +#====== ======= =============== =============== =============================== +create id_resolver * * /usr/sbin/nfs.idmap %k %d 600 + +This will direct all id_resolver requests to the program /usr/sbin/nfs.idmap. +The last parameter, 600, defines how many seconds into the future the key will +expire. This parameter is optional for /usr/sbin/nfs.idmap. When the timeout +is not specified, nfs.idmap will default to 600 seconds. + +id mapper uses for key descriptions: + uid: Find the UID for the given user + gid: Find the GID for the given group + user: Find the user name for the given UID + group: Find the group name for the given GID + +You can handle any of these individually, rather than using the generic upcall +program. If you would like to use your own program for a uid lookup then you +would edit your request-key.conf so it look similar to this: + +#OP TYPE DESCRIPTION CALLOUT INFO PROGRAM ARG1 ARG2 ARG3 ... +#====== ======= =============== =============== =============================== +create id_resolver uid:* * /some/other/program %k %d 600 +create id_resolver * * /usr/sbin/nfs.idmap %k %d 600 + +Notice that the new line was added above the line for the generic program. +request-key will find the first matching line and corresponding program. In +this case, /some/other/program will handle all uid lookups and +/usr/sbin/nfs.idmap will handle gid, user, and group lookups. + +See <file:Documentation/security/keys-request-keys.txt> for more information +about the request-key function. + + +========= +nfs.idmap +========= +nfs.idmap is designed to be called by request-key, and should not be run "by +hand". This program takes two arguments, a serialized key and a key +description. The serialized key is first converted into a key_serial_t, and +then passed as an argument to keyctl_instantiate (both are part of keyutils.h). + +The actual lookups are performed by functions found in nfsidmap.h. nfs.idmap +determines the correct function to call by looking at the first part of the +description string. For example, a uid lookup description will appear as +"uid:user@domain". + +nfs.idmap will return 0 if the key was instantiated, and non-zero otherwise. diff --git a/Documentation/filesystems/nfs/knfsd-stats.txt b/Documentation/filesystems/nfs/knfsd-stats.txt new file mode 100644 index 00000000..64ced514 --- /dev/null +++ b/Documentation/filesystems/nfs/knfsd-stats.txt @@ -0,0 +1,159 @@ + +Kernel NFS Server Statistics +============================ + +This document describes the format and semantics of the statistics +which the kernel NFS server makes available to userspace. These +statistics are available in several text form pseudo files, each of +which is described separately below. + +In most cases you don't need to know these formats, as the nfsstat(8) +program from the nfs-utils distribution provides a helpful command-line +interface for extracting and printing them. + +All the files described here are formatted as a sequence of text lines, +separated by newline '\n' characters. Lines beginning with a hash +'#' character are comments intended for humans and should be ignored +by parsing routines. All other lines contain a sequence of fields +separated by whitespace. + +/proc/fs/nfsd/pool_stats +------------------------ + +This file is available in kernels from 2.6.30 onwards, if the +/proc/fs/nfsd filesystem is mounted (it almost always should be). + +The first line is a comment which describes the fields present in +all the other lines. The other lines present the following data as +a sequence of unsigned decimal numeric fields. One line is shown +for each NFS thread pool. + +All counters are 64 bits wide and wrap naturally. There is no way +to zero these counters, instead applications should do their own +rate conversion. + +pool + The id number of the NFS thread pool to which this line applies. + This number does not change. + + Thread pool ids are a contiguous set of small integers starting + at zero. The maximum value depends on the thread pool mode, but + currently cannot be larger than the number of CPUs in the system. + Note that in the default case there will be a single thread pool + which contains all the nfsd threads and all the CPUs in the system, + and thus this file will have a single line with a pool id of "0". + +packets-arrived + Counts how many NFS packets have arrived. More precisely, this + is the number of times that the network stack has notified the + sunrpc server layer that new data may be available on a transport + (e.g. an NFS or UDP socket or an NFS/RDMA endpoint). + + Depending on the NFS workload patterns and various network stack + effects (such as Large Receive Offload) which can combine packets + on the wire, this may be either more or less than the number + of NFS calls received (which statistic is available elsewhere). + However this is a more accurate and less workload-dependent measure + of how much CPU load is being placed on the sunrpc server layer + due to NFS network traffic. + +sockets-enqueued + Counts how many times an NFS transport is enqueued to wait for + an nfsd thread to service it, i.e. no nfsd thread was considered + available. + + The circumstance this statistic tracks indicates that there was NFS + network-facing work to be done but it couldn't be done immediately, + thus introducing a small delay in servicing NFS calls. The ideal + rate of change for this counter is zero; significantly non-zero + values may indicate a performance limitation. + + This can happen either because there are too few nfsd threads in the + thread pool for the NFS workload (the workload is thread-limited), + or because the NFS workload needs more CPU time than is available in + the thread pool (the workload is CPU-limited). In the former case, + configuring more nfsd threads will probably improve the performance + of the NFS workload. In the latter case, the sunrpc server layer is + already choosing not to wake idle nfsd threads because there are too + many nfsd threads which want to run but cannot, so configuring more + nfsd threads will make no difference whatsoever. The overloads-avoided + statistic (see below) can be used to distinguish these cases. + +threads-woken + Counts how many times an idle nfsd thread is woken to try to + receive some data from an NFS transport. + + This statistic tracks the circumstance where incoming + network-facing NFS work is being handled quickly, which is a good + thing. The ideal rate of change for this counter will be close + to but less than the rate of change of the packets-arrived counter. + +overloads-avoided + Counts how many times the sunrpc server layer chose not to wake an + nfsd thread, despite the presence of idle nfsd threads, because + too many nfsd threads had been recently woken but could not get + enough CPU time to actually run. + + This statistic counts a circumstance where the sunrpc layer + heuristically avoids overloading the CPU scheduler with too many + runnable nfsd threads. The ideal rate of change for this counter + is zero. Significant non-zero values indicate that the workload + is CPU limited. Usually this is associated with heavy CPU usage + on all the CPUs in the nfsd thread pool. + + If a sustained large overloads-avoided rate is detected on a pool, + the top(1) utility should be used to check for the following + pattern of CPU usage on all the CPUs associated with the given + nfsd thread pool. + + - %us ~= 0 (as you're *NOT* running applications on your NFS server) + + - %wa ~= 0 + + - %id ~= 0 + + - %sy + %hi + %si ~= 100 + + If this pattern is seen, configuring more nfsd threads will *not* + improve the performance of the workload. If this patten is not + seen, then something more subtle is wrong. + +threads-timedout + Counts how many times an nfsd thread triggered an idle timeout, + i.e. was not woken to handle any incoming network packets for + some time. + + This statistic counts a circumstance where there are more nfsd + threads configured than can be used by the NFS workload. This is + a clue that the number of nfsd threads can be reduced without + affecting performance. Unfortunately, it's only a clue and not + a strong indication, for a couple of reasons: + + - Currently the rate at which the counter is incremented is quite + slow; the idle timeout is 60 minutes. Unless the NFS workload + remains constant for hours at a time, this counter is unlikely + to be providing information that is still useful. + + - It is usually a wise policy to provide some slack, + i.e. configure a few more nfsds than are currently needed, + to allow for future spikes in load. + + +Note that incoming packets on NFS transports will be dealt with in +one of three ways. An nfsd thread can be woken (threads-woken counts +this case), or the transport can be enqueued for later attention +(sockets-enqueued counts this case), or the packet can be temporarily +deferred because the transport is currently being used by an nfsd +thread. This last case is not very interesting and is not explicitly +counted, but can be inferred from the other counters thus: + +packets-deferred = packets-arrived - ( sockets-enqueued + threads-woken ) + + +More +---- +Descriptions of the other statistics file should go here. + + +Greg Banks <gnb@sgi.com> +26 Mar 2009 diff --git a/Documentation/filesystems/nfs/nfs-rdma.txt b/Documentation/filesystems/nfs/nfs-rdma.txt new file mode 100644 index 00000000..e386f7e4 --- /dev/null +++ b/Documentation/filesystems/nfs/nfs-rdma.txt @@ -0,0 +1,271 @@ +################################################################################ +# # +# NFS/RDMA README # +# # +################################################################################ + + Author: NetApp and Open Grid Computing + Date: May 29, 2008 + +Table of Contents +~~~~~~~~~~~~~~~~~ + - Overview + - Getting Help + - Installation + - Check RDMA and NFS Setup + - NFS/RDMA Setup + +Overview +~~~~~~~~ + + This document describes how to install and setup the Linux NFS/RDMA client + and server software. + + The NFS/RDMA client was first included in Linux 2.6.24. The NFS/RDMA server + was first included in the following release, Linux 2.6.25. + + In our testing, we have obtained excellent performance results (full 10Gbit + wire bandwidth at minimal client CPU) under many workloads. The code passes + the full Connectathon test suite and operates over both Infiniband and iWARP + RDMA adapters. + +Getting Help +~~~~~~~~~~~~ + + If you get stuck, you can ask questions on the + + nfs-rdma-devel@lists.sourceforge.net + + mailing list. + +Installation +~~~~~~~~~~~~ + + These instructions are a step by step guide to building a machine for + use with NFS/RDMA. + + - Install an RDMA device + + Any device supported by the drivers in drivers/infiniband/hw is acceptable. + + Testing has been performed using several Mellanox-based IB cards, the + Ammasso AMS1100 iWARP adapter, and the Chelsio cxgb3 iWARP adapter. + + - Install a Linux distribution and tools + + The first kernel release to contain both the NFS/RDMA client and server was + Linux 2.6.25 Therefore, a distribution compatible with this and subsequent + Linux kernel release should be installed. + + The procedures described in this document have been tested with + distributions from Red Hat's Fedora Project (http://fedora.redhat.com/). + + - Install nfs-utils-1.1.2 or greater on the client + + An NFS/RDMA mount point can be obtained by using the mount.nfs command in + nfs-utils-1.1.2 or greater (nfs-utils-1.1.1 was the first nfs-utils + version with support for NFS/RDMA mounts, but for various reasons we + recommend using nfs-utils-1.1.2 or greater). To see which version of + mount.nfs you are using, type: + + $ /sbin/mount.nfs -V + + If the version is less than 1.1.2 or the command does not exist, + you should install the latest version of nfs-utils. + + Download the latest package from: + + http://www.kernel.org/pub/linux/utils/nfs + + Uncompress the package and follow the installation instructions. + + If you will not need the idmapper and gssd executables (you do not need + these to create an NFS/RDMA enabled mount command), the installation + process can be simplified by disabling these features when running + configure: + + $ ./configure --disable-gss --disable-nfsv4 + + To build nfs-utils you will need the tcp_wrappers package installed. For + more information on this see the package's README and INSTALL files. + + After building the nfs-utils package, there will be a mount.nfs binary in + the utils/mount directory. This binary can be used to initiate NFS v2, v3, + or v4 mounts. To initiate a v4 mount, the binary must be called + mount.nfs4. The standard technique is to create a symlink called + mount.nfs4 to mount.nfs. + + This mount.nfs binary should be installed at /sbin/mount.nfs as follows: + + $ sudo cp utils/mount/mount.nfs /sbin/mount.nfs + + In this location, mount.nfs will be invoked automatically for NFS mounts + by the system mount command. + + NOTE: mount.nfs and therefore nfs-utils-1.1.2 or greater is only needed + on the NFS client machine. You do not need this specific version of + nfs-utils on the server. Furthermore, only the mount.nfs command from + nfs-utils-1.1.2 is needed on the client. + + - Install a Linux kernel with NFS/RDMA + + The NFS/RDMA client and server are both included in the mainline Linux + kernel version 2.6.25 and later. This and other versions of the 2.6 Linux + kernel can be found at: + + ftp://ftp.kernel.org/pub/linux/kernel/v2.6/ + + Download the sources and place them in an appropriate location. + + - Configure the RDMA stack + + Make sure your kernel configuration has RDMA support enabled. Under + Device Drivers -> InfiniBand support, update the kernel configuration + to enable InfiniBand support [NOTE: the option name is misleading. Enabling + InfiniBand support is required for all RDMA devices (IB, iWARP, etc.)]. + + Enable the appropriate IB HCA support (mlx4, mthca, ehca, ipath, etc.) or + iWARP adapter support (amso, cxgb3, etc.). + + If you are using InfiniBand, be sure to enable IP-over-InfiniBand support. + + - Configure the NFS client and server + + Your kernel configuration must also have NFS file system support and/or + NFS server support enabled. These and other NFS related configuration + options can be found under File Systems -> Network File Systems. + + - Build, install, reboot + + The NFS/RDMA code will be enabled automatically if NFS and RDMA + are turned on. The NFS/RDMA client and server are configured via the hidden + SUNRPC_XPRT_RDMA config option that depends on SUNRPC and INFINIBAND. The + value of SUNRPC_XPRT_RDMA will be: + + - N if either SUNRPC or INFINIBAND are N, in this case the NFS/RDMA client + and server will not be built + - M if both SUNRPC and INFINIBAND are on (M or Y) and at least one is M, + in this case the NFS/RDMA client and server will be built as modules + - Y if both SUNRPC and INFINIBAND are Y, in this case the NFS/RDMA client + and server will be built into the kernel + + Therefore, if you have followed the steps above and turned no NFS and RDMA, + the NFS/RDMA client and server will be built. + + Build a new kernel, install it, boot it. + +Check RDMA and NFS Setup +~~~~~~~~~~~~~~~~~~~~~~~~ + + Before configuring the NFS/RDMA software, it is a good idea to test + your new kernel to ensure that the kernel is working correctly. + In particular, it is a good idea to verify that the RDMA stack + is functioning as expected and standard NFS over TCP/IP and/or UDP/IP + is working properly. + + - Check RDMA Setup + + If you built the RDMA components as modules, load them at + this time. For example, if you are using a Mellanox Tavor/Sinai/Arbel + card: + + $ modprobe ib_mthca + $ modprobe ib_ipoib + + If you are using InfiniBand, make sure there is a Subnet Manager (SM) + running on the network. If your IB switch has an embedded SM, you can + use it. Otherwise, you will need to run an SM, such as OpenSM, on one + of your end nodes. + + If an SM is running on your network, you should see the following: + + $ cat /sys/class/infiniband/driverX/ports/1/state + 4: ACTIVE + + where driverX is mthca0, ipath5, ehca3, etc. + + To further test the InfiniBand software stack, use IPoIB (this + assumes you have two IB hosts named host1 and host2): + + host1$ ifconfig ib0 a.b.c.x + host2$ ifconfig ib0 a.b.c.y + host1$ ping a.b.c.y + host2$ ping a.b.c.x + + For other device types, follow the appropriate procedures. + + - Check NFS Setup + + For the NFS components enabled above (client and/or server), + test their functionality over standard Ethernet using TCP/IP or UDP/IP. + +NFS/RDMA Setup +~~~~~~~~~~~~~~ + + We recommend that you use two machines, one to act as the client and + one to act as the server. + + One time configuration: + + - On the server system, configure the /etc/exports file and + start the NFS/RDMA server. + + Exports entries with the following formats have been tested: + + /vol0 192.168.0.47(fsid=0,rw,async,insecure,no_root_squash) + /vol0 192.168.0.0/255.255.255.0(fsid=0,rw,async,insecure,no_root_squash) + + The IP address(es) is(are) the client's IPoIB address for an InfiniBand + HCA or the cleint's iWARP address(es) for an RNIC. + + NOTE: The "insecure" option must be used because the NFS/RDMA client does + not use a reserved port. + + Each time a machine boots: + + - Load and configure the RDMA drivers + + For InfiniBand using a Mellanox adapter: + + $ modprobe ib_mthca + $ modprobe ib_ipoib + $ ifconfig ib0 a.b.c.d + + NOTE: use unique addresses for the client and server + + - Start the NFS server + + If the NFS/RDMA server was built as a module (CONFIG_SUNRPC_XPRT_RDMA=m in + kernel config), load the RDMA transport module: + + $ modprobe svcrdma + + Regardless of how the server was built (module or built-in), start the + server: + + $ /etc/init.d/nfs start + + or + + $ service nfs start + + Instruct the server to listen on the RDMA transport: + + $ echo rdma 20049 > /proc/fs/nfsd/portlist + + - On the client system + + If the NFS/RDMA client was built as a module (CONFIG_SUNRPC_XPRT_RDMA=m in + kernel config), load the RDMA client module: + + $ modprobe xprtrdma.ko + + Regardless of how the client was built (module or built-in), use this + command to mount the NFS/RDMA server: + + $ mount -o rdma,port=20049 <IPoIB-server-name-or-address>:/<export> /mnt + + To verify that the mount is using RDMA, run "cat /proc/mounts" and check + the "proto" field for the given mount. + + Congratulations! You're using NFS/RDMA! diff --git a/Documentation/filesystems/nfs/nfs.txt b/Documentation/filesystems/nfs/nfs.txt new file mode 100644 index 00000000..f50f26ce --- /dev/null +++ b/Documentation/filesystems/nfs/nfs.txt @@ -0,0 +1,98 @@ + +The NFS client +============== + +The NFS version 2 protocol was first documented in RFC1094 (March 1989). +Since then two more major releases of NFS have been published, with NFSv3 +being documented in RFC1813 (June 1995), and NFSv4 in RFC3530 (April +2003). + +The Linux NFS client currently supports all the above published versions, +and work is in progress on adding support for minor version 1 of the NFSv4 +protocol. + +The purpose of this document is to provide information on some of the +upcall interfaces that are used in order to provide the NFS client with +some of the information that it requires in order to fully comply with +the NFS spec. + +The DNS resolver +================ + +NFSv4 allows for one server to refer the NFS client to data that has been +migrated onto another server by means of the special "fs_locations" +attribute. See + http://tools.ietf.org/html/rfc3530#section-6 +and + http://tools.ietf.org/html/draft-ietf-nfsv4-referrals-00 + +The fs_locations information can take the form of either an ip address and +a path, or a DNS hostname and a path. The latter requires the NFS client to +do a DNS lookup in order to mount the new volume, and hence the need for an +upcall to allow userland to provide this service. + +Assuming that the user has the 'rpc_pipefs' filesystem mounted in the usual +/var/lib/nfs/rpc_pipefs, the upcall consists of the following steps: + + (1) The process checks the dns_resolve cache to see if it contains a + valid entry. If so, it returns that entry and exits. + + (2) If no valid entry exists, the helper script '/sbin/nfs_cache_getent' + (may be changed using the 'nfs.cache_getent' kernel boot parameter) + is run, with two arguments: + - the cache name, "dns_resolve" + - the hostname to resolve + + (3) After looking up the corresponding ip address, the helper script + writes the result into the rpc_pipefs pseudo-file + '/var/lib/nfs/rpc_pipefs/cache/dns_resolve/channel' + in the following (text) format: + + "<ip address> <hostname> <ttl>\n" + + Where <ip address> is in the usual IPv4 (123.456.78.90) or IPv6 + (ffee:ddcc:bbaa:9988:7766:5544:3322:1100, ffee::1100, ...) format. + <hostname> is identical to the second argument of the helper + script, and <ttl> is the 'time to live' of this cache entry (in + units of seconds). + + Note: If <ip address> is invalid, say the string "0", then a negative + entry is created, which will cause the kernel to treat the hostname + as having no valid DNS translation. + + + + +A basic sample /sbin/nfs_cache_getent +===================================== + +#!/bin/bash +# +ttl=600 +# +cut=/usr/bin/cut +getent=/usr/bin/getent +rpc_pipefs=/var/lib/nfs/rpc_pipefs +# +die() +{ + echo "Usage: $0 cache_name entry_name" + exit 1 +} + +[ $# -lt 2 ] && die +cachename="$1" +cache_path=${rpc_pipefs}/cache/${cachename}/channel + +case "${cachename}" in + dns_resolve) + name="$2" + result="$(${getent} hosts ${name} | ${cut} -f1 -d\ )" + [ -z "${result}" ] && result="0" + ;; + *) + die + ;; +esac +echo "${result} ${name} ${ttl}" >${cache_path} + diff --git a/Documentation/filesystems/nfs/nfs41-server.txt b/Documentation/filesystems/nfs/nfs41-server.txt new file mode 100644 index 00000000..04884914 --- /dev/null +++ b/Documentation/filesystems/nfs/nfs41-server.txt @@ -0,0 +1,221 @@ +NFSv4.1 Server Implementation + +Server support for minorversion 1 can be controlled using the +/proc/fs/nfsd/versions control file. The string output returned +by reading this file will contain either "+4.1" or "-4.1" +correspondingly. + +Currently, server support for minorversion 1 is disabled by default. +It can be enabled at run time by writing the string "+4.1" to +the /proc/fs/nfsd/versions control file. Note that to write this +control file, the nfsd service must be taken down. Use your user-mode +nfs-utils to set this up; see rpc.nfsd(8) + +(Warning: older servers will interpret "+4.1" and "-4.1" as "+4" and +"-4", respectively. Therefore, code meant to work on both new and old +kernels must turn 4.1 on or off *before* turning support for version 4 +on or off; rpc.nfsd does this correctly.) + +The NFSv4 minorversion 1 (NFSv4.1) implementation in nfsd is based +on RFC 5661. + +From the many new features in NFSv4.1 the current implementation +focuses on the mandatory-to-implement NFSv4.1 Sessions, providing +"exactly once" semantics and better control and throttling of the +resources allocated for each client. + +Other NFSv4.1 features, Parallel NFS operations in particular, +are still under development out of tree. +See http://wiki.linux-nfs.org/wiki/index.php/PNFS_prototype_design +for more information. + +The current implementation is intended for developers only: while it +does support ordinary file operations on clients we have tested against +(including the linux client), it is incomplete in ways which may limit +features unexpectedly, cause known bugs in rare cases, or cause +interoperability problems with future clients. Known issues: + + - gss support is questionable: currently mounts with kerberos + from a linux client are possible, but we aren't really + conformant with the spec (for example, we don't use kerberos + on the backchannel correctly). + - no trunking support: no clients currently take advantage of + trunking, but this is a mandatory feature, and its use is + recommended to clients in a number of places. (E.g. to ensure + timely renewal in case an existing connection's retry timeouts + have gotten too long; see section 8.3 of the RFC.) + Therefore, lack of this feature may cause future clients to + fail. + - Incomplete backchannel support: incomplete backchannel gss + support and no support for BACKCHANNEL_CTL mean that + callbacks (hence delegations and layouts) may not be + available and clients confused by the incomplete + implementation may fail. + - Server reboot recovery is unsupported; if the server reboots, + clients may fail. + - We do not support SSV, which provides security for shared + client-server state (thus preventing unauthorized tampering + with locks and opens, for example). It is mandatory for + servers to support this, though no clients use it yet. + - Mandatory operations which we do not support, such as + DESTROY_CLIENTID, FREE_STATEID, SECINFO_NO_NAME, and + TEST_STATEID, are not currently used by clients, but will be + (and the spec recommends their uses in common cases), and + clients should not be expected to know how to recover from the + case where they are not supported. This will eventually cause + interoperability failures. + +In addition, some limitations are inherited from the current NFSv4 +implementation: + + - Incomplete delegation enforcement: if a file is renamed or + unlinked, a client holding a delegation may continue to + indefinitely allow opens of the file under the old name. + +The table below, taken from the NFSv4.1 document, lists +the operations that are mandatory to implement (REQ), optional +(OPT), and NFSv4.0 operations that are required not to implement (MNI) +in minor version 1. The first column indicates the operations that +are not supported yet by the linux server implementation. + +The OPTIONAL features identified and their abbreviations are as follows: + pNFS Parallel NFS + FDELG File Delegations + DDELG Directory Delegations + +The following abbreviations indicate the linux server implementation status. + I Implemented NFSv4.1 operations. + NS Not Supported. + NS* unimplemented optional feature. + P pNFS features implemented out of tree. + PNS pNFS features that are not supported yet (out of tree). + +Operations + + +----------------------+------------+--------------+----------------+ + | Operation | REQ, REC, | Feature | Definition | + | | OPT, or | (REQ, REC, | | + | | MNI | or OPT) | | + +----------------------+------------+--------------+----------------+ + | ACCESS | REQ | | Section 18.1 | +NS | BACKCHANNEL_CTL | REQ | | Section 18.33 | +NS | BIND_CONN_TO_SESSION | REQ | | Section 18.34 | + | CLOSE | REQ | | Section 18.2 | + | COMMIT | REQ | | Section 18.3 | + | CREATE | REQ | | Section 18.4 | +I | CREATE_SESSION | REQ | | Section 18.36 | +NS*| DELEGPURGE | OPT | FDELG (REQ) | Section 18.5 | + | DELEGRETURN | OPT | FDELG, | Section 18.6 | + | | | DDELG, pNFS | | + | | | (REQ) | | +NS | DESTROY_CLIENTID | REQ | | Section 18.50 | +I | DESTROY_SESSION | REQ | | Section 18.37 | +I | EXCHANGE_ID | REQ | | Section 18.35 | +NS | FREE_STATEID | REQ | | Section 18.38 | + | GETATTR | REQ | | Section 18.7 | +P | GETDEVICEINFO | OPT | pNFS (REQ) | Section 18.40 | +P | GETDEVICELIST | OPT | pNFS (OPT) | Section 18.41 | + | GETFH | REQ | | Section 18.8 | +NS*| GET_DIR_DELEGATION | OPT | DDELG (REQ) | Section 18.39 | +P | LAYOUTCOMMIT | OPT | pNFS (REQ) | Section 18.42 | +P | LAYOUTGET | OPT | pNFS (REQ) | Section 18.43 | +P | LAYOUTRETURN | OPT | pNFS (REQ) | Section 18.44 | + | LINK | OPT | | Section 18.9 | + | LOCK | REQ | | Section 18.10 | + | LOCKT | REQ | | Section 18.11 | + | LOCKU | REQ | | Section 18.12 | + | LOOKUP | REQ | | Section 18.13 | + | LOOKUPP | REQ | | Section 18.14 | + | NVERIFY | REQ | | Section 18.15 | + | OPEN | REQ | | Section 18.16 | +NS*| OPENATTR | OPT | | Section 18.17 | + | OPEN_CONFIRM | MNI | | N/A | + | OPEN_DOWNGRADE | REQ | | Section 18.18 | + | PUTFH | REQ | | Section 18.19 | + | PUTPUBFH | REQ | | Section 18.20 | + | PUTROOTFH | REQ | | Section 18.21 | + | READ | REQ | | Section 18.22 | + | READDIR | REQ | | Section 18.23 | + | READLINK | OPT | | Section 18.24 | + | RECLAIM_COMPLETE | REQ | | Section 18.51 | + | RELEASE_LOCKOWNER | MNI | | N/A | + | REMOVE | REQ | | Section 18.25 | + | RENAME | REQ | | Section 18.26 | + | RENEW | MNI | | N/A | + | RESTOREFH | REQ | | Section 18.27 | + | SAVEFH | REQ | | Section 18.28 | + | SECINFO | REQ | | Section 18.29 | +NS | SECINFO_NO_NAME | REC | pNFS files | Section 18.45, | + | | | layout (REQ) | Section 13.12 | +I | SEQUENCE | REQ | | Section 18.46 | + | SETATTR | REQ | | Section 18.30 | + | SETCLIENTID | MNI | | N/A | + | SETCLIENTID_CONFIRM | MNI | | N/A | +NS | SET_SSV | REQ | | Section 18.47 | +NS | TEST_STATEID | REQ | | Section 18.48 | + | VERIFY | REQ | | Section 18.31 | +NS*| WANT_DELEGATION | OPT | FDELG (OPT) | Section 18.49 | + | WRITE | REQ | | Section 18.32 | + +Callback Operations + + +-------------------------+-----------+-------------+---------------+ + | Operation | REQ, REC, | Feature | Definition | + | | OPT, or | (REQ, REC, | | + | | MNI | or OPT) | | + +-------------------------+-----------+-------------+---------------+ + | CB_GETATTR | OPT | FDELG (REQ) | Section 20.1 | +P | CB_LAYOUTRECALL | OPT | pNFS (REQ) | Section 20.3 | +NS*| CB_NOTIFY | OPT | DDELG (REQ) | Section 20.4 | +P | CB_NOTIFY_DEVICEID | OPT | pNFS (OPT) | Section 20.12 | +NS*| CB_NOTIFY_LOCK | OPT | | Section 20.11 | +NS*| CB_PUSH_DELEG | OPT | FDELG (OPT) | Section 20.5 | + | CB_RECALL | OPT | FDELG, | Section 20.2 | + | | | DDELG, pNFS | | + | | | (REQ) | | +NS*| CB_RECALL_ANY | OPT | FDELG, | Section 20.6 | + | | | DDELG, pNFS | | + | | | (REQ) | | +NS | CB_RECALL_SLOT | REQ | | Section 20.8 | +NS*| CB_RECALLABLE_OBJ_AVAIL | OPT | DDELG, pNFS | Section 20.7 | + | | | (REQ) | | +I | CB_SEQUENCE | OPT | FDELG, | Section 20.9 | + | | | DDELG, pNFS | | + | | | (REQ) | | +NS*| CB_WANTS_CANCELLED | OPT | FDELG, | Section 20.10 | + | | | DDELG, pNFS | | + | | | (REQ) | | + +-------------------------+-----------+-------------+---------------+ + +Implementation notes: + +DELEGPURGE: +* mandatory only for servers that support CLAIM_DELEGATE_PREV and/or + CLAIM_DELEG_PREV_FH (which allows clients to keep delegations that + persist across client reboots). Thus we need not implement this for + now. + +EXCHANGE_ID: +* only SP4_NONE state protection supported +* implementation ids are ignored + +CREATE_SESSION: +* backchannel attributes are ignored +* backchannel security parameters are ignored + +SEQUENCE: +* no support for dynamic slot table renegotiation (optional) + +nfsv4.1 COMPOUND rules: +The following cases aren't supported yet: +* Enforcing of NFS4ERR_NOT_ONLY_OP for: BIND_CONN_TO_SESSION, CREATE_SESSION, + DESTROY_CLIENTID, DESTROY_SESSION, EXCHANGE_ID. +* DESTROY_SESSION MUST be the final operation in the COMPOUND request. + +Nonstandard compound limitations: +* No support for a sessions fore channel RPC compound that requires both a + ca_maxrequestsize request and a ca_maxresponsesize reply, so we may + fail to live up to the promise we made in CREATE_SESSION fore channel + negotiation. +* No more than one IO operation (read, write, readdir) allowed per + compound. diff --git a/Documentation/filesystems/nfs/nfsroot.txt b/Documentation/filesystems/nfs/nfsroot.txt new file mode 100644 index 00000000..90c71c6f --- /dev/null +++ b/Documentation/filesystems/nfs/nfsroot.txt @@ -0,0 +1,294 @@ +Mounting the root filesystem via NFS (nfsroot) +=============================================== + +Written 1996 by Gero Kuhlmann <gero@gkminix.han.de> +Updated 1997 by Martin Mares <mj@atrey.karlin.mff.cuni.cz> +Updated 2006 by Nico Schottelius <nico-kernel-nfsroot@schottelius.org> +Updated 2006 by Horms <horms@verge.net.au> + + + +In order to use a diskless system, such as an X-terminal or printer server +for example, it is necessary for the root filesystem to be present on a +non-disk device. This may be an initramfs (see Documentation/filesystems/ +ramfs-rootfs-initramfs.txt), a ramdisk (see Documentation/initrd.txt) or a +filesystem mounted via NFS. The following text describes on how to use NFS +for the root filesystem. For the rest of this text 'client' means the +diskless system, and 'server' means the NFS server. + + + + +1.) Enabling nfsroot capabilities + ----------------------------- + +In order to use nfsroot, NFS client support needs to be selected as +built-in during configuration. Once this has been selected, the nfsroot +option will become available, which should also be selected. + +In the networking options, kernel level autoconfiguration can be selected, +along with the types of autoconfiguration to support. Selecting all of +DHCP, BOOTP and RARP is safe. + + + + +2.) Kernel command line + ------------------- + +When the kernel has been loaded by a boot loader (see below) it needs to be +told what root fs device to use. And in the case of nfsroot, where to find +both the server and the name of the directory on the server to mount as root. +This can be established using the following kernel command line parameters: + + +root=/dev/nfs + + This is necessary to enable the pseudo-NFS-device. Note that it's not a + real device but just a synonym to tell the kernel to use NFS instead of + a real device. + + +nfsroot=[<server-ip>:]<root-dir>[,<nfs-options>] + + If the `nfsroot' parameter is NOT given on the command line, + the default "/tftpboot/%s" will be used. + + <server-ip> Specifies the IP address of the NFS server. + The default address is determined by the `ip' parameter + (see below). This parameter allows the use of different + servers for IP autoconfiguration and NFS. + + <root-dir> Name of the directory on the server to mount as root. + If there is a "%s" token in the string, it will be + replaced by the ASCII-representation of the client's + IP address. + + <nfs-options> Standard NFS options. All options are separated by commas. + The following defaults are used: + port = as given by server portmap daemon + rsize = 4096 + wsize = 4096 + timeo = 7 + retrans = 3 + acregmin = 3 + acregmax = 60 + acdirmin = 30 + acdirmax = 60 + flags = hard, nointr, noposix, cto, ac + + +ip=<client-ip>:<server-ip>:<gw-ip>:<netmask>:<hostname>:<device>:<autoconf> + + This parameter tells the kernel how to configure IP addresses of devices + and also how to set up the IP routing table. It was originally called + `nfsaddrs', but now the boot-time IP configuration works independently of + NFS, so it was renamed to `ip' and the old name remained as an alias for + compatibility reasons. + + If this parameter is missing from the kernel command line, all fields are + assumed to be empty, and the defaults mentioned below apply. In general + this means that the kernel tries to configure everything using + autoconfiguration. + + The <autoconf> parameter can appear alone as the value to the `ip' + parameter (without all the ':' characters before). If the value is + "ip=off" or "ip=none", no autoconfiguration will take place, otherwise + autoconfiguration will take place. The most common way to use this + is "ip=dhcp". + + <client-ip> IP address of the client. + + Default: Determined using autoconfiguration. + + <server-ip> IP address of the NFS server. If RARP is used to determine + the client address and this parameter is NOT empty only + replies from the specified server are accepted. + + Only required for NFS root. That is autoconfiguration + will not be triggered if it is missing and NFS root is not + in operation. + + Default: Determined using autoconfiguration. + The address of the autoconfiguration server is used. + + <gw-ip> IP address of a gateway if the server is on a different subnet. + + Default: Determined using autoconfiguration. + + <netmask> Netmask for local network interface. If unspecified + the netmask is derived from the client IP address assuming + classful addressing. + + Default: Determined using autoconfiguration. + + <hostname> Name of the client. May be supplied by autoconfiguration, + but its absence will not trigger autoconfiguration. + If specified and DHCP is used, the user provided hostname will + be carried in the DHCP request to hopefully update DNS record. + + Default: Client IP address is used in ASCII notation. + + <device> Name of network device to use. + + Default: If the host only has one device, it is used. + Otherwise the device is determined using + autoconfiguration. This is done by sending + autoconfiguration requests out of all devices, + and using the device that received the first reply. + + <autoconf> Method to use for autoconfiguration. In the case of options + which specify multiple autoconfiguration protocols, + requests are sent using all protocols, and the first one + to reply is used. + + Only autoconfiguration protocols that have been compiled + into the kernel will be used, regardless of the value of + this option. + + off or none: don't use autoconfiguration + (do static IP assignment instead) + on or any: use any protocol available in the kernel + (default) + dhcp: use DHCP + bootp: use BOOTP + rarp: use RARP + both: use both BOOTP and RARP but not DHCP + (old option kept for backwards compatibility) + + Default: any + + +nfsrootdebug + + This parameter enables debugging messages to appear in the kernel + log at boot time so that administrators can verify that the correct + NFS mount options, server address, and root path are passed to the + NFS client. + + +rdinit=<executable file> + + To specify which file contains the program that starts system + initialization, administrators can use this command line parameter. + The default value of this parameter is "/init". If the specified + file exists and the kernel can execute it, root filesystem related + kernel command line parameters, including `nfsroot=', are ignored. + + A description of the process of mounting the root file system can be + found in: + + Documentation/early-userspace/README + + + + +3.) Boot Loader + ---------- + +To get the kernel into memory different approaches can be used. +They depend on various facilities being available: + + +3.1) Booting from a floppy using syslinux + + When building kernels, an easy way to create a boot floppy that uses + syslinux is to use the zdisk or bzdisk make targets which use zimage + and bzimage images respectively. Both targets accept the + FDARGS parameter which can be used to set the kernel command line. + + e.g. + make bzdisk FDARGS="root=/dev/nfs" + + Note that the user running this command will need to have + access to the floppy drive device, /dev/fd0 + + For more information on syslinux, including how to create bootdisks + for prebuilt kernels, see http://syslinux.zytor.com/ + + N.B: Previously it was possible to write a kernel directly to + a floppy using dd, configure the boot device using rdev, and + boot using the resulting floppy. Linux no longer supports this + method of booting. + +3.2) Booting from a cdrom using isolinux + + When building kernels, an easy way to create a bootable cdrom that + uses isolinux is to use the isoimage target which uses a bzimage + image. Like zdisk and bzdisk, this target accepts the FDARGS + parameter which can be used to set the kernel command line. + + e.g. + make isoimage FDARGS="root=/dev/nfs" + + The resulting iso image will be arch/<ARCH>/boot/image.iso + This can be written to a cdrom using a variety of tools including + cdrecord. + + e.g. + cdrecord dev=ATAPI:1,0,0 arch/i386/boot/image.iso + + For more information on isolinux, including how to create bootdisks + for prebuilt kernels, see http://syslinux.zytor.com/ + +3.2) Using LILO + When using LILO all the necessary command line parameters may be + specified using the 'append=' directive in the LILO configuration + file. + + However, to use the 'root=' directive you also need to create + a dummy root device, which may be removed after LILO is run. + + mknod /dev/boot255 c 0 255 + + For information on configuring LILO, please refer to its documentation. + +3.3) Using GRUB + When using GRUB, kernel parameter are simply appended after the kernel + specification: kernel <kernel> <parameters> + +3.4) Using loadlin + loadlin may be used to boot Linux from a DOS command prompt without + requiring a local hard disk to mount as root. This has not been + thoroughly tested by the authors of this document, but in general + it should be possible configure the kernel command line similarly + to the configuration of LILO. + + Please refer to the loadlin documentation for further information. + +3.5) Using a boot ROM + This is probably the most elegant way of booting a diskless client. + With a boot ROM the kernel is loaded using the TFTP protocol. The + authors of this document are not aware of any no commercial boot + ROMs that support booting Linux over the network. However, there + are two free implementations of a boot ROM, netboot-nfs and + etherboot, both of which are available on sunsite.unc.edu, and both + of which contain everything you need to boot a diskless Linux client. + +3.6) Using pxelinux + Pxelinux may be used to boot linux using the PXE boot loader + which is present on many modern network cards. + + When using pxelinux, the kernel image is specified using + "kernel <relative-path-below /tftpboot>". The nfsroot parameters + are passed to the kernel by adding them to the "append" line. + It is common to use serial console in conjunction with pxeliunx, + see Documentation/serial-console.txt for more information. + + For more information on isolinux, including how to create bootdisks + for prebuilt kernels, see http://syslinux.zytor.com/ + + + + +4.) Credits + ------- + + The nfsroot code in the kernel and the RARP support have been written + by Gero Kuhlmann <gero@gkminix.han.de>. + + The rest of the IP layer autoconfiguration code has been written + by Martin Mares <mj@atrey.karlin.mff.cuni.cz>. + + In order to write the initial version of nfsroot I would like to thank + Jens-Uwe Mager <jum@anubis.han.de> for his help. diff --git a/Documentation/filesystems/nfs/pnfs.txt b/Documentation/filesystems/nfs/pnfs.txt new file mode 100644 index 00000000..983e14ab --- /dev/null +++ b/Documentation/filesystems/nfs/pnfs.txt @@ -0,0 +1,55 @@ +Reference counting in pnfs: +========================== + +The are several inter-related caches. We have layouts which can +reference multiple devices, each of which can reference multiple data servers. +Each data server can be referenced by multiple devices. Each device +can be referenced by multiple layouts. To keep all of this straight, +we need to reference count. + + +struct pnfs_layout_hdr +---------------------- +The on-the-wire command LAYOUTGET corresponds to struct +pnfs_layout_segment, usually referred to by the variable name lseg. +Each nfs_inode may hold a pointer to a cache of of these layout +segments in nfsi->layout, of type struct pnfs_layout_hdr. + +We reference the header for the inode pointing to it, across each +outstanding RPC call that references it (LAYOUTGET, LAYOUTRETURN, +LAYOUTCOMMIT), and for each lseg held within. + +Each header is also (when non-empty) put on a list associated with +struct nfs_client (cl_layouts). Being put on this list does not bump +the reference count, as the layout is kept around by the lseg that +keeps it in the list. + +deviceid_cache +-------------- +lsegs reference device ids, which are resolved per nfs_client and +layout driver type. The device ids are held in a RCU cache (struct +nfs4_deviceid_cache). The cache itself is referenced across each +mount. The entries (struct nfs4_deviceid) themselves are held across +the lifetime of each lseg referencing them. + +RCU is used because the deviceid is basically a write once, read many +data structure. The hlist size of 32 buckets needs better +justification, but seems reasonable given that we can have multiple +deviceid's per filesystem, and multiple filesystems per nfs_client. + +The hash code is copied from the nfsd code base. A discussion of +hashing and variations of this algorithm can be found at: +http://groups.google.com/group/comp.lang.c/browse_thread/thread/9522965e2b8d3809 + +data server cache +----------------- +file driver devices refer to data servers, which are kept in a module +level cache. Its reference is held over the lifetime of the deviceid +pointing to it. + +lseg +---- +lseg maintains an extra reference corresponding to the NFS_LSEG_VALID +bit which holds it in the pnfs_layout_hdr's list. When the final lseg +is removed from the pnfs_layout_hdr's list, the NFS_LAYOUT_DESTROYED +bit is set, preventing any new lsegs from being added. diff --git a/Documentation/filesystems/nfs/rpc-cache.txt b/Documentation/filesystems/nfs/rpc-cache.txt new file mode 100644 index 00000000..ebcaaee2 --- /dev/null +++ b/Documentation/filesystems/nfs/rpc-cache.txt @@ -0,0 +1,202 @@ + This document gives a brief introduction to the caching +mechanisms in the sunrpc layer that is used, in particular, +for NFS authentication. + +CACHES +====== +The caching replaces the old exports table and allows for +a wide variety of values to be caches. + +There are a number of caches that are similar in structure though +quite possibly very different in content and use. There is a corpus +of common code for managing these caches. + +Examples of caches that are likely to be needed are: + - mapping from IP address to client name + - mapping from client name and filesystem to export options + - mapping from UID to list of GIDs, to work around NFS's limitation + of 16 gids. + - mappings between local UID/GID and remote UID/GID for sites that + do not have uniform uid assignment + - mapping from network identify to public key for crypto authentication. + +The common code handles such things as: + - general cache lookup with correct locking + - supporting 'NEGATIVE' as well as positive entries + - allowing an EXPIRED time on cache items, and removing + items after they expire, and are no longer in-use. + - making requests to user-space to fill in cache entries + - allowing user-space to directly set entries in the cache + - delaying RPC requests that depend on as-yet incomplete + cache entries, and replaying those requests when the cache entry + is complete. + - clean out old entries as they expire. + +Creating a Cache +---------------- + +1/ A cache needs a datum to store. This is in the form of a + structure definition that must contain a + struct cache_head + as an element, usually the first. + It will also contain a key and some content. + Each cache element is reference counted and contains + expiry and update times for use in cache management. +2/ A cache needs a "cache_detail" structure that + describes the cache. This stores the hash table, some + parameters for cache management, and some operations detailing how + to work with particular cache items. + The operations requires are: + struct cache_head *alloc(void) + This simply allocates appropriate memory and returns + a pointer to the cache_detail embedded within the + structure + void cache_put(struct kref *) + This is called when the last reference to an item is + dropped. The pointer passed is to the 'ref' field + in the cache_head. cache_put should release any + references create by 'cache_init' and, if CACHE_VALID + is set, any references created by cache_update. + It should then release the memory allocated by + 'alloc'. + int match(struct cache_head *orig, struct cache_head *new) + test if the keys in the two structures match. Return + 1 if they do, 0 if they don't. + void init(struct cache_head *orig, struct cache_head *new) + Set the 'key' fields in 'new' from 'orig'. This may + include taking references to shared objects. + void update(struct cache_head *orig, struct cache_head *new) + Set the 'content' fileds in 'new' from 'orig'. + int cache_show(struct seq_file *m, struct cache_detail *cd, + struct cache_head *h) + Optional. Used to provide a /proc file that lists the + contents of a cache. This should show one item, + usually on just one line. + int cache_request(struct cache_detail *cd, struct cache_head *h, + char **bpp, int *blen) + Format a request to be send to user-space for an item + to be instantiated. *bpp is a buffer of size *blen. + bpp should be moved forward over the encoded message, + and *blen should be reduced to show how much free + space remains. Return 0 on success or <0 if not + enough room or other problem. + int cache_parse(struct cache_detail *cd, char *buf, int len) + A message from user space has arrived to fill out a + cache entry. It is in 'buf' of length 'len'. + cache_parse should parse this, find the item in the + cache with sunrpc_cache_lookup, and update the item + with sunrpc_cache_update. + + +3/ A cache needs to be registered using cache_register(). This + includes it on a list of caches that will be regularly + cleaned to discard old data. + +Using a cache +------------- + +To find a value in a cache, call sunrpc_cache_lookup passing a pointer +to the cache_head in a sample item with the 'key' fields filled in. +This will be passed to ->match to identify the target entry. If no +entry is found, a new entry will be create, added to the cache, and +marked as not containing valid data. + +The item returned is typically passed to cache_check which will check +if the data is valid, and may initiate an up-call to get fresh data. +cache_check will return -ENOENT in the entry is negative or if an up +call is needed but not possible, -EAGAIN if an upcall is pending, +or 0 if the data is valid; + +cache_check can be passed a "struct cache_req *". This structure is +typically embedded in the actual request and can be used to create a +deferred copy of the request (struct cache_deferred_req). This is +done when the found cache item is not uptodate, but the is reason to +believe that userspace might provide information soon. When the cache +item does become valid, the deferred copy of the request will be +revisited (->revisit). It is expected that this method will +reschedule the request for processing. + +The value returned by sunrpc_cache_lookup can also be passed to +sunrpc_cache_update to set the content for the item. A second item is +passed which should hold the content. If the item found by _lookup +has valid data, then it is discarded and a new item is created. This +saves any user of an item from worrying about content changing while +it is being inspected. If the item found by _lookup does not contain +valid data, then the content is copied across and CACHE_VALID is set. + +Populating a cache +------------------ + +Each cache has a name, and when the cache is registered, a directory +with that name is created in /proc/net/rpc + +This directory contains a file called 'channel' which is a channel +for communicating between kernel and user for populating the cache. +This directory may later contain other files of interacting +with the cache. + +The 'channel' works a bit like a datagram socket. Each 'write' is +passed as a whole to the cache for parsing and interpretation. +Each cache can treat the write requests differently, but it is +expected that a message written will contain: + - a key + - an expiry time + - a content. +with the intention that an item in the cache with the give key +should be create or updated to have the given content, and the +expiry time should be set on that item. + +Reading from a channel is a bit more interesting. When a cache +lookup fails, or when it succeeds but finds an entry that may soon +expire, a request is lodged for that cache item to be updated by +user-space. These requests appear in the channel file. + +Successive reads will return successive requests. +If there are no more requests to return, read will return EOF, but a +select or poll for read will block waiting for another request to be +added. + +Thus a user-space helper is likely to: + open the channel. + select for readable + read a request + write a response + loop. + +If it dies and needs to be restarted, any requests that have not been +answered will still appear in the file and will be read by the new +instance of the helper. + +Each cache should define a "cache_parse" method which takes a message +written from user-space and processes it. It should return an error +(which propagates back to the write syscall) or 0. + +Each cache should also define a "cache_request" method which +takes a cache item and encodes a request into the buffer +provided. + +Note: If a cache has no active readers on the channel, and has had not +active readers for more than 60 seconds, further requests will not be +added to the channel but instead all lookups that do not find a valid +entry will fail. This is partly for backward compatibility: The +previous nfs exports table was deemed to be authoritative and a +failed lookup meant a definite 'no'. + +request/response format +----------------------- + +While each cache is free to use its own format for requests +and responses over channel, the following is recommended as +appropriate and support routines are available to help: +Each request or response record should be printable ASCII +with precisely one newline character which should be at the end. +Fields within the record should be separated by spaces, normally one. +If spaces, newlines, or nul characters are needed in a field they +much be quoted. two mechanisms are available: +1/ If a field begins '\x' then it must contain an even number of + hex digits, and pairs of these digits provide the bytes in the + field. +2/ otherwise a \ in the field must be followed by 3 octal digits + which give the code for a byte. Other characters are treated + as them selves. At the very least, space, newline, nul, and + '\' must be quoted in this way. diff --git a/Documentation/filesystems/nilfs2.txt b/Documentation/filesystems/nilfs2.txt new file mode 100644 index 00000000..873a2ab2 --- /dev/null +++ b/Documentation/filesystems/nilfs2.txt @@ -0,0 +1,208 @@ +NILFS2 +------ + +NILFS2 is a log-structured file system (LFS) supporting continuous +snapshotting. In addition to versioning capability of the entire file +system, users can even restore files mistakenly overwritten or +destroyed just a few seconds ago. Since NILFS2 can keep consistency +like conventional LFS, it achieves quick recovery after system +crashes. + +NILFS2 creates a number of checkpoints every few seconds or per +synchronous write basis (unless there is no change). Users can select +significant versions among continuously created checkpoints, and can +change them into snapshots which will be preserved until they are +changed back to checkpoints. + +There is no limit on the number of snapshots until the volume gets +full. Each snapshot is mountable as a read-only file system +concurrently with its writable mount, and this feature is convenient +for online backup. + +The userland tools are included in nilfs-utils package, which is +available from the following download page. At least "mkfs.nilfs2", +"mount.nilfs2", "umount.nilfs2", and "nilfs_cleanerd" (so called +cleaner or garbage collector) are required. Details on the tools are +described in the man pages included in the package. + +Project web page: http://www.nilfs.org/en/ +Download page: http://www.nilfs.org/en/download.html +Git tree web page: http://www.nilfs.org/git/ +List info: http://vger.kernel.org/vger-lists.html#linux-nilfs + +Caveats +======= + +Features which NILFS2 does not support yet: + + - atime + - extended attributes + - POSIX ACLs + - quotas + - fsck + - defragmentation + +Mount options +============= + +NILFS2 supports the following mount options: +(*) == default + +barrier(*) This enables/disables the use of write barriers. This +nobarrier requires an IO stack which can support barriers, and + if nilfs gets an error on a barrier write, it will + disable again with a warning. +errors=continue Keep going on a filesystem error. +errors=remount-ro(*) Remount the filesystem read-only on an error. +errors=panic Panic and halt the machine if an error occurs. +cp=n Specify the checkpoint-number of the snapshot to be + mounted. Checkpoints and snapshots are listed by lscp + user command. Only the checkpoints marked as snapshot + are mountable with this option. Snapshot is read-only, + so a read-only mount option must be specified together. +order=relaxed(*) Apply relaxed order semantics that allows modified data + blocks to be written to disk without making a + checkpoint if no metadata update is going. This mode + is equivalent to the ordered data mode of the ext3 + filesystem except for the updates on data blocks still + conserve atomicity. This will improve synchronous + write performance for overwriting. +order=strict Apply strict in-order semantics that preserves sequence + of all file operations including overwriting of data + blocks. That means, it is guaranteed that no + overtaking of events occurs in the recovered file + system after a crash. +norecovery Disable recovery of the filesystem on mount. + This disables every write access on the device for + read-only mounts or snapshots. This option will fail + for r/w mounts on an unclean volume. +discard This enables/disables the use of discard/TRIM commands. +nodiscard(*) The discard/TRIM commands are sent to the underlying + block device when blocks are freed. This is useful + for SSD devices and sparse/thinly-provisioned LUNs. + +NILFS2 usage +============ + +To use nilfs2 as a local file system, simply: + + # mkfs -t nilfs2 /dev/block_device + # mount -t nilfs2 /dev/block_device /dir + +This will also invoke the cleaner through the mount helper program +(mount.nilfs2). + +Checkpoints and snapshots are managed by the following commands. +Their manpages are included in the nilfs-utils package above. + + lscp list checkpoints or snapshots. + mkcp make a checkpoint or a snapshot. + chcp change an existing checkpoint to a snapshot or vice versa. + rmcp invalidate specified checkpoint(s). + +To mount a snapshot, + + # mount -t nilfs2 -r -o cp=<cno> /dev/block_device /snap_dir + +where <cno> is the checkpoint number of the snapshot. + +To unmount the NILFS2 mount point or snapshot, simply: + + # umount /dir + +Then, the cleaner daemon is automatically shut down by the umount +helper program (umount.nilfs2). + +Disk format +=========== + +A nilfs2 volume is equally divided into a number of segments except +for the super block (SB) and segment #0. A segment is the container +of logs. Each log is composed of summary information blocks, payload +blocks, and an optional super root block (SR): + + ______________________________________________________ + | |SB| | Segment | Segment | Segment | ... | Segment | | + |_|__|_|____0____|____1____|____2____|_____|____N____|_| + 0 +1K +4K +8M +16M +24M +(8MB x N) + . . (Typical offsets for 4KB-block) + . . + .______________________. + | log | log |... | log | + |__1__|__2__|____|__m__| + . . + . . + . . + .______________________________. + | Summary | Payload blocks |SR| + |_blocks__|_________________|__| + +The payload blocks are organized per file, and each file consists of +data blocks and B-tree node blocks: + + |<--- File-A --->|<--- File-B --->| + _______________________________________________________________ + | Data blocks | B-tree blocks | Data blocks | B-tree blocks | ... + _|_____________|_______________|_____________|_______________|_ + + +Since only the modified blocks are written in the log, it may have +files without data blocks or B-tree node blocks. + +The organization of the blocks is recorded in the summary information +blocks, which contains a header structure (nilfs_segment_summary), per +file structures (nilfs_finfo), and per block structures (nilfs_binfo): + + _________________________________________________________________________ + | Summary | finfo | binfo | ... | binfo | finfo | binfo | ... | binfo |... + |_blocks__|___A___|_(A,1)_|_____|(A,Na)_|___B___|_(B,1)_|_____|(B,Nb)_|___ + + +The logs include regular files, directory files, symbolic link files +and several meta data files. The mata data files are the files used +to maintain file system meta data. The current version of NILFS2 uses +the following meta data files: + + 1) Inode file (ifile) -- Stores on-disk inodes + 2) Checkpoint file (cpfile) -- Stores checkpoints + 3) Segment usage file (sufile) -- Stores allocation state of segments + 4) Data address translation file -- Maps virtual block numbers to usual + (DAT) block numbers. This file serves to + make on-disk blocks relocatable. + +The following figure shows a typical organization of the logs: + + _________________________________________________________________________ + | Summary | regular file | file | ... | ifile | cpfile | sufile | DAT |SR| + |_blocks__|_or_directory_|_______|_____|_______|________|________|_____|__| + + +To stride over segment boundaries, this sequence of files may be split +into multiple logs. The sequence of logs that should be treated as +logically one log, is delimited with flags marked in the segment +summary. The recovery code of nilfs2 looks this boundary information +to ensure atomicity of updates. + +The super root block is inserted for every checkpoints. It includes +three special inodes, inodes for the DAT, cpfile, and sufile. Inodes +of regular files, directories, symlinks and other special files, are +included in the ifile. The inode of ifile itself is included in the +corresponding checkpoint entry in the cpfile. Thus, the hierarchy +among NILFS2 files can be depicted as follows: + + Super block (SB) + | + v + Super root block (the latest cno=xx) + |-- DAT + |-- sufile + `-- cpfile + |-- ifile (cno=c1) + |-- ifile (cno=c2) ---- file (ino=i1) + : : |-- file (ino=i2) + `-- ifile (cno=xx) |-- file (ino=i3) + : : + `-- file (ino=yy) + ( regular file, directory, or symlink ) + +For detail on the format of each file, please see include/linux/nilfs2_fs.h. diff --git a/Documentation/filesystems/ntfs.txt b/Documentation/filesystems/ntfs.txt new file mode 100644 index 00000000..791af8da --- /dev/null +++ b/Documentation/filesystems/ntfs.txt @@ -0,0 +1,721 @@ +The Linux NTFS filesystem driver +================================ + + +Table of contents +================= + +- Overview +- Web site +- Features +- Supported mount options +- Known bugs and (mis-)features +- Using NTFS volume and stripe sets + - The Device-Mapper driver + - The Software RAID / MD driver + - Limitations when using the MD driver +- ChangeLog + + +Overview +======== + +Linux-NTFS comes with a number of user-space programs known as ntfsprogs. +These include mkntfs, a full-featured ntfs filesystem format utility, +ntfsundelete used for recovering files that were unintentionally deleted +from an NTFS volume and ntfsresize which is used to resize an NTFS partition. +See the web site for more information. + +To mount an NTFS 1.2/3.x (Windows NT4/2000/XP/2003) volume, use the file +system type 'ntfs'. The driver currently supports read-only mode (with no +fault-tolerance, encryption or journalling) and very limited, but safe, write +support. + +For fault tolerance and raid support (i.e. volume and stripe sets), you can +use the kernel's Software RAID / MD driver. See section "Using Software RAID +with NTFS" for details. + + +Web site +======== + +There is plenty of additional information on the linux-ntfs web site +at http://www.linux-ntfs.org/ + +The web site has a lot of additional information, such as a comprehensive +FAQ, documentation on the NTFS on-disk format, information on the Linux-NTFS +userspace utilities, etc. + + +Features +======== + +- This is a complete rewrite of the NTFS driver that used to be in the 2.4 and + earlier kernels. This new driver implements NTFS read support and is + functionally equivalent to the old ntfs driver and it also implements limited + write support. The biggest limitation at present is that files/directories + cannot be created or deleted. See below for the list of write features that + are so far supported. Another limitation is that writing to compressed files + is not implemented at all. Also, neither read nor write access to encrypted + files is so far implemented. +- The new driver has full support for sparse files on NTFS 3.x volumes which + the old driver isn't happy with. +- The new driver supports execution of binaries due to mmap() now being + supported. +- The new driver supports loopback mounting of files on NTFS which is used by + some Linux distributions to enable the user to run Linux from an NTFS + partition by creating a large file while in Windows and then loopback + mounting the file while in Linux and creating a Linux filesystem on it that + is used to install Linux on it. +- A comparison of the two drivers using: + time find . -type f -exec md5sum "{}" \; + run three times in sequence with each driver (after a reboot) on a 1.4GiB + NTFS partition, showed the new driver to be 20% faster in total time elapsed + (from 9:43 minutes on average down to 7:53). The time spent in user space + was unchanged but the time spent in the kernel was decreased by a factor of + 2.5 (from 85 CPU seconds down to 33). +- The driver does not support short file names in general. For backwards + compatibility, we implement access to files using their short file names if + they exist. The driver will not create short file names however, and a + rename will discard any existing short file name. +- The new driver supports exporting of mounted NTFS volumes via NFS. +- The new driver supports async io (aio). +- The new driver supports fsync(2), fdatasync(2), and msync(2). +- The new driver supports readv(2) and writev(2). +- The new driver supports access time updates (including mtime and ctime). +- The new driver supports truncate(2) and open(2) with O_TRUNC. But at present + only very limited support for highly fragmented files, i.e. ones which have + their data attribute split across multiple extents, is included. Another + limitation is that at present truncate(2) will never create sparse files, + since to mark a file sparse we need to modify the directory entry for the + file and we do not implement directory modifications yet. +- The new driver supports write(2) which can both overwrite existing data and + extend the file size so that you can write beyond the existing data. Also, + writing into sparse regions is supported and the holes are filled in with + clusters. But at present only limited support for highly fragmented files, + i.e. ones which have their data attribute split across multiple extents, is + included. Another limitation is that write(2) will never create sparse + files, since to mark a file sparse we need to modify the directory entry for + the file and we do not implement directory modifications yet. + +Supported mount options +======================= + +In addition to the generic mount options described by the manual page for the +mount command (man 8 mount, also see man 5 fstab), the NTFS driver supports the +following mount options: + +iocharset=name Deprecated option. Still supported but please use + nls=name in the future. See description for nls=name. + +nls=name Character set to use when returning file names. + Unlike VFAT, NTFS suppresses names that contain + unconvertible characters. Note that most character + sets contain insufficient characters to represent all + possible Unicode characters that can exist on NTFS. + To be sure you are not missing any files, you are + advised to use nls=utf8 which is capable of + representing all Unicode characters. + +utf8=<bool> Option no longer supported. Currently mapped to + nls=utf8 but please use nls=utf8 in the future and + make sure utf8 is compiled either as module or into + the kernel. See description for nls=name. + +uid= +gid= +umask= Provide default owner, group, and access mode mask. + These options work as documented in mount(8). By + default, the files/directories are owned by root and + he/she has read and write permissions, as well as + browse permission for directories. No one else has any + access permissions. I.e. the mode on all files is by + default rw------- and for directories rwx------, a + consequence of the default fmask=0177 and dmask=0077. + Using a umask of zero will grant all permissions to + everyone, i.e. all files and directories will have mode + rwxrwxrwx. + +fmask= +dmask= Instead of specifying umask which applies both to + files and directories, fmask applies only to files and + dmask only to directories. + +sloppy=<BOOL> If sloppy is specified, ignore unknown mount options. + Otherwise the default behaviour is to abort mount if + any unknown options are found. + +show_sys_files=<BOOL> If show_sys_files is specified, show the system files + in directory listings. Otherwise the default behaviour + is to hide the system files. + Note that even when show_sys_files is specified, "$MFT" + will not be visible due to bugs/mis-features in glibc. + Further, note that irrespective of show_sys_files, all + files are accessible by name, i.e. you can always do + "ls -l \$UpCase" for example to specifically show the + system file containing the Unicode upcase table. + +case_sensitive=<BOOL> If case_sensitive is specified, treat all file names as + case sensitive and create file names in the POSIX + namespace. Otherwise the default behaviour is to treat + file names as case insensitive and to create file names + in the WIN32/LONG name space. Note, the Linux NTFS + driver will never create short file names and will + remove them on rename/delete of the corresponding long + file name. + Note that files remain accessible via their short file + name, if it exists. If case_sensitive, you will need + to provide the correct case of the short file name. + +disable_sparse=<BOOL> If disable_sparse is specified, creation of sparse + regions, i.e. holes, inside files is disabled for the + volume (for the duration of this mount only). By + default, creation of sparse regions is enabled, which + is consistent with the behaviour of traditional Unix + filesystems. + +errors=opt What to do when critical filesystem errors are found. + Following values can be used for "opt": + continue: DEFAULT, try to clean-up as much as + possible, e.g. marking a corrupt inode as + bad so it is no longer accessed, and then + continue. + recover: At present only supported is recovery of + the boot sector from the backup copy. + If read-only mount, the recovery is done + in memory only and not written to disk. + Note that the options are additive, i.e. specifying: + errors=continue,errors=recover + means the driver will attempt to recover and if that + fails it will clean-up as much as possible and + continue. + +mft_zone_multiplier= Set the MFT zone multiplier for the volume (this + setting is not persistent across mounts and can be + changed from mount to mount but cannot be changed on + remount). Values of 1 to 4 are allowed, 1 being the + default. The MFT zone multiplier determines how much + space is reserved for the MFT on the volume. If all + other space is used up, then the MFT zone will be + shrunk dynamically, so this has no impact on the + amount of free space. However, it can have an impact + on performance by affecting fragmentation of the MFT. + In general use the default. If you have a lot of small + files then use a higher value. The values have the + following meaning: + Value MFT zone size (% of volume size) + 1 12.5% + 2 25% + 3 37.5% + 4 50% + Note this option is irrelevant for read-only mounts. + + +Known bugs and (mis-)features +============================= + +- The link count on each directory inode entry is set to 1, due to Linux not + supporting directory hard links. This may well confuse some user space + applications, since the directory names will have the same inode numbers. + This also speeds up ntfs_read_inode() immensely. And we haven't found any + problems with this approach so far. If you find a problem with this, please + let us know. + + +Please send bug reports/comments/feedback/abuse to the Linux-NTFS development +list at sourceforge: linux-ntfs-dev@lists.sourceforge.net + + +Using NTFS volume and stripe sets +================================= + +For support of volume and stripe sets, you can either use the kernel's +Device-Mapper driver or the kernel's Software RAID / MD driver. The former is +the recommended one to use for linear raid. But the latter is required for +raid level 5. For striping and mirroring, either driver should work fine. + + +The Device-Mapper driver +------------------------ + +You will need to create a table of the components of the volume/stripe set and +how they fit together and load this into the kernel using the dmsetup utility +(see man 8 dmsetup). + +Linear volume sets, i.e. linear raid, has been tested and works fine. Even +though untested, there is no reason why stripe sets, i.e. raid level 0, and +mirrors, i.e. raid level 1 should not work, too. Stripes with parity, i.e. +raid level 5, unfortunately cannot work yet because the current version of the +Device-Mapper driver does not support raid level 5. You may be able to use the +Software RAID / MD driver for raid level 5, see the next section for details. + +To create the table describing your volume you will need to know each of its +components and their sizes in sectors, i.e. multiples of 512-byte blocks. + +For NT4 fault tolerant volumes you can obtain the sizes using fdisk. So for +example if one of your partitions is /dev/hda2 you would do: + +$ fdisk -ul /dev/hda + +Disk /dev/hda: 81.9 GB, 81964302336 bytes +255 heads, 63 sectors/track, 9964 cylinders, total 160086528 sectors +Units = sectors of 1 * 512 = 512 bytes + + Device Boot Start End Blocks Id System + /dev/hda1 * 63 4209029 2104483+ 83 Linux + /dev/hda2 4209030 37768814 16779892+ 86 NTFS + /dev/hda3 37768815 46170809 4200997+ 83 Linux + +And you would know that /dev/hda2 has a size of 37768814 - 4209030 + 1 = +33559785 sectors. + +For Win2k and later dynamic disks, you can for example use the ldminfo utility +which is part of the Linux LDM tools (the latest version at the time of +writing is linux-ldm-0.0.8.tar.bz2). You can download it from: + http://www.linux-ntfs.org/ +Simply extract the downloaded archive (tar xvjf linux-ldm-0.0.8.tar.bz2), go +into it (cd linux-ldm-0.0.8) and change to the test directory (cd test). You +will find the precompiled (i386) ldminfo utility there. NOTE: You will not be +able to compile this yourself easily so use the binary version! + +Then you would use ldminfo in dump mode to obtain the necessary information: + +$ ./ldminfo --dump /dev/hda + +This would dump the LDM database found on /dev/hda which describes all of your +dynamic disks and all the volumes on them. At the bottom you will see the +VOLUME DEFINITIONS section which is all you really need. You may need to look +further above to determine which of the disks in the volume definitions is +which device in Linux. Hint: Run ldminfo on each of your dynamic disks and +look at the Disk Id close to the top of the output for each (the PRIVATE HEADER +section). You can then find these Disk Ids in the VBLK DATABASE section in the +<Disk> components where you will get the LDM Name for the disk that is found in +the VOLUME DEFINITIONS section. + +Note you will also need to enable the LDM driver in the Linux kernel. If your +distribution did not enable it, you will need to recompile the kernel with it +enabled. This will create the LDM partitions on each device at boot time. You +would then use those devices (for /dev/hda they would be /dev/hda1, 2, 3, etc) +in the Device-Mapper table. + +You can also bypass using the LDM driver by using the main device (e.g. +/dev/hda) and then using the offsets of the LDM partitions into this device as +the "Start sector of device" when creating the table. Once again ldminfo would +give you the correct information to do this. + +Assuming you know all your devices and their sizes things are easy. + +For a linear raid the table would look like this (note all values are in +512-byte sectors): + +--- cut here --- +# Offset into Size of this Raid type Device Start sector +# volume device of device +0 1028161 linear /dev/hda1 0 +1028161 3903762 linear /dev/hdb2 0 +4931923 2103211 linear /dev/hdc1 0 +--- cut here --- + +For a striped volume, i.e. raid level 0, you will need to know the chunk size +you used when creating the volume. Windows uses 64kiB as the default, so it +will probably be this unless you changes the defaults when creating the array. + +For a raid level 0 the table would look like this (note all values are in +512-byte sectors): + +--- cut here --- +# Offset Size Raid Number Chunk 1st Start 2nd Start +# into of the type of size Device in Device in +# volume volume stripes device device +0 2056320 striped 2 128 /dev/hda1 0 /dev/hdb1 0 +--- cut here --- + +If there are more than two devices, just add each of them to the end of the +line. + +Finally, for a mirrored volume, i.e. raid level 1, the table would look like +this (note all values are in 512-byte sectors): + +--- cut here --- +# Ofs Size Raid Log Number Region Should Number Source Start Target Start +# in of the type type of log size sync? of Device in Device in +# vol volume params mirrors Device Device +0 2056320 mirror core 2 16 nosync 2 /dev/hda1 0 /dev/hdb1 0 +--- cut here --- + +If you are mirroring to multiple devices you can specify further targets at the +end of the line. + +Note the "Should sync?" parameter "nosync" means that the two mirrors are +already in sync which will be the case on a clean shutdown of Windows. If the +mirrors are not clean, you can specify the "sync" option instead of "nosync" +and the Device-Mapper driver will then copy the entirety of the "Source Device" +to the "Target Device" or if you specified multiple target devices to all of +them. + +Once you have your table, save it in a file somewhere (e.g. /etc/ntfsvolume1), +and hand it over to dmsetup to work with, like so: + +$ dmsetup create myvolume1 /etc/ntfsvolume1 + +You can obviously replace "myvolume1" with whatever name you like. + +If it all worked, you will now have the device /dev/device-mapper/myvolume1 +which you can then just use as an argument to the mount command as usual to +mount the ntfs volume. For example: + +$ mount -t ntfs -o ro /dev/device-mapper/myvolume1 /mnt/myvol1 + +(You need to create the directory /mnt/myvol1 first and of course you can use +anything you like instead of /mnt/myvol1 as long as it is an existing +directory.) + +It is advisable to do the mount read-only to see if the volume has been setup +correctly to avoid the possibility of causing damage to the data on the ntfs +volume. + + +The Software RAID / MD driver +----------------------------- + +An alternative to using the Device-Mapper driver is to use the kernel's +Software RAID / MD driver. For which you need to set up your /etc/raidtab +appropriately (see man 5 raidtab). + +Linear volume sets, i.e. linear raid, as well as stripe sets, i.e. raid level +0, have been tested and work fine (though see section "Limitations when using +the MD driver with NTFS volumes" especially if you want to use linear raid). +Even though untested, there is no reason why mirrors, i.e. raid level 1, and +stripes with parity, i.e. raid level 5, should not work, too. + +You have to use the "persistent-superblock 0" option for each raid-disk in the +NTFS volume/stripe you are configuring in /etc/raidtab as the persistent +superblock used by the MD driver would damage the NTFS volume. + +Windows by default uses a stripe chunk size of 64k, so you probably want the +"chunk-size 64k" option for each raid-disk, too. + +For example, if you have a stripe set consisting of two partitions /dev/hda5 +and /dev/hdb1 your /etc/raidtab would look like this: + +raiddev /dev/md0 + raid-level 0 + nr-raid-disks 2 + nr-spare-disks 0 + persistent-superblock 0 + chunk-size 64k + device /dev/hda5 + raid-disk 0 + device /dev/hdb1 + raid-disk 1 + +For linear raid, just change the raid-level above to "raid-level linear", for +mirrors, change it to "raid-level 1", and for stripe sets with parity, change +it to "raid-level 5". + +Note for stripe sets with parity you will also need to tell the MD driver +which parity algorithm to use by specifying the option "parity-algorithm +which", where you need to replace "which" with the name of the algorithm to +use (see man 5 raidtab for available algorithms) and you will have to try the +different available algorithms until you find one that works. Make sure you +are working read-only when playing with this as you may damage your data +otherwise. If you find which algorithm works please let us know (email the +linux-ntfs developers list linux-ntfs-dev@lists.sourceforge.net or drop in on +IRC in channel #ntfs on the irc.freenode.net network) so we can update this +documentation. + +Once the raidtab is setup, run for example raid0run -a to start all devices or +raid0run /dev/md0 to start a particular md device, in this case /dev/md0. + +Then just use the mount command as usual to mount the ntfs volume using for +example: mount -t ntfs -o ro /dev/md0 /mnt/myntfsvolume + +It is advisable to do the mount read-only to see if the md volume has been +setup correctly to avoid the possibility of causing damage to the data on the +ntfs volume. + + +Limitations when using the Software RAID / MD driver +----------------------------------------------------- + +Using the md driver will not work properly if any of your NTFS partitions have +an odd number of sectors. This is especially important for linear raid as all +data after the first partition with an odd number of sectors will be offset by +one or more sectors so if you mount such a partition with write support you +will cause massive damage to the data on the volume which will only become +apparent when you try to use the volume again under Windows. + +So when using linear raid, make sure that all your partitions have an even +number of sectors BEFORE attempting to use it. You have been warned! + +Even better is to simply use the Device-Mapper for linear raid and then you do +not have this problem with odd numbers of sectors. + + +ChangeLog +========= + +Note, a technical ChangeLog aimed at kernel hackers is in fs/ntfs/ChangeLog. + +2.1.30: + - Fix writev() (it kept writing the first segment over and over again + instead of moving onto subsequent segments). + - Fix crash in ntfs_mft_record_alloc() when mapping the new extent mft + record failed. +2.1.29: + - Fix a deadlock when mounting read-write. +2.1.28: + - Fix a deadlock. +2.1.27: + - Implement page migration support so the kernel can move memory used + by NTFS files and directories around for management purposes. + - Add support for writing to sparse files created with Windows XP SP2. + - Many minor improvements and bug fixes. +2.1.26: + - Implement support for sector sizes above 512 bytes (up to the maximum + supported by NTFS which is 4096 bytes). + - Enhance support for NTFS volumes which were supported by Windows but + not by Linux due to invalid attribute list attribute flags. + - A few minor updates and bug fixes. +2.1.25: + - Write support is now extended with write(2) being able to both + overwrite existing file data and to extend files. Also, if a write + to a sparse region occurs, write(2) will fill in the hole. Note, + mmap(2) based writes still do not support writing into holes or + writing beyond the initialized size. + - Write support has a new feature and that is that truncate(2) and + open(2) with O_TRUNC are now implemented thus files can be both made + smaller and larger. + - Note: Both write(2) and truncate(2)/open(2) with O_TRUNC still have + limitations in that they + - only provide limited support for highly fragmented files. + - only work on regular, i.e. uncompressed and unencrypted files. + - never create sparse files although this will change once directory + operations are implemented. + - Lots of bug fixes and enhancements across the board. +2.1.24: + - Support journals ($LogFile) which have been modified by chkdsk. This + means users can boot into Windows after we marked the volume dirty. + The Windows boot will run chkdsk and then reboot. The user can then + immediately boot into Linux rather than having to do a full Windows + boot first before rebooting into Linux and we will recognize such a + journal and empty it as it is clean by definition. + - Support journals ($LogFile) with only one restart page as well as + journals with two different restart pages. We sanity check both and + either use the only sane one or the more recent one of the two in the + case that both are valid. + - Lots of bug fixes and enhancements across the board. +2.1.23: + - Stamp the user space journal, aka transaction log, aka $UsnJrnl, if + it is present and active thus telling Windows and applications using + the transaction log that changes can have happened on the volume + which are not recorded in $UsnJrnl. + - Detect the case when Windows has been hibernated (suspended to disk) + and if this is the case do not allow (re)mounting read-write to + prevent data corruption when you boot back into the suspended + Windows session. + - Implement extension of resident files using the normal file write + code paths, i.e. most very small files can be extended to be a little + bit bigger but not by much. + - Add new mount option "disable_sparse". (See list of mount options + above for details.) + - Improve handling of ntfs volumes with errors and strange boot sectors + in particular. + - Fix various bugs including a nasty deadlock that appeared in recent + kernels (around 2.6.11-2.6.12 timeframe). +2.1.22: + - Improve handling of ntfs volumes with errors. + - Fix various bugs and race conditions. +2.1.21: + - Fix several race conditions and various other bugs. + - Many internal cleanups, code reorganization, optimizations, and mft + and index record writing code rewritten to fit in with the changes. + - Update Documentation/filesystems/ntfs.txt with instructions on how to + use the Device-Mapper driver with NTFS ftdisk/LDM raid. +2.1.20: + - Fix two stupid bugs introduced in 2.1.18 release. +2.1.19: + - Minor bugfix in handling of the default upcase table. + - Many internal cleanups and improvements. Many thanks to Linus + Torvalds and Al Viro for the help and advice with the sparse + annotations and cleanups. +2.1.18: + - Fix scheduling latencies at mount time. (Ingo Molnar) + - Fix endianness bug in a little traversed portion of the attribute + lookup code. +2.1.17: + - Fix bugs in mount time error code paths. +2.1.16: + - Implement access time updates (including mtime and ctime). + - Implement fsync(2), fdatasync(2), and msync(2) system calls. + - Enable the readv(2) and writev(2) system calls. + - Enable access via the asynchronous io (aio) API by adding support for + the aio_read(3) and aio_write(3) functions. +2.1.15: + - Invalidate quotas when (re)mounting read-write. + NOTE: This now only leave user space journalling on the side. (See + note for version 2.1.13, below.) +2.1.14: + - Fix an NFSd caused deadlock reported by several users. +2.1.13: + - Implement writing of inodes (access time updates are not implemented + yet so mounting with -o noatime,nodiratime is enforced). + - Enable writing out of resident files so you can now overwrite any + uncompressed, unencrypted, nonsparse file as long as you do not + change the file size. + - Add housekeeping of ntfs system files so that ntfsfix no longer needs + to be run after writing to an NTFS volume. + NOTE: This still leaves quota tracking and user space journalling on + the side but they should not cause data corruption. In the worst + case the charged quotas will be out of date ($Quota) and some + userspace applications might get confused due to the out of date + userspace journal ($UsnJrnl). +2.1.12: + - Fix the second fix to the decompression engine from the 2.1.9 release + and some further internals cleanups. +2.1.11: + - Driver internal cleanups. +2.1.10: + - Force read-only (re)mounting of volumes with unsupported volume + flags and various cleanups. +2.1.9: + - Fix two bugs in handling of corner cases in the decompression engine. +2.1.8: + - Read the $MFT mirror and compare it to the $MFT and if the two do not + match, force a read-only mount and do not allow read-write remounts. + - Read and parse the $LogFile journal and if it indicates that the + volume was not shutdown cleanly, force a read-only mount and do not + allow read-write remounts. If the $LogFile indicates a clean + shutdown and a read-write (re)mount is requested, empty $LogFile to + ensure that Windows cannot cause data corruption by replaying a stale + journal after Linux has written to the volume. + - Improve time handling so that the NTFS time is fully preserved when + converted to kernel time and only up to 99 nano-seconds are lost when + kernel time is converted to NTFS time. +2.1.7: + - Enable NFS exporting of mounted NTFS volumes. +2.1.6: + - Fix minor bug in handling of compressed directories that fixes the + erroneous "du" and "stat" output people reported. +2.1.5: + - Minor bug fix in attribute list attribute handling that fixes the + I/O errors on "ls" of certain fragmented files found by at least two + people running Windows XP. +2.1.4: + - Minor update allowing compilation with all gcc versions (well, the + ones the kernel can be compiled with anyway). +2.1.3: + - Major bug fixes for reading files and volumes in corner cases which + were being hit by Windows 2k/XP users. +2.1.2: + - Major bug fixes alleviating the hangs in statfs experienced by some + users. +2.1.1: + - Update handling of compressed files so people no longer get the + frequently reported warning messages about initialized_size != + data_size. +2.1.0: + - Add configuration option for developmental write support. + - Initial implementation of file overwriting. (Writes to resident files + are not written out to disk yet, so avoid writing to files smaller + than about 1kiB.) + - Intercept/abort changes in file size as they are not implemented yet. +2.0.25: + - Minor bugfixes in error code paths and small cleanups. +2.0.24: + - Small internal cleanups. + - Support for sendfile system call. (Christoph Hellwig) +2.0.23: + - Massive internal locking changes to mft record locking. Fixes + various race conditions and deadlocks. + - Fix ntfs over loopback for compressed files by adding an + optimization barrier. (gcc was screwing up otherwise ?) + Thanks go to Christoph Hellwig for pointing these two out: + - Remove now unused function fs/ntfs/malloc.h::vmalloc_nofs(). + - Fix ntfs_free() for ia64 and parisc. +2.0.22: + - Small internal cleanups. +2.0.21: + These only affect 32-bit architectures: + - Check for, and refuse to mount too large volumes (maximum is 2TiB). + - Check for, and refuse to open too large files and directories + (maximum is 16TiB). +2.0.20: + - Support non-resident directory index bitmaps. This means we now cope + with huge directories without problems. + - Fix a page leak that manifested itself in some cases when reading + directory contents. + - Internal cleanups. +2.0.19: + - Fix race condition and improvements in block i/o interface. + - Optimization when reading compressed files. +2.0.18: + - Fix race condition in reading of compressed files. +2.0.17: + - Cleanups and optimizations. +2.0.16: + - Fix stupid bug introduced in 2.0.15 in new attribute inode API. + - Big internal cleanup replacing the mftbmp access hacks by using the + new attribute inode API instead. +2.0.15: + - Bug fix in parsing of remount options. + - Internal changes implementing attribute (fake) inodes allowing all + attribute i/o to go via the page cache and to use all the normal + vfs/mm functionality. +2.0.14: + - Internal changes improving run list merging code and minor locking + change to not rely on BKL in ntfs_statfs(). +2.0.13: + - Internal changes towards using iget5_locked() in preparation for + fake inodes and small cleanups to ntfs_volume structure. +2.0.12: + - Internal cleanups in address space operations made possible by the + changes introduced in the previous release. +2.0.11: + - Internal updates and cleanups introducing the first step towards + fake inode based attribute i/o. +2.0.10: + - Microsoft says that the maximum number of inodes is 2^32 - 1. Update + the driver accordingly to only use 32-bits to store inode numbers on + 32-bit architectures. This improves the speed of the driver a little. +2.0.9: + - Change decompression engine to use a single buffer. This should not + affect performance except perhaps on the most heavy i/o on SMP + systems when accessing multiple compressed files from multiple + devices simultaneously. + - Minor updates and cleanups. +2.0.8: + - Remove now obsolete show_inodes and posix mount option(s). + - Restore show_sys_files mount option. + - Add new mount option case_sensitive, to determine if the driver + treats file names as case sensitive or not. + - Mostly drop support for short file names (for backwards compatibility + we only support accessing files via their short file name if one + exists). + - Fix dcache aliasing issues wrt short/long file names. + - Cleanups and minor fixes. +2.0.7: + - Just cleanups. +2.0.6: + - Major bugfix to make compatible with other kernel changes. This fixes + the hangs/oopses on umount. + - Locking cleanup in directory operations (remove BKL usage). +2.0.5: + - Major buffer overflow bug fix. + - Minor cleanups and updates for kernel 2.5.12. +2.0.4: + - Cleanups and updates for kernel 2.5.11. +2.0.3: + - Small bug fixes, cleanups, and performance improvements. +2.0.2: + - Use default fmask of 0177 so that files are no executable by default. + If you want owner executable files, just use fmask=0077. + - Update for kernel 2.5.9 but preserve backwards compatibility with + kernel 2.5.7. + - Minor bug fixes, cleanups, and updates. +2.0.1: + - Minor updates, primarily set the executable bit by default on files + so they can be executed. +2.0.0: + - Started ChangeLog. + diff --git a/Documentation/filesystems/ocfs2.txt b/Documentation/filesystems/ocfs2.txt new file mode 100644 index 00000000..7618a287 --- /dev/null +++ b/Documentation/filesystems/ocfs2.txt @@ -0,0 +1,102 @@ +OCFS2 filesystem +================== +OCFS2 is a general purpose extent based shared disk cluster file +system with many similarities to ext3. It supports 64 bit inode +numbers, and has automatically extending metadata groups which may +also make it attractive for non-clustered use. + +You'll want to install the ocfs2-tools package in order to at least +get "mount.ocfs2" and "ocfs2_hb_ctl". + +Project web page: http://oss.oracle.com/projects/ocfs2 +Tools web page: http://oss.oracle.com/projects/ocfs2-tools +OCFS2 mailing lists: http://oss.oracle.com/projects/ocfs2/mailman/ + +All code copyright 2005 Oracle except when otherwise noted. + +CREDITS: +Lots of code taken from ext3 and other projects. + +Authors in alphabetical order: +Joel Becker <joel.becker@oracle.com> +Zach Brown <zach.brown@oracle.com> +Mark Fasheh <mfasheh@suse.com> +Kurt Hackel <kurt.hackel@oracle.com> +Tao Ma <tao.ma@oracle.com> +Sunil Mushran <sunil.mushran@oracle.com> +Manish Singh <manish.singh@oracle.com> +Tiger Yang <tiger.yang@oracle.com> + +Caveats +======= +Features which OCFS2 does not support yet: + - Directory change notification (F_NOTIFY) + - Distributed Caching (F_SETLEASE/F_GETLEASE/break_lease) + +Mount options +============= + +OCFS2 supports the following mount options: +(*) == default + +barrier=1 This enables/disables barriers. barrier=0 disables it, + barrier=1 enables it. +errors=remount-ro(*) Remount the filesystem read-only on an error. +errors=panic Panic and halt the machine if an error occurs. +intr (*) Allow signals to interrupt cluster operations. +nointr Do not allow signals to interrupt cluster + operations. +noatime Do not update access time. +relatime(*) Update atime if the previous atime is older than + mtime or ctime +strictatime Always update atime, but the minimum update interval + is specified by atime_quantum. +atime_quantum=60(*) OCFS2 will not update atime unless this number + of seconds has passed since the last update. + Set to zero to always update atime. This option need + work with strictatime. +data=ordered (*) All data are forced directly out to the main file + system prior to its metadata being committed to the + journal. +data=writeback Data ordering is not preserved, data may be written + into the main file system after its metadata has been + committed to the journal. +preferred_slot=0(*) During mount, try to use this filesystem slot first. If + it is in use by another node, the first empty one found + will be chosen. Invalid values will be ignored. +commit=nrsec (*) Ocfs2 can be told to sync all its data and metadata + every 'nrsec' seconds. The default value is 5 seconds. + This means that if you lose your power, you will lose + as much as the latest 5 seconds of work (your + filesystem will not be damaged though, thanks to the + journaling). This default value (or any low value) + will hurt performance, but it's good for data-safety. + Setting it to 0 will have the same effect as leaving + it at the default (5 seconds). + Setting it to very large values will improve + performance. +localalloc=8(*) Allows custom localalloc size in MB. If the value is too + large, the fs will silently revert it to the default. +localflocks This disables cluster aware flock. +inode64 Indicates that Ocfs2 is allowed to create inodes at + any location in the filesystem, including those which + will result in inode numbers occupying more than 32 + bits of significance. +user_xattr (*) Enables Extended User Attributes. +nouser_xattr Disables Extended User Attributes. +acl Enables POSIX Access Control Lists support. +noacl (*) Disables POSIX Access Control Lists support. +resv_level=2 (*) Set how aggressive allocation reservations will be. + Valid values are between 0 (reservations off) to 8 + (maximum space for reservations). +dir_resv_level= (*) By default, directory reservations will scale with file + reservations - users should rarely need to change this + value. If allocation reservations are turned off, this + option will have no effect. +coherency=full (*) Disallow concurrent O_DIRECT writes, cluster inode + lock will be taken to force other nodes drop cache, + therefore full cluster coherency is guaranteed even + for O_DIRECT writes. +coherency=buffered Allow concurrent O_DIRECT writes without EX lock among + nodes, which gains high performance at risk of getting + stale data on other nodes. diff --git a/Documentation/filesystems/omfs.txt b/Documentation/filesystems/omfs.txt new file mode 100644 index 00000000..1d0d41ff --- /dev/null +++ b/Documentation/filesystems/omfs.txt @@ -0,0 +1,106 @@ +Optimized MPEG Filesystem (OMFS) + +Overview +======== + +OMFS is a filesystem created by SonicBlue for use in the ReplayTV DVR +and Rio Karma MP3 player. The filesystem is extent-based, utilizing +block sizes from 2k to 8k, with hash-based directories. This +filesystem driver may be used to read and write disks from these +devices. + +Note, it is not recommended that this FS be used in place of a general +filesystem for your own streaming media device. Native Linux filesystems +will likely perform better. + +More information is available at: + + http://linux-karma.sf.net/ + +Various utilities, including mkomfs and omfsck, are included with +omfsprogs, available at: + + http://bobcopeland.com/karma/ + +Instructions are included in its README. + +Options +======= + +OMFS supports the following mount-time options: + + uid=n - make all files owned by specified user + gid=n - make all files owned by specified group + umask=xxx - set permission umask to xxx + fmask=xxx - set umask to xxx for files + dmask=xxx - set umask to xxx for directories + +Disk format +=========== + +OMFS discriminates between "sysblocks" and normal data blocks. The sysblock +group consists of super block information, file metadata, directory structures, +and extents. Each sysblock has a header containing CRCs of the entire +sysblock, and may be mirrored in successive blocks on the disk. A sysblock may +have a smaller size than a data block, but since they are both addressed by the +same 64-bit block number, any remaining space in the smaller sysblock is +unused. + +Sysblock header information: + +struct omfs_header { + __be64 h_self; /* FS block where this is located */ + __be32 h_body_size; /* size of useful data after header */ + __be16 h_crc; /* crc-ccitt of body_size bytes */ + char h_fill1[2]; + u8 h_version; /* version, always 1 */ + char h_type; /* OMFS_INODE_X */ + u8 h_magic; /* OMFS_IMAGIC */ + u8 h_check_xor; /* XOR of header bytes before this */ + __be32 h_fill2; +}; + +Files and directories are both represented by omfs_inode: + +struct omfs_inode { + struct omfs_header i_head; /* header */ + __be64 i_parent; /* parent containing this inode */ + __be64 i_sibling; /* next inode in hash bucket */ + __be64 i_ctime; /* ctime, in milliseconds */ + char i_fill1[35]; + char i_type; /* OMFS_[DIR,FILE] */ + __be32 i_fill2; + char i_fill3[64]; + char i_name[OMFS_NAMELEN]; /* filename */ + __be64 i_size; /* size of file, in bytes */ +}; + +Directories in OMFS are implemented as a large hash table. Filenames are +hashed then prepended into the bucket list beginning at OMFS_DIR_START. +Lookup requires hashing the filename, then seeking across i_sibling pointers +until a match is found on i_name. Empty buckets are represented by block +pointers with all-1s (~0). + +A file is an omfs_inode structure followed by an extent table beginning at +OMFS_EXTENT_START: + +struct omfs_extent_entry { + __be64 e_cluster; /* start location of a set of blocks */ + __be64 e_blocks; /* number of blocks after e_cluster */ +}; + +struct omfs_extent { + __be64 e_next; /* next extent table location */ + __be32 e_extent_count; /* total # extents in this table */ + __be32 e_fill; + struct omfs_extent_entry e_entry; /* start of extent entries */ +}; + +Each extent holds the block offset followed by number of blocks allocated to +the extent. The final extent in each table is a terminator with e_cluster +being ~0 and e_blocks being ones'-complement of the total number of blocks +in the table. + +If this table overflows, a continuation inode is written and pointed to by +e_next. These have a header but lack the rest of the inode structure. + diff --git a/Documentation/filesystems/path-lookup.txt b/Documentation/filesystems/path-lookup.txt new file mode 100644 index 00000000..3571667c --- /dev/null +++ b/Documentation/filesystems/path-lookup.txt @@ -0,0 +1,382 @@ +Path walking and name lookup locking +==================================== + +Path resolution is the finding a dentry corresponding to a path name string, by +performing a path walk. Typically, for every open(), stat() etc., the path name +will be resolved. Paths are resolved by walking the namespace tree, starting +with the first component of the pathname (eg. root or cwd) with a known dentry, +then finding the child of that dentry, which is named the next component in the +path string. Then repeating the lookup from the child dentry and finding its +child with the next element, and so on. + +Since it is a frequent operation for workloads like multiuser environments and +web servers, it is important to optimize this code. + +Path walking synchronisation history: +Prior to 2.5.10, dcache_lock was acquired in d_lookup (dcache hash lookup) and +thus in every component during path look-up. Since 2.5.10 onwards, fast-walk +algorithm changed this by holding the dcache_lock at the beginning and walking +as many cached path component dentries as possible. This significantly +decreases the number of acquisition of dcache_lock. However it also increases +the lock hold time significantly and affects performance in large SMP machines. +Since 2.5.62 kernel, dcache has been using a new locking model that uses RCU to +make dcache look-up lock-free. + +All the above algorithms required taking a lock and reference count on the +dentry that was looked up, so that may be used as the basis for walking the +next path element. This is inefficient and unscalable. It is inefficient +because of the locks and atomic operations required for every dentry element +slows things down. It is not scalable because many parallel applications that +are path-walk intensive tend to do path lookups starting from a common dentry +(usually, the root "/" or current working directory). So contention on these +common path elements causes lock and cacheline queueing. + +Since 2.6.38, RCU is used to make a significant part of the entire path walk +(including dcache look-up) completely "store-free" (so, no locks, atomics, or +even stores into cachelines of common dentries). This is known as "rcu-walk" +path walking. + +Path walking overview +===================== + +A name string specifies a start (root directory, cwd, fd-relative) and a +sequence of elements (directory entry names), which together refer to a path in +the namespace. A path is represented as a (dentry, vfsmount) tuple. The name +elements are sub-strings, separated by '/'. + +Name lookups will want to find a particular path that a name string refers to +(usually the final element, or parent of final element). This is done by taking +the path given by the name's starting point (which we know in advance -- eg. +current->fs->cwd or current->fs->root) as the first parent of the lookup. Then +iteratively for each subsequent name element, look up the child of the current +parent with the given name and if it is not the desired entry, make it the +parent for the next lookup. + +A parent, of course, must be a directory, and we must have appropriate +permissions on the parent inode to be able to walk into it. + +Turning the child into a parent for the next lookup requires more checks and +procedures. Symlinks essentially substitute the symlink name for the target +name in the name string, and require some recursive path walking. Mount points +must be followed into (thus changing the vfsmount that subsequent path elements +refer to), switching from the mount point path to the root of the particular +mounted vfsmount. These behaviours are variously modified depending on the +exact path walking flags. + +Path walking then must, broadly, do several particular things: +- find the start point of the walk; +- perform permissions and validity checks on inodes; +- perform dcache hash name lookups on (parent, name element) tuples; +- traverse mount points; +- traverse symlinks; +- lookup and create missing parts of the path on demand. + +Safe store-free look-up of dcache hash table +============================================ + +Dcache name lookup +------------------ +In order to lookup a dcache (parent, name) tuple, we take a hash on the tuple +and use that to select a bucket in the dcache-hash table. The list of entries +in that bucket is then walked, and we do a full comparison of each entry +against our (parent, name) tuple. + +The hash lists are RCU protected, so list walking is not serialised with +concurrent updates (insertion, deletion from the hash). This is a standard RCU +list application with the exception of renames, which will be covered below. + +Parent and name members of a dentry, as well as its membership in the dcache +hash, and its inode are protected by the per-dentry d_lock spinlock. A +reference is taken on the dentry (while the fields are verified under d_lock), +and this stabilises its d_inode pointer and actual inode. This gives a stable +point to perform the next step of our path walk against. + +These members are also protected by d_seq seqlock, although this offers +read-only protection and no durability of results, so care must be taken when +using d_seq for synchronisation (see seqcount based lookups, below). + +Renames +------- +Back to the rename case. In usual RCU protected lists, the only operations that +will happen to an object is insertion, and then eventually removal from the +list. The object will not be reused until an RCU grace period is complete. +This ensures the RCU list traversal primitives can run over the object without +problems (see RCU documentation for how this works). + +However when a dentry is renamed, its hash value can change, requiring it to be +moved to a new hash list. Allocating and inserting a new alias would be +expensive and also problematic for directory dentries. Latency would be far to +high to wait for a grace period after removing the dentry and before inserting +it in the new hash bucket. So what is done is to insert the dentry into the +new list immediately. + +However, when the dentry's list pointers are updated to point to objects in the +new list before waiting for a grace period, this can result in a concurrent RCU +lookup of the old list veering off into the new (incorrect) list and missing +the remaining dentries on the list. + +There is no fundamental problem with walking down the wrong list, because the +dentry comparisons will never match. However it is fatal to miss a matching +dentry. So a seqlock is used to detect when a rename has occurred, and so the +lookup can be retried. + + 1 2 3 + +---+ +---+ +---+ +hlist-->| N-+->| N-+->| N-+-> +head <--+-P |<-+-P |<-+-P | + +---+ +---+ +---+ + +Rename of dentry 2 may require it deleted from the above list, and inserted +into a new list. Deleting 2 gives the following list. + + 1 3 + +---+ +---+ (don't worry, the longer pointers do not +hlist-->| N-+-------->| N-+-> impose a measurable performance overhead +head <--+-P |<--------+-P | on modern CPUs) + +---+ +---+ + ^ 2 ^ + | +---+ | + | | N-+----+ + +----+-P | + +---+ + +This is a standard RCU-list deletion, which leaves the deleted object's +pointers intact, so a concurrent list walker that is currently looking at +object 2 will correctly continue to object 3 when it is time to traverse the +next object. + +However, when inserting object 2 onto a new list, we end up with this: + + 1 3 + +---+ +---+ +hlist-->| N-+-------->| N-+-> +head <--+-P |<--------+-P | + +---+ +---+ + 2 + +---+ + | N-+----> + <----+-P | + +---+ + +Because we didn't wait for a grace period, there may be a concurrent lookup +still at 2. Now when it follows 2's 'next' pointer, it will walk off into +another list without ever having checked object 3. + +A related, but distinctly different, issue is that of rename atomicity versus +lookup operations. If a file is renamed from 'A' to 'B', a lookup must only +find either 'A' or 'B'. So if a lookup of 'A' returns NULL, a subsequent lookup +of 'B' must succeed (note the reverse is not true). + +Between deleting the dentry from the old hash list, and inserting it on the new +hash list, a lookup may find neither 'A' nor 'B' matching the dentry. The same +rename seqlock is also used to cover this race in much the same way, by +retrying a negative lookup result if a rename was in progress. + +Seqcount based lookups +---------------------- +In refcount based dcache lookups, d_lock is used to serialise access to +the dentry, stabilising it while comparing its name and parent and then +taking a reference count (the reference count then gives a stable place to +start the next part of the path walk from). + +As explained above, we would like to do path walking without taking locks or +reference counts on intermediate dentries along the path. To do this, a per +dentry seqlock (d_seq) is used to take a "coherent snapshot" of what the dentry +looks like (its name, parent, and inode). That snapshot is then used to start +the next part of the path walk. When loading the coherent snapshot under d_seq, +care must be taken to load the members up-front, and use those pointers rather +than reloading from the dentry later on (otherwise we'd have interesting things +like d_inode going NULL underneath us, if the name was unlinked). + +Also important is to avoid performing any destructive operations (pretty much: +no non-atomic stores to shared data), and to recheck the seqcount when we are +"done" with the operation. Retry or abort if the seqcount does not match. +Avoiding destructive or changing operations means we can easily unwind from +failure. + +What this means is that a caller, provided they are holding RCU lock to +protect the dentry object from disappearing, can perform a seqcount based +lookup which does not increment the refcount on the dentry or write to +it in any way. This returned dentry can be used for subsequent operations, +provided that d_seq is rechecked after that operation is complete. + +Inodes are also rcu freed, so the seqcount lookup dentry's inode may also be +queried for permissions. + +With this two parts of the puzzle, we can do path lookups without taking +locks or refcounts on dentry elements. + +RCU-walk path walking design +============================ + +Path walking code now has two distinct modes, ref-walk and rcu-walk. ref-walk +is the traditional[*] way of performing dcache lookups using d_lock to +serialise concurrent modifications to the dentry and take a reference count on +it. ref-walk is simple and obvious, and may sleep, take locks, etc while path +walking is operating on each dentry. rcu-walk uses seqcount based dentry +lookups, and can perform lookup of intermediate elements without any stores to +shared data in the dentry or inode. rcu-walk can not be applied to all cases, +eg. if the filesystem must sleep or perform non trivial operations, rcu-walk +must be switched to ref-walk mode. + +[*] RCU is still used for the dentry hash lookup in ref-walk, but not the full + path walk. + +Where ref-walk uses a stable, refcounted ``parent'' to walk the remaining +path string, rcu-walk uses a d_seq protected snapshot. When looking up a +child of this parent snapshot, we open d_seq critical section on the child +before closing d_seq critical section on the parent. This gives an interlocking +ladder of snapshots to walk down. + + + proc 101 + /----------------\ + / comm: "vi" \ + / fs.root: dentry0 \ + \ fs.cwd: dentry2 / + \ / + \----------------/ + +So when vi wants to open("/home/npiggin/test.c", O_RDWR), then it will +start from current->fs->root, which is a pinned dentry. Alternatively, +"./test.c" would start from cwd; both names refer to the same path in +the context of proc101. + + dentry 0 + +---------------------+ rcu-walk begins here, we note d_seq, check the + | name: "/" | inode's permission, and then look up the next + | inode: 10 | path element which is "home"... + | children:"home", ...| + +---------------------+ + | + dentry 1 V + +---------------------+ ... which brings us here. We find dentry1 via + | name: "home" | hash lookup, then note d_seq and compare name + | inode: 678 | string and parent pointer. When we have a match, + | children:"npiggin" | we now recheck the d_seq of dentry0. Then we + +---------------------+ check inode and look up the next element. + | + dentry2 V + +---------------------+ Note: if dentry0 is now modified, lookup is + | name: "npiggin" | not necessarily invalid, so we need only keep a + | inode: 543 | parent for d_seq verification, and grandparents + | children:"a.c", ... | can be forgotten. + +---------------------+ + | + dentry3 V + +---------------------+ At this point we have our destination dentry. + | name: "a.c" | We now take its d_lock, verify d_seq of this + | inode: 14221 | dentry. If that checks out, we can increment + | children:NULL | its refcount because we're holding d_lock. + +---------------------+ + +Taking a refcount on a dentry from rcu-walk mode, by taking its d_lock, +re-checking its d_seq, and then incrementing its refcount is called +"dropping rcu" or dropping from rcu-walk into ref-walk mode. + +It is, in some sense, a bit of a house of cards. If the seqcount check of the +parent snapshot fails, the house comes down, because we had closed the d_seq +section on the grandparent, so we have nothing left to stand on. In that case, +the path walk must be fully restarted (which we do in ref-walk mode, to avoid +live locks). It is costly to have a full restart, but fortunately they are +quite rare. + +When we reach a point where sleeping is required, or a filesystem callout +requires ref-walk, then instead of restarting the walk, we attempt to drop rcu +at the last known good dentry we have. Avoiding a full restart in ref-walk in +these cases is fundamental for performance and scalability because blocking +operations such as creates and unlinks are not uncommon. + +The detailed design for rcu-walk is like this: +* LOOKUP_RCU is set in nd->flags, which distinguishes rcu-walk from ref-walk. +* Take the RCU lock for the entire path walk, starting with the acquiring + of the starting path (eg. root/cwd/fd-path). So now dentry refcounts are + not required for dentry persistence. +* synchronize_rcu is called when unregistering a filesystem, so we can + access d_ops and i_ops during rcu-walk. +* Similarly take the vfsmount lock for the entire path walk. So now mnt + refcounts are not required for persistence. Also we are free to perform mount + lookups, and to assume dentry mount points and mount roots are stable up and + down the path. +* Have a per-dentry seqlock to protect the dentry name, parent, and inode, + so we can load this tuple atomically, and also check whether any of its + members have changed. +* Dentry lookups (based on parent, candidate string tuple) recheck the parent + sequence after the child is found in case anything changed in the parent + during the path walk. +* inode is also RCU protected so we can load d_inode and use the inode for + limited things. +* i_mode, i_uid, i_gid can be tested for exec permissions during path walk. +* i_op can be loaded. +* When the destination dentry is reached, drop rcu there (ie. take d_lock, + verify d_seq, increment refcount). +* If seqlock verification fails anywhere along the path, do a full restart + of the path lookup in ref-walk mode. -ECHILD tends to be used (for want of + a better errno) to signal an rcu-walk failure. + +The cases where rcu-walk cannot continue are: +* NULL dentry (ie. any uncached path element) +* Following links + +It may be possible eventually to make following links rcu-walk aware. + +Uncached path elements will always require dropping to ref-walk mode, at the +very least because i_mutex needs to be grabbed, and objects allocated. + +Final note: +"store-free" path walking is not strictly store free. We take vfsmount lock +and refcounts (both of which can be made per-cpu), and we also store to the +stack (which is essentially CPU-local), and we also have to take locks and +refcount on final dentry. + +The point is that shared data, where practically possible, is not locked +or stored into. The result is massive improvements in performance and +scalability of path resolution. + + +Interesting statistics +====================== + +The following table gives rcu lookup statistics for a few simple workloads +(2s12c24t Westmere, debian non-graphical system). Ungraceful are attempts to +drop rcu that fail due to d_seq failure and requiring the entire path lookup +again. Other cases are successful rcu-drops that are required before the final +element, nodentry for missing dentry, revalidate for filesystem revalidate +routine requiring rcu drop, permission for permission check requiring drop, +and link for symlink traversal requiring drop. + + rcu-lookups restart nodentry link revalidate permission +bootup 47121 0 4624 1010 10283 7852 +dbench 25386793 0 6778659(26.7%) 55 549 1156 +kbuild 2696672 10 64442(2.3%) 108764(4.0%) 1 1590 +git diff 39605 0 28 2 0 106 +vfstest 24185492 4945 708725(2.9%) 1076136(4.4%) 0 2651 + +What this shows is that failed rcu-walk lookups, ie. ones that are restarted +entirely with ref-walk, are quite rare. Even the "vfstest" case which +specifically has concurrent renames/mkdir/rmdir/ creat/unlink/etc to exercise +such races is not showing a huge amount of restarts. + +Dropping from rcu-walk to ref-walk mean that we have encountered a dentry where +the reference count needs to be taken for some reason. This is either because +we have reached the target of the path walk, or because we have encountered a +condition that can't be resolved in rcu-walk mode. Ideally, we drop rcu-walk +only when we have reached the target dentry, so the other statistics show where +this does not happen. + +Note that a graceful drop from rcu-walk mode due to something such as the +dentry not existing (which can be common) is not necessarily a failure of +rcu-walk scheme, because some elements of the path may have been walked in +rcu-walk mode. The further we get from common path elements (such as cwd or +root), the less contended the dentry is likely to be. The closer we are to +common path elements, the more likely they will exist in dentry cache. + + +Papers and other documentation on dcache locking +================================================ + +1. Scaling dcache with RCU (http://linuxjournal.com/article.php?sid=7124). + +2. http://lse.sourceforge.net/locking/dcache/dcache.html + + diff --git a/Documentation/filesystems/pohmelfs/design_notes.txt b/Documentation/filesystems/pohmelfs/design_notes.txt new file mode 100644 index 00000000..dcf83358 --- /dev/null +++ b/Documentation/filesystems/pohmelfs/design_notes.txt @@ -0,0 +1,71 @@ +POHMELFS: Parallel Optimized Host Message Exchange Layered File System. + + Evgeniy Polyakov <zbr@ioremap.net> + +Homepage: http://www.ioremap.net/projects/pohmelfs + +POHMELFS first began as a network filesystem with coherent local data and +metadata caches but is now evolving into a parallel distributed filesystem. + +Main features of this FS include: + * Locally coherent cache for data and metadata with (potentially) byte-range locks. + Since all Linux filesystems lock the whole inode during writing, algorithm + is very simple and does not use byte-ranges, although they are sent in + locking messages. + * Completely async processing of all events except creation of hard and symbolic + links, and rename events. + Object creation and data reading and writing are processed asynchronously. + * Flexible object architecture optimized for network processing. + Ability to create long paths to objects and remove arbitrarily huge + directories with a single network command. + (like removing the whole kernel tree via a single network command). + * Very high performance. + * Fast and scalable multithreaded userspace server. Being in userspace it works + with any underlying filesystem and still is much faster than async in-kernel NFS one. + * Client is able to switch between different servers (if one goes down, client + automatically reconnects to second and so on). + * Transactions support. Full failover for all operations. + Resending transactions to different servers on timeout or error. + * Read request (data read, directory listing, lookup requests) balancing between multiple servers. + * Write requests are replicated to multiple servers and completed only when all of them are acked. + * Ability to add and/or remove servers from the working set at run-time. + * Strong authentification and possible data encryption in network channel. + * Extended attributes support. + +POHMELFS is based on transactions, which are potentially long-standing objects that live +in the client's memory. Each transaction contains all the information needed to process a given +command (or set of commands, which is frequently used during data writing: single transactions +can contain creation and data writing commands). Transactions are committed by all the servers +to which they are sent and, in case of failures, are eventually resent or dropped with an error. +For example, reading will return an error if no servers are available. + +POHMELFS uses a asynchronous approach to data processing. Courtesy of transactions, it is +possible to detach replies from requests and, if the command requires data to be received, the +caller sleeps waiting for it. Thus, it is possible to issue multiple read commands to different +servers and async threads will pick up replies in parallel, find appropriate transactions in the +system and put the data where it belongs (like the page or inode cache). + +The main feature of POHMELFS is writeback data and the metadata cache. +Only a few non-performance critical operations use the write-through cache and +are synchronous: hard and symbolic link creation, and object rename. Creation, +removal of objects and data writing are asynchronous and are sent to +the server during system writeback. Only one writer at a time is allowed for any +given inode, which is guarded by an appropriate locking protocol. +Because of this feature, POHMELFS is extremely fast at metadata intensive +workloads and can fully utilize the bandwidth to the servers when doing bulk +data transfers. + +POHMELFS clients operate with a working set of servers and are capable of balancing read-only +operations (like lookups or directory listings) between them according to IO priorities. +Administrators can add or remove servers from the set at run-time via special commands (described +in Documentation/pohmelfs/info.txt file). Writes are replicated to all servers, which are connected +with write permission turned on. IO priority and permissions can be changed in run-time. + +POHMELFS is capable of full data channel encryption and/or strong crypto hashing. +One can select any kernel supported cipher, encryption mode, hash type and operation mode +(hmac or digest). It is also possible to use both or neither (default). Crypto configuration +is checked during mount time and, if the server does not support it, appropriate capabilities +will be disabled or mount will fail (if 'crypto_fail_unsupported' mount option is specified). +Crypto performance heavily depends on the number of crypto threads, which asynchronously perform +crypto operations and send the resulting data to server or submit it up the stack. This number +can be controlled via a mount option. diff --git a/Documentation/filesystems/pohmelfs/info.txt b/Documentation/filesystems/pohmelfs/info.txt new file mode 100644 index 00000000..db2e4139 --- /dev/null +++ b/Documentation/filesystems/pohmelfs/info.txt @@ -0,0 +1,99 @@ +POHMELFS usage information. + +Mount options. +All but index, number of crypto threads and maximum IO size can changed via remount. + +idx=%u + Each mountpoint is associated with a special index via this option. + Administrator can add or remove servers from the given index, so all mounts, + which were attached to it, are updated. + Default it is 0. + +trans_scan_timeout=%u + This timeout, expressed in milliseconds, specifies time to scan transaction + trees looking for stale requests, which have to be resent, or if number of + retries exceed specified limit, dropped with error. + Default is 5 seconds. + +drop_scan_timeout=%u + Internal timeout, expressed in milliseconds, which specifies how frequently + inodes marked to be dropped are freed. It also specifies how frequently + the system checks that servers have to be added or removed from current working set. + Default is 1 second. + +wait_on_page_timeout=%u + Number of milliseconds to wait for reply from remote server for data reading command. + If this timeout is exceeded, reading returns an error. + Default is 5 seconds. + +trans_retries=%u + This is the number of times that a transaction will be resent to a server that did + not answer for the last @trans_scan_timeout milliseconds. + When the number of resends exceeds this limit, the transaction is completed with error. + Default is 5 resends. + +crypto_thread_num=%u + Number of crypto processing threads. Threads are used both for RX and TX traffic. + Default is 2, or no threads if crypto operations are not supported. + +trans_max_pages=%u + Maximum number of pages in a single transaction. This parameter also controls + the number of pages, allocated for crypto processing (each crypto thread has + pool of pages, the number of which is equal to 'trans_max_pages'. + Default is 100 pages. + +crypto_fail_unsupported + If specified, mount will fail if the server does not support requested crypto operations. + By default mount will disable non-matching crypto operations. + +mcache_timeout=%u + Maximum number of milliseconds to wait for the mcache objects to be processed. + Mcache includes locks (given lock should be granted by server), attributes (they should be + fully received in the given timeframe). + Default is 5 seconds. + +Usage examples. + +Add server server1.net:1025 into the working set with index $idx +with appropriate hash algorithm and key file and cipher algorithm, mode and key file: +$cfg A add -a server1.net -p 1025 -i $idx -K $hash_key -k $cipher_key + +Mount filesystem with given index $idx to /mnt mountpoint. +Client will connect to all servers specified in the working set via previous command: +mount -t pohmel -o idx=$idx q /mnt + +Change permissions to read-only (-I 1 option, '-I 2' - write-only, 3 - rw): +$cfg A modify -a server1.net -p 1025 -i $idx -I 1 + +Change IO priority to 123 (node with the highest priority gets read requests). +$cfg A modify -a server1.net -p 1025 -i $idx -P 123 + +One can check currect status of all connections in the mountstats file: +# cat /proc/$PID/mountstats +... +device none mounted on /mnt with fstype pohmel +idx addr(:port) socket_type protocol active priority permissions +0 server1.net:1026 1 6 1 250 1 +0 server2.net:1025 1 6 1 123 3 + +Server installation. + +Creating a server, which listens at port 1025 and 0.0.0.0 address. +Working root directory (note, that server chroots there, so you have to have appropriate permissions) +is set to /mnt, server will negotiate hash/cipher with client, in case client requested it, there +are appropriate key files. +Number of working threads is set to 10. + +# ./fserver -a 0.0.0.0 -p 1025 -r /mnt -w 10 -K hash_key -k cipher_key + + -A 6 - listen on ipv6 address. Default: Disabled. + -r root - path to root directory. Default: /tmp. + -a addr - listen address. Default: 0.0.0.0. + -p port - listen port. Default: 1025. + -w workers - number of workers per connected client. Default: 1. + -K file - hash key size. Default: none. + -k file - cipher key size. Default: none. + -h - this help. + +Number of worker threads specifies how many workers will be created for each client. +Bulk single-client transafers usually are better handled with smaller number (like 1-3). diff --git a/Documentation/filesystems/pohmelfs/network_protocol.txt b/Documentation/filesystems/pohmelfs/network_protocol.txt new file mode 100644 index 00000000..65e03dd4 --- /dev/null +++ b/Documentation/filesystems/pohmelfs/network_protocol.txt @@ -0,0 +1,227 @@ +POHMELFS network protocol. + +Basic structure used in network communication is following command: + +struct netfs_cmd +{ + __u16 cmd; /* Command number */ + __u16 csize; /* Attached crypto information size */ + __u16 cpad; /* Attached padding size */ + __u16 ext; /* External flags */ + __u32 size; /* Size of the attached data */ + __u32 trans; /* Transaction id */ + __u64 id; /* Object ID to operate on. Used for feedback.*/ + __u64 start; /* Start of the object. */ + __u64 iv; /* IV sequence */ + __u8 data[0]; +}; + +Commands can be embedded into transaction command (which in turn has own command), +so one can extend protocol as needed without breaking backward compatibility as long +as old commands are supported. All string lengths include tail 0 byte. + +All commands are transferred over the network in big-endian. CPU endianess is used at the end peers. + +@cmd - command number, which specifies command to be processed. Following + commands are used currently: + + NETFS_READDIR = 1, /* Read directory for given inode number */ + NETFS_READ_PAGE, /* Read data page from the server */ + NETFS_WRITE_PAGE, /* Write data page to the server */ + NETFS_CREATE, /* Create directory entry */ + NETFS_REMOVE, /* Remove directory entry */ + NETFS_LOOKUP, /* Lookup single object */ + NETFS_LINK, /* Create a link */ + NETFS_TRANS, /* Transaction */ + NETFS_OPEN, /* Open intent */ + NETFS_INODE_INFO, /* Metadata cache coherency synchronization message */ + NETFS_PAGE_CACHE, /* Page cache invalidation message */ + NETFS_READ_PAGES, /* Read multiple contiguous pages in one go */ + NETFS_RENAME, /* Rename object */ + NETFS_CAPABILITIES, /* Capabilities of the client, for example supported crypto */ + NETFS_LOCK, /* Distributed lock message */ + NETFS_XATTR_SET, /* Set extended attribute */ + NETFS_XATTR_GET, /* Get extended attribute */ + +@ext - external flags. Used by different commands to specify some extra arguments + like partial size of the embedded objects or creation flags. + +@size - size of the attached data. For NETFS_READ_PAGE and NETFS_READ_PAGES no data is attached, + but size of the requested data is incorporated here. It does not include size of the command + header (struct netfs_cmd) itself. + +@id - id of the object this command operates on. Each command can use it for own purpose. + +@start - start of the object this command operates on. Each command can use it for own purpose. + +@csize, @cpad - size and padding size of the (attached if needed) crypto information. + +Command specifications. + +@NETFS_READDIR +This command is used to sync content of the remote dir to the client. + +@ext - length of the path to object. +@size - the same. +@id - local inode number of the directory to read. +@start - zero. + + +@NETFS_READ_PAGE +This command is used to read data from remote server. +Data size does not exceed local page cache size. + +@id - inode number. +@start - first byte offset. +@size - number of bytes to read plus length of the path to object. +@ext - object path length. + + +@NETFS_CREATE +Used to create object. +It does not require that all directories on top of the object were +already created, it will create them automatically. Each object has +associated @netfs_path_entry data structure, which contains creation +mode (permissions and type) and length of the name as long as name itself. + +@start - 0 +@size - size of the all data structures needed to create a path +@id - local inode number +@ext - 0 + + +@NETFS_REMOVE +Used to remove object. + +@ext - length of the path to object. +@size - the same. +@id - local inode number. +@start - zero. + + +@NETFS_LOOKUP +Lookup information about object on server. + +@ext - length of the path to object. +@size - the same. +@id - local inode number of the directory to look object in. +@start - local inode number of the object to look at. + + +@NETFS_LINK +Create hard of symlink. +Command is sent as "object_path|target_path". + +@size - size of the above string. +@id - parent local inode number. +@start - 1 for symlink, 0 for hardlink. +@ext - size of the "object_path" above. + + +@NETFS_TRANS +Transaction header. + +@size - incorporates all embedded command sizes including theirs header sizes. +@start - transaction generation number - unique id used to find transaction. +@ext - transaction flags. Unused at the moment. +@id - 0. + + +@NETFS_OPEN +Open intent for given transaction. + +@id - local inode number. +@start - 0. +@size - path length to the object. +@ext - open flags (O_RDWR and so on). + + +@NETFS_INODE_INFO +Metadata update command. +It is sent to servers when attributes of the object are changed and received +when data or metadata were updated. It operates with the following structure: + +struct netfs_inode_info +{ + unsigned int mode; + unsigned int nlink; + unsigned int uid; + unsigned int gid; + unsigned int blocksize; + unsigned int padding; + __u64 ino; + __u64 blocks; + __u64 rdev; + __u64 size; + __u64 version; +}; + +It effectively mirrors stat(2) returned data. + + +@ext - path length to the object. +@size - the same plus size of the netfs_inode_info structure. +@id - local inode number. +@start - 0. + + +@NETFS_PAGE_CACHE +Command is only received by clients. It contains information about +page to be marked as not up-to-date. + +@id - client's inode number. +@start - last byte of the page to be invalidated. If it is not equal to + current inode size, it will be vmtruncated(). +@size - 0 +@ext - 0 + + +@NETFS_READ_PAGES +Used to read multiple contiguous pages in one go. + +@start - first byte of the contiguous region to read. +@size - contains of two fields: lower 8 bits are used to represent page cache shift + used by client, another 3 bytes are used to get number of pages. +@id - local inode number. +@ext - path length to the object. + + +@NETFS_RENAME +Used to rename object. +Attached data is formed into following string: "old_path|new_path". + +@id - local inode number. +@start - parent inode number. +@size - length of the above string. +@ext - length of the old path part. + + +@NETFS_CAPABILITIES +Used to exchange crypto capabilities with server. +If crypto capabilities are not supported by server, then client will disable it +or fail (if 'crypto_fail_unsupported' mount options was specified). + +@id - superblock index. Used to specify crypto information for group of servers. +@size - size of the attached capabilities structure. +@start - 0. +@size - 0. +@scsize - 0. + +@NETFS_LOCK +Used to send lock request/release messages. Although it sends byte range request +and is capable of flushing pages based on that, it is not used, since all Linux +filesystems lock the whole inode. + +@id - lock generation number. +@start - start of the locked range. +@size - size of the locked range. +@ext - lock type: read/write. Not used actually. 15'th bit is used to determine, + if it is lock request (1) or release (0). + +@NETFS_XATTR_SET +@NETFS_XATTR_GET +Used to set/get extended attributes for given inode. +@id - attribute generation number or xattr setting type +@start - size of the attribute (request or attached) +@size - name length, path len and data size for given attribute +@ext - path length for given object diff --git a/Documentation/filesystems/porting b/Documentation/filesystems/porting new file mode 100644 index 00000000..6e299548 --- /dev/null +++ b/Documentation/filesystems/porting @@ -0,0 +1,409 @@ +Changes since 2.5.0: + +--- +[recommended] + +New helpers: sb_bread(), sb_getblk(), sb_find_get_block(), set_bh(), + sb_set_blocksize() and sb_min_blocksize(). + +Use them. + +(sb_find_get_block() replaces 2.4's get_hash_table()) + +--- +[recommended] + +New methods: ->alloc_inode() and ->destroy_inode(). + +Remove inode->u.foo_inode_i +Declare + struct foo_inode_info { + /* fs-private stuff */ + struct inode vfs_inode; + }; + static inline struct foo_inode_info *FOO_I(struct inode *inode) + { + return list_entry(inode, struct foo_inode_info, vfs_inode); + } + +Use FOO_I(inode) instead of &inode->u.foo_inode_i; + +Add foo_alloc_inode() and foo_destroy_inode() - the former should allocate +foo_inode_info and return the address of ->vfs_inode, the latter should free +FOO_I(inode) (see in-tree filesystems for examples). + +Make them ->alloc_inode and ->destroy_inode in your super_operations. + +Keep in mind that now you need explicit initialization of private data +typically between calling iget_locked() and unlocking the inode. + +At some point that will become mandatory. + +--- +[mandatory] + +Change of file_system_type method (->read_super to ->get_sb) + +->read_super() is no more. Ditto for DECLARE_FSTYPE and DECLARE_FSTYPE_DEV. + +Turn your foo_read_super() into a function that would return 0 in case of +success and negative number in case of error (-EINVAL unless you have more +informative error value to report). Call it foo_fill_super(). Now declare + +int foo_get_sb(struct file_system_type *fs_type, + int flags, const char *dev_name, void *data, struct vfsmount *mnt) +{ + return get_sb_bdev(fs_type, flags, dev_name, data, foo_fill_super, + mnt); +} + +(or similar with s/bdev/nodev/ or s/bdev/single/, depending on the kind of +filesystem). + +Replace DECLARE_FSTYPE... with explicit initializer and have ->get_sb set as +foo_get_sb. + +--- +[mandatory] + +Locking change: ->s_vfs_rename_sem is taken only by cross-directory renames. +Most likely there is no need to change anything, but if you relied on +global exclusion between renames for some internal purpose - you need to +change your internal locking. Otherwise exclusion warranties remain the +same (i.e. parents and victim are locked, etc.). + +--- +[informational] + +Now we have the exclusion between ->lookup() and directory removal (by +->rmdir() and ->rename()). If you used to need that exclusion and do +it by internal locking (most of filesystems couldn't care less) - you +can relax your locking. + +--- +[mandatory] + +->lookup(), ->truncate(), ->create(), ->unlink(), ->mknod(), ->mkdir(), +->rmdir(), ->link(), ->lseek(), ->symlink(), ->rename() +and ->readdir() are called without BKL now. Grab it on entry, drop upon return +- that will guarantee the same locking you used to have. If your method or its +parts do not need BKL - better yet, now you can shift lock_kernel() and +unlock_kernel() so that they would protect exactly what needs to be +protected. + +--- +[mandatory] + +BKL is also moved from around sb operations. ->write_super() Is now called +without BKL held. BKL should have been shifted into individual fs sb_op +functions. If you don't need it, remove it. + +--- +[informational] + +check for ->link() target not being a directory is done by callers. Feel +free to drop it... + +--- +[informational] + +->link() callers hold ->i_mutex on the object we are linking to. Some of your +problems might be over... + +--- +[mandatory] + +new file_system_type method - kill_sb(superblock). If you are converting +an existing filesystem, set it according to ->fs_flags: + FS_REQUIRES_DEV - kill_block_super + FS_LITTER - kill_litter_super + neither - kill_anon_super +FS_LITTER is gone - just remove it from fs_flags. + +--- +[mandatory] + + FS_SINGLE is gone (actually, that had happened back when ->get_sb() +went in - and hadn't been documented ;-/). Just remove it from fs_flags +(and see ->get_sb() entry for other actions). + +--- +[mandatory] + +->setattr() is called without BKL now. Caller _always_ holds ->i_mutex, so +watch for ->i_mutex-grabbing code that might be used by your ->setattr(). +Callers of notify_change() need ->i_mutex now. + +--- +[recommended] + +New super_block field "struct export_operations *s_export_op" for +explicit support for exporting, e.g. via NFS. The structure is fully +documented at its declaration in include/linux/fs.h, and in +Documentation/filesystems/nfs/Exporting. + +Briefly it allows for the definition of decode_fh and encode_fh operations +to encode and decode filehandles, and allows the filesystem to use +a standard helper function for decode_fh, and provide file-system specific +support for this helper, particularly get_parent. + +It is planned that this will be required for exporting once the code +settles down a bit. + +[mandatory] + +s_export_op is now required for exporting a filesystem. +isofs, ext2, ext3, resierfs, fat +can be used as examples of very different filesystems. + +--- +[mandatory] + +iget4() and the read_inode2 callback have been superseded by iget5_locked() +which has the following prototype, + + struct inode *iget5_locked(struct super_block *sb, unsigned long ino, + int (*test)(struct inode *, void *), + int (*set)(struct inode *, void *), + void *data); + +'test' is an additional function that can be used when the inode +number is not sufficient to identify the actual file object. 'set' +should be a non-blocking function that initializes those parts of a +newly created inode to allow the test function to succeed. 'data' is +passed as an opaque value to both test and set functions. + +When the inode has been created by iget5_locked(), it will be returned with the +I_NEW flag set and will still be locked. The filesystem then needs to finalize +the initialization. Once the inode is initialized it must be unlocked by +calling unlock_new_inode(). + +The filesystem is responsible for setting (and possibly testing) i_ino +when appropriate. There is also a simpler iget_locked function that +just takes the superblock and inode number as arguments and does the +test and set for you. + +e.g. + inode = iget_locked(sb, ino); + if (inode->i_state & I_NEW) { + err = read_inode_from_disk(inode); + if (err < 0) { + iget_failed(inode); + return err; + } + unlock_new_inode(inode); + } + +Note that if the process of setting up a new inode fails, then iget_failed() +should be called on the inode to render it dead, and an appropriate error +should be passed back to the caller. + +--- +[recommended] + +->getattr() finally getting used. See instances in nfs, minix, etc. + +--- +[mandatory] + +->revalidate() is gone. If your filesystem had it - provide ->getattr() +and let it call whatever you had as ->revlidate() + (for symlinks that +had ->revalidate()) add calls in ->follow_link()/->readlink(). + +--- +[mandatory] + +->d_parent changes are not protected by BKL anymore. Read access is safe +if at least one of the following is true: + * filesystem has no cross-directory rename() + * we know that parent had been locked (e.g. we are looking at +->d_parent of ->lookup() argument). + * we are called from ->rename(). + * the child's ->d_lock is held +Audit your code and add locking if needed. Notice that any place that is +not protected by the conditions above is risky even in the old tree - you +had been relying on BKL and that's prone to screwups. Old tree had quite +a few holes of that kind - unprotected access to ->d_parent leading to +anything from oops to silent memory corruption. + +--- +[mandatory] + + FS_NOMOUNT is gone. If you use it - just set MS_NOUSER in flags +(see rootfs for one kind of solution and bdev/socket/pipe for another). + +--- +[recommended] + + Use bdev_read_only(bdev) instead of is_read_only(kdev). The latter +is still alive, but only because of the mess in drivers/s390/block/dasd.c. +As soon as it gets fixed is_read_only() will die. + +--- +[mandatory] + +->permission() is called without BKL now. Grab it on entry, drop upon +return - that will guarantee the same locking you used to have. If +your method or its parts do not need BKL - better yet, now you can +shift lock_kernel() and unlock_kernel() so that they would protect +exactly what needs to be protected. + +--- +[mandatory] + +->statfs() is now called without BKL held. BKL should have been +shifted into individual fs sb_op functions where it's not clear that +it's safe to remove it. If you don't need it, remove it. + +--- +[mandatory] + + is_read_only() is gone; use bdev_read_only() instead. + +--- +[mandatory] + + destroy_buffers() is gone; use invalidate_bdev(). + +--- +[mandatory] + + fsync_dev() is gone; use fsync_bdev(). NOTE: lvm breakage is +deliberate; as soon as struct block_device * is propagated in a reasonable +way by that code fixing will become trivial; until then nothing can be +done. + +[mandatory] + + block truncatation on error exit from ->write_begin, and ->direct_IO +moved from generic methods (block_write_begin, cont_write_begin, +nobh_write_begin, blockdev_direct_IO*) to callers. Take a look at +ext2_write_failed and callers for an example. + +[mandatory] + + ->truncate is going away. The whole truncate sequence needs to be +implemented in ->setattr, which is now mandatory for filesystems +implementing on-disk size changes. Start with a copy of the old inode_setattr +and vmtruncate, and the reorder the vmtruncate + foofs_vmtruncate sequence to +be in order of zeroing blocks using block_truncate_page or similar helpers, +size update and on finally on-disk truncation which should not fail. +inode_change_ok now includes the size checks for ATTR_SIZE and must be called +in the beginning of ->setattr unconditionally. + +[mandatory] + + ->clear_inode() and ->delete_inode() are gone; ->evict_inode() should +be used instead. It gets called whenever the inode is evicted, whether it has +remaining links or not. Caller does *not* evict the pagecache or inode-associated +metadata buffers; getting rid of those is responsibility of method, as it had +been for ->delete_inode(). + + ->drop_inode() returns int now; it's called on final iput() with +inode->i_lock held and it returns true if filesystems wants the inode to be +dropped. As before, generic_drop_inode() is still the default and it's been +updated appropriately. generic_delete_inode() is also alive and it consists +simply of return 1. Note that all actual eviction work is done by caller after +->drop_inode() returns. + + clear_inode() is gone; use end_writeback() instead. As before, it must +be called exactly once on each call of ->evict_inode() (as it used to be for +each call of ->delete_inode()). Unlike before, if you are using inode-associated +metadata buffers (i.e. mark_buffer_dirty_inode()), it's your responsibility to +call invalidate_inode_buffers() before end_writeback(). + No async writeback (and thus no calls of ->write_inode()) will happen +after end_writeback() returns, so actions that should not overlap with ->write_inode() +(e.g. freeing on-disk inode if i_nlink is 0) ought to be done after that call. + + NOTE: checking i_nlink in the beginning of ->write_inode() and bailing out +if it's zero is not *and* *never* *had* *been* enough. Final unlink() and iput() +may happen while the inode is in the middle of ->write_inode(); e.g. if you blindly +free the on-disk inode, you may end up doing that while ->write_inode() is writing +to it. + +--- +[mandatory] + + .d_delete() now only advises the dcache as to whether or not to cache +unreferenced dentries, and is now only called when the dentry refcount goes to +0. Even on 0 refcount transition, it must be able to tolerate being called 0, +1, or more times (eg. constant, idempotent). + +--- +[mandatory] + + .d_compare() calling convention and locking rules are significantly +changed. Read updated documentation in Documentation/filesystems/vfs.txt (and +look at examples of other filesystems) for guidance. + +--- +[mandatory] + + .d_hash() calling convention and locking rules are significantly +changed. Read updated documentation in Documentation/filesystems/vfs.txt (and +look at examples of other filesystems) for guidance. + +--- +[mandatory] + dcache_lock is gone, replaced by fine grained locks. See fs/dcache.c +for details of what locks to replace dcache_lock with in order to protect +particular things. Most of the time, a filesystem only needs ->d_lock, which +protects *all* the dcache state of a given dentry. + +-- +[mandatory] + + Filesystems must RCU-free their inodes, if they can have been accessed +via rcu-walk path walk (basically, if the file can have had a path name in the +vfs namespace). + + i_dentry and i_rcu share storage in a union, and the vfs expects +i_dentry to be reinitialized before it is freed, so an: + + INIT_LIST_HEAD(&inode->i_dentry); + +must be done in the RCU callback. + +-- +[recommended] + vfs now tries to do path walking in "rcu-walk mode", which avoids +atomic operations and scalability hazards on dentries and inodes (see +Documentation/filesystems/path-lookup.txt). d_hash and d_compare changes +(above) are examples of the changes required to support this. For more complex +filesystem callbacks, the vfs drops out of rcu-walk mode before the fs call, so +no changes are required to the filesystem. However, this is costly and loses +the benefits of rcu-walk mode. We will begin to add filesystem callbacks that +are rcu-walk aware, shown below. Filesystems should take advantage of this +where possible. + +-- +[mandatory] + d_revalidate is a callback that is made on every path element (if +the filesystem provides it), which requires dropping out of rcu-walk mode. This +may now be called in rcu-walk mode (nd->flags & LOOKUP_RCU). -ECHILD should be +returned if the filesystem cannot handle rcu-walk. See +Documentation/filesystems/vfs.txt for more details. + + permission and check_acl are inode permission checks that are called +on many or all directory inodes on the way down a path walk (to check for +exec permission). These must now be rcu-walk aware (flags & IPERM_FLAG_RCU). +See Documentation/filesystems/vfs.txt for more details. + +-- +[mandatory] + In ->fallocate() you must check the mode option passed in. If your +filesystem does not support hole punching (deallocating space in the middle of a +file) you must return -EOPNOTSUPP if FALLOC_FL_PUNCH_HOLE is set in mode. +Currently you can only have FALLOC_FL_PUNCH_HOLE with FALLOC_FL_KEEP_SIZE set, +so the i_size should not change when hole punching, even when puching the end of +a file off. + +-- +[mandatory] + +-- +[mandatory] + ->get_sb() is gone. Switch to use of ->mount(). Typically it's just +a matter of switching from calling get_sb_... to mount_... and changing the +function type. If you were doing it manually, just switch from setting ->mnt_root +to some pointer to returning that pointer. On errors return ERR_PTR(...). diff --git a/Documentation/filesystems/proc.txt b/Documentation/filesystems/proc.txt new file mode 100644 index 00000000..db3b1aba --- /dev/null +++ b/Documentation/filesystems/proc.txt @@ -0,0 +1,1544 @@ +------------------------------------------------------------------------------ + T H E /proc F I L E S Y S T E M +------------------------------------------------------------------------------ +/proc/sys Terrehon Bowden <terrehon@pacbell.net> October 7 1999 + Bodo Bauer <bb@ricochet.net> + +2.4.x update Jorge Nerin <comandante@zaralinux.com> November 14 2000 +move /proc/sys Shen Feng <shen@cn.fujitsu.com> April 1 2009 +------------------------------------------------------------------------------ +Version 1.3 Kernel version 2.2.12 + Kernel version 2.4.0-test11-pre4 +------------------------------------------------------------------------------ +fixes/update part 1.1 Stefani Seibold <stefani@seibold.net> June 9 2009 + +Table of Contents +----------------- + + 0 Preface + 0.1 Introduction/Credits + 0.2 Legal Stuff + + 1 Collecting System Information + 1.1 Process-Specific Subdirectories + 1.2 Kernel data + 1.3 IDE devices in /proc/ide + 1.4 Networking info in /proc/net + 1.5 SCSI info + 1.6 Parallel port info in /proc/parport + 1.7 TTY info in /proc/tty + 1.8 Miscellaneous kernel statistics in /proc/stat + 1.9 Ext4 file system parameters + + 2 Modifying System Parameters + + 3 Per-Process Parameters + 3.1 /proc/<pid>/oom_adj & /proc/<pid>/oom_score_adj - Adjust the oom-killer + score + 3.2 /proc/<pid>/oom_score - Display current oom-killer score + 3.3 /proc/<pid>/io - Display the IO accounting fields + 3.4 /proc/<pid>/coredump_filter - Core dump filtering settings + 3.5 /proc/<pid>/mountinfo - Information about mounts + 3.6 /proc/<pid>/comm & /proc/<pid>/task/<tid>/comm + + +------------------------------------------------------------------------------ +Preface +------------------------------------------------------------------------------ + +0.1 Introduction/Credits +------------------------ + +This documentation is part of a soon (or so we hope) to be released book on +the SuSE Linux distribution. As there is no complete documentation for the +/proc file system and we've used many freely available sources to write these +chapters, it seems only fair to give the work back to the Linux community. +This work is based on the 2.2.* kernel version and the upcoming 2.4.*. I'm +afraid it's still far from complete, but we hope it will be useful. As far as +we know, it is the first 'all-in-one' document about the /proc file system. It +is focused on the Intel x86 hardware, so if you are looking for PPC, ARM, +SPARC, AXP, etc., features, you probably won't find what you are looking for. +It also only covers IPv4 networking, not IPv6 nor other protocols - sorry. But +additions and patches are welcome and will be added to this document if you +mail them to Bodo. + +We'd like to thank Alan Cox, Rik van Riel, and Alexey Kuznetsov and a lot of +other people for help compiling this documentation. We'd also like to extend a +special thank you to Andi Kleen for documentation, which we relied on heavily +to create this document, as well as the additional information he provided. +Thanks to everybody else who contributed source or docs to the Linux kernel +and helped create a great piece of software... :) + +If you have any comments, corrections or additions, please don't hesitate to +contact Bodo Bauer at bb@ricochet.net. We'll be happy to add them to this +document. + +The latest version of this document is available online at +http://tldp.org/LDP/Linux-Filesystem-Hierarchy/html/proc.html + +If the above direction does not works for you, you could try the kernel +mailing list at linux-kernel@vger.kernel.org and/or try to reach me at +comandante@zaralinux.com. + +0.2 Legal Stuff +--------------- + +We don't guarantee the correctness of this document, and if you come to us +complaining about how you screwed up your system because of incorrect +documentation, we won't feel responsible... + +------------------------------------------------------------------------------ +CHAPTER 1: COLLECTING SYSTEM INFORMATION +------------------------------------------------------------------------------ + +------------------------------------------------------------------------------ +In This Chapter +------------------------------------------------------------------------------ +* Investigating the properties of the pseudo file system /proc and its + ability to provide information on the running Linux system +* Examining /proc's structure +* Uncovering various information about the kernel and the processes running + on the system +------------------------------------------------------------------------------ + + +The proc file system acts as an interface to internal data structures in the +kernel. It can be used to obtain information about the system and to change +certain kernel parameters at runtime (sysctl). + +First, we'll take a look at the read-only parts of /proc. In Chapter 2, we +show you how you can use /proc/sys to change settings. + +1.1 Process-Specific Subdirectories +----------------------------------- + +The directory /proc contains (among other things) one subdirectory for each +process running on the system, which is named after the process ID (PID). + +The link self points to the process reading the file system. Each process +subdirectory has the entries listed in Table 1-1. + + +Table 1-1: Process specific entries in /proc +.............................................................................. + File Content + clear_refs Clears page referenced bits shown in smaps output + cmdline Command line arguments + cpu Current and last cpu in which it was executed (2.4)(smp) + cwd Link to the current working directory + environ Values of environment variables + exe Link to the executable of this process + fd Directory, which contains all file descriptors + maps Memory maps to executables and library files (2.4) + mem Memory held by this process + root Link to the root directory of this process + stat Process status + statm Process memory status information + status Process status in human readable form + wchan If CONFIG_KALLSYMS is set, a pre-decoded wchan + pagemap Page table + stack Report full stack trace, enable via CONFIG_STACKTRACE + smaps a extension based on maps, showing the memory consumption of + each mapping +.............................................................................. + +For example, to get the status information of a process, all you have to do is +read the file /proc/PID/status: + + >cat /proc/self/status + Name: cat + State: R (running) + Tgid: 5452 + Pid: 5452 + PPid: 743 + TracerPid: 0 (2.4) + Uid: 501 501 501 501 + Gid: 100 100 100 100 + FDSize: 256 + Groups: 100 14 16 + VmPeak: 5004 kB + VmSize: 5004 kB + VmLck: 0 kB + VmHWM: 476 kB + VmRSS: 476 kB + VmData: 156 kB + VmStk: 88 kB + VmExe: 68 kB + VmLib: 1412 kB + VmPTE: 20 kb + VmSwap: 0 kB + Threads: 1 + SigQ: 0/28578 + SigPnd: 0000000000000000 + ShdPnd: 0000000000000000 + SigBlk: 0000000000000000 + SigIgn: 0000000000000000 + SigCgt: 0000000000000000 + CapInh: 00000000fffffeff + CapPrm: 0000000000000000 + CapEff: 0000000000000000 + CapBnd: ffffffffffffffff + voluntary_ctxt_switches: 0 + nonvoluntary_ctxt_switches: 1 + +This shows you nearly the same information you would get if you viewed it with +the ps command. In fact, ps uses the proc file system to obtain its +information. But you get a more detailed view of the process by reading the +file /proc/PID/status. It fields are described in table 1-2. + +The statm file contains more detailed information about the process +memory usage. Its seven fields are explained in Table 1-3. The stat file +contains details information about the process itself. Its fields are +explained in Table 1-4. + +(for SMP CONFIG users) +For making accounting scalable, RSS related information are handled in +asynchronous manner and the vaule may not be very precise. To see a precise +snapshot of a moment, you can see /proc/<pid>/smaps file and scan page table. +It's slow but very precise. + +Table 1-2: Contents of the status files (as of 2.6.30-rc7) +.............................................................................. + Field Content + Name filename of the executable + State state (R is running, S is sleeping, D is sleeping + in an uninterruptible wait, Z is zombie, + T is traced or stopped) + Tgid thread group ID + Pid process id + PPid process id of the parent process + TracerPid PID of process tracing this process (0 if not) + Uid Real, effective, saved set, and file system UIDs + Gid Real, effective, saved set, and file system GIDs + FDSize number of file descriptor slots currently allocated + Groups supplementary group list + VmPeak peak virtual memory size + VmSize total program size + VmLck locked memory size + VmHWM peak resident set size ("high water mark") + VmRSS size of memory portions + VmData size of data, stack, and text segments + VmStk size of data, stack, and text segments + VmExe size of text segment + VmLib size of shared library code + VmPTE size of page table entries + VmSwap size of swap usage (the number of referred swapents) + Threads number of threads + SigQ number of signals queued/max. number for queue + SigPnd bitmap of pending signals for the thread + ShdPnd bitmap of shared pending signals for the process + SigBlk bitmap of blocked signals + SigIgn bitmap of ignored signals + SigCgt bitmap of catched signals + CapInh bitmap of inheritable capabilities + CapPrm bitmap of permitted capabilities + CapEff bitmap of effective capabilities + CapBnd bitmap of capabilities bounding set + Cpus_allowed mask of CPUs on which this process may run + Cpus_allowed_list Same as previous, but in "list format" + Mems_allowed mask of memory nodes allowed to this process + Mems_allowed_list Same as previous, but in "list format" + voluntary_ctxt_switches number of voluntary context switches + nonvoluntary_ctxt_switches number of non voluntary context switches +.............................................................................. + +Table 1-3: Contents of the statm files (as of 2.6.8-rc3) +.............................................................................. + Field Content + size total program size (pages) (same as VmSize in status) + resident size of memory portions (pages) (same as VmRSS in status) + shared number of pages that are shared (i.e. backed by a file) + trs number of pages that are 'code' (not including libs; broken, + includes data segment) + lrs number of pages of library (always 0 on 2.6) + drs number of pages of data/stack (including libs; broken, + includes library text) + dt number of dirty pages (always 0 on 2.6) +.............................................................................. + + +Table 1-4: Contents of the stat files (as of 2.6.30-rc7) +.............................................................................. + Field Content + pid process id + tcomm filename of the executable + state state (R is running, S is sleeping, D is sleeping in an + uninterruptible wait, Z is zombie, T is traced or stopped) + ppid process id of the parent process + pgrp pgrp of the process + sid session id + tty_nr tty the process uses + tty_pgrp pgrp of the tty + flags task flags + min_flt number of minor faults + cmin_flt number of minor faults with child's + maj_flt number of major faults + cmaj_flt number of major faults with child's + utime user mode jiffies + stime kernel mode jiffies + cutime user mode jiffies with child's + cstime kernel mode jiffies with child's + priority priority level + nice nice level + num_threads number of threads + it_real_value (obsolete, always 0) + start_time time the process started after system boot + vsize virtual memory size + rss resident set memory size + rsslim current limit in bytes on the rss + start_code address above which program text can run + end_code address below which program text can run + start_stack address of the start of the stack + esp current value of ESP + eip current value of EIP + pending bitmap of pending signals + blocked bitmap of blocked signals + sigign bitmap of ignored signals + sigcatch bitmap of catched signals + wchan address where process went to sleep + 0 (place holder) + 0 (place holder) + exit_signal signal to send to parent thread on exit + task_cpu which CPU the task is scheduled on + rt_priority realtime priority + policy scheduling policy (man sched_setscheduler) + blkio_ticks time spent waiting for block IO + gtime guest time of the task in jiffies + cgtime guest time of the task children in jiffies +.............................................................................. + +The /proc/PID/maps file containing the currently mapped memory regions and +their access permissions. + +The format is: + +address perms offset dev inode pathname + +08048000-08049000 r-xp 00000000 03:00 8312 /opt/test +08049000-0804a000 rw-p 00001000 03:00 8312 /opt/test +0804a000-0806b000 rw-p 00000000 00:00 0 [heap] +a7cb1000-a7cb2000 ---p 00000000 00:00 0 +a7cb2000-a7eb2000 rw-p 00000000 00:00 0 +a7eb2000-a7eb3000 ---p 00000000 00:00 0 +a7eb3000-a7ed5000 rw-p 00000000 00:00 0 +a7ed5000-a8008000 r-xp 00000000 03:00 4222 /lib/libc.so.6 +a8008000-a800a000 r--p 00133000 03:00 4222 /lib/libc.so.6 +a800a000-a800b000 rw-p 00135000 03:00 4222 /lib/libc.so.6 +a800b000-a800e000 rw-p 00000000 00:00 0 +a800e000-a8022000 r-xp 00000000 03:00 14462 /lib/libpthread.so.0 +a8022000-a8023000 r--p 00013000 03:00 14462 /lib/libpthread.so.0 +a8023000-a8024000 rw-p 00014000 03:00 14462 /lib/libpthread.so.0 +a8024000-a8027000 rw-p 00000000 00:00 0 +a8027000-a8043000 r-xp 00000000 03:00 8317 /lib/ld-linux.so.2 +a8043000-a8044000 r--p 0001b000 03:00 8317 /lib/ld-linux.so.2 +a8044000-a8045000 rw-p 0001c000 03:00 8317 /lib/ld-linux.so.2 +aff35000-aff4a000 rw-p 00000000 00:00 0 [stack] +ffffe000-fffff000 r-xp 00000000 00:00 0 [vdso] + +where "address" is the address space in the process that it occupies, "perms" +is a set of permissions: + + r = read + w = write + x = execute + s = shared + p = private (copy on write) + +"offset" is the offset into the mapping, "dev" is the device (major:minor), and +"inode" is the inode on that device. 0 indicates that no inode is associated +with the memory region, as the case would be with BSS (uninitialized data). +The "pathname" shows the name associated file for this mapping. If the mapping +is not associated with a file: + + [heap] = the heap of the program + [stack] = the stack of the main process + [vdso] = the "virtual dynamic shared object", + the kernel system call handler + + or if empty, the mapping is anonymous. + + +The /proc/PID/smaps is an extension based on maps, showing the memory +consumption for each of the process's mappings. For each of mappings there +is a series of lines such as the following: + +08048000-080bc000 r-xp 00000000 03:02 13130 /bin/bash +Size: 1084 kB +Rss: 892 kB +Pss: 374 kB +Shared_Clean: 892 kB +Shared_Dirty: 0 kB +Private_Clean: 0 kB +Private_Dirty: 0 kB +Referenced: 892 kB +Anonymous: 0 kB +Swap: 0 kB +KernelPageSize: 4 kB +MMUPageSize: 4 kB +Locked: 374 kB + +The first of these lines shows the same information as is displayed for the +mapping in /proc/PID/maps. The remaining lines show the size of the mapping +(size), the amount of the mapping that is currently resident in RAM (RSS), the +process' proportional share of this mapping (PSS), the number of clean and +dirty private pages in the mapping. Note that even a page which is part of a +MAP_SHARED mapping, but has only a single pte mapped, i.e. is currently used +by only one process, is accounted as private and not as shared. "Referenced" +indicates the amount of memory currently marked as referenced or accessed. +"Anonymous" shows the amount of memory that does not belong to any file. Even +a mapping associated with a file may contain anonymous pages: when MAP_PRIVATE +and a page is modified, the file page is replaced by a private anonymous copy. +"Swap" shows how much would-be-anonymous memory is also used, but out on +swap. + +This file is only present if the CONFIG_MMU kernel configuration option is +enabled. + +The /proc/PID/clear_refs is used to reset the PG_Referenced and ACCESSED/YOUNG +bits on both physical and virtual pages associated with a process. +To clear the bits for all the pages associated with the process + > echo 1 > /proc/PID/clear_refs + +To clear the bits for the anonymous pages associated with the process + > echo 2 > /proc/PID/clear_refs + +To clear the bits for the file mapped pages associated with the process + > echo 3 > /proc/PID/clear_refs +Any other value written to /proc/PID/clear_refs will have no effect. + +The /proc/pid/pagemap gives the PFN, which can be used to find the pageflags +using /proc/kpageflags and number of times a page is mapped using +/proc/kpagecount. For detailed explanation, see Documentation/vm/pagemap.txt. + +1.2 Kernel data +--------------- + +Similar to the process entries, the kernel data files give information about +the running kernel. The files used to obtain this information are contained in +/proc and are listed in Table 1-5. Not all of these will be present in your +system. It depends on the kernel configuration and the loaded modules, which +files are there, and which are missing. + +Table 1-5: Kernel info in /proc +.............................................................................. + File Content + apm Advanced power management info + buddyinfo Kernel memory allocator information (see text) (2.5) + bus Directory containing bus specific information + cmdline Kernel command line + cpuinfo Info about the CPU + devices Available devices (block and character) + dma Used DMS channels + filesystems Supported filesystems + driver Various drivers grouped here, currently rtc (2.4) + execdomains Execdomains, related to security (2.4) + fb Frame Buffer devices (2.4) + fs File system parameters, currently nfs/exports (2.4) + ide Directory containing info about the IDE subsystem + interrupts Interrupt usage + iomem Memory map (2.4) + ioports I/O port usage + irq Masks for irq to cpu affinity (2.4)(smp?) + isapnp ISA PnP (Plug&Play) Info (2.4) + kcore Kernel core image (can be ELF or A.OUT(deprecated in 2.4)) + kmsg Kernel messages + ksyms Kernel symbol table + loadavg Load average of last 1, 5 & 15 minutes + locks Kernel locks + meminfo Memory info + misc Miscellaneous + modules List of loaded modules + mounts Mounted filesystems + net Networking info (see text) + pagetypeinfo Additional page allocator information (see text) (2.5) + partitions Table of partitions known to the system + pci Deprecated info of PCI bus (new way -> /proc/bus/pci/, + decoupled by lspci (2.4) + rtc Real time clock + scsi SCSI info (see text) + slabinfo Slab pool info + softirqs softirq usage + stat Overall statistics + swaps Swap space utilization + sys See chapter 2 + sysvipc Info of SysVIPC Resources (msg, sem, shm) (2.4) + tty Info of tty drivers + uptime System uptime + version Kernel version + video bttv info of video resources (2.4) + vmallocinfo Show vmalloced areas +.............................................................................. + +You can, for example, check which interrupts are currently in use and what +they are used for by looking in the file /proc/interrupts: + + > cat /proc/interrupts + CPU0 + 0: 8728810 XT-PIC timer + 1: 895 XT-PIC keyboard + 2: 0 XT-PIC cascade + 3: 531695 XT-PIC aha152x + 4: 2014133 XT-PIC serial + 5: 44401 XT-PIC pcnet_cs + 8: 2 XT-PIC rtc + 11: 8 XT-PIC i82365 + 12: 182918 XT-PIC PS/2 Mouse + 13: 1 XT-PIC fpu + 14: 1232265 XT-PIC ide0 + 15: 7 XT-PIC ide1 + NMI: 0 + +In 2.4.* a couple of lines where added to this file LOC & ERR (this time is the +output of a SMP machine): + + > cat /proc/interrupts + + CPU0 CPU1 + 0: 1243498 1214548 IO-APIC-edge timer + 1: 8949 8958 IO-APIC-edge keyboard + 2: 0 0 XT-PIC cascade + 5: 11286 10161 IO-APIC-edge soundblaster + 8: 1 0 IO-APIC-edge rtc + 9: 27422 27407 IO-APIC-edge 3c503 + 12: 113645 113873 IO-APIC-edge PS/2 Mouse + 13: 0 0 XT-PIC fpu + 14: 22491 24012 IO-APIC-edge ide0 + 15: 2183 2415 IO-APIC-edge ide1 + 17: 30564 30414 IO-APIC-level eth0 + 18: 177 164 IO-APIC-level bttv + NMI: 2457961 2457959 + LOC: 2457882 2457881 + ERR: 2155 + +NMI is incremented in this case because every timer interrupt generates a NMI +(Non Maskable Interrupt) which is used by the NMI Watchdog to detect lockups. + +LOC is the local interrupt counter of the internal APIC of every CPU. + +ERR is incremented in the case of errors in the IO-APIC bus (the bus that +connects the CPUs in a SMP system. This means that an error has been detected, +the IO-APIC automatically retry the transmission, so it should not be a big +problem, but you should read the SMP-FAQ. + +In 2.6.2* /proc/interrupts was expanded again. This time the goal was for +/proc/interrupts to display every IRQ vector in use by the system, not +just those considered 'most important'. The new vectors are: + + THR -- interrupt raised when a machine check threshold counter + (typically counting ECC corrected errors of memory or cache) exceeds + a configurable threshold. Only available on some systems. + + TRM -- a thermal event interrupt occurs when a temperature threshold + has been exceeded for the CPU. This interrupt may also be generated + when the temperature drops back to normal. + + SPU -- a spurious interrupt is some interrupt that was raised then lowered + by some IO device before it could be fully processed by the APIC. Hence + the APIC sees the interrupt but does not know what device it came from. + For this case the APIC will generate the interrupt with a IRQ vector + of 0xff. This might also be generated by chipset bugs. + + RES, CAL, TLB -- rescheduling, call and TLB flush interrupts are + sent from one CPU to another per the needs of the OS. Typically, + their statistics are used by kernel developers and interested users to + determine the occurrence of interrupts of the given type. + +The above IRQ vectors are displayed only when relevant. For example, +the threshold vector does not exist on x86_64 platforms. Others are +suppressed when the system is a uniprocessor. As of this writing, only +i386 and x86_64 platforms support the new IRQ vector displays. + +Of some interest is the introduction of the /proc/irq directory to 2.4. +It could be used to set IRQ to CPU affinity, this means that you can "hook" an +IRQ to only one CPU, or to exclude a CPU of handling IRQs. The contents of the +irq subdir is one subdir for each IRQ, and two files; default_smp_affinity and +prof_cpu_mask. + +For example + > ls /proc/irq/ + 0 10 12 14 16 18 2 4 6 8 prof_cpu_mask + 1 11 13 15 17 19 3 5 7 9 default_smp_affinity + > ls /proc/irq/0/ + smp_affinity + +smp_affinity is a bitmask, in which you can specify which CPUs can handle the +IRQ, you can set it by doing: + + > echo 1 > /proc/irq/10/smp_affinity + +This means that only the first CPU will handle the IRQ, but you can also echo +5 which means that only the first and fourth CPU can handle the IRQ. + +The contents of each smp_affinity file is the same by default: + + > cat /proc/irq/0/smp_affinity + ffffffff + +There is an alternate interface, smp_affinity_list which allows specifying +a cpu range instead of a bitmask: + + > cat /proc/irq/0/smp_affinity_list + 1024-1031 + +The default_smp_affinity mask applies to all non-active IRQs, which are the +IRQs which have not yet been allocated/activated, and hence which lack a +/proc/irq/[0-9]* directory. + +The node file on an SMP system shows the node to which the device using the IRQ +reports itself as being attached. This hardware locality information does not +include information about any possible driver locality preference. + +prof_cpu_mask specifies which CPUs are to be profiled by the system wide +profiler. Default value is ffffffff (all cpus if there are only 32 of them). + +The way IRQs are routed is handled by the IO-APIC, and it's Round Robin +between all the CPUs which are allowed to handle it. As usual the kernel has +more info than you and does a better job than you, so the defaults are the +best choice for almost everyone. [Note this applies only to those IO-APIC's +that support "Round Robin" interrupt distribution.] + +There are three more important subdirectories in /proc: net, scsi, and sys. +The general rule is that the contents, or even the existence of these +directories, depend on your kernel configuration. If SCSI is not enabled, the +directory scsi may not exist. The same is true with the net, which is there +only when networking support is present in the running kernel. + +The slabinfo file gives information about memory usage at the slab level. +Linux uses slab pools for memory management above page level in version 2.2. +Commonly used objects have their own slab pool (such as network buffers, +directory cache, and so on). + +.............................................................................. + +> cat /proc/buddyinfo + +Node 0, zone DMA 0 4 5 4 4 3 ... +Node 0, zone Normal 1 0 0 1 101 8 ... +Node 0, zone HighMem 2 0 0 1 1 0 ... + +External fragmentation is a problem under some workloads, and buddyinfo is a +useful tool for helping diagnose these problems. Buddyinfo will give you a +clue as to how big an area you can safely allocate, or why a previous +allocation failed. + +Each column represents the number of pages of a certain order which are +available. In this case, there are 0 chunks of 2^0*PAGE_SIZE available in +ZONE_DMA, 4 chunks of 2^1*PAGE_SIZE in ZONE_DMA, 101 chunks of 2^4*PAGE_SIZE +available in ZONE_NORMAL, etc... + +More information relevant to external fragmentation can be found in +pagetypeinfo. + +> cat /proc/pagetypeinfo +Page block order: 9 +Pages per block: 512 + +Free pages count per migrate type at order 0 1 2 3 4 5 6 7 8 9 10 +Node 0, zone DMA, type Unmovable 0 0 0 1 1 1 1 1 1 1 0 +Node 0, zone DMA, type Reclaimable 0 0 0 0 0 0 0 0 0 0 0 +Node 0, zone DMA, type Movable 1 1 2 1 2 1 1 0 1 0 2 +Node 0, zone DMA, type Reserve 0 0 0 0 0 0 0 0 0 1 0 +Node 0, zone DMA, type Isolate 0 0 0 0 0 0 0 0 0 0 0 +Node 0, zone DMA32, type Unmovable 103 54 77 1 1 1 11 8 7 1 9 +Node 0, zone DMA32, type Reclaimable 0 0 2 1 0 0 0 0 1 0 0 +Node 0, zone DMA32, type Movable 169 152 113 91 77 54 39 13 6 1 452 +Node 0, zone DMA32, type Reserve 1 2 2 2 2 0 1 1 1 1 0 +Node 0, zone DMA32, type Isolate 0 0 0 0 0 0 0 0 0 0 0 + +Number of blocks type Unmovable Reclaimable Movable Reserve Isolate +Node 0, zone DMA 2 0 5 1 0 +Node 0, zone DMA32 41 6 967 2 0 + +Fragmentation avoidance in the kernel works by grouping pages of different +migrate types into the same contiguous regions of memory called page blocks. +A page block is typically the size of the default hugepage size e.g. 2MB on +X86-64. By keeping pages grouped based on their ability to move, the kernel +can reclaim pages within a page block to satisfy a high-order allocation. + +The pagetypinfo begins with information on the size of a page block. It +then gives the same type of information as buddyinfo except broken down +by migrate-type and finishes with details on how many page blocks of each +type exist. + +If min_free_kbytes has been tuned correctly (recommendations made by hugeadm +from libhugetlbfs http://sourceforge.net/projects/libhugetlbfs/), one can +make an estimate of the likely number of huge pages that can be allocated +at a given point in time. All the "Movable" blocks should be allocatable +unless memory has been mlock()'d. Some of the Reclaimable blocks should +also be allocatable although a lot of filesystem metadata may have to be +reclaimed to achieve this. + +.............................................................................. + +meminfo: + +Provides information about distribution and utilization of memory. This +varies by architecture and compile options. The following is from a +16GB PIII, which has highmem enabled. You may not have all of these fields. + +> cat /proc/meminfo + +The "Locked" indicates whether the mapping is locked in memory or not. + + +MemTotal: 16344972 kB +MemFree: 13634064 kB +Buffers: 3656 kB +Cached: 1195708 kB +SwapCached: 0 kB +Active: 891636 kB +Inactive: 1077224 kB +HighTotal: 15597528 kB +HighFree: 13629632 kB +LowTotal: 747444 kB +LowFree: 4432 kB +SwapTotal: 0 kB +SwapFree: 0 kB +Dirty: 968 kB +Writeback: 0 kB +AnonPages: 861800 kB +Mapped: 280372 kB +Slab: 284364 kB +SReclaimable: 159856 kB +SUnreclaim: 124508 kB +PageTables: 24448 kB +NFS_Unstable: 0 kB +Bounce: 0 kB +WritebackTmp: 0 kB +CommitLimit: 7669796 kB +Committed_AS: 100056 kB +VmallocTotal: 112216 kB +VmallocUsed: 428 kB +VmallocChunk: 111088 kB + + MemTotal: Total usable ram (i.e. physical ram minus a few reserved + bits and the kernel binary code) + MemFree: The sum of LowFree+HighFree + Buffers: Relatively temporary storage for raw disk blocks + shouldn't get tremendously large (20MB or so) + Cached: in-memory cache for files read from the disk (the + pagecache). Doesn't include SwapCached + SwapCached: Memory that once was swapped out, is swapped back in but + still also is in the swapfile (if memory is needed it + doesn't need to be swapped out AGAIN because it is already + in the swapfile. This saves I/O) + Active: Memory that has been used more recently and usually not + reclaimed unless absolutely necessary. + Inactive: Memory which has been less recently used. It is more + eligible to be reclaimed for other purposes + HighTotal: + HighFree: Highmem is all memory above ~860MB of physical memory + Highmem areas are for use by userspace programs, or + for the pagecache. The kernel must use tricks to access + this memory, making it slower to access than lowmem. + LowTotal: + LowFree: Lowmem is memory which can be used for everything that + highmem can be used for, but it is also available for the + kernel's use for its own data structures. Among many + other things, it is where everything from the Slab is + allocated. Bad things happen when you're out of lowmem. + SwapTotal: total amount of swap space available + SwapFree: Memory which has been evicted from RAM, and is temporarily + on the disk + Dirty: Memory which is waiting to get written back to the disk + Writeback: Memory which is actively being written back to the disk + AnonPages: Non-file backed pages mapped into userspace page tables + Mapped: files which have been mmaped, such as libraries + Slab: in-kernel data structures cache +SReclaimable: Part of Slab, that might be reclaimed, such as caches + SUnreclaim: Part of Slab, that cannot be reclaimed on memory pressure + PageTables: amount of memory dedicated to the lowest level of page + tables. +NFS_Unstable: NFS pages sent to the server, but not yet committed to stable + storage + Bounce: Memory used for block device "bounce buffers" +WritebackTmp: Memory used by FUSE for temporary writeback buffers + CommitLimit: Based on the overcommit ratio ('vm.overcommit_ratio'), + this is the total amount of memory currently available to + be allocated on the system. This limit is only adhered to + if strict overcommit accounting is enabled (mode 2 in + 'vm.overcommit_memory'). + The CommitLimit is calculated with the following formula: + CommitLimit = ('vm.overcommit_ratio' * Physical RAM) + Swap + For example, on a system with 1G of physical RAM and 7G + of swap with a `vm.overcommit_ratio` of 30 it would + yield a CommitLimit of 7.3G. + For more details, see the memory overcommit documentation + in vm/overcommit-accounting. +Committed_AS: The amount of memory presently allocated on the system. + The committed memory is a sum of all of the memory which + has been allocated by processes, even if it has not been + "used" by them as of yet. A process which malloc()'s 1G + of memory, but only touches 300M of it will only show up + as using 300M of memory even if it has the address space + allocated for the entire 1G. This 1G is memory which has + been "committed" to by the VM and can be used at any time + by the allocating application. With strict overcommit + enabled on the system (mode 2 in 'vm.overcommit_memory'), + allocations which would exceed the CommitLimit (detailed + above) will not be permitted. This is useful if one needs + to guarantee that processes will not fail due to lack of + memory once that memory has been successfully allocated. +VmallocTotal: total size of vmalloc memory area + VmallocUsed: amount of vmalloc area which is used +VmallocChunk: largest contiguous block of vmalloc area which is free + +.............................................................................. + +vmallocinfo: + +Provides information about vmalloced/vmaped areas. One line per area, +containing the virtual address range of the area, size in bytes, +caller information of the creator, and optional information depending +on the kind of area : + + pages=nr number of pages + phys=addr if a physical address was specified + ioremap I/O mapping (ioremap() and friends) + vmalloc vmalloc() area + vmap vmap()ed pages + user VM_USERMAP area + vpages buffer for pages pointers was vmalloced (huge area) + N<node>=nr (Only on NUMA kernels) + Number of pages allocated on memory node <node> + +> cat /proc/vmallocinfo +0xffffc20000000000-0xffffc20000201000 2101248 alloc_large_system_hash+0x204 ... + /0x2c0 pages=512 vmalloc N0=128 N1=128 N2=128 N3=128 +0xffffc20000201000-0xffffc20000302000 1052672 alloc_large_system_hash+0x204 ... + /0x2c0 pages=256 vmalloc N0=64 N1=64 N2=64 N3=64 +0xffffc20000302000-0xffffc20000304000 8192 acpi_tb_verify_table+0x21/0x4f... + phys=7fee8000 ioremap +0xffffc20000304000-0xffffc20000307000 12288 acpi_tb_verify_table+0x21/0x4f... + phys=7fee7000 ioremap +0xffffc2000031d000-0xffffc2000031f000 8192 init_vdso_vars+0x112/0x210 +0xffffc2000031f000-0xffffc2000032b000 49152 cramfs_uncompress_init+0x2e ... + /0x80 pages=11 vmalloc N0=3 N1=3 N2=2 N3=3 +0xffffc2000033a000-0xffffc2000033d000 12288 sys_swapon+0x640/0xac0 ... + pages=2 vmalloc N1=2 +0xffffc20000347000-0xffffc2000034c000 20480 xt_alloc_table_info+0xfe ... + /0x130 [x_tables] pages=4 vmalloc N0=4 +0xffffffffa0000000-0xffffffffa000f000 61440 sys_init_module+0xc27/0x1d00 ... + pages=14 vmalloc N2=14 +0xffffffffa000f000-0xffffffffa0014000 20480 sys_init_module+0xc27/0x1d00 ... + pages=4 vmalloc N1=4 +0xffffffffa0014000-0xffffffffa0017000 12288 sys_init_module+0xc27/0x1d00 ... + pages=2 vmalloc N1=2 +0xffffffffa0017000-0xffffffffa0022000 45056 sys_init_module+0xc27/0x1d00 ... + pages=10 vmalloc N0=10 + +.............................................................................. + +softirqs: + +Provides counts of softirq handlers serviced since boot time, for each cpu. + +> cat /proc/softirqs + CPU0 CPU1 CPU2 CPU3 + HI: 0 0 0 0 + TIMER: 27166 27120 27097 27034 + NET_TX: 0 0 0 17 + NET_RX: 42 0 0 39 + BLOCK: 0 0 107 1121 + TASKLET: 0 0 0 290 + SCHED: 27035 26983 26971 26746 + HRTIMER: 0 0 0 0 + RCU: 1678 1769 2178 2250 + + +1.3 IDE devices in /proc/ide +---------------------------- + +The subdirectory /proc/ide contains information about all IDE devices of which +the kernel is aware. There is one subdirectory for each IDE controller, the +file drivers and a link for each IDE device, pointing to the device directory +in the controller specific subtree. + +The file drivers contains general information about the drivers used for the +IDE devices: + + > cat /proc/ide/drivers + ide-cdrom version 4.53 + ide-disk version 1.08 + +More detailed information can be found in the controller specific +subdirectories. These are named ide0, ide1 and so on. Each of these +directories contains the files shown in table 1-6. + + +Table 1-6: IDE controller info in /proc/ide/ide? +.............................................................................. + File Content + channel IDE channel (0 or 1) + config Configuration (only for PCI/IDE bridge) + mate Mate name + model Type/Chipset of IDE controller +.............................................................................. + +Each device connected to a controller has a separate subdirectory in the +controllers directory. The files listed in table 1-7 are contained in these +directories. + + +Table 1-7: IDE device information +.............................................................................. + File Content + cache The cache + capacity Capacity of the medium (in 512Byte blocks) + driver driver and version + geometry physical and logical geometry + identify device identify block + media media type + model device identifier + settings device setup + smart_thresholds IDE disk management thresholds + smart_values IDE disk management values +.............................................................................. + +The most interesting file is settings. This file contains a nice overview of +the drive parameters: + + # cat /proc/ide/ide0/hda/settings + name value min max mode + ---- ----- --- --- ---- + bios_cyl 526 0 65535 rw + bios_head 255 0 255 rw + bios_sect 63 0 63 rw + breada_readahead 4 0 127 rw + bswap 0 0 1 r + file_readahead 72 0 2097151 rw + io_32bit 0 0 3 rw + keepsettings 0 0 1 rw + max_kb_per_request 122 1 127 rw + multcount 0 0 8 rw + nice1 1 0 1 rw + nowerr 0 0 1 rw + pio_mode write-only 0 255 w + slow 0 0 1 rw + unmaskirq 0 0 1 rw + using_dma 0 0 1 rw + + +1.4 Networking info in /proc/net +-------------------------------- + +The subdirectory /proc/net follows the usual pattern. Table 1-8 shows the +additional values you get for IP version 6 if you configure the kernel to +support this. Table 1-9 lists the files and their meaning. + + +Table 1-8: IPv6 info in /proc/net +.............................................................................. + File Content + udp6 UDP sockets (IPv6) + tcp6 TCP sockets (IPv6) + raw6 Raw device statistics (IPv6) + igmp6 IP multicast addresses, which this host joined (IPv6) + if_inet6 List of IPv6 interface addresses + ipv6_route Kernel routing table for IPv6 + rt6_stats Global IPv6 routing tables statistics + sockstat6 Socket statistics (IPv6) + snmp6 Snmp data (IPv6) +.............................................................................. + + +Table 1-9: Network info in /proc/net +.............................................................................. + File Content + arp Kernel ARP table + dev network devices with statistics + dev_mcast the Layer2 multicast groups a device is listening too + (interface index, label, number of references, number of bound + addresses). + dev_stat network device status + ip_fwchains Firewall chain linkage + ip_fwnames Firewall chain names + ip_masq Directory containing the masquerading tables + ip_masquerade Major masquerading table + netstat Network statistics + raw raw device statistics + route Kernel routing table + rpc Directory containing rpc info + rt_cache Routing cache + snmp SNMP data + sockstat Socket statistics + tcp TCP sockets + tr_rif Token ring RIF routing table + udp UDP sockets + unix UNIX domain sockets + wireless Wireless interface data (Wavelan etc) + igmp IP multicast addresses, which this host joined + psched Global packet scheduler parameters. + netlink List of PF_NETLINK sockets + ip_mr_vifs List of multicast virtual interfaces + ip_mr_cache List of multicast routing cache +.............................................................................. + +You can use this information to see which network devices are available in +your system and how much traffic was routed over those devices: + + > cat /proc/net/dev + Inter-|Receive |[... + face |bytes packets errs drop fifo frame compressed multicast|[... + lo: 908188 5596 0 0 0 0 0 0 [... + ppp0:15475140 20721 410 0 0 410 0 0 [... + eth0: 614530 7085 0 0 0 0 0 1 [... + + ...] Transmit + ...] bytes packets errs drop fifo colls carrier compressed + ...] 908188 5596 0 0 0 0 0 0 + ...] 1375103 17405 0 0 0 0 0 0 + ...] 1703981 5535 0 0 0 3 0 0 + +In addition, each Channel Bond interface has its own directory. For +example, the bond0 device will have a directory called /proc/net/bond0/. +It will contain information that is specific to that bond, such as the +current slaves of the bond, the link status of the slaves, and how +many times the slaves link has failed. + +1.5 SCSI info +------------- + +If you have a SCSI host adapter in your system, you'll find a subdirectory +named after the driver for this adapter in /proc/scsi. You'll also see a list +of all recognized SCSI devices in /proc/scsi: + + >cat /proc/scsi/scsi + Attached devices: + Host: scsi0 Channel: 00 Id: 00 Lun: 00 + Vendor: IBM Model: DGHS09U Rev: 03E0 + Type: Direct-Access ANSI SCSI revision: 03 + Host: scsi0 Channel: 00 Id: 06 Lun: 00 + Vendor: PIONEER Model: CD-ROM DR-U06S Rev: 1.04 + Type: CD-ROM ANSI SCSI revision: 02 + + +The directory named after the driver has one file for each adapter found in +the system. These files contain information about the controller, including +the used IRQ and the IO address range. The amount of information shown is +dependent on the adapter you use. The example shows the output for an Adaptec +AHA-2940 SCSI adapter: + + > cat /proc/scsi/aic7xxx/0 + + Adaptec AIC7xxx driver version: 5.1.19/3.2.4 + Compile Options: + TCQ Enabled By Default : Disabled + AIC7XXX_PROC_STATS : Disabled + AIC7XXX_RESET_DELAY : 5 + Adapter Configuration: + SCSI Adapter: Adaptec AHA-294X Ultra SCSI host adapter + Ultra Wide Controller + PCI MMAPed I/O Base: 0xeb001000 + Adapter SEEPROM Config: SEEPROM found and used. + Adaptec SCSI BIOS: Enabled + IRQ: 10 + SCBs: Active 0, Max Active 2, + Allocated 15, HW 16, Page 255 + Interrupts: 160328 + BIOS Control Word: 0x18b6 + Adapter Control Word: 0x005b + Extended Translation: Enabled + Disconnect Enable Flags: 0xffff + Ultra Enable Flags: 0x0001 + Tag Queue Enable Flags: 0x0000 + Ordered Queue Tag Flags: 0x0000 + Default Tag Queue Depth: 8 + Tagged Queue By Device array for aic7xxx host instance 0: + {255,255,255,255,255,255,255,255,255,255,255,255,255,255,255,255} + Actual queue depth per device for aic7xxx host instance 0: + {1,1,1,1,1,1,1,1,1,1,1,1,1,1,1,1} + Statistics: + (scsi0:0:0:0) + Device using Wide/Sync transfers at 40.0 MByte/sec, offset 8 + Transinfo settings: current(12/8/1/0), goal(12/8/1/0), user(12/15/1/0) + Total transfers 160151 (74577 reads and 85574 writes) + (scsi0:0:6:0) + Device using Narrow/Sync transfers at 5.0 MByte/sec, offset 15 + Transinfo settings: current(50/15/0/0), goal(50/15/0/0), user(50/15/0/0) + Total transfers 0 (0 reads and 0 writes) + + +1.6 Parallel port info in /proc/parport +--------------------------------------- + +The directory /proc/parport contains information about the parallel ports of +your system. It has one subdirectory for each port, named after the port +number (0,1,2,...). + +These directories contain the four files shown in Table 1-10. + + +Table 1-10: Files in /proc/parport +.............................................................................. + File Content + autoprobe Any IEEE-1284 device ID information that has been acquired. + devices list of the device drivers using that port. A + will appear by the + name of the device currently using the port (it might not appear + against any). + hardware Parallel port's base address, IRQ line and DMA channel. + irq IRQ that parport is using for that port. This is in a separate + file to allow you to alter it by writing a new value in (IRQ + number or none). +.............................................................................. + +1.7 TTY info in /proc/tty +------------------------- + +Information about the available and actually used tty's can be found in the +directory /proc/tty.You'll find entries for drivers and line disciplines in +this directory, as shown in Table 1-11. + + +Table 1-11: Files in /proc/tty +.............................................................................. + File Content + drivers list of drivers and their usage + ldiscs registered line disciplines + driver/serial usage statistic and status of single tty lines +.............................................................................. + +To see which tty's are currently in use, you can simply look into the file +/proc/tty/drivers: + + > cat /proc/tty/drivers + pty_slave /dev/pts 136 0-255 pty:slave + pty_master /dev/ptm 128 0-255 pty:master + pty_slave /dev/ttyp 3 0-255 pty:slave + pty_master /dev/pty 2 0-255 pty:master + serial /dev/cua 5 64-67 serial:callout + serial /dev/ttyS 4 64-67 serial + /dev/tty0 /dev/tty0 4 0 system:vtmaster + /dev/ptmx /dev/ptmx 5 2 system + /dev/console /dev/console 5 1 system:console + /dev/tty /dev/tty 5 0 system:/dev/tty + unknown /dev/tty 4 1-63 console + + +1.8 Miscellaneous kernel statistics in /proc/stat +------------------------------------------------- + +Various pieces of information about kernel activity are available in the +/proc/stat file. All of the numbers reported in this file are aggregates +since the system first booted. For a quick look, simply cat the file: + + > cat /proc/stat + cpu 2255 34 2290 22625563 6290 127 456 0 0 + cpu0 1132 34 1441 11311718 3675 127 438 0 0 + cpu1 1123 0 849 11313845 2614 0 18 0 0 + intr 114930548 113199788 3 0 5 263 0 4 [... lots more numbers ...] + ctxt 1990473 + btime 1062191376 + processes 2915 + procs_running 1 + procs_blocked 0 + softirq 183433 0 21755 12 39 1137 231 21459 2263 + +The very first "cpu" line aggregates the numbers in all of the other "cpuN" +lines. These numbers identify the amount of time the CPU has spent performing +different kinds of work. Time units are in USER_HZ (typically hundredths of a +second). The meanings of the columns are as follows, from left to right: + +- user: normal processes executing in user mode +- nice: niced processes executing in user mode +- system: processes executing in kernel mode +- idle: twiddling thumbs +- iowait: waiting for I/O to complete +- irq: servicing interrupts +- softirq: servicing softirqs +- steal: involuntary wait +- guest: running a normal guest +- guest_nice: running a niced guest + +The "intr" line gives counts of interrupts serviced since boot time, for each +of the possible system interrupts. The first column is the total of all +interrupts serviced; each subsequent column is the total for that particular +interrupt. + +The "ctxt" line gives the total number of context switches across all CPUs. + +The "btime" line gives the time at which the system booted, in seconds since +the Unix epoch. + +The "processes" line gives the number of processes and threads created, which +includes (but is not limited to) those created by calls to the fork() and +clone() system calls. + +The "procs_running" line gives the total number of threads that are +running or ready to run (i.e., the total number of runnable threads). + +The "procs_blocked" line gives the number of processes currently blocked, +waiting for I/O to complete. + +The "softirq" line gives counts of softirqs serviced since boot time, for each +of the possible system softirqs. The first column is the total of all +softirqs serviced; each subsequent column is the total for that particular +softirq. + + +1.9 Ext4 file system parameters +------------------------------ + +Information about mounted ext4 file systems can be found in +/proc/fs/ext4. Each mounted filesystem will have a directory in +/proc/fs/ext4 based on its device name (i.e., /proc/fs/ext4/hdc or +/proc/fs/ext4/dm-0). The files in each per-device directory are shown +in Table 1-12, below. + +Table 1-12: Files in /proc/fs/ext4/<devname> +.............................................................................. + File Content + mb_groups details of multiblock allocator buddy cache of free blocks +.............................................................................. + +2.0 /proc/consoles +------------------ +Shows registered system console lines. + +To see which character device lines are currently used for the system console +/dev/console, you may simply look into the file /proc/consoles: + + > cat /proc/consoles + tty0 -WU (ECp) 4:7 + ttyS0 -W- (Ep) 4:64 + +The columns are: + + device name of the device + operations R = can do read operations + W = can do write operations + U = can do unblank + flags E = it is enabled + C = it is preferred console + B = it is primary boot console + p = it is used for printk buffer + b = it is not a TTY but a Braille device + a = it is safe to use when cpu is offline + major:minor major and minor number of the device separated by a colon + +------------------------------------------------------------------------------ +Summary +------------------------------------------------------------------------------ +The /proc file system serves information about the running system. It not only +allows access to process data but also allows you to request the kernel status +by reading files in the hierarchy. + +The directory structure of /proc reflects the types of information and makes +it easy, if not obvious, where to look for specific data. +------------------------------------------------------------------------------ + +------------------------------------------------------------------------------ +CHAPTER 2: MODIFYING SYSTEM PARAMETERS +------------------------------------------------------------------------------ + +------------------------------------------------------------------------------ +In This Chapter +------------------------------------------------------------------------------ +* Modifying kernel parameters by writing into files found in /proc/sys +* Exploring the files which modify certain parameters +* Review of the /proc/sys file tree +------------------------------------------------------------------------------ + + +A very interesting part of /proc is the directory /proc/sys. This is not only +a source of information, it also allows you to change parameters within the +kernel. Be very careful when attempting this. You can optimize your system, +but you can also cause it to crash. Never alter kernel parameters on a +production system. Set up a development machine and test to make sure that +everything works the way you want it to. You may have no alternative but to +reboot the machine once an error has been made. + +To change a value, simply echo the new value into the file. An example is +given below in the section on the file system data. You need to be root to do +this. You can create your own boot script to perform this every time your +system boots. + +The files in /proc/sys can be used to fine tune and monitor miscellaneous and +general things in the operation of the Linux kernel. Since some of the files +can inadvertently disrupt your system, it is advisable to read both +documentation and source before actually making adjustments. In any case, be +very careful when writing to any of these files. The entries in /proc may +change slightly between the 2.1.* and the 2.2 kernel, so if there is any doubt +review the kernel documentation in the directory /usr/src/linux/Documentation. +This chapter is heavily based on the documentation included in the pre 2.2 +kernels, and became part of it in version 2.2.1 of the Linux kernel. + +Please see: Documentation/sysctls/ directory for descriptions of these +entries. + +------------------------------------------------------------------------------ +Summary +------------------------------------------------------------------------------ +Certain aspects of kernel behavior can be modified at runtime, without the +need to recompile the kernel, or even to reboot the system. The files in the +/proc/sys tree can not only be read, but also modified. You can use the echo +command to write value into these files, thereby changing the default settings +of the kernel. +------------------------------------------------------------------------------ + +------------------------------------------------------------------------------ +CHAPTER 3: PER-PROCESS PARAMETERS +------------------------------------------------------------------------------ + +3.1 /proc/<pid>/oom_adj & /proc/<pid>/oom_score_adj- Adjust the oom-killer score +-------------------------------------------------------------------------------- + +These file can be used to adjust the badness heuristic used to select which +process gets killed in out of memory conditions. + +The badness heuristic assigns a value to each candidate task ranging from 0 +(never kill) to 1000 (always kill) to determine which process is targeted. The +units are roughly a proportion along that range of allowed memory the process +may allocate from based on an estimation of its current memory and swap use. +For example, if a task is using all allowed memory, its badness score will be +1000. If it is using half of its allowed memory, its score will be 500. + +There is an additional factor included in the badness score: root +processes are given 3% extra memory over other tasks. + +The amount of "allowed" memory depends on the context in which the oom killer +was called. If it is due to the memory assigned to the allocating task's cpuset +being exhausted, the allowed memory represents the set of mems assigned to that +cpuset. If it is due to a mempolicy's node(s) being exhausted, the allowed +memory represents the set of mempolicy nodes. If it is due to a memory +limit (or swap limit) being reached, the allowed memory is that configured +limit. Finally, if it is due to the entire system being out of memory, the +allowed memory represents all allocatable resources. + +The value of /proc/<pid>/oom_score_adj is added to the badness score before it +is used to determine which task to kill. Acceptable values range from -1000 +(OOM_SCORE_ADJ_MIN) to +1000 (OOM_SCORE_ADJ_MAX). This allows userspace to +polarize the preference for oom killing either by always preferring a certain +task or completely disabling it. The lowest possible value, -1000, is +equivalent to disabling oom killing entirely for that task since it will always +report a badness score of 0. + +Consequently, it is very simple for userspace to define the amount of memory to +consider for each task. Setting a /proc/<pid>/oom_score_adj value of +500, for +example, is roughly equivalent to allowing the remainder of tasks sharing the +same system, cpuset, mempolicy, or memory controller resources to use at least +50% more memory. A value of -500, on the other hand, would be roughly +equivalent to discounting 50% of the task's allowed memory from being considered +as scoring against the task. + +For backwards compatibility with previous kernels, /proc/<pid>/oom_adj may also +be used to tune the badness score. Its acceptable values range from -16 +(OOM_ADJUST_MIN) to +15 (OOM_ADJUST_MAX) and a special value of -17 +(OOM_DISABLE) to disable oom killing entirely for that task. Its value is +scaled linearly with /proc/<pid>/oom_score_adj. + +Writing to /proc/<pid>/oom_score_adj or /proc/<pid>/oom_adj will change the +other with its scaled value. + +The value of /proc/<pid>/oom_score_adj may be reduced no lower than the last +value set by a CAP_SYS_RESOURCE process. To reduce the value any lower +requires CAP_SYS_RESOURCE. + +NOTICE: /proc/<pid>/oom_adj is deprecated and will be removed, please see +Documentation/feature-removal-schedule.txt. + +Caveat: when a parent task is selected, the oom killer will sacrifice any first +generation children with separate address spaces instead, if possible. This +avoids servers and important system daemons from being killed and loses the +minimal amount of work. + + +3.2 /proc/<pid>/oom_score - Display current oom-killer score +------------------------------------------------------------- + +This file can be used to check the current score used by the oom-killer is for +any given <pid>. Use it together with /proc/<pid>/oom_adj to tune which +process should be killed in an out-of-memory situation. + + +3.3 /proc/<pid>/io - Display the IO accounting fields +------------------------------------------------------- + +This file contains IO statistics for each running process + +Example +------- + +test:/tmp # dd if=/dev/zero of=/tmp/test.dat & +[1] 3828 + +test:/tmp # cat /proc/3828/io +rchar: 323934931 +wchar: 323929600 +syscr: 632687 +syscw: 632675 +read_bytes: 0 +write_bytes: 323932160 +cancelled_write_bytes: 0 + + +Description +----------- + +rchar +----- + +I/O counter: chars read +The number of bytes which this task has caused to be read from storage. This +is simply the sum of bytes which this process passed to read() and pread(). +It includes things like tty IO and it is unaffected by whether or not actual +physical disk IO was required (the read might have been satisfied from +pagecache) + + +wchar +----- + +I/O counter: chars written +The number of bytes which this task has caused, or shall cause to be written +to disk. Similar caveats apply here as with rchar. + + +syscr +----- + +I/O counter: read syscalls +Attempt to count the number of read I/O operations, i.e. syscalls like read() +and pread(). + + +syscw +----- + +I/O counter: write syscalls +Attempt to count the number of write I/O operations, i.e. syscalls like +write() and pwrite(). + + +read_bytes +---------- + +I/O counter: bytes read +Attempt to count the number of bytes which this process really did cause to +be fetched from the storage layer. Done at the submit_bio() level, so it is +accurate for block-backed filesystems. <please add status regarding NFS and +CIFS at a later time> + + +write_bytes +----------- + +I/O counter: bytes written +Attempt to count the number of bytes which this process caused to be sent to +the storage layer. This is done at page-dirtying time. + + +cancelled_write_bytes +--------------------- + +The big inaccuracy here is truncate. If a process writes 1MB to a file and +then deletes the file, it will in fact perform no writeout. But it will have +been accounted as having caused 1MB of write. +In other words: The number of bytes which this process caused to not happen, +by truncating pagecache. A task can cause "negative" IO too. If this task +truncates some dirty pagecache, some IO which another task has been accounted +for (in its write_bytes) will not be happening. We _could_ just subtract that +from the truncating task's write_bytes, but there is information loss in doing +that. + + +Note +---- + +At its current implementation state, this is a bit racy on 32-bit machines: if +process A reads process B's /proc/pid/io while process B is updating one of +those 64-bit counters, process A could see an intermediate result. + + +More information about this can be found within the taskstats documentation in +Documentation/accounting. + +3.4 /proc/<pid>/coredump_filter - Core dump filtering settings +--------------------------------------------------------------- +When a process is dumped, all anonymous memory is written to a core file as +long as the size of the core file isn't limited. But sometimes we don't want +to dump some memory segments, for example, huge shared memory. Conversely, +sometimes we want to save file-backed memory segments into a core file, not +only the individual files. + +/proc/<pid>/coredump_filter allows you to customize which memory segments +will be dumped when the <pid> process is dumped. coredump_filter is a bitmask +of memory types. If a bit of the bitmask is set, memory segments of the +corresponding memory type are dumped, otherwise they are not dumped. + +The following 7 memory types are supported: + - (bit 0) anonymous private memory + - (bit 1) anonymous shared memory + - (bit 2) file-backed private memory + - (bit 3) file-backed shared memory + - (bit 4) ELF header pages in file-backed private memory areas (it is + effective only if the bit 2 is cleared) + - (bit 5) hugetlb private memory + - (bit 6) hugetlb shared memory + + Note that MMIO pages such as frame buffer are never dumped and vDSO pages + are always dumped regardless of the bitmask status. + + Note bit 0-4 doesn't effect any hugetlb memory. hugetlb memory are only + effected by bit 5-6. + +Default value of coredump_filter is 0x23; this means all anonymous memory +segments and hugetlb private memory are dumped. + +If you don't want to dump all shared memory segments attached to pid 1234, +write 0x21 to the process's proc file. + + $ echo 0x21 > /proc/1234/coredump_filter + +When a new process is created, the process inherits the bitmask status from its +parent. It is useful to set up coredump_filter before the program runs. +For example: + + $ echo 0x7 > /proc/self/coredump_filter + $ ./some_program + +3.5 /proc/<pid>/mountinfo - Information about mounts +-------------------------------------------------------- + +This file contains lines of the form: + +36 35 98:0 /mnt1 /mnt2 rw,noatime master:1 - ext3 /dev/root rw,errors=continue +(1)(2)(3) (4) (5) (6) (7) (8) (9) (10) (11) + +(1) mount ID: unique identifier of the mount (may be reused after umount) +(2) parent ID: ID of parent (or of self for the top of the mount tree) +(3) major:minor: value of st_dev for files on filesystem +(4) root: root of the mount within the filesystem +(5) mount point: mount point relative to the process's root +(6) mount options: per mount options +(7) optional fields: zero or more fields of the form "tag[:value]" +(8) separator: marks the end of the optional fields +(9) filesystem type: name of filesystem of the form "type[.subtype]" +(10) mount source: filesystem specific information or "none" +(11) super options: per super block options + +Parsers should ignore all unrecognised optional fields. Currently the +possible optional fields are: + +shared:X mount is shared in peer group X +master:X mount is slave to peer group X +propagate_from:X mount is slave and receives propagation from peer group X (*) +unbindable mount is unbindable + +(*) X is the closest dominant peer group under the process's root. If +X is the immediate master of the mount, or if there's no dominant peer +group under the same root, then only the "master:X" field is present +and not the "propagate_from:X" field. + +For more information on mount propagation see: + + Documentation/filesystems/sharedsubtree.txt + + +3.6 /proc/<pid>/comm & /proc/<pid>/task/<tid>/comm +-------------------------------------------------------- +These files provide a method to access a tasks comm value. It also allows for +a task to set its own or one of its thread siblings comm value. The comm value +is limited in size compared to the cmdline value, so writing anything longer +then the kernel's TASK_COMM_LEN (currently 16 chars) will result in a truncated +comm value. diff --git a/Documentation/filesystems/quota.txt b/Documentation/filesystems/quota.txt new file mode 100644 index 00000000..5e8de25b --- /dev/null +++ b/Documentation/filesystems/quota.txt @@ -0,0 +1,65 @@ + +Quota subsystem +=============== + +Quota subsystem allows system administrator to set limits on used space and +number of used inodes (inode is a filesystem structure which is associated with +each file or directory) for users and/or groups. For both used space and number +of used inodes there are actually two limits. The first one is called softlimit +and the second one hardlimit. An user can never exceed a hardlimit for any +resource (unless he has CAP_SYS_RESOURCE capability). User is allowed to exceed +softlimit but only for limited period of time. This period is called "grace +period" or "grace time". When grace time is over, user is not able to allocate +more space/inodes until he frees enough of them to get below softlimit. + +Quota limits (and amount of grace time) are set independently for each +filesystem. + +For more details about quota design, see the documentation in quota-tools package +(http://sourceforge.net/projects/linuxquota). + +Quota netlink interface +======================= +When user exceeds a softlimit, runs out of grace time or reaches hardlimit, +quota subsystem traditionally printed a message to the controlling terminal of +the process which caused the excess. This method has the disadvantage that +when user is using a graphical desktop he usually cannot see the message. +Thus quota netlink interface has been designed to pass information about +the above events to userspace. There they can be captured by an application +and processed accordingly. + +The interface uses generic netlink framework (see +http://lwn.net/Articles/208755/ and http://people.suug.ch/~tgr/libnl/ for more +details about this layer). The name of the quota generic netlink interface +is "VFS_DQUOT". Definitions of constants below are in <linux/quota.h>. + Currently, the interface supports only one message type QUOTA_NL_C_WARNING. +This command is used to send a notification about any of the above mentioned +events. Each message has six attributes. These are (type of the argument is +in parentheses): + QUOTA_NL_A_QTYPE (u32) + - type of quota being exceeded (one of USRQUOTA, GRPQUOTA) + QUOTA_NL_A_EXCESS_ID (u64) + - UID/GID (depends on quota type) of user / group whose limit + is being exceeded. + QUOTA_NL_A_CAUSED_ID (u64) + - UID of a user who caused the event + QUOTA_NL_A_WARNING (u32) + - what kind of limit is exceeded: + QUOTA_NL_IHARDWARN - inode hardlimit + QUOTA_NL_ISOFTLONGWARN - inode softlimit is exceeded longer + than given grace period + QUOTA_NL_ISOFTWARN - inode softlimit + QUOTA_NL_BHARDWARN - space (block) hardlimit + QUOTA_NL_BSOFTLONGWARN - space (block) softlimit is exceeded + longer than given grace period. + QUOTA_NL_BSOFTWARN - space (block) softlimit + - four warnings are also defined for the event when user stops + exceeding some limit: + QUOTA_NL_IHARDBELOW - inode hardlimit + QUOTA_NL_ISOFTBELOW - inode softlimit + QUOTA_NL_BHARDBELOW - space (block) hardlimit + QUOTA_NL_BSOFTBELOW - space (block) softlimit + QUOTA_NL_A_DEV_MAJOR (u32) + - major number of a device with the affected filesystem + QUOTA_NL_A_DEV_MINOR (u32) + - minor number of a device with the affected filesystem diff --git a/Documentation/filesystems/ramfs-rootfs-initramfs.txt b/Documentation/filesystems/ramfs-rootfs-initramfs.txt new file mode 100644 index 00000000..a8273d5f --- /dev/null +++ b/Documentation/filesystems/ramfs-rootfs-initramfs.txt @@ -0,0 +1,355 @@ +ramfs, rootfs and initramfs +October 17, 2005 +Rob Landley <rob@landley.net> +============================= + +What is ramfs? +-------------- + +Ramfs is a very simple filesystem that exports Linux's disk caching +mechanisms (the page cache and dentry cache) as a dynamically resizable +RAM-based filesystem. + +Normally all files are cached in memory by Linux. Pages of data read from +backing store (usually the block device the filesystem is mounted on) are kept +around in case it's needed again, but marked as clean (freeable) in case the +Virtual Memory system needs the memory for something else. Similarly, data +written to files is marked clean as soon as it has been written to backing +store, but kept around for caching purposes until the VM reallocates the +memory. A similar mechanism (the dentry cache) greatly speeds up access to +directories. + +With ramfs, there is no backing store. Files written into ramfs allocate +dentries and page cache as usual, but there's nowhere to write them to. +This means the pages are never marked clean, so they can't be freed by the +VM when it's looking to recycle memory. + +The amount of code required to implement ramfs is tiny, because all the +work is done by the existing Linux caching infrastructure. Basically, +you're mounting the disk cache as a filesystem. Because of this, ramfs is not +an optional component removable via menuconfig, since there would be negligible +space savings. + +ramfs and ramdisk: +------------------ + +The older "ram disk" mechanism created a synthetic block device out of +an area of RAM and used it as backing store for a filesystem. This block +device was of fixed size, so the filesystem mounted on it was of fixed +size. Using a ram disk also required unnecessarily copying memory from the +fake block device into the page cache (and copying changes back out), as well +as creating and destroying dentries. Plus it needed a filesystem driver +(such as ext2) to format and interpret this data. + +Compared to ramfs, this wastes memory (and memory bus bandwidth), creates +unnecessary work for the CPU, and pollutes the CPU caches. (There are tricks +to avoid this copying by playing with the page tables, but they're unpleasantly +complicated and turn out to be about as expensive as the copying anyway.) +More to the point, all the work ramfs is doing has to happen _anyway_, +since all file access goes through the page and dentry caches. The RAM +disk is simply unnecessary; ramfs is internally much simpler. + +Another reason ramdisks are semi-obsolete is that the introduction of +loopback devices offered a more flexible and convenient way to create +synthetic block devices, now from files instead of from chunks of memory. +See losetup (8) for details. + +ramfs and tmpfs: +---------------- + +One downside of ramfs is you can keep writing data into it until you fill +up all memory, and the VM can't free it because the VM thinks that files +should get written to backing store (rather than swap space), but ramfs hasn't +got any backing store. Because of this, only root (or a trusted user) should +be allowed write access to a ramfs mount. + +A ramfs derivative called tmpfs was created to add size limits, and the ability +to write the data to swap space. Normal users can be allowed write access to +tmpfs mounts. See Documentation/filesystems/tmpfs.txt for more information. + +What is rootfs? +--------------- + +Rootfs is a special instance of ramfs (or tmpfs, if that's enabled), which is +always present in 2.6 systems. You can't unmount rootfs for approximately the +same reason you can't kill the init process; rather than having special code +to check for and handle an empty list, it's smaller and simpler for the kernel +to just make sure certain lists can't become empty. + +Most systems just mount another filesystem over rootfs and ignore it. The +amount of space an empty instance of ramfs takes up is tiny. + +What is initramfs? +------------------ + +All 2.6 Linux kernels contain a gzipped "cpio" format archive, which is +extracted into rootfs when the kernel boots up. After extracting, the kernel +checks to see if rootfs contains a file "init", and if so it executes it as PID +1. If found, this init process is responsible for bringing the system the +rest of the way up, including locating and mounting the real root device (if +any). If rootfs does not contain an init program after the embedded cpio +archive is extracted into it, the kernel will fall through to the older code +to locate and mount a root partition, then exec some variant of /sbin/init +out of that. + +All this differs from the old initrd in several ways: + + - The old initrd was always a separate file, while the initramfs archive is + linked into the linux kernel image. (The directory linux-*/usr is devoted + to generating this archive during the build.) + + - The old initrd file was a gzipped filesystem image (in some file format, + such as ext2, that needed a driver built into the kernel), while the new + initramfs archive is a gzipped cpio archive (like tar only simpler, + see cpio(1) and Documentation/early-userspace/buffer-format.txt). The + kernel's cpio extraction code is not only extremely small, it's also + __init text and data that can be discarded during the boot process. + + - The program run by the old initrd (which was called /initrd, not /init) did + some setup and then returned to the kernel, while the init program from + initramfs is not expected to return to the kernel. (If /init needs to hand + off control it can overmount / with a new root device and exec another init + program. See the switch_root utility, below.) + + - When switching another root device, initrd would pivot_root and then + umount the ramdisk. But initramfs is rootfs: you can neither pivot_root + rootfs, nor unmount it. Instead delete everything out of rootfs to + free up the space (find -xdev / -exec rm '{}' ';'), overmount rootfs + with the new root (cd /newmount; mount --move . /; chroot .), attach + stdin/stdout/stderr to the new /dev/console, and exec the new init. + + Since this is a remarkably persnickety process (and involves deleting + commands before you can run them), the klibc package introduced a helper + program (utils/run_init.c) to do all this for you. Most other packages + (such as busybox) have named this command "switch_root". + +Populating initramfs: +--------------------- + +The 2.6 kernel build process always creates a gzipped cpio format initramfs +archive and links it into the resulting kernel binary. By default, this +archive is empty (consuming 134 bytes on x86). + +The config option CONFIG_INITRAMFS_SOURCE (in General Setup in menuconfig, +and living in usr/Kconfig) can be used to specify a source for the +initramfs archive, which will automatically be incorporated into the +resulting binary. This option can point to an existing gzipped cpio +archive, a directory containing files to be archived, or a text file +specification such as the following example: + + dir /dev 755 0 0 + nod /dev/console 644 0 0 c 5 1 + nod /dev/loop0 644 0 0 b 7 0 + dir /bin 755 1000 1000 + slink /bin/sh busybox 777 0 0 + file /bin/busybox initramfs/busybox 755 0 0 + dir /proc 755 0 0 + dir /sys 755 0 0 + dir /mnt 755 0 0 + file /init initramfs/init.sh 755 0 0 + +Run "usr/gen_init_cpio" (after the kernel build) to get a usage message +documenting the above file format. + +One advantage of the configuration file is that root access is not required to +set permissions or create device nodes in the new archive. (Note that those +two example "file" entries expect to find files named "init.sh" and "busybox" in +a directory called "initramfs", under the linux-2.6.* directory. See +Documentation/early-userspace/README for more details.) + +The kernel does not depend on external cpio tools. If you specify a +directory instead of a configuration file, the kernel's build infrastructure +creates a configuration file from that directory (usr/Makefile calls +scripts/gen_initramfs_list.sh), and proceeds to package up that directory +using the config file (by feeding it to usr/gen_init_cpio, which is created +from usr/gen_init_cpio.c). The kernel's build-time cpio creation code is +entirely self-contained, and the kernel's boot-time extractor is also +(obviously) self-contained. + +The one thing you might need external cpio utilities installed for is creating +or extracting your own preprepared cpio files to feed to the kernel build +(instead of a config file or directory). + +The following command line can extract a cpio image (either by the above script +or by the kernel build) back into its component files: + + cpio -i -d -H newc -F initramfs_data.cpio --no-absolute-filenames + +The following shell script can create a prebuilt cpio archive you can +use in place of the above config file: + + #!/bin/sh + + # Copyright 2006 Rob Landley <rob@landley.net> and TimeSys Corporation. + # Licensed under GPL version 2 + + if [ $# -ne 2 ] + then + echo "usage: mkinitramfs directory imagename.cpio.gz" + exit 1 + fi + + if [ -d "$1" ] + then + echo "creating $2 from $1" + (cd "$1"; find . | cpio -o -H newc | gzip) > "$2" + else + echo "First argument must be a directory" + exit 1 + fi + +Note: The cpio man page contains some bad advice that will break your initramfs +archive if you follow it. It says "A typical way to generate the list +of filenames is with the find command; you should give find the -depth option +to minimize problems with permissions on directories that are unwritable or not +searchable." Don't do this when creating initramfs.cpio.gz images, it won't +work. The Linux kernel cpio extractor won't create files in a directory that +doesn't exist, so the directory entries must go before the files that go in +those directories. The above script gets them in the right order. + +External initramfs images: +-------------------------- + +If the kernel has initrd support enabled, an external cpio.gz archive can also +be passed into a 2.6 kernel in place of an initrd. In this case, the kernel +will autodetect the type (initramfs, not initrd) and extract the external cpio +archive into rootfs before trying to run /init. + +This has the memory efficiency advantages of initramfs (no ramdisk block +device) but the separate packaging of initrd (which is nice if you have +non-GPL code you'd like to run from initramfs, without conflating it with +the GPL licensed Linux kernel binary). + +It can also be used to supplement the kernel's built-in initramfs image. The +files in the external archive will overwrite any conflicting files in +the built-in initramfs archive. Some distributors also prefer to customize +a single kernel image with task-specific initramfs images, without recompiling. + +Contents of initramfs: +---------------------- + +An initramfs archive is a complete self-contained root filesystem for Linux. +If you don't already understand what shared libraries, devices, and paths +you need to get a minimal root filesystem up and running, here are some +references: +http://www.tldp.org/HOWTO/Bootdisk-HOWTO/ +http://www.tldp.org/HOWTO/From-PowerUp-To-Bash-Prompt-HOWTO.html +http://www.linuxfromscratch.org/lfs/view/stable/ + +The "klibc" package (http://www.kernel.org/pub/linux/libs/klibc) is +designed to be a tiny C library to statically link early userspace +code against, along with some related utilities. It is BSD licensed. + +I use uClibc (http://www.uclibc.org) and busybox (http://www.busybox.net) +myself. These are LGPL and GPL, respectively. (A self-contained initramfs +package is planned for the busybox 1.3 release.) + +In theory you could use glibc, but that's not well suited for small embedded +uses like this. (A "hello world" program statically linked against glibc is +over 400k. With uClibc it's 7k. Also note that glibc dlopens libnss to do +name lookups, even when otherwise statically linked.) + +A good first step is to get initramfs to run a statically linked "hello world" +program as init, and test it under an emulator like qemu (www.qemu.org) or +User Mode Linux, like so: + + cat > hello.c << EOF + #include <stdio.h> + #include <unistd.h> + + int main(int argc, char *argv[]) + { + printf("Hello world!\n"); + sleep(999999999); + } + EOF + gcc -static hello.c -o init + echo init | cpio -o -H newc | gzip > test.cpio.gz + # Testing external initramfs using the initrd loading mechanism. + qemu -kernel /boot/vmlinuz -initrd test.cpio.gz /dev/zero + +When debugging a normal root filesystem, it's nice to be able to boot with +"init=/bin/sh". The initramfs equivalent is "rdinit=/bin/sh", and it's +just as useful. + +Why cpio rather than tar? +------------------------- + +This decision was made back in December, 2001. The discussion started here: + + http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1538.html + +And spawned a second thread (specifically on tar vs cpio), starting here: + + http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1587.html + +The quick and dirty summary version (which is no substitute for reading +the above threads) is: + +1) cpio is a standard. It's decades old (from the AT&T days), and already + widely used on Linux (inside RPM, Red Hat's device driver disks). Here's + a Linux Journal article about it from 1996: + + http://www.linuxjournal.com/article/1213 + + It's not as popular as tar because the traditional cpio command line tools + require _truly_hideous_ command line arguments. But that says nothing + either way about the archive format, and there are alternative tools, + such as: + + http://freshmeat.net/projects/afio/ + +2) The cpio archive format chosen by the kernel is simpler and cleaner (and + thus easier to create and parse) than any of the (literally dozens of) + various tar archive formats. The complete initramfs archive format is + explained in buffer-format.txt, created in usr/gen_init_cpio.c, and + extracted in init/initramfs.c. All three together come to less than 26k + total of human-readable text. + +3) The GNU project standardizing on tar is approximately as relevant as + Windows standardizing on zip. Linux is not part of either, and is free + to make its own technical decisions. + +4) Since this is a kernel internal format, it could easily have been + something brand new. The kernel provides its own tools to create and + extract this format anyway. Using an existing standard was preferable, + but not essential. + +5) Al Viro made the decision (quote: "tar is ugly as hell and not going to be + supported on the kernel side"): + + http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1540.html + + explained his reasoning: + + http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1550.html + http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1638.html + + and, most importantly, designed and implemented the initramfs code. + +Future directions: +------------------ + +Today (2.6.16), initramfs is always compiled in, but not always used. The +kernel falls back to legacy boot code that is reached only if initramfs does +not contain an /init program. The fallback is legacy code, there to ensure a +smooth transition and allowing early boot functionality to gradually move to +"early userspace" (I.E. initramfs). + +The move to early userspace is necessary because finding and mounting the real +root device is complex. Root partitions can span multiple devices (raid or +separate journal). They can be out on the network (requiring dhcp, setting a +specific MAC address, logging into a server, etc). They can live on removable +media, with dynamically allocated major/minor numbers and persistent naming +issues requiring a full udev implementation to sort out. They can be +compressed, encrypted, copy-on-write, loopback mounted, strangely partitioned, +and so on. + +This kind of complexity (which inevitably includes policy) is rightly handled +in userspace. Both klibc and busybox/uClibc are working on simple initramfs +packages to drop into a kernel build. + +The klibc package has now been accepted into Andrew Morton's 2.6.17-mm tree. +The kernel's current early boot code (partition detection, etc) will probably +be migrated into a default initramfs, automatically created and used by the +kernel build. diff --git a/Documentation/filesystems/relay.txt b/Documentation/filesystems/relay.txt new file mode 100644 index 00000000..510b7226 --- /dev/null +++ b/Documentation/filesystems/relay.txt @@ -0,0 +1,494 @@ +relay interface (formerly relayfs) +================================== + +The relay interface provides a means for kernel applications to +efficiently log and transfer large quantities of data from the kernel +to userspace via user-defined 'relay channels'. + +A 'relay channel' is a kernel->user data relay mechanism implemented +as a set of per-cpu kernel buffers ('channel buffers'), each +represented as a regular file ('relay file') in user space. Kernel +clients write into the channel buffers using efficient write +functions; these automatically log into the current cpu's channel +buffer. User space applications mmap() or read() from the relay files +and retrieve the data as it becomes available. The relay files +themselves are files created in a host filesystem, e.g. debugfs, and +are associated with the channel buffers using the API described below. + +The format of the data logged into the channel buffers is completely +up to the kernel client; the relay interface does however provide +hooks which allow kernel clients to impose some structure on the +buffer data. The relay interface doesn't implement any form of data +filtering - this also is left to the kernel client. The purpose is to +keep things as simple as possible. + +This document provides an overview of the relay interface API. The +details of the function parameters are documented along with the +functions in the relay interface code - please see that for details. + +Semantics +========= + +Each relay channel has one buffer per CPU, each buffer has one or more +sub-buffers. Messages are written to the first sub-buffer until it is +too full to contain a new message, in which case it it is written to +the next (if available). Messages are never split across sub-buffers. +At this point, userspace can be notified so it empties the first +sub-buffer, while the kernel continues writing to the next. + +When notified that a sub-buffer is full, the kernel knows how many +bytes of it are padding i.e. unused space occurring because a complete +message couldn't fit into a sub-buffer. Userspace can use this +knowledge to copy only valid data. + +After copying it, userspace can notify the kernel that a sub-buffer +has been consumed. + +A relay channel can operate in a mode where it will overwrite data not +yet collected by userspace, and not wait for it to be consumed. + +The relay channel itself does not provide for communication of such +data between userspace and kernel, allowing the kernel side to remain +simple and not impose a single interface on userspace. It does +provide a set of examples and a separate helper though, described +below. + +The read() interface both removes padding and internally consumes the +read sub-buffers; thus in cases where read(2) is being used to drain +the channel buffers, special-purpose communication between kernel and +user isn't necessary for basic operation. + +One of the major goals of the relay interface is to provide a low +overhead mechanism for conveying kernel data to userspace. While the +read() interface is easy to use, it's not as efficient as the mmap() +approach; the example code attempts to make the tradeoff between the +two approaches as small as possible. + +klog and relay-apps example code +================================ + +The relay interface itself is ready to use, but to make things easier, +a couple simple utility functions and a set of examples are provided. + +The relay-apps example tarball, available on the relay sourceforge +site, contains a set of self-contained examples, each consisting of a +pair of .c files containing boilerplate code for each of the user and +kernel sides of a relay application. When combined these two sets of +boilerplate code provide glue to easily stream data to disk, without +having to bother with mundane housekeeping chores. + +The 'klog debugging functions' patch (klog.patch in the relay-apps +tarball) provides a couple of high-level logging functions to the +kernel which allow writing formatted text or raw data to a channel, +regardless of whether a channel to write into exists or not, or even +whether the relay interface is compiled into the kernel or not. These +functions allow you to put unconditional 'trace' statements anywhere +in the kernel or kernel modules; only when there is a 'klog handler' +registered will data actually be logged (see the klog and kleak +examples for details). + +It is of course possible to use the relay interface from scratch, +i.e. without using any of the relay-apps example code or klog, but +you'll have to implement communication between userspace and kernel, +allowing both to convey the state of buffers (full, empty, amount of +padding). The read() interface both removes padding and internally +consumes the read sub-buffers; thus in cases where read(2) is being +used to drain the channel buffers, special-purpose communication +between kernel and user isn't necessary for basic operation. Things +such as buffer-full conditions would still need to be communicated via +some channel though. + +klog and the relay-apps examples can be found in the relay-apps +tarball on http://relayfs.sourceforge.net + +The relay interface user space API +================================== + +The relay interface implements basic file operations for user space +access to relay channel buffer data. Here are the file operations +that are available and some comments regarding their behavior: + +open() enables user to open an _existing_ channel buffer. + +mmap() results in channel buffer being mapped into the caller's + memory space. Note that you can't do a partial mmap - you + must map the entire file, which is NRBUF * SUBBUFSIZE. + +read() read the contents of a channel buffer. The bytes read are + 'consumed' by the reader, i.e. they won't be available + again to subsequent reads. If the channel is being used + in no-overwrite mode (the default), it can be read at any + time even if there's an active kernel writer. If the + channel is being used in overwrite mode and there are + active channel writers, results may be unpredictable - + users should make sure that all logging to the channel has + ended before using read() with overwrite mode. Sub-buffer + padding is automatically removed and will not be seen by + the reader. + +sendfile() transfer data from a channel buffer to an output file + descriptor. Sub-buffer padding is automatically removed + and will not be seen by the reader. + +poll() POLLIN/POLLRDNORM/POLLERR supported. User applications are + notified when sub-buffer boundaries are crossed. + +close() decrements the channel buffer's refcount. When the refcount + reaches 0, i.e. when no process or kernel client has the + buffer open, the channel buffer is freed. + +In order for a user application to make use of relay files, the +host filesystem must be mounted. For example, + + mount -t debugfs debugfs /sys/kernel/debug + +NOTE: the host filesystem doesn't need to be mounted for kernel + clients to create or use channels - it only needs to be + mounted when user space applications need access to the buffer + data. + + +The relay interface kernel API +============================== + +Here's a summary of the API the relay interface provides to in-kernel clients: + +TBD(curr. line MT:/API/) + channel management functions: + + relay_open(base_filename, parent, subbuf_size, n_subbufs, + callbacks, private_data) + relay_close(chan) + relay_flush(chan) + relay_reset(chan) + + channel management typically called on instigation of userspace: + + relay_subbufs_consumed(chan, cpu, subbufs_consumed) + + write functions: + + relay_write(chan, data, length) + __relay_write(chan, data, length) + relay_reserve(chan, length) + + callbacks: + + subbuf_start(buf, subbuf, prev_subbuf, prev_padding) + buf_mapped(buf, filp) + buf_unmapped(buf, filp) + create_buf_file(filename, parent, mode, buf, is_global) + remove_buf_file(dentry) + + helper functions: + + relay_buf_full(buf) + subbuf_start_reserve(buf, length) + + +Creating a channel +------------------ + +relay_open() is used to create a channel, along with its per-cpu +channel buffers. Each channel buffer will have an associated file +created for it in the host filesystem, which can be and mmapped or +read from in user space. The files are named basename0...basenameN-1 +where N is the number of online cpus, and by default will be created +in the root of the filesystem (if the parent param is NULL). If you +want a directory structure to contain your relay files, you should +create it using the host filesystem's directory creation function, +e.g. debugfs_create_dir(), and pass the parent directory to +relay_open(). Users are responsible for cleaning up any directory +structure they create, when the channel is closed - again the host +filesystem's directory removal functions should be used for that, +e.g. debugfs_remove(). + +In order for a channel to be created and the host filesystem's files +associated with its channel buffers, the user must provide definitions +for two callback functions, create_buf_file() and remove_buf_file(). +create_buf_file() is called once for each per-cpu buffer from +relay_open() and allows the user to create the file which will be used +to represent the corresponding channel buffer. The callback should +return the dentry of the file created to represent the channel buffer. +remove_buf_file() must also be defined; it's responsible for deleting +the file(s) created in create_buf_file() and is called during +relay_close(). + +Here are some typical definitions for these callbacks, in this case +using debugfs: + +/* + * create_buf_file() callback. Creates relay file in debugfs. + */ +static struct dentry *create_buf_file_handler(const char *filename, + struct dentry *parent, + int mode, + struct rchan_buf *buf, + int *is_global) +{ + return debugfs_create_file(filename, mode, parent, buf, + &relay_file_operations); +} + +/* + * remove_buf_file() callback. Removes relay file from debugfs. + */ +static int remove_buf_file_handler(struct dentry *dentry) +{ + debugfs_remove(dentry); + + return 0; +} + +/* + * relay interface callbacks + */ +static struct rchan_callbacks relay_callbacks = +{ + .create_buf_file = create_buf_file_handler, + .remove_buf_file = remove_buf_file_handler, +}; + +And an example relay_open() invocation using them: + + chan = relay_open("cpu", NULL, SUBBUF_SIZE, N_SUBBUFS, &relay_callbacks, NULL); + +If the create_buf_file() callback fails, or isn't defined, channel +creation and thus relay_open() will fail. + +The total size of each per-cpu buffer is calculated by multiplying the +number of sub-buffers by the sub-buffer size passed into relay_open(). +The idea behind sub-buffers is that they're basically an extension of +double-buffering to N buffers, and they also allow applications to +easily implement random-access-on-buffer-boundary schemes, which can +be important for some high-volume applications. The number and size +of sub-buffers is completely dependent on the application and even for +the same application, different conditions will warrant different +values for these parameters at different times. Typically, the right +values to use are best decided after some experimentation; in general, +though, it's safe to assume that having only 1 sub-buffer is a bad +idea - you're guaranteed to either overwrite data or lose events +depending on the channel mode being used. + +The create_buf_file() implementation can also be defined in such a way +as to allow the creation of a single 'global' buffer instead of the +default per-cpu set. This can be useful for applications interested +mainly in seeing the relative ordering of system-wide events without +the need to bother with saving explicit timestamps for the purpose of +merging/sorting per-cpu files in a postprocessing step. + +To have relay_open() create a global buffer, the create_buf_file() +implementation should set the value of the is_global outparam to a +non-zero value in addition to creating the file that will be used to +represent the single buffer. In the case of a global buffer, +create_buf_file() and remove_buf_file() will be called only once. The +normal channel-writing functions, e.g. relay_write(), can still be +used - writes from any cpu will transparently end up in the global +buffer - but since it is a global buffer, callers should make sure +they use the proper locking for such a buffer, either by wrapping +writes in a spinlock, or by copying a write function from relay.h and +creating a local version that internally does the proper locking. + +The private_data passed into relay_open() allows clients to associate +user-defined data with a channel, and is immediately available +(including in create_buf_file()) via chan->private_data or +buf->chan->private_data. + +Buffer-only channels +-------------------- + +These channels have no files associated and can be created with +relay_open(NULL, NULL, ...). Such channels are useful in scenarios such +as when doing early tracing in the kernel, before the VFS is up. In these +cases, one may open a buffer-only channel and then call +relay_late_setup_files() when the kernel is ready to handle files, +to expose the buffered data to the userspace. + +Channel 'modes' +--------------- + +relay channels can be used in either of two modes - 'overwrite' or +'no-overwrite'. The mode is entirely determined by the implementation +of the subbuf_start() callback, as described below. The default if no +subbuf_start() callback is defined is 'no-overwrite' mode. If the +default mode suits your needs, and you plan to use the read() +interface to retrieve channel data, you can ignore the details of this +section, as it pertains mainly to mmap() implementations. + +In 'overwrite' mode, also known as 'flight recorder' mode, writes +continuously cycle around the buffer and will never fail, but will +unconditionally overwrite old data regardless of whether it's actually +been consumed. In no-overwrite mode, writes will fail, i.e. data will +be lost, if the number of unconsumed sub-buffers equals the total +number of sub-buffers in the channel. It should be clear that if +there is no consumer or if the consumer can't consume sub-buffers fast +enough, data will be lost in either case; the only difference is +whether data is lost from the beginning or the end of a buffer. + +As explained above, a relay channel is made of up one or more +per-cpu channel buffers, each implemented as a circular buffer +subdivided into one or more sub-buffers. Messages are written into +the current sub-buffer of the channel's current per-cpu buffer via the +write functions described below. Whenever a message can't fit into +the current sub-buffer, because there's no room left for it, the +client is notified via the subbuf_start() callback that a switch to a +new sub-buffer is about to occur. The client uses this callback to 1) +initialize the next sub-buffer if appropriate 2) finalize the previous +sub-buffer if appropriate and 3) return a boolean value indicating +whether or not to actually move on to the next sub-buffer. + +To implement 'no-overwrite' mode, the userspace client would provide +an implementation of the subbuf_start() callback something like the +following: + +static int subbuf_start(struct rchan_buf *buf, + void *subbuf, + void *prev_subbuf, + unsigned int prev_padding) +{ + if (prev_subbuf) + *((unsigned *)prev_subbuf) = prev_padding; + + if (relay_buf_full(buf)) + return 0; + + subbuf_start_reserve(buf, sizeof(unsigned int)); + + return 1; +} + +If the current buffer is full, i.e. all sub-buffers remain unconsumed, +the callback returns 0 to indicate that the buffer switch should not +occur yet, i.e. until the consumer has had a chance to read the +current set of ready sub-buffers. For the relay_buf_full() function +to make sense, the consumer is responsible for notifying the relay +interface when sub-buffers have been consumed via +relay_subbufs_consumed(). Any subsequent attempts to write into the +buffer will again invoke the subbuf_start() callback with the same +parameters; only when the consumer has consumed one or more of the +ready sub-buffers will relay_buf_full() return 0, in which case the +buffer switch can continue. + +The implementation of the subbuf_start() callback for 'overwrite' mode +would be very similar: + +static int subbuf_start(struct rchan_buf *buf, + void *subbuf, + void *prev_subbuf, + unsigned int prev_padding) +{ + if (prev_subbuf) + *((unsigned *)prev_subbuf) = prev_padding; + + subbuf_start_reserve(buf, sizeof(unsigned int)); + + return 1; +} + +In this case, the relay_buf_full() check is meaningless and the +callback always returns 1, causing the buffer switch to occur +unconditionally. It's also meaningless for the client to use the +relay_subbufs_consumed() function in this mode, as it's never +consulted. + +The default subbuf_start() implementation, used if the client doesn't +define any callbacks, or doesn't define the subbuf_start() callback, +implements the simplest possible 'no-overwrite' mode, i.e. it does +nothing but return 0. + +Header information can be reserved at the beginning of each sub-buffer +by calling the subbuf_start_reserve() helper function from within the +subbuf_start() callback. This reserved area can be used to store +whatever information the client wants. In the example above, room is +reserved in each sub-buffer to store the padding count for that +sub-buffer. This is filled in for the previous sub-buffer in the +subbuf_start() implementation; the padding value for the previous +sub-buffer is passed into the subbuf_start() callback along with a +pointer to the previous sub-buffer, since the padding value isn't +known until a sub-buffer is filled. The subbuf_start() callback is +also called for the first sub-buffer when the channel is opened, to +give the client a chance to reserve space in it. In this case the +previous sub-buffer pointer passed into the callback will be NULL, so +the client should check the value of the prev_subbuf pointer before +writing into the previous sub-buffer. + +Writing to a channel +-------------------- + +Kernel clients write data into the current cpu's channel buffer using +relay_write() or __relay_write(). relay_write() is the main logging +function - it uses local_irqsave() to protect the buffer and should be +used if you might be logging from interrupt context. If you know +you'll never be logging from interrupt context, you can use +__relay_write(), which only disables preemption. These functions +don't return a value, so you can't determine whether or not they +failed - the assumption is that you wouldn't want to check a return +value in the fast logging path anyway, and that they'll always succeed +unless the buffer is full and no-overwrite mode is being used, in +which case you can detect a failed write in the subbuf_start() +callback by calling the relay_buf_full() helper function. + +relay_reserve() is used to reserve a slot in a channel buffer which +can be written to later. This would typically be used in applications +that need to write directly into a channel buffer without having to +stage data in a temporary buffer beforehand. Because the actual write +may not happen immediately after the slot is reserved, applications +using relay_reserve() can keep a count of the number of bytes actually +written, either in space reserved in the sub-buffers themselves or as +a separate array. See the 'reserve' example in the relay-apps tarball +at http://relayfs.sourceforge.net for an example of how this can be +done. Because the write is under control of the client and is +separated from the reserve, relay_reserve() doesn't protect the buffer +at all - it's up to the client to provide the appropriate +synchronization when using relay_reserve(). + +Closing a channel +----------------- + +The client calls relay_close() when it's finished using the channel. +The channel and its associated buffers are destroyed when there are no +longer any references to any of the channel buffers. relay_flush() +forces a sub-buffer switch on all the channel buffers, and can be used +to finalize and process the last sub-buffers before the channel is +closed. + +Misc +---- + +Some applications may want to keep a channel around and re-use it +rather than open and close a new channel for each use. relay_reset() +can be used for this purpose - it resets a channel to its initial +state without reallocating channel buffer memory or destroying +existing mappings. It should however only be called when it's safe to +do so, i.e. when the channel isn't currently being written to. + +Finally, there are a couple of utility callbacks that can be used for +different purposes. buf_mapped() is called whenever a channel buffer +is mmapped from user space and buf_unmapped() is called when it's +unmapped. The client can use this notification to trigger actions +within the kernel application, such as enabling/disabling logging to +the channel. + + +Resources +========= + +For news, example code, mailing list, etc. see the relay interface homepage: + + http://relayfs.sourceforge.net + + +Credits +======= + +The ideas and specs for the relay interface came about as a result of +discussions on tracing involving the following: + +Michel Dagenais <michel.dagenais@polymtl.ca> +Richard Moore <richardj_moore@uk.ibm.com> +Bob Wisniewski <bob@watson.ibm.com> +Karim Yaghmour <karim@opersys.com> +Tom Zanussi <zanussi@us.ibm.com> + +Also thanks to Hubertus Franke for a lot of useful suggestions and bug +reports. diff --git a/Documentation/filesystems/romfs.txt b/Documentation/filesystems/romfs.txt new file mode 100644 index 00000000..e2b07cc9 --- /dev/null +++ b/Documentation/filesystems/romfs.txt @@ -0,0 +1,186 @@ +ROMFS - ROM FILE SYSTEM + +This is a quite dumb, read only filesystem, mainly for initial RAM +disks of installation disks. It has grown up by the need of having +modules linked at boot time. Using this filesystem, you get a very +similar feature, and even the possibility of a small kernel, with a +file system which doesn't take up useful memory from the router +functions in the basement of your office. + +For comparison, both the older minix and xiafs (the latter is now +defunct) filesystems, compiled as module need more than 20000 bytes, +while romfs is less than a page, about 4000 bytes (assuming i586 +code). Under the same conditions, the msdos filesystem would need +about 30K (and does not support device nodes or symlinks), while the +nfs module with nfsroot is about 57K. Furthermore, as a bit unfair +comparison, an actual rescue disk used up 3202 blocks with ext2, while +with romfs, it needed 3079 blocks. + +To create such a file system, you'll need a user program named +genromfs. It is available on http://romfs.sourceforge.net/ + +As the name suggests, romfs could be also used (space-efficiently) on +various read-only media, like (E)EPROM disks if someone will have the +motivation.. :) + +However, the main purpose of romfs is to have a very small kernel, +which has only this filesystem linked in, and then can load any module +later, with the current module utilities. It can also be used to run +some program to decide if you need SCSI devices, and even IDE or +floppy drives can be loaded later if you use the "initrd"--initial +RAM disk--feature of the kernel. This would not be really news +flash, but with romfs, you can even spare off your ext2 or minix or +maybe even affs filesystem until you really know that you need it. + +For example, a distribution boot disk can contain only the cd disk +drivers (and possibly the SCSI drivers), and the ISO 9660 filesystem +module. The kernel can be small enough, since it doesn't have other +filesystems, like the quite large ext2fs module, which can then be +loaded off the CD at a later stage of the installation. Another use +would be for a recovery disk, when you are reinstalling a workstation +from the network, and you will have all the tools/modules available +from a nearby server, so you don't want to carry two disks for this +purpose, just because it won't fit into ext2. + +romfs operates on block devices as you can expect, and the underlying +structure is very simple. Every accessible structure begins on 16 +byte boundaries for fast access. The minimum space a file will take +is 32 bytes (this is an empty file, with a less than 16 character +name). The maximum overhead for any non-empty file is the header, and +the 16 byte padding for the name and the contents, also 16+14+15 = 45 +bytes. This is quite rare however, since most file names are longer +than 3 bytes, and shorter than 15 bytes. + +The layout of the filesystem is the following: + +offset content + + +---+---+---+---+ + 0 | - | r | o | m | \ + +---+---+---+---+ The ASCII representation of those bytes + 4 | 1 | f | s | - | / (i.e. "-rom1fs-") + +---+---+---+---+ + 8 | full size | The number of accessible bytes in this fs. + +---+---+---+---+ + 12 | checksum | The checksum of the FIRST 512 BYTES. + +---+---+---+---+ + 16 | volume name | The zero terminated name of the volume, + : : padded to 16 byte boundary. + +---+---+---+---+ + xx | file | + : headers : + +Every multi byte value (32 bit words, I'll use the longwords term from +now on) must be in big endian order. + +The first eight bytes identify the filesystem, even for the casual +inspector. After that, in the 3rd longword, it contains the number of +bytes accessible from the start of this filesystem. The 4th longword +is the checksum of the first 512 bytes (or the number of bytes +accessible, whichever is smaller). The applied algorithm is the same +as in the AFFS filesystem, namely a simple sum of the longwords +(assuming bigendian quantities again). For details, please consult +the source. This algorithm was chosen because although it's not quite +reliable, it does not require any tables, and it is very simple. + +The following bytes are now part of the file system; each file header +must begin on a 16 byte boundary. + +offset content + + +---+---+---+---+ + 0 | next filehdr|X| The offset of the next file header + +---+---+---+---+ (zero if no more files) + 4 | spec.info | Info for directories/hard links/devices + +---+---+---+---+ + 8 | size | The size of this file in bytes + +---+---+---+---+ + 12 | checksum | Covering the meta data, including the file + +---+---+---+---+ name, and padding + 16 | file name | The zero terminated name of the file, + : : padded to 16 byte boundary + +---+---+---+---+ + xx | file data | + : : + +Since the file headers begin always at a 16 byte boundary, the lowest +4 bits would be always zero in the next filehdr pointer. These four +bits are used for the mode information. Bits 0..2 specify the type of +the file; while bit 4 shows if the file is executable or not. The +permissions are assumed to be world readable, if this bit is not set, +and world executable if it is; except the character and block devices, +they are never accessible for other than owner. The owner of every +file is user and group 0, this should never be a problem for the +intended use. The mapping of the 8 possible values to file types is +the following: + + mapping spec.info means + 0 hard link link destination [file header] + 1 directory first file's header + 2 regular file unused, must be zero [MBZ] + 3 symbolic link unused, MBZ (file data is the link content) + 4 block device 16/16 bits major/minor number + 5 char device - " - + 6 socket unused, MBZ + 7 fifo unused, MBZ + +Note that hard links are specifically marked in this filesystem, but +they will behave as you can expect (i.e. share the inode number). +Note also that it is your responsibility to not create hard link +loops, and creating all the . and .. links for directories. This is +normally done correctly by the genromfs program. Please refrain from +using the executable bits for special purposes on the socket and fifo +special files, they may have other uses in the future. Additionally, +please remember that only regular files, and symlinks are supposed to +have a nonzero size field; they contain the number of bytes available +directly after the (padded) file name. + +Another thing to note is that romfs works on file headers and data +aligned to 16 byte boundaries, but most hardware devices and the block +device drivers are unable to cope with smaller than block-sized data. +To overcome this limitation, the whole size of the file system must be +padded to an 1024 byte boundary. + +If you have any problems or suggestions concerning this file system, +please contact me. However, think twice before wanting me to add +features and code, because the primary and most important advantage of +this file system is the small code. On the other hand, don't be +alarmed, I'm not getting that much romfs related mail. Now I can +understand why Avery wrote poems in the ARCnet docs to get some more +feedback. :) + +romfs has also a mailing list, and to date, it hasn't received any +traffic, so you are welcome to join it to discuss your ideas. :) + +It's run by ezmlm, so you can subscribe to it by sending a message +to romfs-subscribe@shadow.banki.hu, the content is irrelevant. + +Pending issues: + +- Permissions and owner information are pretty essential features of a +Un*x like system, but romfs does not provide the full possibilities. +I have never found this limiting, but others might. + +- The file system is read only, so it can be very small, but in case +one would want to write _anything_ to a file system, he still needs +a writable file system, thus negating the size advantages. Possible +solutions: implement write access as a compile-time option, or a new, +similarly small writable filesystem for RAM disks. + +- Since the files are only required to have alignment on a 16 byte +boundary, it is currently possibly suboptimal to read or execute files +from the filesystem. It might be resolved by reordering file data to +have most of it (i.e. except the start and the end) laying at "natural" +boundaries, thus it would be possible to directly map a big portion of +the file contents to the mm subsystem. + +- Compression might be an useful feature, but memory is quite a +limiting factor in my eyes. + +- Where it is used? + +- Does it work on other architectures than intel and motorola? + + +Have fun, +Janos Farkas <chexum@shadow.banki.hu> diff --git a/Documentation/filesystems/seq_file.txt b/Documentation/filesystems/seq_file.txt new file mode 100644 index 00000000..a1e2e0dd --- /dev/null +++ b/Documentation/filesystems/seq_file.txt @@ -0,0 +1,292 @@ +The seq_file interface + + Copyright 2003 Jonathan Corbet <corbet@lwn.net> + This file is originally from the LWN.net Driver Porting series at + http://lwn.net/Articles/driver-porting/ + + +There are numerous ways for a device driver (or other kernel component) to +provide information to the user or system administrator. One useful +technique is the creation of virtual files, in debugfs, /proc or elsewhere. +Virtual files can provide human-readable output that is easy to get at +without any special utility programs; they can also make life easier for +script writers. It is not surprising that the use of virtual files has +grown over the years. + +Creating those files correctly has always been a bit of a challenge, +however. It is not that hard to make a virtual file which returns a +string. But life gets trickier if the output is long - anything greater +than an application is likely to read in a single operation. Handling +multiple reads (and seeks) requires careful attention to the reader's +position within the virtual file - that position is, likely as not, in the +middle of a line of output. The kernel has traditionally had a number of +implementations that got this wrong. + +The 2.6 kernel contains a set of functions (implemented by Alexander Viro) +which are designed to make it easy for virtual file creators to get it +right. + +The seq_file interface is available via <linux/seq_file.h>. There are +three aspects to seq_file: + + * An iterator interface which lets a virtual file implementation + step through the objects it is presenting. + + * Some utility functions for formatting objects for output without + needing to worry about things like output buffers. + + * A set of canned file_operations which implement most operations on + the virtual file. + +We'll look at the seq_file interface via an extremely simple example: a +loadable module which creates a file called /proc/sequence. The file, when +read, simply produces a set of increasing integer values, one per line. The +sequence will continue until the user loses patience and finds something +better to do. The file is seekable, in that one can do something like the +following: + + dd if=/proc/sequence of=out1 count=1 + dd if=/proc/sequence skip=1 of=out2 count=1 + +Then concatenate the output files out1 and out2 and get the right +result. Yes, it is a thoroughly useless module, but the point is to show +how the mechanism works without getting lost in other details. (Those +wanting to see the full source for this module can find it at +http://lwn.net/Articles/22359/). + + +The iterator interface + +Modules implementing a virtual file with seq_file must implement a simple +iterator object that allows stepping through the data of interest. +Iterators must be able to move to a specific position - like the file they +implement - but the interpretation of that position is up to the iterator +itself. A seq_file implementation that is formatting firewall rules, for +example, could interpret position N as the Nth rule in the chain. +Positioning can thus be done in whatever way makes the most sense for the +generator of the data, which need not be aware of how a position translates +to an offset in the virtual file. The one obvious exception is that a +position of zero should indicate the beginning of the file. + +The /proc/sequence iterator just uses the count of the next number it +will output as its position. + +Four functions must be implemented to make the iterator work. The first, +called start() takes a position as an argument and returns an iterator +which will start reading at that position. For our simple sequence example, +the start() function looks like: + + static void *ct_seq_start(struct seq_file *s, loff_t *pos) + { + loff_t *spos = kmalloc(sizeof(loff_t), GFP_KERNEL); + if (! spos) + return NULL; + *spos = *pos; + return spos; + } + +The entire data structure for this iterator is a single loff_t value +holding the current position. There is no upper bound for the sequence +iterator, but that will not be the case for most other seq_file +implementations; in most cases the start() function should check for a +"past end of file" condition and return NULL if need be. + +For more complicated applications, the private field of the seq_file +structure can be used. There is also a special value which can be returned +by the start() function called SEQ_START_TOKEN; it can be used if you wish +to instruct your show() function (described below) to print a header at the +top of the output. SEQ_START_TOKEN should only be used if the offset is +zero, however. + +The next function to implement is called, amazingly, next(); its job is to +move the iterator forward to the next position in the sequence. The +example module can simply increment the position by one; more useful +modules will do what is needed to step through some data structure. The +next() function returns a new iterator, or NULL if the sequence is +complete. Here's the example version: + + static void *ct_seq_next(struct seq_file *s, void *v, loff_t *pos) + { + loff_t *spos = v; + *pos = ++*spos; + return spos; + } + +The stop() function is called when iteration is complete; its job, of +course, is to clean up. If dynamic memory is allocated for the iterator, +stop() is the place to free it. + + static void ct_seq_stop(struct seq_file *s, void *v) + { + kfree(v); + } + +Finally, the show() function should format the object currently pointed to +by the iterator for output. The example module's show() function is: + + static int ct_seq_show(struct seq_file *s, void *v) + { + loff_t *spos = v; + seq_printf(s, "%lld\n", (long long)*spos); + return 0; + } + +If all is well, the show() function should return zero. A negative error +code in the usual manner indicates that something went wrong; it will be +passed back to user space. This function can also return SEQ_SKIP, which +causes the current item to be skipped; if the show() function has already +generated output before returning SEQ_SKIP, that output will be dropped. + +We will look at seq_printf() in a moment. But first, the definition of the +seq_file iterator is finished by creating a seq_operations structure with +the four functions we have just defined: + + static const struct seq_operations ct_seq_ops = { + .start = ct_seq_start, + .next = ct_seq_next, + .stop = ct_seq_stop, + .show = ct_seq_show + }; + +This structure will be needed to tie our iterator to the /proc file in +a little bit. + +It's worth noting that the iterator value returned by start() and +manipulated by the other functions is considered to be completely opaque by +the seq_file code. It can thus be anything that is useful in stepping +through the data to be output. Counters can be useful, but it could also be +a direct pointer into an array or linked list. Anything goes, as long as +the programmer is aware that things can happen between calls to the +iterator function. However, the seq_file code (by design) will not sleep +between the calls to start() and stop(), so holding a lock during that time +is a reasonable thing to do. The seq_file code will also avoid taking any +other locks while the iterator is active. + + +Formatted output + +The seq_file code manages positioning within the output created by the +iterator and getting it into the user's buffer. But, for that to work, that +output must be passed to the seq_file code. Some utility functions have +been defined which make this task easy. + +Most code will simply use seq_printf(), which works pretty much like +printk(), but which requires the seq_file pointer as an argument. It is +common to ignore the return value from seq_printf(), but a function +producing complicated output may want to check that value and quit if +something non-zero is returned; an error return means that the seq_file +buffer has been filled and further output will be discarded. + +For straight character output, the following functions may be used: + + int seq_putc(struct seq_file *m, char c); + int seq_puts(struct seq_file *m, const char *s); + int seq_escape(struct seq_file *m, const char *s, const char *esc); + +The first two output a single character and a string, just like one would +expect. seq_escape() is like seq_puts(), except that any character in s +which is in the string esc will be represented in octal form in the output. + +There is also a pair of functions for printing filenames: + + int seq_path(struct seq_file *m, struct path *path, char *esc); + int seq_path_root(struct seq_file *m, struct path *path, + struct path *root, char *esc) + +Here, path indicates the file of interest, and esc is a set of characters +which should be escaped in the output. A call to seq_path() will output +the path relative to the current process's filesystem root. If a different +root is desired, it can be used with seq_path_root(). Note that, if it +turns out that path cannot be reached from root, the value of root will be +changed in seq_file_root() to a root which *does* work. + + +Making it all work + +So far, we have a nice set of functions which can produce output within the +seq_file system, but we have not yet turned them into a file that a user +can see. Creating a file within the kernel requires, of course, the +creation of a set of file_operations which implement the operations on that +file. The seq_file interface provides a set of canned operations which do +most of the work. The virtual file author still must implement the open() +method, however, to hook everything up. The open function is often a single +line, as in the example module: + + static int ct_open(struct inode *inode, struct file *file) + { + return seq_open(file, &ct_seq_ops); + } + +Here, the call to seq_open() takes the seq_operations structure we created +before, and gets set up to iterate through the virtual file. + +On a successful open, seq_open() stores the struct seq_file pointer in +file->private_data. If you have an application where the same iterator can +be used for more than one file, you can store an arbitrary pointer in the +private field of the seq_file structure; that value can then be retrieved +by the iterator functions. + +The other operations of interest - read(), llseek(), and release() - are +all implemented by the seq_file code itself. So a virtual file's +file_operations structure will look like: + + static const struct file_operations ct_file_ops = { + .owner = THIS_MODULE, + .open = ct_open, + .read = seq_read, + .llseek = seq_lseek, + .release = seq_release + }; + +There is also a seq_release_private() which passes the contents of the +seq_file private field to kfree() before releasing the structure. + +The final step is the creation of the /proc file itself. In the example +code, that is done in the initialization code in the usual way: + + static int ct_init(void) + { + struct proc_dir_entry *entry; + + proc_create("sequence", 0, NULL, &ct_file_ops); + return 0; + } + + module_init(ct_init); + +And that is pretty much it. + + +seq_list + +If your file will be iterating through a linked list, you may find these +routines useful: + + struct list_head *seq_list_start(struct list_head *head, + loff_t pos); + struct list_head *seq_list_start_head(struct list_head *head, + loff_t pos); + struct list_head *seq_list_next(void *v, struct list_head *head, + loff_t *ppos); + +These helpers will interpret pos as a position within the list and iterate +accordingly. Your start() and next() functions need only invoke the +seq_list_* helpers with a pointer to the appropriate list_head structure. + + +The extra-simple version + +For extremely simple virtual files, there is an even easier interface. A +module can define only the show() function, which should create all the +output that the virtual file will contain. The file's open() method then +calls: + + int single_open(struct file *file, + int (*show)(struct seq_file *m, void *p), + void *data); + +When output time comes, the show() function will be called once. The data +value given to single_open() can be found in the private field of the +seq_file structure. When using single_open(), the programmer should use +single_release() instead of seq_release() in the file_operations structure +to avoid a memory leak. diff --git a/Documentation/filesystems/sharedsubtree.txt b/Documentation/filesystems/sharedsubtree.txt new file mode 100644 index 00000000..4ede421c --- /dev/null +++ b/Documentation/filesystems/sharedsubtree.txt @@ -0,0 +1,939 @@ +Shared Subtrees +--------------- + +Contents: + 1) Overview + 2) Features + 3) Setting mount states + 4) Use-case + 5) Detailed semantics + 6) Quiz + 7) FAQ + 8) Implementation + + +1) Overview +----------- + +Consider the following situation: + +A process wants to clone its own namespace, but still wants to access the CD +that got mounted recently. Shared subtree semantics provide the necessary +mechanism to accomplish the above. + +It provides the necessary building blocks for features like per-user-namespace +and versioned filesystem. + +2) Features +----------- + +Shared subtree provides four different flavors of mounts; struct vfsmount to be +precise + + a. shared mount + b. slave mount + c. private mount + d. unbindable mount + + +2a) A shared mount can be replicated to as many mountpoints and all the +replicas continue to be exactly same. + + Here is an example: + + Let's say /mnt has a mount that is shared. + mount --make-shared /mnt + + Note: mount(8) command now supports the --make-shared flag, + so the sample 'smount' program is no longer needed and has been + removed. + + # mount --bind /mnt /tmp + The above command replicates the mount at /mnt to the mountpoint /tmp + and the contents of both the mounts remain identical. + + #ls /mnt + a b c + + #ls /tmp + a b c + + Now let's say we mount a device at /tmp/a + # mount /dev/sd0 /tmp/a + + #ls /tmp/a + t1 t2 t3 + + #ls /mnt/a + t1 t2 t3 + + Note that the mount has propagated to the mount at /mnt as well. + + And the same is true even when /dev/sd0 is mounted on /mnt/a. The + contents will be visible under /tmp/a too. + + +2b) A slave mount is like a shared mount except that mount and umount events + only propagate towards it. + + All slave mounts have a master mount which is a shared. + + Here is an example: + + Let's say /mnt has a mount which is shared. + # mount --make-shared /mnt + + Let's bind mount /mnt to /tmp + # mount --bind /mnt /tmp + + the new mount at /tmp becomes a shared mount and it is a replica of + the mount at /mnt. + + Now let's make the mount at /tmp; a slave of /mnt + # mount --make-slave /tmp + + let's mount /dev/sd0 on /mnt/a + # mount /dev/sd0 /mnt/a + + #ls /mnt/a + t1 t2 t3 + + #ls /tmp/a + t1 t2 t3 + + Note the mount event has propagated to the mount at /tmp + + However let's see what happens if we mount something on the mount at /tmp + + # mount /dev/sd1 /tmp/b + + #ls /tmp/b + s1 s2 s3 + + #ls /mnt/b + + Note how the mount event has not propagated to the mount at + /mnt + + +2c) A private mount does not forward or receive propagation. + + This is the mount we are familiar with. Its the default type. + + +2d) A unbindable mount is a unbindable private mount + + let's say we have a mount at /mnt and we make is unbindable + + # mount --make-unbindable /mnt + + Let's try to bind mount this mount somewhere else. + # mount --bind /mnt /tmp + mount: wrong fs type, bad option, bad superblock on /mnt, + or too many mounted file systems + + Binding a unbindable mount is a invalid operation. + + +3) Setting mount states + + The mount command (util-linux package) can be used to set mount + states: + + mount --make-shared mountpoint + mount --make-slave mountpoint + mount --make-private mountpoint + mount --make-unbindable mountpoint + + +4) Use cases +------------ + + A) A process wants to clone its own namespace, but still wants to + access the CD that got mounted recently. + + Solution: + + The system administrator can make the mount at /cdrom shared + mount --bind /cdrom /cdrom + mount --make-shared /cdrom + + Now any process that clones off a new namespace will have a + mount at /cdrom which is a replica of the same mount in the + parent namespace. + + So when a CD is inserted and mounted at /cdrom that mount gets + propagated to the other mount at /cdrom in all the other clone + namespaces. + + B) A process wants its mounts invisible to any other process, but + still be able to see the other system mounts. + + Solution: + + To begin with, the administrator can mark the entire mount tree + as shareable. + + mount --make-rshared / + + A new process can clone off a new namespace. And mark some part + of its namespace as slave + + mount --make-rslave /myprivatetree + + Hence forth any mounts within the /myprivatetree done by the + process will not show up in any other namespace. However mounts + done in the parent namespace under /myprivatetree still shows + up in the process's namespace. + + + Apart from the above semantics this feature provides the + building blocks to solve the following problems: + + C) Per-user namespace + + The above semantics allows a way to share mounts across + namespaces. But namespaces are associated with processes. If + namespaces are made first class objects with user API to + associate/disassociate a namespace with userid, then each user + could have his/her own namespace and tailor it to his/her + requirements. Offcourse its needs support from PAM. + + D) Versioned files + + If the entire mount tree is visible at multiple locations, then + a underlying versioning file system can return different + version of the file depending on the path used to access that + file. + + An example is: + + mount --make-shared / + mount --rbind / /view/v1 + mount --rbind / /view/v2 + mount --rbind / /view/v3 + mount --rbind / /view/v4 + + and if /usr has a versioning filesystem mounted, then that + mount appears at /view/v1/usr, /view/v2/usr, /view/v3/usr and + /view/v4/usr too + + A user can request v3 version of the file /usr/fs/namespace.c + by accessing /view/v3/usr/fs/namespace.c . The underlying + versioning filesystem can then decipher that v3 version of the + filesystem is being requested and return the corresponding + inode. + +5) Detailed semantics: +------------------- + The section below explains the detailed semantics of + bind, rbind, move, mount, umount and clone-namespace operations. + + Note: the word 'vfsmount' and the noun 'mount' have been used + to mean the same thing, throughout this document. + +5a) Mount states + + A given mount can be in one of the following states + 1) shared + 2) slave + 3) shared and slave + 4) private + 5) unbindable + + A 'propagation event' is defined as event generated on a vfsmount + that leads to mount or unmount actions in other vfsmounts. + + A 'peer group' is defined as a group of vfsmounts that propagate + events to each other. + + (1) Shared mounts + + A 'shared mount' is defined as a vfsmount that belongs to a + 'peer group'. + + For example: + mount --make-shared /mnt + mount --bind /mnt /tmp + + The mount at /mnt and that at /tmp are both shared and belong + to the same peer group. Anything mounted or unmounted under + /mnt or /tmp reflect in all the other mounts of its peer + group. + + + (2) Slave mounts + + A 'slave mount' is defined as a vfsmount that receives + propagation events and does not forward propagation events. + + A slave mount as the name implies has a master mount from which + mount/unmount events are received. Events do not propagate from + the slave mount to the master. Only a shared mount can be made + a slave by executing the following command + + mount --make-slave mount + + A shared mount that is made as a slave is no more shared unless + modified to become shared. + + (3) Shared and Slave + + A vfsmount can be both shared as well as slave. This state + indicates that the mount is a slave of some vfsmount, and + has its own peer group too. This vfsmount receives propagation + events from its master vfsmount, and also forwards propagation + events to its 'peer group' and to its slave vfsmounts. + + Strictly speaking, the vfsmount is shared having its own + peer group, and this peer-group is a slave of some other + peer group. + + Only a slave vfsmount can be made as 'shared and slave' by + either executing the following command + mount --make-shared mount + or by moving the slave vfsmount under a shared vfsmount. + + (4) Private mount + + A 'private mount' is defined as vfsmount that does not + receive or forward any propagation events. + + (5) Unbindable mount + + A 'unbindable mount' is defined as vfsmount that does not + receive or forward any propagation events and cannot + be bind mounted. + + + State diagram: + The state diagram below explains the state transition of a mount, + in response to various commands. + ------------------------------------------------------------------------ + | |make-shared | make-slave | make-private |make-unbindab| + --------------|------------|--------------|--------------|-------------| + |shared |shared |*slave/private| private | unbindable | + | | | | | | + |-------------|------------|--------------|--------------|-------------| + |slave |shared | **slave | private | unbindable | + | |and slave | | | | + |-------------|------------|--------------|--------------|-------------| + |shared |shared | slave | private | unbindable | + |and slave |and slave | | | | + |-------------|------------|--------------|--------------|-------------| + |private |shared | **private | private | unbindable | + |-------------|------------|--------------|--------------|-------------| + |unbindable |shared |**unbindable | private | unbindable | + ------------------------------------------------------------------------ + + * if the shared mount is the only mount in its peer group, making it + slave, makes it private automatically. Note that there is no master to + which it can be slaved to. + + ** slaving a non-shared mount has no effect on the mount. + + Apart from the commands listed below, the 'move' operation also changes + the state of a mount depending on type of the destination mount. Its + explained in section 5d. + +5b) Bind semantics + + Consider the following command + + mount --bind A/a B/b + + where 'A' is the source mount, 'a' is the dentry in the mount 'A', 'B' + is the destination mount and 'b' is the dentry in the destination mount. + + The outcome depends on the type of mount of 'A' and 'B'. The table + below contains quick reference. + --------------------------------------------------------------------------- + | BIND MOUNT OPERATION | + |************************************************************************** + |source(A)->| shared | private | slave | unbindable | + | dest(B) | | | | | + | | | | | | | + | v | | | | | + |************************************************************************** + | shared | shared | shared | shared & slave | invalid | + | | | | | | + |non-shared| shared | private | slave | invalid | + *************************************************************************** + + Details: + + 1. 'A' is a shared mount and 'B' is a shared mount. A new mount 'C' + which is clone of 'A', is created. Its root dentry is 'a' . 'C' is + mounted on mount 'B' at dentry 'b'. Also new mount 'C1', 'C2', 'C3' ... + are created and mounted at the dentry 'b' on all mounts where 'B' + propagates to. A new propagation tree containing 'C1',..,'Cn' is + created. This propagation tree is identical to the propagation tree of + 'B'. And finally the peer-group of 'C' is merged with the peer group + of 'A'. + + 2. 'A' is a private mount and 'B' is a shared mount. A new mount 'C' + which is clone of 'A', is created. Its root dentry is 'a'. 'C' is + mounted on mount 'B' at dentry 'b'. Also new mount 'C1', 'C2', 'C3' ... + are created and mounted at the dentry 'b' on all mounts where 'B' + propagates to. A new propagation tree is set containing all new mounts + 'C', 'C1', .., 'Cn' with exactly the same configuration as the + propagation tree for 'B'. + + 3. 'A' is a slave mount of mount 'Z' and 'B' is a shared mount. A new + mount 'C' which is clone of 'A', is created. Its root dentry is 'a' . + 'C' is mounted on mount 'B' at dentry 'b'. Also new mounts 'C1', 'C2', + 'C3' ... are created and mounted at the dentry 'b' on all mounts where + 'B' propagates to. A new propagation tree containing the new mounts + 'C','C1',.. 'Cn' is created. This propagation tree is identical to the + propagation tree for 'B'. And finally the mount 'C' and its peer group + is made the slave of mount 'Z'. In other words, mount 'C' is in the + state 'slave and shared'. + + 4. 'A' is a unbindable mount and 'B' is a shared mount. This is a + invalid operation. + + 5. 'A' is a private mount and 'B' is a non-shared(private or slave or + unbindable) mount. A new mount 'C' which is clone of 'A', is created. + Its root dentry is 'a'. 'C' is mounted on mount 'B' at dentry 'b'. + + 6. 'A' is a shared mount and 'B' is a non-shared mount. A new mount 'C' + which is a clone of 'A' is created. Its root dentry is 'a'. 'C' is + mounted on mount 'B' at dentry 'b'. 'C' is made a member of the + peer-group of 'A'. + + 7. 'A' is a slave mount of mount 'Z' and 'B' is a non-shared mount. A + new mount 'C' which is a clone of 'A' is created. Its root dentry is + 'a'. 'C' is mounted on mount 'B' at dentry 'b'. Also 'C' is set as a + slave mount of 'Z'. In other words 'A' and 'C' are both slave mounts of + 'Z'. All mount/unmount events on 'Z' propagates to 'A' and 'C'. But + mount/unmount on 'A' do not propagate anywhere else. Similarly + mount/unmount on 'C' do not propagate anywhere else. + + 8. 'A' is a unbindable mount and 'B' is a non-shared mount. This is a + invalid operation. A unbindable mount cannot be bind mounted. + +5c) Rbind semantics + + rbind is same as bind. Bind replicates the specified mount. Rbind + replicates all the mounts in the tree belonging to the specified mount. + Rbind mount is bind mount applied to all the mounts in the tree. + + If the source tree that is rbind has some unbindable mounts, + then the subtree under the unbindable mount is pruned in the new + location. + + eg: let's say we have the following mount tree. + + A + / \ + B C + / \ / \ + D E F G + + Let's say all the mount except the mount C in the tree are + of a type other than unbindable. + + If this tree is rbound to say Z + + We will have the following tree at the new location. + + Z + | + A' + / + B' Note how the tree under C is pruned + / \ in the new location. + D' E' + + + +5d) Move semantics + + Consider the following command + + mount --move A B/b + + where 'A' is the source mount, 'B' is the destination mount and 'b' is + the dentry in the destination mount. + + The outcome depends on the type of the mount of 'A' and 'B'. The table + below is a quick reference. + --------------------------------------------------------------------------- + | MOVE MOUNT OPERATION | + |************************************************************************** + | source(A)->| shared | private | slave | unbindable | + | dest(B) | | | | | + | | | | | | | + | v | | | | | + |************************************************************************** + | shared | shared | shared |shared and slave| invalid | + | | | | | | + |non-shared| shared | private | slave | unbindable | + *************************************************************************** + NOTE: moving a mount residing under a shared mount is invalid. + + Details follow: + + 1. 'A' is a shared mount and 'B' is a shared mount. The mount 'A' is + mounted on mount 'B' at dentry 'b'. Also new mounts 'A1', 'A2'...'An' + are created and mounted at dentry 'b' on all mounts that receive + propagation from mount 'B'. A new propagation tree is created in the + exact same configuration as that of 'B'. This new propagation tree + contains all the new mounts 'A1', 'A2'... 'An'. And this new + propagation tree is appended to the already existing propagation tree + of 'A'. + + 2. 'A' is a private mount and 'B' is a shared mount. The mount 'A' is + mounted on mount 'B' at dentry 'b'. Also new mount 'A1', 'A2'... 'An' + are created and mounted at dentry 'b' on all mounts that receive + propagation from mount 'B'. The mount 'A' becomes a shared mount and a + propagation tree is created which is identical to that of + 'B'. This new propagation tree contains all the new mounts 'A1', + 'A2'... 'An'. + + 3. 'A' is a slave mount of mount 'Z' and 'B' is a shared mount. The + mount 'A' is mounted on mount 'B' at dentry 'b'. Also new mounts 'A1', + 'A2'... 'An' are created and mounted at dentry 'b' on all mounts that + receive propagation from mount 'B'. A new propagation tree is created + in the exact same configuration as that of 'B'. This new propagation + tree contains all the new mounts 'A1', 'A2'... 'An'. And this new + propagation tree is appended to the already existing propagation tree of + 'A'. Mount 'A' continues to be the slave mount of 'Z' but it also + becomes 'shared'. + + 4. 'A' is a unbindable mount and 'B' is a shared mount. The operation + is invalid. Because mounting anything on the shared mount 'B' can + create new mounts that get mounted on the mounts that receive + propagation from 'B'. And since the mount 'A' is unbindable, cloning + it to mount at other mountpoints is not possible. + + 5. 'A' is a private mount and 'B' is a non-shared(private or slave or + unbindable) mount. The mount 'A' is mounted on mount 'B' at dentry 'b'. + + 6. 'A' is a shared mount and 'B' is a non-shared mount. The mount 'A' + is mounted on mount 'B' at dentry 'b'. Mount 'A' continues to be a + shared mount. + + 7. 'A' is a slave mount of mount 'Z' and 'B' is a non-shared mount. + The mount 'A' is mounted on mount 'B' at dentry 'b'. Mount 'A' + continues to be a slave mount of mount 'Z'. + + 8. 'A' is a unbindable mount and 'B' is a non-shared mount. The mount + 'A' is mounted on mount 'B' at dentry 'b'. Mount 'A' continues to be a + unbindable mount. + +5e) Mount semantics + + Consider the following command + + mount device B/b + + 'B' is the destination mount and 'b' is the dentry in the destination + mount. + + The above operation is the same as bind operation with the exception + that the source mount is always a private mount. + + +5f) Unmount semantics + + Consider the following command + + umount A + + where 'A' is a mount mounted on mount 'B' at dentry 'b'. + + If mount 'B' is shared, then all most-recently-mounted mounts at dentry + 'b' on mounts that receive propagation from mount 'B' and does not have + sub-mounts within them are unmounted. + + Example: Let's say 'B1', 'B2', 'B3' are shared mounts that propagate to + each other. + + let's say 'A1', 'A2', 'A3' are first mounted at dentry 'b' on mount + 'B1', 'B2' and 'B3' respectively. + + let's say 'C1', 'C2', 'C3' are next mounted at the same dentry 'b' on + mount 'B1', 'B2' and 'B3' respectively. + + if 'C1' is unmounted, all the mounts that are most-recently-mounted on + 'B1' and on the mounts that 'B1' propagates-to are unmounted. + + 'B1' propagates to 'B2' and 'B3'. And the most recently mounted mount + on 'B2' at dentry 'b' is 'C2', and that of mount 'B3' is 'C3'. + + So all 'C1', 'C2' and 'C3' should be unmounted. + + If any of 'C2' or 'C3' has some child mounts, then that mount is not + unmounted, but all other mounts are unmounted. However if 'C1' is told + to be unmounted and 'C1' has some sub-mounts, the umount operation is + failed entirely. + +5g) Clone Namespace + + A cloned namespace contains all the mounts as that of the parent + namespace. + + Let's say 'A' and 'B' are the corresponding mounts in the parent and the + child namespace. + + If 'A' is shared, then 'B' is also shared and 'A' and 'B' propagate to + each other. + + If 'A' is a slave mount of 'Z', then 'B' is also the slave mount of + 'Z'. + + If 'A' is a private mount, then 'B' is a private mount too. + + If 'A' is unbindable mount, then 'B' is a unbindable mount too. + + +6) Quiz + + A. What is the result of the following command sequence? + + mount --bind /mnt /mnt + mount --make-shared /mnt + mount --bind /mnt /tmp + mount --move /tmp /mnt/1 + + what should be the contents of /mnt /mnt/1 /mnt/1/1 should be? + Should they all be identical? or should /mnt and /mnt/1 be + identical only? + + + B. What is the result of the following command sequence? + + mount --make-rshared / + mkdir -p /v/1 + mount --rbind / /v/1 + + what should be the content of /v/1/v/1 be? + + + C. What is the result of the following command sequence? + + mount --bind /mnt /mnt + mount --make-shared /mnt + mkdir -p /mnt/1/2/3 /mnt/1/test + mount --bind /mnt/1 /tmp + mount --make-slave /mnt + mount --make-shared /mnt + mount --bind /mnt/1/2 /tmp1 + mount --make-slave /mnt + + At this point we have the first mount at /tmp and + its root dentry is 1. Let's call this mount 'A' + And then we have a second mount at /tmp1 with root + dentry 2. Let's call this mount 'B' + Next we have a third mount at /mnt with root dentry + mnt. Let's call this mount 'C' + + 'B' is the slave of 'A' and 'C' is a slave of 'B' + A -> B -> C + + at this point if we execute the following command + + mount --bind /bin /tmp/test + + The mount is attempted on 'A' + + will the mount propagate to 'B' and 'C' ? + + what would be the contents of + /mnt/1/test be? + +7) FAQ + + Q1. Why is bind mount needed? How is it different from symbolic links? + symbolic links can get stale if the destination mount gets + unmounted or moved. Bind mounts continue to exist even if the + other mount is unmounted or moved. + + Q2. Why can't the shared subtree be implemented using exportfs? + + exportfs is a heavyweight way of accomplishing part of what + shared subtree can do. I cannot imagine a way to implement the + semantics of slave mount using exportfs? + + Q3 Why is unbindable mount needed? + + Let's say we want to replicate the mount tree at multiple + locations within the same subtree. + + if one rbind mounts a tree within the same subtree 'n' times + the number of mounts created is an exponential function of 'n'. + Having unbindable mount can help prune the unneeded bind + mounts. Here is a example. + + step 1: + let's say the root tree has just two directories with + one vfsmount. + root + / \ + tmp usr + + And we want to replicate the tree at multiple + mountpoints under /root/tmp + + step2: + mount --make-shared /root + + mkdir -p /tmp/m1 + + mount --rbind /root /tmp/m1 + + the new tree now looks like this: + + root + / \ + tmp usr + / + m1 + / \ + tmp usr + / + m1 + + it has two vfsmounts + + step3: + mkdir -p /tmp/m2 + mount --rbind /root /tmp/m2 + + the new tree now looks like this: + + root + / \ + tmp usr + / \ + m1 m2 + / \ / \ + tmp usr tmp usr + / \ / + m1 m2 m1 + / \ / \ + tmp usr tmp usr + / / \ + m1 m1 m2 + / \ + tmp usr + / \ + m1 m2 + + it has 6 vfsmounts + + step 4: + mkdir -p /tmp/m3 + mount --rbind /root /tmp/m3 + + I wont' draw the tree..but it has 24 vfsmounts + + + at step i the number of vfsmounts is V[i] = i*V[i-1]. + This is an exponential function. And this tree has way more + mounts than what we really needed in the first place. + + One could use a series of umount at each step to prune + out the unneeded mounts. But there is a better solution. + Unclonable mounts come in handy here. + + step 1: + let's say the root tree has just two directories with + one vfsmount. + root + / \ + tmp usr + + How do we set up the same tree at multiple locations under + /root/tmp + + step2: + mount --bind /root/tmp /root/tmp + + mount --make-rshared /root + mount --make-unbindable /root/tmp + + mkdir -p /tmp/m1 + + mount --rbind /root /tmp/m1 + + the new tree now looks like this: + + root + / \ + tmp usr + / + m1 + / \ + tmp usr + + step3: + mkdir -p /tmp/m2 + mount --rbind /root /tmp/m2 + + the new tree now looks like this: + + root + / \ + tmp usr + / \ + m1 m2 + / \ / \ + tmp usr tmp usr + + step4: + + mkdir -p /tmp/m3 + mount --rbind /root /tmp/m3 + + the new tree now looks like this: + + root + / \ + tmp usr + / \ \ + m1 m2 m3 + / \ / \ / \ + tmp usr tmp usr tmp usr + +8) Implementation + +8A) Datastructure + + 4 new fields are introduced to struct vfsmount + ->mnt_share + ->mnt_slave_list + ->mnt_slave + ->mnt_master + + ->mnt_share links together all the mount to/from which this vfsmount + send/receives propagation events. + + ->mnt_slave_list links all the mounts to which this vfsmount propagates + to. + + ->mnt_slave links together all the slaves that its master vfsmount + propagates to. + + ->mnt_master points to the master vfsmount from which this vfsmount + receives propagation. + + ->mnt_flags takes two more flags to indicate the propagation status of + the vfsmount. MNT_SHARE indicates that the vfsmount is a shared + vfsmount. MNT_UNCLONABLE indicates that the vfsmount cannot be + replicated. + + All the shared vfsmounts in a peer group form a cyclic list through + ->mnt_share. + + All vfsmounts with the same ->mnt_master form on a cyclic list anchored + in ->mnt_master->mnt_slave_list and going through ->mnt_slave. + + ->mnt_master can point to arbitrary (and possibly different) members + of master peer group. To find all immediate slaves of a peer group + you need to go through _all_ ->mnt_slave_list of its members. + Conceptually it's just a single set - distribution among the + individual lists does not affect propagation or the way propagation + tree is modified by operations. + + All vfsmounts in a peer group have the same ->mnt_master. If it is + non-NULL, they form a contiguous (ordered) segment of slave list. + + A example propagation tree looks as shown in the figure below. + [ NOTE: Though it looks like a forest, if we consider all the shared + mounts as a conceptual entity called 'pnode', it becomes a tree] + + + A <--> B <--> C <---> D + /|\ /| |\ + / F G J K H I + / + E<-->K + /|\ + M L N + + In the above figure A,B,C and D all are shared and propagate to each + other. 'A' has got 3 slave mounts 'E' 'F' and 'G' 'C' has got 2 slave + mounts 'J' and 'K' and 'D' has got two slave mounts 'H' and 'I'. + 'E' is also shared with 'K' and they propagate to each other. And + 'K' has 3 slaves 'M', 'L' and 'N' + + A's ->mnt_share links with the ->mnt_share of 'B' 'C' and 'D' + + A's ->mnt_slave_list links with ->mnt_slave of 'E', 'K', 'F' and 'G' + + E's ->mnt_share links with ->mnt_share of K + 'E', 'K', 'F', 'G' have their ->mnt_master point to struct + vfsmount of 'A' + 'M', 'L', 'N' have their ->mnt_master point to struct vfsmount of 'K' + K's ->mnt_slave_list links with ->mnt_slave of 'M', 'L' and 'N' + + C's ->mnt_slave_list links with ->mnt_slave of 'J' and 'K' + J and K's ->mnt_master points to struct vfsmount of C + and finally D's ->mnt_slave_list links with ->mnt_slave of 'H' and 'I' + 'H' and 'I' have their ->mnt_master pointing to struct vfsmount of 'D'. + + + NOTE: The propagation tree is orthogonal to the mount tree. + +8B Locking: + + ->mnt_share, ->mnt_slave, ->mnt_slave_list, ->mnt_master are protected + by namespace_sem (exclusive for modifications, shared for reading). + + Normally we have ->mnt_flags modifications serialized by vfsmount_lock. + There are two exceptions: do_add_mount() and clone_mnt(). + The former modifies a vfsmount that has not been visible in any shared + data structures yet. + The latter holds namespace_sem and the only references to vfsmount + are in lists that can't be traversed without namespace_sem. + +8C Algorithm: + + The crux of the implementation resides in rbind/move operation. + + The overall algorithm breaks the operation into 3 phases: (look at + attach_recursive_mnt() and propagate_mnt()) + + 1. prepare phase. + 2. commit phases. + 3. abort phases. + + Prepare phase: + + for each mount in the source tree: + a) Create the necessary number of mount trees to + be attached to each of the mounts that receive + propagation from the destination mount. + b) Do not attach any of the trees to its destination. + However note down its ->mnt_parent and ->mnt_mountpoint + c) Link all the new mounts to form a propagation tree that + is identical to the propagation tree of the destination + mount. + + If this phase is successful, there should be 'n' new + propagation trees; where 'n' is the number of mounts in the + source tree. Go to the commit phase + + Also there should be 'm' new mount trees, where 'm' is + the number of mounts to which the destination mount + propagates to. + + if any memory allocations fail, go to the abort phase. + + Commit phase + attach each of the mount trees to their corresponding + destination mounts. + + Abort phase + delete all the newly created trees. + + NOTE: all the propagation related functionality resides in the file + pnode.c + + +------------------------------------------------------------------------ + +version 0.1 (created the initial document, Ram Pai linuxram@us.ibm.com) +version 0.2 (Incorporated comments from Al Viro) diff --git a/Documentation/filesystems/spufs.txt b/Documentation/filesystems/spufs.txt new file mode 100644 index 00000000..1343d118 --- /dev/null +++ b/Documentation/filesystems/spufs.txt @@ -0,0 +1,521 @@ +SPUFS(2) Linux Programmer's Manual SPUFS(2) + + + +NAME + spufs - the SPU file system + + +DESCRIPTION + The SPU file system is used on PowerPC machines that implement the Cell + Broadband Engine Architecture in order to access Synergistic Processor + Units (SPUs). + + The file system provides a name space similar to posix shared memory or + message queues. Users that have write permissions on the file system + can use spu_create(2) to establish SPU contexts in the spufs root. + + Every SPU context is represented by a directory containing a predefined + set of files. These files can be used for manipulating the state of the + logical SPU. Users can change permissions on those files, but not actu- + ally add or remove files. + + +MOUNT OPTIONS + uid=<uid> + set the user owning the mount point, the default is 0 (root). + + gid=<gid> + set the group owning the mount point, the default is 0 (root). + + +FILES + The files in spufs mostly follow the standard behavior for regular sys- + tem calls like read(2) or write(2), but often support only a subset of + the operations supported on regular file systems. This list details the + supported operations and the deviations from the behaviour in the + respective man pages. + + All files that support the read(2) operation also support readv(2) and + all files that support the write(2) operation also support writev(2). + All files support the access(2) and stat(2) family of operations, but + only the st_mode, st_nlink, st_uid and st_gid fields of struct stat + contain reliable information. + + All files support the chmod(2)/fchmod(2) and chown(2)/fchown(2) opera- + tions, but will not be able to grant permissions that contradict the + possible operations, e.g. read access on the wbox file. + + The current set of files is: + + + /mem + the contents of the local storage memory of the SPU. This can be + accessed like a regular shared memory file and contains both code and + data in the address space of the SPU. The possible operations on an + open mem file are: + + read(2), pread(2), write(2), pwrite(2), lseek(2) + These operate as documented, with the exception that seek(2), + write(2) and pwrite(2) are not supported beyond the end of the + file. The file size is the size of the local storage of the SPU, + which normally is 256 kilobytes. + + mmap(2) + Mapping mem into the process address space gives access to the + SPU local storage within the process address space. Only + MAP_SHARED mappings are allowed. + + + /mbox + The first SPU to CPU communication mailbox. This file is read-only and + can be read in units of 32 bits. The file can only be used in non- + blocking mode and it even poll() will not block on it. The possible + operations on an open mbox file are: + + read(2) + If a count smaller than four is requested, read returns -1 and + sets errno to EINVAL. If there is no data available in the mail + box, the return value is set to -1 and errno becomes EAGAIN. + When data has been read successfully, four bytes are placed in + the data buffer and the value four is returned. + + + /ibox + The second SPU to CPU communication mailbox. This file is similar to + the first mailbox file, but can be read in blocking I/O mode, and the + poll family of system calls can be used to wait for it. The possible + operations on an open ibox file are: + + read(2) + If a count smaller than four is requested, read returns -1 and + sets errno to EINVAL. If there is no data available in the mail + box and the file descriptor has been opened with O_NONBLOCK, the + return value is set to -1 and errno becomes EAGAIN. + + If there is no data available in the mail box and the file + descriptor has been opened without O_NONBLOCK, the call will + block until the SPU writes to its interrupt mailbox channel. + When data has been read successfully, four bytes are placed in + the data buffer and the value four is returned. + + poll(2) + Poll on the ibox file returns (POLLIN | POLLRDNORM) whenever + data is available for reading. + + + /wbox + The CPU to SPU communation mailbox. It is write-only and can be written + in units of 32 bits. If the mailbox is full, write() will block and + poll can be used to wait for it becoming empty again. The possible + operations on an open wbox file are: write(2) If a count smaller than + four is requested, write returns -1 and sets errno to EINVAL. If there + is no space available in the mail box and the file descriptor has been + opened with O_NONBLOCK, the return value is set to -1 and errno becomes + EAGAIN. + + If there is no space available in the mail box and the file descriptor + has been opened without O_NONBLOCK, the call will block until the SPU + reads from its PPE mailbox channel. When data has been read success- + fully, four bytes are placed in the data buffer and the value four is + returned. + + poll(2) + Poll on the ibox file returns (POLLOUT | POLLWRNORM) whenever + space is available for writing. + + + /mbox_stat + /ibox_stat + /wbox_stat + Read-only files that contain the length of the current queue, i.e. how + many words can be read from mbox or ibox or how many words can be + written to wbox without blocking. The files can be read only in 4-byte + units and return a big-endian binary integer number. The possible + operations on an open *box_stat file are: + + read(2) + If a count smaller than four is requested, read returns -1 and + sets errno to EINVAL. Otherwise, a four byte value is placed in + the data buffer, containing the number of elements that can be + read from (for mbox_stat and ibox_stat) or written to (for + wbox_stat) the respective mail box without blocking or resulting + in EAGAIN. + + + /npc + /decr + /decr_status + /spu_tag_mask + /event_mask + /srr0 + Internal registers of the SPU. The representation is an ASCII string + with the numeric value of the next instruction to be executed. These + can be used in read/write mode for debugging, but normal operation of + programs should not rely on them because access to any of them except + npc requires an SPU context save and is therefore very inefficient. + + The contents of these files are: + + npc Next Program Counter + + decr SPU Decrementer + + decr_status Decrementer Status + + spu_tag_mask MFC tag mask for SPU DMA + + event_mask Event mask for SPU interrupts + + srr0 Interrupt Return address register + + + The possible operations on an open npc, decr, decr_status, + spu_tag_mask, event_mask or srr0 file are: + + read(2) + When the count supplied to the read call is shorter than the + required length for the pointer value plus a newline character, + subsequent reads from the same file descriptor will result in + completing the string, regardless of changes to the register by + a running SPU task. When a complete string has been read, all + subsequent read operations will return zero bytes and a new file + descriptor needs to be opened to read the value again. + + write(2) + A write operation on the file results in setting the register to + the value given in the string. The string is parsed from the + beginning to the first non-numeric character or the end of the + buffer. Subsequent writes to the same file descriptor overwrite + the previous setting. + + + /fpcr + This file gives access to the Floating Point Status and Control Regis- + ter as a four byte long file. The operations on the fpcr file are: + + read(2) + If a count smaller than four is requested, read returns -1 and + sets errno to EINVAL. Otherwise, a four byte value is placed in + the data buffer, containing the current value of the fpcr regis- + ter. + + write(2) + If a count smaller than four is requested, write returns -1 and + sets errno to EINVAL. Otherwise, a four byte value is copied + from the data buffer, updating the value of the fpcr register. + + + /signal1 + /signal2 + The two signal notification channels of an SPU. These are read-write + files that operate on a 32 bit word. Writing to one of these files + triggers an interrupt on the SPU. The value written to the signal + files can be read from the SPU through a channel read or from host user + space through the file. After the value has been read by the SPU, it + is reset to zero. The possible operations on an open signal1 or sig- + nal2 file are: + + read(2) + If a count smaller than four is requested, read returns -1 and + sets errno to EINVAL. Otherwise, a four byte value is placed in + the data buffer, containing the current value of the specified + signal notification register. + + write(2) + If a count smaller than four is requested, write returns -1 and + sets errno to EINVAL. Otherwise, a four byte value is copied + from the data buffer, updating the value of the specified signal + notification register. The signal notification register will + either be replaced with the input data or will be updated to the + bitwise OR or the old value and the input data, depending on the + contents of the signal1_type, or signal2_type respectively, + file. + + + /signal1_type + /signal2_type + These two files change the behavior of the signal1 and signal2 notifi- + cation files. The contain a numerical ASCII string which is read as + either "1" or "0". In mode 0 (overwrite), the hardware replaces the + contents of the signal channel with the data that is written to it. in + mode 1 (logical OR), the hardware accumulates the bits that are subse- + quently written to it. The possible operations on an open signal1_type + or signal2_type file are: + + read(2) + When the count supplied to the read call is shorter than the + required length for the digit plus a newline character, subse- + quent reads from the same file descriptor will result in com- + pleting the string. When a complete string has been read, all + subsequent read operations will return zero bytes and a new file + descriptor needs to be opened to read the value again. + + write(2) + A write operation on the file results in setting the register to + the value given in the string. The string is parsed from the + beginning to the first non-numeric character or the end of the + buffer. Subsequent writes to the same file descriptor overwrite + the previous setting. + + +EXAMPLES + /etc/fstab entry + none /spu spufs gid=spu 0 0 + + +AUTHORS + Arnd Bergmann <arndb@de.ibm.com>, Mark Nutter <mnutter@us.ibm.com>, + Ulrich Weigand <Ulrich.Weigand@de.ibm.com> + +SEE ALSO + capabilities(7), close(2), spu_create(2), spu_run(2), spufs(7) + + + +Linux 2005-09-28 SPUFS(2) + +------------------------------------------------------------------------------ + +SPU_RUN(2) Linux Programmer's Manual SPU_RUN(2) + + + +NAME + spu_run - execute an spu context + + +SYNOPSIS + #include <sys/spu.h> + + int spu_run(int fd, unsigned int *npc, unsigned int *event); + +DESCRIPTION + The spu_run system call is used on PowerPC machines that implement the + Cell Broadband Engine Architecture in order to access Synergistic Pro- + cessor Units (SPUs). It uses the fd that was returned from spu_cre- + ate(2) to address a specific SPU context. When the context gets sched- + uled to a physical SPU, it starts execution at the instruction pointer + passed in npc. + + Execution of SPU code happens synchronously, meaning that spu_run does + not return while the SPU is still running. If there is a need to exe- + cute SPU code in parallel with other code on either the main CPU or + other SPUs, you need to create a new thread of execution first, e.g. + using the pthread_create(3) call. + + When spu_run returns, the current value of the SPU instruction pointer + is written back to npc, so you can call spu_run again without updating + the pointers. + + event can be a NULL pointer or point to an extended status code that + gets filled when spu_run returns. It can be one of the following con- + stants: + + SPE_EVENT_DMA_ALIGNMENT + A DMA alignment error + + SPE_EVENT_SPE_DATA_SEGMENT + A DMA segmentation error + + SPE_EVENT_SPE_DATA_STORAGE + A DMA storage error + + If NULL is passed as the event argument, these errors will result in a + signal delivered to the calling process. + +RETURN VALUE + spu_run returns the value of the spu_status register or -1 to indicate + an error and set errno to one of the error codes listed below. The + spu_status register value contains a bit mask of status codes and + optionally a 14 bit code returned from the stop-and-signal instruction + on the SPU. The bit masks for the status codes are: + + 0x02 SPU was stopped by stop-and-signal. + + 0x04 SPU was stopped by halt. + + 0x08 SPU is waiting for a channel. + + 0x10 SPU is in single-step mode. + + 0x20 SPU has tried to execute an invalid instruction. + + 0x40 SPU has tried to access an invalid channel. + + 0x3fff0000 + The bits masked with this value contain the code returned from + stop-and-signal. + + There are always one or more of the lower eight bits set or an error + code is returned from spu_run. + +ERRORS + EAGAIN or EWOULDBLOCK + fd is in non-blocking mode and spu_run would block. + + EBADF fd is not a valid file descriptor. + + EFAULT npc is not a valid pointer or status is neither NULL nor a valid + pointer. + + EINTR A signal occurred while spu_run was in progress. The npc value + has been updated to the new program counter value if necessary. + + EINVAL fd is not a file descriptor returned from spu_create(2). + + ENOMEM Insufficient memory was available to handle a page fault result- + ing from an MFC direct memory access. + + ENOSYS the functionality is not provided by the current system, because + either the hardware does not provide SPUs or the spufs module is + not loaded. + + +NOTES + spu_run is meant to be used from libraries that implement a more + abstract interface to SPUs, not to be used from regular applications. + See http://www.bsc.es/projects/deepcomputing/linuxoncell/ for the rec- + ommended libraries. + + +CONFORMING TO + This call is Linux specific and only implemented by the ppc64 architec- + ture. Programs using this system call are not portable. + + +BUGS + The code does not yet fully implement all features lined out here. + + +AUTHOR + Arnd Bergmann <arndb@de.ibm.com> + +SEE ALSO + capabilities(7), close(2), spu_create(2), spufs(7) + + + +Linux 2005-09-28 SPU_RUN(2) + +------------------------------------------------------------------------------ + +SPU_CREATE(2) Linux Programmer's Manual SPU_CREATE(2) + + + +NAME + spu_create - create a new spu context + + +SYNOPSIS + #include <sys/types.h> + #include <sys/spu.h> + + int spu_create(const char *pathname, int flags, mode_t mode); + +DESCRIPTION + The spu_create system call is used on PowerPC machines that implement + the Cell Broadband Engine Architecture in order to access Synergistic + Processor Units (SPUs). It creates a new logical context for an SPU in + pathname and returns a handle to associated with it. pathname must + point to a non-existing directory in the mount point of the SPU file + system (spufs). When spu_create is successful, a directory gets cre- + ated on pathname and it is populated with files. + + The returned file handle can only be passed to spu_run(2) or closed, + other operations are not defined on it. When it is closed, all associ- + ated directory entries in spufs are removed. When the last file handle + pointing either inside of the context directory or to this file + descriptor is closed, the logical SPU context is destroyed. + + The parameter flags can be zero or any bitwise or'd combination of the + following constants: + + SPU_RAWIO + Allow mapping of some of the hardware registers of the SPU into + user space. This flag requires the CAP_SYS_RAWIO capability, see + capabilities(7). + + The mode parameter specifies the permissions used for creating the new + directory in spufs. mode is modified with the user's umask(2) value + and then used for both the directory and the files contained in it. The + file permissions mask out some more bits of mode because they typically + support only read or write access. See stat(2) for a full list of the + possible mode values. + + +RETURN VALUE + spu_create returns a new file descriptor. It may return -1 to indicate + an error condition and set errno to one of the error codes listed + below. + + +ERRORS + EACCESS + The current user does not have write access on the spufs mount + point. + + EEXIST An SPU context already exists at the given path name. + + EFAULT pathname is not a valid string pointer in the current address + space. + + EINVAL pathname is not a directory in the spufs mount point. + + ELOOP Too many symlinks were found while resolving pathname. + + EMFILE The process has reached its maximum open file limit. + + ENAMETOOLONG + pathname was too long. + + ENFILE The system has reached the global open file limit. + + ENOENT Part of pathname could not be resolved. + + ENOMEM The kernel could not allocate all resources required. + + ENOSPC There are not enough SPU resources available to create a new + context or the user specific limit for the number of SPU con- + texts has been reached. + + ENOSYS the functionality is not provided by the current system, because + either the hardware does not provide SPUs or the spufs module is + not loaded. + + ENOTDIR + A part of pathname is not a directory. + + + +NOTES + spu_create is meant to be used from libraries that implement a more + abstract interface to SPUs, not to be used from regular applications. + See http://www.bsc.es/projects/deepcomputing/linuxoncell/ for the rec- + ommended libraries. + + +FILES + pathname must point to a location beneath the mount point of spufs. By + convention, it gets mounted in /spu. + + +CONFORMING TO + This call is Linux specific and only implemented by the ppc64 architec- + ture. Programs using this system call are not portable. + + +BUGS + The code does not yet fully implement all features lined out here. + + +AUTHOR + Arnd Bergmann <arndb@de.ibm.com> + +SEE ALSO + capabilities(7), close(2), spu_run(2), spufs(7) + + + +Linux 2005-09-28 SPU_CREATE(2) diff --git a/Documentation/filesystems/squashfs.txt b/Documentation/filesystems/squashfs.txt new file mode 100644 index 00000000..d4d41465 --- /dev/null +++ b/Documentation/filesystems/squashfs.txt @@ -0,0 +1,257 @@ +SQUASHFS 4.0 FILESYSTEM +======================= + +Squashfs is a compressed read-only filesystem for Linux. +It uses zlib/lzo compression to compress files, inodes and directories. +Inodes in the system are very small and all blocks are packed to minimise +data overhead. Block sizes greater than 4K are supported up to a maximum +of 1Mbytes (default block size 128K). + +Squashfs is intended for general read-only filesystem use, for archival +use (i.e. in cases where a .tar.gz file may be used), and in constrained +block device/memory systems (e.g. embedded systems) where low overhead is +needed. + +Mailing list: squashfs-devel@lists.sourceforge.net +Web site: www.squashfs.org + +1. FILESYSTEM FEATURES +---------------------- + +Squashfs filesystem features versus Cramfs: + + Squashfs Cramfs + +Max filesystem size: 2^64 256 MiB +Max file size: ~ 2 TiB 16 MiB +Max files: unlimited unlimited +Max directories: unlimited unlimited +Max entries per directory: unlimited unlimited +Max block size: 1 MiB 4 KiB +Metadata compression: yes no +Directory indexes: yes no +Sparse file support: yes no +Tail-end packing (fragments): yes no +Exportable (NFS etc.): yes no +Hard link support: yes no +"." and ".." in readdir: yes no +Real inode numbers: yes no +32-bit uids/gids: yes no +File creation time: yes no +Xattr support: yes no +ACL support: no no + +Squashfs compresses data, inodes and directories. In addition, inode and +directory data are highly compacted, and packed on byte boundaries. Each +compressed inode is on average 8 bytes in length (the exact length varies on +file type, i.e. regular file, directory, symbolic link, and block/char device +inodes have different sizes). + +2. USING SQUASHFS +----------------- + +As squashfs is a read-only filesystem, the mksquashfs program must be used to +create populated squashfs filesystems. This and other squashfs utilities +can be obtained from http://www.squashfs.org. Usage instructions can be +obtained from this site also. + + +3. SQUASHFS FILESYSTEM DESIGN +----------------------------- + +A squashfs filesystem consists of a maximum of nine parts, packed together on a +byte alignment: + + --------------- + | superblock | + |---------------| + | compression | + | options | + |---------------| + | datablocks | + | & fragments | + |---------------| + | inode table | + |---------------| + | directory | + | table | + |---------------| + | fragment | + | table | + |---------------| + | export | + | table | + |---------------| + | uid/gid | + | lookup table | + |---------------| + | xattr | + | table | + --------------- + +Compressed data blocks are written to the filesystem as files are read from +the source directory, and checked for duplicates. Once all file data has been +written the completed inode, directory, fragment, export and uid/gid lookup +tables are written. + +3.1 Compression options +----------------------- + +Compressors can optionally support compression specific options (e.g. +dictionary size). If non-default compression options have been used, then +these are stored here. + +3.2 Inodes +---------- + +Metadata (inodes and directories) are compressed in 8Kbyte blocks. Each +compressed block is prefixed by a two byte length, the top bit is set if the +block is uncompressed. A block will be uncompressed if the -noI option is set, +or if the compressed block was larger than the uncompressed block. + +Inodes are packed into the metadata blocks, and are not aligned to block +boundaries, therefore inodes overlap compressed blocks. Inodes are identified +by a 48-bit number which encodes the location of the compressed metadata block +containing the inode, and the byte offset into that block where the inode is +placed (<block, offset>). + +To maximise compression there are different inodes for each file type +(regular file, directory, device, etc.), the inode contents and length +varying with the type. + +To further maximise compression, two types of regular file inode and +directory inode are defined: inodes optimised for frequently occurring +regular files and directories, and extended types where extra +information has to be stored. + +3.3 Directories +--------------- + +Like inodes, directories are packed into compressed metadata blocks, stored +in a directory table. Directories are accessed using the start address of +the metablock containing the directory and the offset into the +decompressed block (<block, offset>). + +Directories are organised in a slightly complex way, and are not simply +a list of file names. The organisation takes advantage of the +fact that (in most cases) the inodes of the files will be in the same +compressed metadata block, and therefore, can share the start block. +Directories are therefore organised in a two level list, a directory +header containing the shared start block value, and a sequence of directory +entries, each of which share the shared start block. A new directory header +is written once/if the inode start block changes. The directory +header/directory entry list is repeated as many times as necessary. + +Directories are sorted, and can contain a directory index to speed up +file lookup. Directory indexes store one entry per metablock, each entry +storing the index/filename mapping to the first directory header +in each metadata block. Directories are sorted in alphabetical order, +and at lookup the index is scanned linearly looking for the first filename +alphabetically larger than the filename being looked up. At this point the +location of the metadata block the filename is in has been found. +The general idea of the index is ensure only one metadata block needs to be +decompressed to do a lookup irrespective of the length of the directory. +This scheme has the advantage that it doesn't require extra memory overhead +and doesn't require much extra storage on disk. + +3.4 File data +------------- + +Regular files consist of a sequence of contiguous compressed blocks, and/or a +compressed fragment block (tail-end packed block). The compressed size +of each datablock is stored in a block list contained within the +file inode. + +To speed up access to datablocks when reading 'large' files (256 Mbytes or +larger), the code implements an index cache that caches the mapping from +block index to datablock location on disk. + +The index cache allows Squashfs to handle large files (up to 1.75 TiB) while +retaining a simple and space-efficient block list on disk. The cache +is split into slots, caching up to eight 224 GiB files (128 KiB blocks). +Larger files use multiple slots, with 1.75 TiB files using all 8 slots. +The index cache is designed to be memory efficient, and by default uses +16 KiB. + +3.5 Fragment lookup table +------------------------- + +Regular files can contain a fragment index which is mapped to a fragment +location on disk and compressed size using a fragment lookup table. This +fragment lookup table is itself stored compressed into metadata blocks. +A second index table is used to locate these. This second index table for +speed of access (and because it is small) is read at mount time and cached +in memory. + +3.6 Uid/gid lookup table +------------------------ + +For space efficiency regular files store uid and gid indexes, which are +converted to 32-bit uids/gids using an id look up table. This table is +stored compressed into metadata blocks. A second index table is used to +locate these. This second index table for speed of access (and because it +is small) is read at mount time and cached in memory. + +3.7 Export table +---------------- + +To enable Squashfs filesystems to be exportable (via NFS etc.) filesystems +can optionally (disabled with the -no-exports Mksquashfs option) contain +an inode number to inode disk location lookup table. This is required to +enable Squashfs to map inode numbers passed in filehandles to the inode +location on disk, which is necessary when the export code reinstantiates +expired/flushed inodes. + +This table is stored compressed into metadata blocks. A second index table is +used to locate these. This second index table for speed of access (and because +it is small) is read at mount time and cached in memory. + +3.8 Xattr table +--------------- + +The xattr table contains extended attributes for each inode. The xattrs +for each inode are stored in a list, each list entry containing a type, +name and value field. The type field encodes the xattr prefix +("user.", "trusted." etc) and it also encodes how the name/value fields +should be interpreted. Currently the type indicates whether the value +is stored inline (in which case the value field contains the xattr value), +or if it is stored out of line (in which case the value field stores a +reference to where the actual value is stored). This allows large values +to be stored out of line improving scanning and lookup performance and it +also allows values to be de-duplicated, the value being stored once, and +all other occurrences holding an out of line reference to that value. + +The xattr lists are packed into compressed 8K metadata blocks. +To reduce overhead in inodes, rather than storing the on-disk +location of the xattr list inside each inode, a 32-bit xattr id +is stored. This xattr id is mapped into the location of the xattr +list using a second xattr id lookup table. + +4. TODOS AND OUTSTANDING ISSUES +------------------------------- + +4.1 Todo list +------------- + +Implement ACL support. + +4.2 Squashfs internal cache +--------------------------- + +Blocks in Squashfs are compressed. To avoid repeatedly decompressing +recently accessed data Squashfs uses two small metadata and fragment caches. + +The cache is not used for file datablocks, these are decompressed and cached in +the page-cache in the normal way. The cache is used to temporarily cache +fragment and metadata blocks which have been read as a result of a metadata +(i.e. inode or directory) or fragment access. Because metadata and fragments +are packed together into blocks (to gain greater compression) the read of a +particular piece of metadata or fragment will retrieve other metadata/fragments +which have been packed with it, these because of locality-of-reference may be +read in the near future. Temporarily caching them ensures they are available +for near future access without requiring an additional read and decompress. + +In the future this internal cache may be replaced with an implementation which +uses the kernel page cache. Because the page cache operates on page sized +units this may introduce additional complexity in terms of locking and +associated race conditions. diff --git a/Documentation/filesystems/sysfs-pci.txt b/Documentation/filesystems/sysfs-pci.txt new file mode 100644 index 00000000..74eaac26 --- /dev/null +++ b/Documentation/filesystems/sysfs-pci.txt @@ -0,0 +1,120 @@ +Accessing PCI device resources through sysfs +-------------------------------------------- + +sysfs, usually mounted at /sys, provides access to PCI resources on platforms +that support it. For example, a given bus might look like this: + + /sys/devices/pci0000:17 + |-- 0000:17:00.0 + | |-- class + | |-- config + | |-- device + | |-- enable + | |-- irq + | |-- local_cpus + | |-- remove + | |-- resource + | |-- resource0 + | |-- resource1 + | |-- resource2 + | |-- rom + | |-- subsystem_device + | |-- subsystem_vendor + | `-- vendor + `-- ... + +The topmost element describes the PCI domain and bus number. In this case, +the domain number is 0000 and the bus number is 17 (both values are in hex). +This bus contains a single function device in slot 0. The domain and bus +numbers are reproduced for convenience. Under the device directory are several +files, each with their own function. + + file function + ---- -------- + class PCI class (ascii, ro) + config PCI config space (binary, rw) + device PCI device (ascii, ro) + enable Whether the device is enabled (ascii, rw) + irq IRQ number (ascii, ro) + local_cpus nearby CPU mask (cpumask, ro) + remove remove device from kernel's list (ascii, wo) + resource PCI resource host addresses (ascii, ro) + resource0..N PCI resource N, if present (binary, mmap, rw[1]) + resource0_wc..N_wc PCI WC map resource N, if prefetchable (binary, mmap) + rom PCI ROM resource, if present (binary, ro) + subsystem_device PCI subsystem device (ascii, ro) + subsystem_vendor PCI subsystem vendor (ascii, ro) + vendor PCI vendor (ascii, ro) + + ro - read only file + rw - file is readable and writable + wo - write only file + mmap - file is mmapable + ascii - file contains ascii text + binary - file contains binary data + cpumask - file contains a cpumask type + +[1] rw for RESOURCE_IO (I/O port) regions only + +The read only files are informational, writes to them will be ignored, with +the exception of the 'rom' file. Writable files can be used to perform +actions on the device (e.g. changing config space, detaching a device). +mmapable files are available via an mmap of the file at offset 0 and can be +used to do actual device programming from userspace. Note that some platforms +don't support mmapping of certain resources, so be sure to check the return +value from any attempted mmap. The most notable of these are I/O port +resources, which also provide read/write access. + +The 'enable' file provides a counter that indicates how many times the device +has been enabled. If the 'enable' file currently returns '4', and a '1' is +echoed into it, it will then return '5'. Echoing a '0' into it will decrease +the count. Even when it returns to 0, though, some of the initialisation +may not be reversed. + +The 'rom' file is special in that it provides read-only access to the device's +ROM file, if available. It's disabled by default, however, so applications +should write the string "1" to the file to enable it before attempting a read +call, and disable it following the access by writing "0" to the file. Note +that the device must be enabled for a rom read to return data successfully. +In the event a driver is not bound to the device, it can be enabled using the +'enable' file, documented above. + +The 'remove' file is used to remove the PCI device, by writing a non-zero +integer to the file. This does not involve any kind of hot-plug functionality, +e.g. powering off the device. The device is removed from the kernel's list of +PCI devices, the sysfs directory for it is removed, and the device will be +removed from any drivers attached to it. Removal of PCI root buses is +disallowed. + +Accessing legacy resources through sysfs +---------------------------------------- + +Legacy I/O port and ISA memory resources are also provided in sysfs if the +underlying platform supports them. They're located in the PCI class hierarchy, +e.g. + + /sys/class/pci_bus/0000:17/ + |-- bridge -> ../../../devices/pci0000:17 + |-- cpuaffinity + |-- legacy_io + `-- legacy_mem + +The legacy_io file is a read/write file that can be used by applications to +do legacy port I/O. The application should open the file, seek to the desired +port (e.g. 0x3e8) and do a read or a write of 1, 2 or 4 bytes. The legacy_mem +file should be mmapped with an offset corresponding to the memory offset +desired, e.g. 0xa0000 for the VGA frame buffer. The application can then +simply dereference the returned pointer (after checking for errors of course) +to access legacy memory space. + +Supporting PCI access on new platforms +-------------------------------------- + +In order to support PCI resource mapping as described above, Linux platform +code must define HAVE_PCI_MMAP and provide a pci_mmap_page_range function. +Platforms are free to only support subsets of the mmap functionality, but +useful return codes should be provided. + +Legacy resources are protected by the HAVE_PCI_LEGACY define. Platforms +wishing to support legacy functionality should define it and provide +pci_legacy_read, pci_legacy_write and pci_mmap_legacy_page_range functions. diff --git a/Documentation/filesystems/sysfs-tagging.txt b/Documentation/filesystems/sysfs-tagging.txt new file mode 100644 index 00000000..caaaf126 --- /dev/null +++ b/Documentation/filesystems/sysfs-tagging.txt @@ -0,0 +1,42 @@ +Sysfs tagging +------------- + +(Taken almost verbatim from Eric Biederman's netns tagging patch +commit msg) + +The problem. Network devices show up in sysfs and with the network +namespace active multiple devices with the same name can show up in +the same directory, ouch! + +To avoid that problem and allow existing applications in network +namespaces to see the same interface that is currently presented in +sysfs, sysfs now has tagging directory support. + +By using the network namespace pointers as tags to separate out the +the sysfs directory entries we ensure that we don't have conflicts +in the directories and applications only see a limited set of +the network devices. + +Each sysfs directory entry may be tagged with zero or one +namespaces. A sysfs_dirent is augmented with a void *s_ns. If a +directory entry is tagged, then sysfs_dirent->s_flags will have a +flag between KOBJ_NS_TYPE_NONE and KOBJ_NS_TYPES, and s_ns will +point to the namespace to which it belongs. + +Each sysfs superblock's sysfs_super_info contains an array void +*ns[KOBJ_NS_TYPES]. When a a task in a tagging namespace +kobj_nstype first mounts sysfs, a new superblock is created. It +will be differentiated from other sysfs mounts by having its +s_fs_info->ns[kobj_nstype] set to the new namespace. Note that +through bind mounting and mounts propagation, a task can easily view +the contents of other namespaces' sysfs mounts. Therefore, when a +namespace exits, it will call kobj_ns_exit() to invalidate any +sysfs_dirent->s_ns pointers pointing to it. + +Users of this interface: +- define a type in the kobj_ns_type enumeration. +- call kobj_ns_type_register() with its kobj_ns_type_operations which has + - current_ns() which returns current's namespace + - netlink_ns() which returns a socket's namespace + - initial_ns() which returns the initial namesapce +- call kobj_ns_exit() when an individual tag is no longer valid diff --git a/Documentation/filesystems/sysfs.txt b/Documentation/filesystems/sysfs.txt new file mode 100644 index 00000000..597f728e --- /dev/null +++ b/Documentation/filesystems/sysfs.txt @@ -0,0 +1,372 @@ + +sysfs - _The_ filesystem for exporting kernel objects. + +Patrick Mochel <mochel@osdl.org> +Mike Murphy <mamurph@cs.clemson.edu> + +Revised: 15 July 2010 +Original: 10 January 2003 + + +What it is: +~~~~~~~~~~~ + +sysfs is a ram-based filesystem initially based on ramfs. It provides +a means to export kernel data structures, their attributes, and the +linkages between them to userspace. + +sysfs is tied inherently to the kobject infrastructure. Please read +Documentation/kobject.txt for more information concerning the kobject +interface. + + +Using sysfs +~~~~~~~~~~~ + +sysfs is always compiled in if CONFIG_SYSFS is defined. You can access +it by doing: + + mount -t sysfs sysfs /sys + + +Directory Creation +~~~~~~~~~~~~~~~~~~ + +For every kobject that is registered with the system, a directory is +created for it in sysfs. That directory is created as a subdirectory +of the kobject's parent, expressing internal object hierarchies to +userspace. Top-level directories in sysfs represent the common +ancestors of object hierarchies; i.e. the subsystems the objects +belong to. + +Sysfs internally stores a pointer to the kobject that implements a +directory in the sysfs_dirent object associated with the directory. In +the past this kobject pointer has been used by sysfs to do reference +counting directly on the kobject whenever the file is opened or closed. +With the current sysfs implementation the kobject reference count is +only modified directly by the function sysfs_schedule_callback(). + + +Attributes +~~~~~~~~~~ + +Attributes can be exported for kobjects in the form of regular files in +the filesystem. Sysfs forwards file I/O operations to methods defined +for the attributes, providing a means to read and write kernel +attributes. + +Attributes should be ASCII text files, preferably with only one value +per file. It is noted that it may not be efficient to contain only one +value per file, so it is socially acceptable to express an array of +values of the same type. + +Mixing types, expressing multiple lines of data, and doing fancy +formatting of data is heavily frowned upon. Doing these things may get +you publicly humiliated and your code rewritten without notice. + + +An attribute definition is simply: + +struct attribute { + char * name; + struct module *owner; + mode_t mode; +}; + + +int sysfs_create_file(struct kobject * kobj, const struct attribute * attr); +void sysfs_remove_file(struct kobject * kobj, const struct attribute * attr); + + +A bare attribute contains no means to read or write the value of the +attribute. Subsystems are encouraged to define their own attribute +structure and wrapper functions for adding and removing attributes for +a specific object type. + +For example, the driver model defines struct device_attribute like: + +struct device_attribute { + struct attribute attr; + ssize_t (*show)(struct device *dev, struct device_attribute *attr, + char *buf); + ssize_t (*store)(struct device *dev, struct device_attribute *attr, + const char *buf, size_t count); +}; + +int device_create_file(struct device *, const struct device_attribute *); +void device_remove_file(struct device *, const struct device_attribute *); + +It also defines this helper for defining device attributes: + +#define DEVICE_ATTR(_name, _mode, _show, _store) \ +struct device_attribute dev_attr_##_name = __ATTR(_name, _mode, _show, _store) + +For example, declaring + +static DEVICE_ATTR(foo, S_IWUSR | S_IRUGO, show_foo, store_foo); + +is equivalent to doing: + +static struct device_attribute dev_attr_foo = { + .attr = { + .name = "foo", + .mode = S_IWUSR | S_IRUGO, + .show = show_foo, + .store = store_foo, + }, +}; + + +Subsystem-Specific Callbacks +~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + +When a subsystem defines a new attribute type, it must implement a +set of sysfs operations for forwarding read and write calls to the +show and store methods of the attribute owners. + +struct sysfs_ops { + ssize_t (*show)(struct kobject *, struct attribute *, char *); + ssize_t (*store)(struct kobject *, struct attribute *, const char *, size_t); +}; + +[ Subsystems should have already defined a struct kobj_type as a +descriptor for this type, which is where the sysfs_ops pointer is +stored. See the kobject documentation for more information. ] + +When a file is read or written, sysfs calls the appropriate method +for the type. The method then translates the generic struct kobject +and struct attribute pointers to the appropriate pointer types, and +calls the associated methods. + + +To illustrate: + +#define to_dev(obj) container_of(obj, struct device, kobj) +#define to_dev_attr(_attr) container_of(_attr, struct device_attribute, attr) + +static ssize_t dev_attr_show(struct kobject *kobj, struct attribute *attr, + char *buf) +{ + struct device_attribute *dev_attr = to_dev_attr(attr); + struct device *dev = to_dev(kobj); + ssize_t ret = -EIO; + + if (dev_attr->show) + ret = dev_attr->show(dev, dev_attr, buf); + if (ret >= (ssize_t)PAGE_SIZE) { + print_symbol("dev_attr_show: %s returned bad count\n", + (unsigned long)dev_attr->show); + } + return ret; +} + + + +Reading/Writing Attribute Data +~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ + +To read or write attributes, show() or store() methods must be +specified when declaring the attribute. The method types should be as +simple as those defined for device attributes: + +ssize_t (*show)(struct device *dev, struct device_attribute *attr, char *buf); +ssize_t (*store)(struct device *dev, struct device_attribute *attr, + const char *buf, size_t count); + +IOW, they should take only an object, an attribute, and a buffer as parameters. + + +sysfs allocates a buffer of size (PAGE_SIZE) and passes it to the +method. Sysfs will call the method exactly once for each read or +write. This forces the following behavior on the method +implementations: + +- On read(2), the show() method should fill the entire buffer. + Recall that an attribute should only be exporting one value, or an + array of similar values, so this shouldn't be that expensive. + + This allows userspace to do partial reads and forward seeks + arbitrarily over the entire file at will. If userspace seeks back to + zero or does a pread(2) with an offset of '0' the show() method will + be called again, rearmed, to fill the buffer. + +- On write(2), sysfs expects the entire buffer to be passed during the + first write. Sysfs then passes the entire buffer to the store() + method. + + When writing sysfs files, userspace processes should first read the + entire file, modify the values it wishes to change, then write the + entire buffer back. + + Attribute method implementations should operate on an identical + buffer when reading and writing values. + +Other notes: + +- Writing causes the show() method to be rearmed regardless of current + file position. + +- The buffer will always be PAGE_SIZE bytes in length. On i386, this + is 4096. + +- show() methods should return the number of bytes printed into the + buffer. This is the return value of scnprintf(). + +- show() should always use scnprintf(). + +- store() should return the number of bytes used from the buffer. If the + entire buffer has been used, just return the count argument. + +- show() or store() can always return errors. If a bad value comes + through, be sure to return an error. + +- The object passed to the methods will be pinned in memory via sysfs + referencing counting its embedded object. However, the physical + entity (e.g. device) the object represents may not be present. Be + sure to have a way to check this, if necessary. + + +A very simple (and naive) implementation of a device attribute is: + +static ssize_t show_name(struct device *dev, struct device_attribute *attr, + char *buf) +{ + return scnprintf(buf, PAGE_SIZE, "%s\n", dev->name); +} + +static ssize_t store_name(struct device *dev, struct device_attribute *attr, + const char *buf, size_t count) +{ + snprintf(dev->name, sizeof(dev->name), "%.*s", + (int)min(count, sizeof(dev->name) - 1), buf); + return count; +} + +static DEVICE_ATTR(name, S_IRUGO, show_name, store_name); + + +(Note that the real implementation doesn't allow userspace to set the +name for a device.) + + +Top Level Directory Layout +~~~~~~~~~~~~~~~~~~~~~~~~~~ + +The sysfs directory arrangement exposes the relationship of kernel +data structures. + +The top level sysfs directory looks like: + +block/ +bus/ +class/ +dev/ +devices/ +firmware/ +net/ +fs/ + +devices/ contains a filesystem representation of the device tree. It maps +directly to the internal kernel device tree, which is a hierarchy of +struct device. + +bus/ contains flat directory layout of the various bus types in the +kernel. Each bus's directory contains two subdirectories: + + devices/ + drivers/ + +devices/ contains symlinks for each device discovered in the system +that point to the device's directory under root/. + +drivers/ contains a directory for each device driver that is loaded +for devices on that particular bus (this assumes that drivers do not +span multiple bus types). + +fs/ contains a directory for some filesystems. Currently each +filesystem wanting to export attributes must create its own hierarchy +below fs/ (see ./fuse.txt for an example). + +dev/ contains two directories char/ and block/. Inside these two +directories there are symlinks named <major>:<minor>. These symlinks +point to the sysfs directory for the given device. /sys/dev provides a +quick way to lookup the sysfs interface for a device from the result of +a stat(2) operation. + +More information can driver-model specific features can be found in +Documentation/driver-model/. + + +TODO: Finish this section. + + +Current Interfaces +~~~~~~~~~~~~~~~~~~ + +The following interface layers currently exist in sysfs: + + +- devices (include/linux/device.h) +---------------------------------- +Structure: + +struct device_attribute { + struct attribute attr; + ssize_t (*show)(struct device *dev, struct device_attribute *attr, + char *buf); + ssize_t (*store)(struct device *dev, struct device_attribute *attr, + const char *buf, size_t count); +}; + +Declaring: + +DEVICE_ATTR(_name, _mode, _show, _store); + +Creation/Removal: + +int device_create_file(struct device *dev, const struct device_attribute * attr); +void device_remove_file(struct device *dev, const struct device_attribute * attr); + + +- bus drivers (include/linux/device.h) +-------------------------------------- +Structure: + +struct bus_attribute { + struct attribute attr; + ssize_t (*show)(struct bus_type *, char * buf); + ssize_t (*store)(struct bus_type *, const char * buf, size_t count); +}; + +Declaring: + +BUS_ATTR(_name, _mode, _show, _store) + +Creation/Removal: + +int bus_create_file(struct bus_type *, struct bus_attribute *); +void bus_remove_file(struct bus_type *, struct bus_attribute *); + + +- device drivers (include/linux/device.h) +----------------------------------------- + +Structure: + +struct driver_attribute { + struct attribute attr; + ssize_t (*show)(struct device_driver *, char * buf); + ssize_t (*store)(struct device_driver *, const char * buf, + size_t count); +}; + +Declaring: + +DRIVER_ATTR(_name, _mode, _show, _store) + +Creation/Removal: + +int driver_create_file(struct device_driver *, const struct driver_attribute *); +void driver_remove_file(struct device_driver *, const struct driver_attribute *); + + diff --git a/Documentation/filesystems/sysv-fs.txt b/Documentation/filesystems/sysv-fs.txt new file mode 100644 index 00000000..253b50d1 --- /dev/null +++ b/Documentation/filesystems/sysv-fs.txt @@ -0,0 +1,197 @@ +It implements all of + - Xenix FS, + - SystemV/386 FS, + - Coherent FS. + +To install: +* Answer the 'System V and Coherent filesystem support' question with 'y' + when configuring the kernel. +* To mount a disk or a partition, use + mount [-r] -t sysv device mountpoint + The file system type names + -t sysv + -t xenix + -t coherent + may be used interchangeably, but the last two will eventually disappear. + +Bugs in the present implementation: +- Coherent FS: + - The "free list interleave" n:m is currently ignored. + - Only file systems with no filesystem name and no pack name are recognized. + (See Coherent "man mkfs" for a description of these features.) +- SystemV Release 2 FS: + The superblock is only searched in the blocks 9, 15, 18, which + corresponds to the beginning of track 1 on floppy disks. No support + for this FS on hard disk yet. + + +These filesystems are rather similar. Here is a comparison with Minix FS: + +* Linux fdisk reports on partitions + - Minix FS 0x81 Linux/Minix + - Xenix FS ?? + - SystemV FS ?? + - Coherent FS 0x08 AIX bootable + +* Size of a block or zone (data allocation unit on disk) + - Minix FS 1024 + - Xenix FS 1024 (also 512 ??) + - SystemV FS 1024 (also 512 and 2048) + - Coherent FS 512 + +* General layout: all have one boot block, one super block and + separate areas for inodes and for directories/data. + On SystemV Release 2 FS (e.g. Microport) the first track is reserved and + all the block numbers (including the super block) are offset by one track. + +* Byte ordering of "short" (16 bit entities) on disk: + - Minix FS little endian 0 1 + - Xenix FS little endian 0 1 + - SystemV FS little endian 0 1 + - Coherent FS little endian 0 1 + Of course, this affects only the file system, not the data of files on it! + +* Byte ordering of "long" (32 bit entities) on disk: + - Minix FS little endian 0 1 2 3 + - Xenix FS little endian 0 1 2 3 + - SystemV FS little endian 0 1 2 3 + - Coherent FS PDP-11 2 3 0 1 + Of course, this affects only the file system, not the data of files on it! + +* Inode on disk: "short", 0 means non-existent, the root dir ino is: + - Minix FS 1 + - Xenix FS, SystemV FS, Coherent FS 2 + +* Maximum number of hard links to a file: + - Minix FS 250 + - Xenix FS ?? + - SystemV FS ?? + - Coherent FS >=10000 + +* Free inode management: + - Minix FS a bitmap + - Xenix FS, SystemV FS, Coherent FS + There is a cache of a certain number of free inodes in the super-block. + When it is exhausted, new free inodes are found using a linear search. + +* Free block management: + - Minix FS a bitmap + - Xenix FS, SystemV FS, Coherent FS + Free blocks are organized in a "free list". Maybe a misleading term, + since it is not true that every free block contains a pointer to + the next free block. Rather, the free blocks are organized in chunks + of limited size, and every now and then a free block contains pointers + to the free blocks pertaining to the next chunk; the first of these + contains pointers and so on. The list terminates with a "block number" + 0 on Xenix FS and SystemV FS, with a block zeroed out on Coherent FS. + +* Super-block location: + - Minix FS block 1 = bytes 1024..2047 + - Xenix FS block 1 = bytes 1024..2047 + - SystemV FS bytes 512..1023 + - Coherent FS block 1 = bytes 512..1023 + +* Super-block layout: + - Minix FS + unsigned short s_ninodes; + unsigned short s_nzones; + unsigned short s_imap_blocks; + unsigned short s_zmap_blocks; + unsigned short s_firstdatazone; + unsigned short s_log_zone_size; + unsigned long s_max_size; + unsigned short s_magic; + - Xenix FS, SystemV FS, Coherent FS + unsigned short s_firstdatazone; + unsigned long s_nzones; + unsigned short s_fzone_count; + unsigned long s_fzones[NICFREE]; + unsigned short s_finode_count; + unsigned short s_finodes[NICINOD]; + char s_flock; + char s_ilock; + char s_modified; + char s_rdonly; + unsigned long s_time; + short s_dinfo[4]; -- SystemV FS only + unsigned long s_free_zones; + unsigned short s_free_inodes; + short s_dinfo[4]; -- Xenix FS only + unsigned short s_interleave_m,s_interleave_n; -- Coherent FS only + char s_fname[6]; + char s_fpack[6]; + then they differ considerably: + Xenix FS + char s_clean; + char s_fill[371]; + long s_magic; + long s_type; + SystemV FS + long s_fill[12 or 14]; + long s_state; + long s_magic; + long s_type; + Coherent FS + unsigned long s_unique; + Note that Coherent FS has no magic. + +* Inode layout: + - Minix FS + unsigned short i_mode; + unsigned short i_uid; + unsigned long i_size; + unsigned long i_time; + unsigned char i_gid; + unsigned char i_nlinks; + unsigned short i_zone[7+1+1]; + - Xenix FS, SystemV FS, Coherent FS + unsigned short i_mode; + unsigned short i_nlink; + unsigned short i_uid; + unsigned short i_gid; + unsigned long i_size; + unsigned char i_zone[3*(10+1+1+1)]; + unsigned long i_atime; + unsigned long i_mtime; + unsigned long i_ctime; + +* Regular file data blocks are organized as + - Minix FS + 7 direct blocks + 1 indirect block (pointers to blocks) + 1 double-indirect block (pointer to pointers to blocks) + - Xenix FS, SystemV FS, Coherent FS + 10 direct blocks + 1 indirect block (pointers to blocks) + 1 double-indirect block (pointer to pointers to blocks) + 1 triple-indirect block (pointer to pointers to pointers to blocks) + +* Inode size, inodes per block + - Minix FS 32 32 + - Xenix FS 64 16 + - SystemV FS 64 16 + - Coherent FS 64 8 + +* Directory entry on disk + - Minix FS + unsigned short inode; + char name[14/30]; + - Xenix FS, SystemV FS, Coherent FS + unsigned short inode; + char name[14]; + +* Dir entry size, dir entries per block + - Minix FS 16/32 64/32 + - Xenix FS 16 64 + - SystemV FS 16 64 + - Coherent FS 16 32 + +* How to implement symbolic links such that the host fsck doesn't scream: + - Minix FS normal + - Xenix FS kludge: as regular files with chmod 1000 + - SystemV FS ?? + - Coherent FS kludge: as regular files with chmod 1000 + + +Notation: We often speak of a "block" but mean a zone (the allocation unit) +and not the disk driver's notion of "block". diff --git a/Documentation/filesystems/tmpfs.txt b/Documentation/filesystems/tmpfs.txt new file mode 100644 index 00000000..98ef5512 --- /dev/null +++ b/Documentation/filesystems/tmpfs.txt @@ -0,0 +1,148 @@ +Tmpfs is a file system which keeps all files in virtual memory. + + +Everything in tmpfs is temporary in the sense that no files will be +created on your hard drive. If you unmount a tmpfs instance, +everything stored therein is lost. + +tmpfs puts everything into the kernel internal caches and grows and +shrinks to accommodate the files it contains and is able to swap +unneeded pages out to swap space. It has maximum size limits which can +be adjusted on the fly via 'mount -o remount ...' + +If you compare it to ramfs (which was the template to create tmpfs) +you gain swapping and limit checking. Another similar thing is the RAM +disk (/dev/ram*), which simulates a fixed size hard disk in physical +RAM, where you have to create an ordinary filesystem on top. Ramdisks +cannot swap and you do not have the possibility to resize them. + +Since tmpfs lives completely in the page cache and on swap, all tmpfs +pages currently in memory will show up as cached. It will not show up +as shared or something like that. Further on you can check the actual +RAM+swap use of a tmpfs instance with df(1) and du(1). + + +tmpfs has the following uses: + +1) There is always a kernel internal mount which you will not see at + all. This is used for shared anonymous mappings and SYSV shared + memory. + + This mount does not depend on CONFIG_TMPFS. If CONFIG_TMPFS is not + set, the user visible part of tmpfs is not build. But the internal + mechanisms are always present. + +2) glibc 2.2 and above expects tmpfs to be mounted at /dev/shm for + POSIX shared memory (shm_open, shm_unlink). Adding the following + line to /etc/fstab should take care of this: + + tmpfs /dev/shm tmpfs defaults 0 0 + + Remember to create the directory that you intend to mount tmpfs on + if necessary. + + This mount is _not_ needed for SYSV shared memory. The internal + mount is used for that. (In the 2.3 kernel versions it was + necessary to mount the predecessor of tmpfs (shm fs) to use SYSV + shared memory) + +3) Some people (including me) find it very convenient to mount it + e.g. on /tmp and /var/tmp and have a big swap partition. And now + loop mounts of tmpfs files do work, so mkinitrd shipped by most + distributions should succeed with a tmpfs /tmp. + +4) And probably a lot more I do not know about :-) + + +tmpfs has three mount options for sizing: + +size: The limit of allocated bytes for this tmpfs instance. The + default is half of your physical RAM without swap. If you + oversize your tmpfs instances the machine will deadlock + since the OOM handler will not be able to free that memory. +nr_blocks: The same as size, but in blocks of PAGE_CACHE_SIZE. +nr_inodes: The maximum number of inodes for this instance. The default + is half of the number of your physical RAM pages, or (on a + machine with highmem) the number of lowmem RAM pages, + whichever is the lower. + +These parameters accept a suffix k, m or g for kilo, mega and giga and +can be changed on remount. The size parameter also accepts a suffix % +to limit this tmpfs instance to that percentage of your physical RAM: +the default, when neither size nor nr_blocks is specified, is size=50% + +If nr_blocks=0 (or size=0), blocks will not be limited in that instance; +if nr_inodes=0, inodes will not be limited. It is generally unwise to +mount with such options, since it allows any user with write access to +use up all the memory on the machine; but enhances the scalability of +that instance in a system with many cpus making intensive use of it. + + +tmpfs has a mount option to set the NUMA memory allocation policy for +all files in that instance (if CONFIG_NUMA is enabled) - which can be +adjusted on the fly via 'mount -o remount ...' + +mpol=default use the process allocation policy + (see set_mempolicy(2)) +mpol=prefer:Node prefers to allocate memory from the given Node +mpol=bind:NodeList allocates memory only from nodes in NodeList +mpol=interleave prefers to allocate from each node in turn +mpol=interleave:NodeList allocates from each node of NodeList in turn +mpol=local prefers to allocate memory from the local node + +NodeList format is a comma-separated list of decimal numbers and ranges, +a range being two hyphen-separated decimal numbers, the smallest and +largest node numbers in the range. For example, mpol=bind:0-3,5,7,9-15 + +A memory policy with a valid NodeList will be saved, as specified, for +use at file creation time. When a task allocates a file in the file +system, the mount option memory policy will be applied with a NodeList, +if any, modified by the calling task's cpuset constraints +[See Documentation/cgroups/cpusets.txt] and any optional flags, listed +below. If the resulting NodeLists is the empty set, the effective memory +policy for the file will revert to "default" policy. + +NUMA memory allocation policies have optional flags that can be used in +conjunction with their modes. These optional flags can be specified +when tmpfs is mounted by appending them to the mode before the NodeList. +See Documentation/vm/numa_memory_policy.txt for a list of all available +memory allocation policy mode flags and their effect on memory policy. + + =static is equivalent to MPOL_F_STATIC_NODES + =relative is equivalent to MPOL_F_RELATIVE_NODES + +For example, mpol=bind=static:NodeList, is the equivalent of an +allocation policy of MPOL_BIND | MPOL_F_STATIC_NODES. + +Note that trying to mount a tmpfs with an mpol option will fail if the +running kernel does not support NUMA; and will fail if its nodelist +specifies a node which is not online. If your system relies on that +tmpfs being mounted, but from time to time runs a kernel built without +NUMA capability (perhaps a safe recovery kernel), or with fewer nodes +online, then it is advisable to omit the mpol option from automatic +mount options. It can be added later, when the tmpfs is already mounted +on MountPoint, by 'mount -o remount,mpol=Policy:NodeList MountPoint'. + + +To specify the initial root directory you can use the following mount +options: + +mode: The permissions as an octal number +uid: The user id +gid: The group id + +These options do not have any effect on remount. You can change these +parameters with chmod(1), chown(1) and chgrp(1) on a mounted filesystem. + + +So 'mount -t tmpfs -o size=10G,nr_inodes=10k,mode=700 tmpfs /mytmpfs' +will give you tmpfs instance on /mytmpfs which can allocate 10GB +RAM/SWAP in 10240 inodes and it is only accessible by root. + + +Author: + Christoph Rohland <cr@sap.com>, 1.12.01 +Updated: + Hugh Dickins, 4 June 2007 +Updated: + KOSAKI Motohiro, 16 Mar 2010 diff --git a/Documentation/filesystems/ubifs.txt b/Documentation/filesystems/ubifs.txt new file mode 100644 index 00000000..8e4fab63 --- /dev/null +++ b/Documentation/filesystems/ubifs.txt @@ -0,0 +1,147 @@ +Introduction +============= + +UBIFS file-system stands for UBI File System. UBI stands for "Unsorted +Block Images". UBIFS is a flash file system, which means it is designed +to work with flash devices. It is important to understand, that UBIFS +is completely different to any traditional file-system in Linux, like +Ext2, XFS, JFS, etc. UBIFS represents a separate class of file-systems +which work with MTD devices, not block devices. The other Linux +file-system of this class is JFFS2. + +To make it more clear, here is a small comparison of MTD devices and +block devices. + +1 MTD devices represent flash devices and they consist of eraseblocks of + rather large size, typically about 128KiB. Block devices consist of + small blocks, typically 512 bytes. +2 MTD devices support 3 main operations - read from some offset within an + eraseblock, write to some offset within an eraseblock, and erase a whole + eraseblock. Block devices support 2 main operations - read a whole + block and write a whole block. +3 The whole eraseblock has to be erased before it becomes possible to + re-write its contents. Blocks may be just re-written. +4 Eraseblocks become worn out after some number of erase cycles - + typically 100K-1G for SLC NAND and NOR flashes, and 1K-10K for MLC + NAND flashes. Blocks do not have the wear-out property. +5 Eraseblocks may become bad (only on NAND flashes) and software should + deal with this. Blocks on hard drives typically do not become bad, + because hardware has mechanisms to substitute bad blocks, at least in + modern LBA disks. + +It should be quite obvious why UBIFS is very different to traditional +file-systems. + +UBIFS works on top of UBI. UBI is a separate software layer which may be +found in drivers/mtd/ubi. UBI is basically a volume management and +wear-leveling layer. It provides so called UBI volumes which is a higher +level abstraction than a MTD device. The programming model of UBI devices +is very similar to MTD devices - they still consist of large eraseblocks, +they have read/write/erase operations, but UBI devices are devoid of +limitations like wear and bad blocks (items 4 and 5 in the above list). + +In a sense, UBIFS is a next generation of JFFS2 file-system, but it is +very different and incompatible to JFFS2. The following are the main +differences. + +* JFFS2 works on top of MTD devices, UBIFS depends on UBI and works on + top of UBI volumes. +* JFFS2 does not have on-media index and has to build it while mounting, + which requires full media scan. UBIFS maintains the FS indexing + information on the flash media and does not require full media scan, + so it mounts many times faster than JFFS2. +* JFFS2 is a write-through file-system, while UBIFS supports write-back, + which makes UBIFS much faster on writes. + +Similarly to JFFS2, UBIFS supports on-the-flight compression which makes +it possible to fit quite a lot of data to the flash. + +Similarly to JFFS2, UBIFS is tolerant of unclean reboots and power-cuts. +It does not need stuff like fsck.ext2. UBIFS automatically replays its +journal and recovers from crashes, ensuring that the on-flash data +structures are consistent. + +UBIFS scales logarithmically (most of the data structures it uses are +trees), so the mount time and memory consumption do not linearly depend +on the flash size, like in case of JFFS2. This is because UBIFS +maintains the FS index on the flash media. However, UBIFS depends on +UBI, which scales linearly. So overall UBI/UBIFS stack scales linearly. +Nevertheless, UBI/UBIFS scales considerably better than JFFS2. + +The authors of UBIFS believe, that it is possible to develop UBI2 which +would scale logarithmically as well. UBI2 would support the same API as UBI, +but it would be binary incompatible to UBI. So UBIFS would not need to be +changed to use UBI2 + + +Mount options +============= + +(*) == default. + +bulk_read read more in one go to take advantage of flash + media that read faster sequentially +no_bulk_read (*) do not bulk-read +no_chk_data_crc (*) skip checking of CRCs on data nodes in order to + improve read performance. Use this option only + if the flash media is highly reliable. The effect + of this option is that corruption of the contents + of a file can go unnoticed. +chk_data_crc do not skip checking CRCs on data nodes +compr=none override default compressor and set it to "none" +compr=lzo override default compressor and set it to "lzo" +compr=zlib override default compressor and set it to "zlib" + + +Quick usage instructions +======================== + +The UBI volume to mount is specified using "ubiX_Y" or "ubiX:NAME" syntax, +where "X" is UBI device number, "Y" is UBI volume number, and "NAME" is +UBI volume name. + +Mount volume 0 on UBI device 0 to /mnt/ubifs: +$ mount -t ubifs ubi0_0 /mnt/ubifs + +Mount "rootfs" volume of UBI device 0 to /mnt/ubifs ("rootfs" is volume +name): +$ mount -t ubifs ubi0:rootfs /mnt/ubifs + +The following is an example of the kernel boot arguments to attach mtd0 +to UBI and mount volume "rootfs": +ubi.mtd=0 root=ubi0:rootfs rootfstype=ubifs + + +Module Parameters for Debugging +=============================== + +When UBIFS has been compiled with debugging enabled, there are 2 module +parameters that are available to control aspects of testing and debugging. + +debug_chks Selects extra checks that UBIFS can do while running: + + Check Flag value + + General checks 1 + Check Tree Node Cache (TNC) 2 + Check indexing tree size 4 + Check orphan area 8 + Check old indexing tree 16 + Check LEB properties (lprops) 32 + Check leaf nodes and inodes 64 + +debug_tsts Selects a mode of testing, as follows: + + Test mode Flag value + + Failure mode for recovery testing 4 + +For example, set debug_chks to 3 to enable general and TNC checks. + + +References +========== + +UBIFS documentation and FAQ/HOWTO at the MTD web site: +http://www.linux-mtd.infradead.org/doc/ubifs.html +http://www.linux-mtd.infradead.org/faq/ubifs.html diff --git a/Documentation/filesystems/udf.txt b/Documentation/filesystems/udf.txt new file mode 100644 index 00000000..902b95d0 --- /dev/null +++ b/Documentation/filesystems/udf.txt @@ -0,0 +1,82 @@ +* +* Documentation/filesystems/udf.txt +* +UDF Filesystem version 0.9.8.1 + +If you encounter problems with reading UDF discs using this driver, +please report them to linux_udf@hpesjro.fc.hp.com, which is the +developer's list. + +Write support requires a block driver which supports writing. Currently +dvd+rw drives and media support true random sector writes, and so a udf +filesystem on such devices can be directly mounted read/write. CD-RW +media however, does not support this. Instead the media can be formatted +for packet mode using the utility cdrwtool, then the pktcdvd driver can +be bound to the underlying cd device to provide the required buffering +and read-modify-write cycles to allow the filesystem random sector writes +while providing the hardware with only full packet writes. While not +required for dvd+rw media, use of the pktcdvd driver often enhances +performance due to very poor read-modify-write support supplied internally +by drive firmware. + +------------------------------------------------------------------------------- +The following mount options are supported: + + gid= Set the default group. + umask= Set the default umask. + mode= Set the default file permissions. + dmode= Set the default directory permissions. + uid= Set the default user. + bs= Set the block size. + unhide Show otherwise hidden files. + undelete Show deleted files in lists. + adinicb Embed data in the inode (default) + noadinicb Don't embed data in the inode + shortad Use short ad's + longad Use long ad's (default) + nostrict Unset strict conformance + iocharset= Set the NLS character set + +The uid= and gid= options need a bit more explaining. They will accept a +decimal numeric value which will be used as the default ID for that mount. +They will also accept the string "ignore" and "forget". For files on the disk +that are owned by nobody ( -1 ), they will instead look as if they are owned +by the default ID. The ignore option causes the default ID to override all +IDs on the disk, not just -1. The forget option causes all IDs to be written +to disk as -1, so when the media is later remounted, they will appear to be +owned by whatever default ID it is mounted with at that time. + +For typical desktop use of removable media, you should set the ID to that +of the interactively logged on user, and also specify both the forget and +ignore options. This way the interactive user will always see the files +on the disk as belonging to him. + +The remaining are for debugging and disaster recovery: + + novrs Skip volume sequence recognition + +The following expect a offset from 0. + + session= Set the CDROM session (default= last session) + anchor= Override standard anchor location. (default= 256) + volume= Override the VolumeDesc location. (unused) + partition= Override the PartitionDesc location. (unused) + lastblock= Set the last block of the filesystem/ + +The following expect a offset from the partition root. + + fileset= Override the fileset block location. (unused) + rootdir= Override the root directory location. (unused) + WARNING: overriding the rootdir to a non-directory may + yield highly unpredictable results. +------------------------------------------------------------------------------- + + +For the latest version and toolset see: + http://linux-udf.sourceforge.net/ + +Documentation on UDF and ECMA 167 is available FREE from: + http://www.osta.org/ + http://www.ecma-international.org/ + +Ben Fennema <bfennema@falcon.csc.calpoly.edu> diff --git a/Documentation/filesystems/ufs.txt b/Documentation/filesystems/ufs.txt new file mode 100644 index 00000000..7a602ade --- /dev/null +++ b/Documentation/filesystems/ufs.txt @@ -0,0 +1,60 @@ +USING UFS +========= + +mount -t ufs -o ufstype=type_of_ufs device dir + + +UFS OPTIONS +=========== + +ufstype=type_of_ufs + UFS is a file system widely used in different operating systems. + The problem are differences among implementations. Features of + some implementations are undocumented, so its hard to recognize + type of ufs automatically. That's why user must specify type of + ufs manually by mount option ufstype. Possible values are: + + old old format of ufs + default value, supported as read-only + + 44bsd used in FreeBSD, NetBSD, OpenBSD + supported as read-write + + ufs2 used in FreeBSD 5.x + supported as read-write + + 5xbsd synonym for ufs2 + + sun used in SunOS (Solaris) + supported as read-write + + sunx86 used in SunOS for Intel (Solarisx86) + supported as read-write + + hp used in HP-UX + supported as read-only + + nextstep + used in NextStep + supported as read-only + + nextstep-cd + used for NextStep CDROMs (block_size == 2048) + supported as read-only + + openstep + used in OpenStep + supported as read-only + + +POSSIBLE PROBLEMS +================= + +See next section, if you have any. + + +BUG REPORTS +=========== + +Any ufs bug report you can send to daniel.pirkl@email.cz or +to dushistov@mail.ru (do not send partition tables bug reports). diff --git a/Documentation/filesystems/vfat.txt b/Documentation/filesystems/vfat.txt new file mode 100644 index 00000000..ead764b2 --- /dev/null +++ b/Documentation/filesystems/vfat.txt @@ -0,0 +1,295 @@ +USING VFAT +---------------------------------------------------------------------- +To use the vfat filesystem, use the filesystem type 'vfat'. i.e. + mount -t vfat /dev/fd0 /mnt + +No special partition formatter is required. mkdosfs will work fine +if you want to format from within Linux. + +VFAT MOUNT OPTIONS +---------------------------------------------------------------------- +uid=### -- Set the owner of all files on this filesystem. + The default is the uid of current process. + +gid=### -- Set the group of all files on this filesystem. + The default is the gid of current process. + +umask=### -- The permission mask (for files and directories, see umask(1)). + The default is the umask of current process. + +dmask=### -- The permission mask for the directory. + The default is the umask of current process. + +fmask=### -- The permission mask for files. + The default is the umask of current process. + +allow_utime=### -- This option controls the permission check of mtime/atime. + + 20 - If current process is in group of file's group ID, + you can change timestamp. + 2 - Other users can change timestamp. + + The default is set from `dmask' option. (If the directory is + writable, utime(2) is also allowed. I.e. ~dmask & 022) + + Normally utime(2) checks current process is owner of + the file, or it has CAP_FOWNER capability. But FAT + filesystem doesn't have uid/gid on disk, so normal + check is too unflexible. With this option you can + relax it. + +codepage=### -- Sets the codepage number for converting to shortname + characters on FAT filesystem. + By default, FAT_DEFAULT_CODEPAGE setting is used. + +iocharset=<name> -- Character set to use for converting between the + encoding is used for user visible filename and 16 bit + Unicode characters. Long filenames are stored on disk + in Unicode format, but Unix for the most part doesn't + know how to deal with Unicode. + By default, FAT_DEFAULT_IOCHARSET setting is used. + + There is also an option of doing UTF-8 translations + with the utf8 option. + + NOTE: "iocharset=utf8" is not recommended. If unsure, + you should consider the following option instead. + +utf8=<bool> -- UTF-8 is the filesystem safe version of Unicode that + is used by the console. It can be enabled for the + filesystem with this option. If 'uni_xlate' gets set, + UTF-8 gets disabled. + +uni_xlate=<bool> -- Translate unhandled Unicode characters to special + escaped sequences. This would let you backup and + restore filenames that are created with any Unicode + characters. Until Linux supports Unicode for real, + this gives you an alternative. Without this option, + a '?' is used when no translation is possible. The + escape character is ':' because it is otherwise + illegal on the vfat filesystem. The escape sequence + that gets used is ':' and the four digits of hexadecimal + unicode. + +nonumtail=<bool> -- When creating 8.3 aliases, normally the alias will + end in '~1' or tilde followed by some number. If this + option is set, then if the filename is + "longfilename.txt" and "longfile.txt" does not + currently exist in the directory, 'longfile.txt' will + be the short alias instead of 'longfi~1.txt'. + +usefree -- Use the "free clusters" value stored on FSINFO. It'll + be used to determine number of free clusters without + scanning disk. But it's not used by default, because + recent Windows don't update it correctly in some + case. If you are sure the "free clusters" on FSINFO is + correct, by this option you can avoid scanning disk. + +quiet -- Stops printing certain warning messages. + +check=s|r|n -- Case sensitivity checking setting. + s: strict, case sensitive + r: relaxed, case insensitive + n: normal, default setting, currently case insensitive + +nocase -- This was deprecated for vfat. Use shortname=win95 instead. + +shortname=lower|win95|winnt|mixed + -- Shortname display/create setting. + lower: convert to lowercase for display, + emulate the Windows 95 rule for create. + win95: emulate the Windows 95 rule for display/create. + winnt: emulate the Windows NT rule for display/create. + mixed: emulate the Windows NT rule for display, + emulate the Windows 95 rule for create. + Default setting is `mixed'. + +tz=UTC -- Interpret timestamps as UTC rather than local time. + This option disables the conversion of timestamps + between local time (as used by Windows on FAT) and UTC + (which Linux uses internally). This is particularly + useful when mounting devices (like digital cameras) + that are set to UTC in order to avoid the pitfalls of + local time. + +showexec -- If set, the execute permission bits of the file will be + allowed only if the extension part of the name is .EXE, + .COM, or .BAT. Not set by default. + +debug -- Can be set, but unused by the current implementation. + +sys_immutable -- If set, ATTR_SYS attribute on FAT is handled as + IMMUTABLE flag on Linux. Not set by default. + +flush -- If set, the filesystem will try to flush to disk more + early than normal. Not set by default. + +rodir -- FAT has the ATTR_RO (read-only) attribute. On Windows, + the ATTR_RO of the directory will just be ignored, + and is used only by applications as a flag (e.g. it's set + for the customized folder). + + If you want to use ATTR_RO as read-only flag even for + the directory, set this option. + +errors=panic|continue|remount-ro + -- specify FAT behavior on critical errors: panic, continue + without doing anything or remount the partition in + read-only mode (default behavior). + +<bool>: 0,1,yes,no,true,false + +TODO +---------------------------------------------------------------------- +* Need to get rid of the raw scanning stuff. Instead, always use + a get next directory entry approach. The only thing left that uses + raw scanning is the directory renaming code. + + +POSSIBLE PROBLEMS +---------------------------------------------------------------------- +* vfat_valid_longname does not properly checked reserved names. +* When a volume name is the same as a directory name in the root + directory of the filesystem, the directory name sometimes shows + up as an empty file. +* autoconv option does not work correctly. + +BUG REPORTS +---------------------------------------------------------------------- +If you have trouble with the VFAT filesystem, mail bug reports to +chaffee@bmrc.cs.berkeley.edu. Please specify the filename +and the operation that gave you trouble. + +TEST SUITE +---------------------------------------------------------------------- +If you plan to make any modifications to the vfat filesystem, please +get the test suite that comes with the vfat distribution at + + http://web.archive.org/web/*/http://bmrc.berkeley.edu/ + people/chaffee/vfat.html + +This tests quite a few parts of the vfat filesystem and additional +tests for new features or untested features would be appreciated. + +NOTES ON THE STRUCTURE OF THE VFAT FILESYSTEM +---------------------------------------------------------------------- +(This documentation was provided by Galen C. Hunt <gchunt@cs.rochester.edu> + and lightly annotated by Gordon Chaffee). + +This document presents a very rough, technical overview of my +knowledge of the extended FAT file system used in Windows NT 3.5 and +Windows 95. I don't guarantee that any of the following is correct, +but it appears to be so. + +The extended FAT file system is almost identical to the FAT +file system used in DOS versions up to and including 6.223410239847 +:-). The significant change has been the addition of long file names. +These names support up to 255 characters including spaces and lower +case characters as opposed to the traditional 8.3 short names. + +Here is the description of the traditional FAT entry in the current +Windows 95 filesystem: + + struct directory { // Short 8.3 names + unsigned char name[8]; // file name + unsigned char ext[3]; // file extension + unsigned char attr; // attribute byte + unsigned char lcase; // Case for base and extension + unsigned char ctime_ms; // Creation time, milliseconds + unsigned char ctime[2]; // Creation time + unsigned char cdate[2]; // Creation date + unsigned char adate[2]; // Last access date + unsigned char reserved[2]; // reserved values (ignored) + unsigned char time[2]; // time stamp + unsigned char date[2]; // date stamp + unsigned char start[2]; // starting cluster number + unsigned char size[4]; // size of the file + }; + +The lcase field specifies if the base and/or the extension of an 8.3 +name should be capitalized. This field does not seem to be used by +Windows 95 but it is used by Windows NT. The case of filenames is not +completely compatible from Windows NT to Windows 95. It is not completely +compatible in the reverse direction, however. Filenames that fit in +the 8.3 namespace and are written on Windows NT to be lowercase will +show up as uppercase on Windows 95. + +Note that the "start" and "size" values are actually little +endian integer values. The descriptions of the fields in this +structure are public knowledge and can be found elsewhere. + +With the extended FAT system, Microsoft has inserted extra +directory entries for any files with extended names. (Any name which +legally fits within the old 8.3 encoding scheme does not have extra +entries.) I call these extra entries slots. Basically, a slot is a +specially formatted directory entry which holds up to 13 characters of +a file's extended name. Think of slots as additional labeling for the +directory entry of the file to which they correspond. Microsoft +prefers to refer to the 8.3 entry for a file as its alias and the +extended slot directory entries as the file name. + +The C structure for a slot directory entry follows: + + struct slot { // Up to 13 characters of a long name + unsigned char id; // sequence number for slot + unsigned char name0_4[10]; // first 5 characters in name + unsigned char attr; // attribute byte + unsigned char reserved; // always 0 + unsigned char alias_checksum; // checksum for 8.3 alias + unsigned char name5_10[12]; // 6 more characters in name + unsigned char start[2]; // starting cluster number + unsigned char name11_12[4]; // last 2 characters in name + }; + +If the layout of the slots looks a little odd, it's only +because of Microsoft's efforts to maintain compatibility with old +software. The slots must be disguised to prevent old software from +panicking. To this end, a number of measures are taken: + + 1) The attribute byte for a slot directory entry is always set + to 0x0f. This corresponds to an old directory entry with + attributes of "hidden", "system", "read-only", and "volume + label". Most old software will ignore any directory + entries with the "volume label" bit set. Real volume label + entries don't have the other three bits set. + + 2) The starting cluster is always set to 0, an impossible + value for a DOS file. + +Because the extended FAT system is backward compatible, it is +possible for old software to modify directory entries. Measures must +be taken to ensure the validity of slots. An extended FAT system can +verify that a slot does in fact belong to an 8.3 directory entry by +the following: + + 1) Positioning. Slots for a file always immediately proceed + their corresponding 8.3 directory entry. In addition, each + slot has an id which marks its order in the extended file + name. Here is a very abbreviated view of an 8.3 directory + entry and its corresponding long name slots for the file + "My Big File.Extension which is long": + + <proceeding files...> + <slot #3, id = 0x43, characters = "h is long"> + <slot #2, id = 0x02, characters = "xtension whic"> + <slot #1, id = 0x01, characters = "My Big File.E"> + <directory entry, name = "MYBIGFIL.EXT"> + + Note that the slots are stored from last to first. Slots + are numbered from 1 to N. The Nth slot is or'ed with 0x40 + to mark it as the last one. + + 2) Checksum. Each slot has an "alias_checksum" value. The + checksum is calculated from the 8.3 name using the + following algorithm: + + for (sum = i = 0; i < 11; i++) { + sum = (((sum&1)<<7)|((sum&0xfe)>>1)) + name[i] + } + + 3) If there is free space in the final slot, a Unicode NULL (0x0000) + is stored after the final character. After that, all unused + characters in the final slot are set to Unicode 0xFFFF. + +Finally, note that the extended name is stored in Unicode. Each Unicode +character takes two bytes. diff --git a/Documentation/filesystems/vfs.txt b/Documentation/filesystems/vfs.txt new file mode 100644 index 00000000..88b9f551 --- /dev/null +++ b/Documentation/filesystems/vfs.txt @@ -0,0 +1,1097 @@ + + Overview of the Linux Virtual File System + + Original author: Richard Gooch <rgooch@atnf.csiro.au> + + Last updated on June 24, 2007. + + Copyright (C) 1999 Richard Gooch + Copyright (C) 2005 Pekka Enberg + + This file is released under the GPLv2. + + +Introduction +============ + +The Virtual File System (also known as the Virtual Filesystem Switch) +is the software layer in the kernel that provides the filesystem +interface to userspace programs. It also provides an abstraction +within the kernel which allows different filesystem implementations to +coexist. + +VFS system calls open(2), stat(2), read(2), write(2), chmod(2) and so +on are called from a process context. Filesystem locking is described +in the document Documentation/filesystems/Locking. + + +Directory Entry Cache (dcache) +------------------------------ + +The VFS implements the open(2), stat(2), chmod(2), and similar system +calls. The pathname argument that is passed to them is used by the VFS +to search through the directory entry cache (also known as the dentry +cache or dcache). This provides a very fast look-up mechanism to +translate a pathname (filename) into a specific dentry. Dentries live +in RAM and are never saved to disc: they exist only for performance. + +The dentry cache is meant to be a view into your entire filespace. As +most computers cannot fit all dentries in the RAM at the same time, +some bits of the cache are missing. In order to resolve your pathname +into a dentry, the VFS may have to resort to creating dentries along +the way, and then loading the inode. This is done by looking up the +inode. + + +The Inode Object +---------------- + +An individual dentry usually has a pointer to an inode. Inodes are +filesystem objects such as regular files, directories, FIFOs and other +beasts. They live either on the disc (for block device filesystems) +or in the memory (for pseudo filesystems). Inodes that live on the +disc are copied into the memory when required and changes to the inode +are written back to disc. A single inode can be pointed to by multiple +dentries (hard links, for example, do this). + +To look up an inode requires that the VFS calls the lookup() method of +the parent directory inode. This method is installed by the specific +filesystem implementation that the inode lives in. Once the VFS has +the required dentry (and hence the inode), we can do all those boring +things like open(2) the file, or stat(2) it to peek at the inode +data. The stat(2) operation is fairly simple: once the VFS has the +dentry, it peeks at the inode data and passes some of it back to +userspace. + + +The File Object +--------------- + +Opening a file requires another operation: allocation of a file +structure (this is the kernel-side implementation of file +descriptors). The freshly allocated file structure is initialized with +a pointer to the dentry and a set of file operation member functions. +These are taken from the inode data. The open() file method is then +called so the specific filesystem implementation can do its work. You +can see that this is another switch performed by the VFS. The file +structure is placed into the file descriptor table for the process. + +Reading, writing and closing files (and other assorted VFS operations) +is done by using the userspace file descriptor to grab the appropriate +file structure, and then calling the required file structure method to +do whatever is required. For as long as the file is open, it keeps the +dentry in use, which in turn means that the VFS inode is still in use. + + +Registering and Mounting a Filesystem +===================================== + +To register and unregister a filesystem, use the following API +functions: + + #include <linux/fs.h> + + extern int register_filesystem(struct file_system_type *); + extern int unregister_filesystem(struct file_system_type *); + +The passed struct file_system_type describes your filesystem. When a +request is made to mount a filesystem onto a directory in your namespace, +the VFS will call the appropriate mount() method for the specific +filesystem. New vfsmount referring to the tree returned by ->mount() +will be attached to the mountpoint, so that when pathname resolution +reaches the mountpoint it will jump into the root of that vfsmount. + +You can see all filesystems that are registered to the kernel in the +file /proc/filesystems. + + +struct file_system_type +----------------------- + +This describes the filesystem. As of kernel 2.6.39, the following +members are defined: + +struct file_system_type { + const char *name; + int fs_flags; + struct dentry (*mount) (struct file_system_type *, int, + const char *, void *); + void (*kill_sb) (struct super_block *); + struct module *owner; + struct file_system_type * next; + struct list_head fs_supers; + struct lock_class_key s_lock_key; + struct lock_class_key s_umount_key; +}; + + name: the name of the filesystem type, such as "ext2", "iso9660", + "msdos" and so on + + fs_flags: various flags (i.e. FS_REQUIRES_DEV, FS_NO_DCACHE, etc.) + + mount: the method to call when a new instance of this + filesystem should be mounted + + kill_sb: the method to call when an instance of this filesystem + should be shut down + + owner: for internal VFS use: you should initialize this to THIS_MODULE in + most cases. + + next: for internal VFS use: you should initialize this to NULL + + s_lock_key, s_umount_key: lockdep-specific + +The mount() method has the following arguments: + + struct file_system_type *fs_type: describes the filesystem, partly initialized + by the specific filesystem code + + int flags: mount flags + + const char *dev_name: the device name we are mounting. + + void *data: arbitrary mount options, usually comes as an ASCII + string (see "Mount Options" section) + +The mount() method must return the root dentry of the tree requested by +caller. An active reference to its superblock must be grabbed and the +superblock must be locked. On failure it should return ERR_PTR(error). + +The arguments match those of mount(2) and their interpretation +depends on filesystem type. E.g. for block filesystems, dev_name is +interpreted as block device name, that device is opened and if it +contains a suitable filesystem image the method creates and initializes +struct super_block accordingly, returning its root dentry to caller. + +->mount() may choose to return a subtree of existing filesystem - it +doesn't have to create a new one. The main result from the caller's +point of view is a reference to dentry at the root of (sub)tree to +be attached; creation of new superblock is a common side effect. + +The most interesting member of the superblock structure that the +mount() method fills in is the "s_op" field. This is a pointer to +a "struct super_operations" which describes the next level of the +filesystem implementation. + +Usually, a filesystem uses one of the generic mount() implementations +and provides a fill_super() callback instead. The generic variants are: + + mount_bdev: mount a filesystem residing on a block device + + mount_nodev: mount a filesystem that is not backed by a device + + mount_single: mount a filesystem which shares the instance between + all mounts + +A fill_super() callback implementation has the following arguments: + + struct super_block *sb: the superblock structure. The callback + must initialize this properly. + + void *data: arbitrary mount options, usually comes as an ASCII + string (see "Mount Options" section) + + int silent: whether or not to be silent on error + + +The Superblock Object +===================== + +A superblock object represents a mounted filesystem. + + +struct super_operations +----------------------- + +This describes how the VFS can manipulate the superblock of your +filesystem. As of kernel 2.6.22, the following members are defined: + +struct super_operations { + struct inode *(*alloc_inode)(struct super_block *sb); + void (*destroy_inode)(struct inode *); + + void (*dirty_inode) (struct inode *, int flags); + int (*write_inode) (struct inode *, int); + void (*drop_inode) (struct inode *); + void (*delete_inode) (struct inode *); + void (*put_super) (struct super_block *); + void (*write_super) (struct super_block *); + int (*sync_fs)(struct super_block *sb, int wait); + int (*freeze_fs) (struct super_block *); + int (*unfreeze_fs) (struct super_block *); + int (*statfs) (struct dentry *, struct kstatfs *); + int (*remount_fs) (struct super_block *, int *, char *); + void (*clear_inode) (struct inode *); + void (*umount_begin) (struct super_block *); + + int (*show_options)(struct seq_file *, struct vfsmount *); + + ssize_t (*quota_read)(struct super_block *, int, char *, size_t, loff_t); + ssize_t (*quota_write)(struct super_block *, int, const char *, size_t, loff_t); +}; + +All methods are called without any locks being held, unless otherwise +noted. This means that most methods can block safely. All methods are +only called from a process context (i.e. not from an interrupt handler +or bottom half). + + alloc_inode: this method is called by inode_alloc() to allocate memory + for struct inode and initialize it. If this function is not + defined, a simple 'struct inode' is allocated. Normally + alloc_inode will be used to allocate a larger structure which + contains a 'struct inode' embedded within it. + + destroy_inode: this method is called by destroy_inode() to release + resources allocated for struct inode. It is only required if + ->alloc_inode was defined and simply undoes anything done by + ->alloc_inode. + + dirty_inode: this method is called by the VFS to mark an inode dirty. + + write_inode: this method is called when the VFS needs to write an + inode to disc. The second parameter indicates whether the write + should be synchronous or not, not all filesystems check this flag. + + drop_inode: called when the last access to the inode is dropped, + with the inode->i_lock spinlock held. + + This method should be either NULL (normal UNIX filesystem + semantics) or "generic_delete_inode" (for filesystems that do not + want to cache inodes - causing "delete_inode" to always be + called regardless of the value of i_nlink) + + The "generic_delete_inode()" behavior is equivalent to the + old practice of using "force_delete" in the put_inode() case, + but does not have the races that the "force_delete()" approach + had. + + delete_inode: called when the VFS wants to delete an inode + + put_super: called when the VFS wishes to free the superblock + (i.e. unmount). This is called with the superblock lock held + + write_super: called when the VFS superblock needs to be written to + disc. This method is optional + + sync_fs: called when VFS is writing out all dirty data associated with + a superblock. The second parameter indicates whether the method + should wait until the write out has been completed. Optional. + + freeze_fs: called when VFS is locking a filesystem and + forcing it into a consistent state. This method is currently + used by the Logical Volume Manager (LVM). + + unfreeze_fs: called when VFS is unlocking a filesystem and making it writable + again. + + statfs: called when the VFS needs to get filesystem statistics. + + remount_fs: called when the filesystem is remounted. This is called + with the kernel lock held + + clear_inode: called then the VFS clears the inode. Optional + + umount_begin: called when the VFS is unmounting a filesystem. + + show_options: called by the VFS to show mount options for + /proc/<pid>/mounts. (see "Mount Options" section) + + quota_read: called by the VFS to read from filesystem quota file. + + quota_write: called by the VFS to write to filesystem quota file. + +Whoever sets up the inode is responsible for filling in the "i_op" field. This +is a pointer to a "struct inode_operations" which describes the methods that +can be performed on individual inodes. + + +The Inode Object +================ + +An inode object represents an object within the filesystem. + + +struct inode_operations +----------------------- + +This describes how the VFS can manipulate an inode in your +filesystem. As of kernel 2.6.22, the following members are defined: + +struct inode_operations { + int (*create) (struct inode *,struct dentry *,int, struct nameidata *); + struct dentry * (*lookup) (struct inode *,struct dentry *, struct nameidata *); + int (*link) (struct dentry *,struct inode *,struct dentry *); + int (*unlink) (struct inode *,struct dentry *); + int (*symlink) (struct inode *,struct dentry *,const char *); + int (*mkdir) (struct inode *,struct dentry *,int); + int (*rmdir) (struct inode *,struct dentry *); + int (*mknod) (struct inode *,struct dentry *,int,dev_t); + int (*rename) (struct inode *, struct dentry *, + struct inode *, struct dentry *); + int (*readlink) (struct dentry *, char __user *,int); + void * (*follow_link) (struct dentry *, struct nameidata *); + void (*put_link) (struct dentry *, struct nameidata *, void *); + void (*truncate) (struct inode *); + int (*permission) (struct inode *, int, unsigned int); + int (*check_acl)(struct inode *, int, unsigned int); + int (*setattr) (struct dentry *, struct iattr *); + int (*getattr) (struct vfsmount *mnt, struct dentry *, struct kstat *); + int (*setxattr) (struct dentry *, const char *,const void *,size_t,int); + ssize_t (*getxattr) (struct dentry *, const char *, void *, size_t); + ssize_t (*listxattr) (struct dentry *, char *, size_t); + int (*removexattr) (struct dentry *, const char *); + void (*truncate_range)(struct inode *, loff_t, loff_t); +}; + +Again, all methods are called without any locks being held, unless +otherwise noted. + + create: called by the open(2) and creat(2) system calls. Only + required if you want to support regular files. The dentry you + get should not have an inode (i.e. it should be a negative + dentry). Here you will probably call d_instantiate() with the + dentry and the newly created inode + + lookup: called when the VFS needs to look up an inode in a parent + directory. The name to look for is found in the dentry. This + method must call d_add() to insert the found inode into the + dentry. The "i_count" field in the inode structure should be + incremented. If the named inode does not exist a NULL inode + should be inserted into the dentry (this is called a negative + dentry). Returning an error code from this routine must only + be done on a real error, otherwise creating inodes with system + calls like create(2), mknod(2), mkdir(2) and so on will fail. + If you wish to overload the dentry methods then you should + initialise the "d_dop" field in the dentry; this is a pointer + to a struct "dentry_operations". + This method is called with the directory inode semaphore held + + link: called by the link(2) system call. Only required if you want + to support hard links. You will probably need to call + d_instantiate() just as you would in the create() method + + unlink: called by the unlink(2) system call. Only required if you + want to support deleting inodes + + symlink: called by the symlink(2) system call. Only required if you + want to support symlinks. You will probably need to call + d_instantiate() just as you would in the create() method + + mkdir: called by the mkdir(2) system call. Only required if you want + to support creating subdirectories. You will probably need to + call d_instantiate() just as you would in the create() method + + rmdir: called by the rmdir(2) system call. Only required if you want + to support deleting subdirectories + + mknod: called by the mknod(2) system call to create a device (char, + block) inode or a named pipe (FIFO) or socket. Only required + if you want to support creating these types of inodes. You + will probably need to call d_instantiate() just as you would + in the create() method + + rename: called by the rename(2) system call to rename the object to + have the parent and name given by the second inode and dentry. + + readlink: called by the readlink(2) system call. Only required if + you want to support reading symbolic links + + follow_link: called by the VFS to follow a symbolic link to the + inode it points to. Only required if you want to support + symbolic links. This method returns a void pointer cookie + that is passed to put_link(). + + put_link: called by the VFS to release resources allocated by + follow_link(). The cookie returned by follow_link() is passed + to this method as the last parameter. It is used by + filesystems such as NFS where page cache is not stable + (i.e. page that was installed when the symbolic link walk + started might not be in the page cache at the end of the + walk). + + truncate: Deprecated. This will not be called if ->setsize is defined. + Called by the VFS to change the size of a file. The + i_size field of the inode is set to the desired size by the + VFS before this method is called. This method is called by + the truncate(2) system call and related functionality. + + Note: ->truncate and vmtruncate are deprecated. Do not add new + instances/calls of these. Filesystems should be converted to do their + truncate sequence via ->setattr(). + + permission: called by the VFS to check for access rights on a POSIX-like + filesystem. + + May be called in rcu-walk mode (flags & IPERM_FLAG_RCU). If in rcu-walk + mode, the filesystem must check the permission without blocking or + storing to the inode. + + If a situation is encountered that rcu-walk cannot handle, return + -ECHILD and it will be called again in ref-walk mode. + + setattr: called by the VFS to set attributes for a file. This method + is called by chmod(2) and related system calls. + + getattr: called by the VFS to get attributes of a file. This method + is called by stat(2) and related system calls. + + setxattr: called by the VFS to set an extended attribute for a file. + Extended attribute is a name:value pair associated with an + inode. This method is called by setxattr(2) system call. + + getxattr: called by the VFS to retrieve the value of an extended + attribute name. This method is called by getxattr(2) function + call. + + listxattr: called by the VFS to list all extended attributes for a + given file. This method is called by listxattr(2) system call. + + removexattr: called by the VFS to remove an extended attribute from + a file. This method is called by removexattr(2) system call. + + truncate_range: a method provided by the underlying filesystem to truncate a + range of blocks , i.e. punch a hole somewhere in a file. + + +The Address Space Object +======================== + +The address space object is used to group and manage pages in the page +cache. It can be used to keep track of the pages in a file (or +anything else) and also track the mapping of sections of the file into +process address spaces. + +There are a number of distinct yet related services that an +address-space can provide. These include communicating memory +pressure, page lookup by address, and keeping track of pages tagged as +Dirty or Writeback. + +The first can be used independently to the others. The VM can try to +either write dirty pages in order to clean them, or release clean +pages in order to reuse them. To do this it can call the ->writepage +method on dirty pages, and ->releasepage on clean pages with +PagePrivate set. Clean pages without PagePrivate and with no external +references will be released without notice being given to the +address_space. + +To achieve this functionality, pages need to be placed on an LRU with +lru_cache_add and mark_page_active needs to be called whenever the +page is used. + +Pages are normally kept in a radix tree index by ->index. This tree +maintains information about the PG_Dirty and PG_Writeback status of +each page, so that pages with either of these flags can be found +quickly. + +The Dirty tag is primarily used by mpage_writepages - the default +->writepages method. It uses the tag to find dirty pages to call +->writepage on. If mpage_writepages is not used (i.e. the address +provides its own ->writepages) , the PAGECACHE_TAG_DIRTY tag is +almost unused. write_inode_now and sync_inode do use it (through +__sync_single_inode) to check if ->writepages has been successful in +writing out the whole address_space. + +The Writeback tag is used by filemap*wait* and sync_page* functions, +via filemap_fdatawait_range, to wait for all writeback to +complete. While waiting ->sync_page (if defined) will be called on +each page that is found to require writeback. + +An address_space handler may attach extra information to a page, +typically using the 'private' field in the 'struct page'. If such +information is attached, the PG_Private flag should be set. This will +cause various VM routines to make extra calls into the address_space +handler to deal with that data. + +An address space acts as an intermediate between storage and +application. Data is read into the address space a whole page at a +time, and provided to the application either by copying of the page, +or by memory-mapping the page. +Data is written into the address space by the application, and then +written-back to storage typically in whole pages, however the +address_space has finer control of write sizes. + +The read process essentially only requires 'readpage'. The write +process is more complicated and uses write_begin/write_end or +set_page_dirty to write data into the address_space, and writepage, +sync_page, and writepages to writeback data to storage. + +Adding and removing pages to/from an address_space is protected by the +inode's i_mutex. + +When data is written to a page, the PG_Dirty flag should be set. It +typically remains set until writepage asks for it to be written. This +should clear PG_Dirty and set PG_Writeback. It can be actually +written at any point after PG_Dirty is clear. Once it is known to be +safe, PG_Writeback is cleared. + +Writeback makes use of a writeback_control structure... + +struct address_space_operations +------------------------------- + +This describes how the VFS can manipulate mapping of a file to page cache in +your filesystem. As of kernel 2.6.22, the following members are defined: + +struct address_space_operations { + int (*writepage)(struct page *page, struct writeback_control *wbc); + int (*readpage)(struct file *, struct page *); + int (*sync_page)(struct page *); + int (*writepages)(struct address_space *, struct writeback_control *); + int (*set_page_dirty)(struct page *page); + int (*readpages)(struct file *filp, struct address_space *mapping, + struct list_head *pages, unsigned nr_pages); + int (*write_begin)(struct file *, struct address_space *mapping, + loff_t pos, unsigned len, unsigned flags, + struct page **pagep, void **fsdata); + int (*write_end)(struct file *, struct address_space *mapping, + loff_t pos, unsigned len, unsigned copied, + struct page *page, void *fsdata); + sector_t (*bmap)(struct address_space *, sector_t); + int (*invalidatepage) (struct page *, unsigned long); + int (*releasepage) (struct page *, int); + void (*freepage)(struct page *); + ssize_t (*direct_IO)(int, struct kiocb *, const struct iovec *iov, + loff_t offset, unsigned long nr_segs); + struct page* (*get_xip_page)(struct address_space *, sector_t, + int); + /* migrate the contents of a page to the specified target */ + int (*migratepage) (struct page *, struct page *); + int (*launder_page) (struct page *); + int (*error_remove_page) (struct mapping *mapping, struct page *page); +}; + + writepage: called by the VM to write a dirty page to backing store. + This may happen for data integrity reasons (i.e. 'sync'), or + to free up memory (flush). The difference can be seen in + wbc->sync_mode. + The PG_Dirty flag has been cleared and PageLocked is true. + writepage should start writeout, should set PG_Writeback, + and should make sure the page is unlocked, either synchronously + or asynchronously when the write operation completes. + + If wbc->sync_mode is WB_SYNC_NONE, ->writepage doesn't have to + try too hard if there are problems, and may choose to write out + other pages from the mapping if that is easier (e.g. due to + internal dependencies). If it chooses not to start writeout, it + should return AOP_WRITEPAGE_ACTIVATE so that the VM will not keep + calling ->writepage on that page. + + See the file "Locking" for more details. + + readpage: called by the VM to read a page from backing store. + The page will be Locked when readpage is called, and should be + unlocked and marked uptodate once the read completes. + If ->readpage discovers that it needs to unlock the page for + some reason, it can do so, and then return AOP_TRUNCATED_PAGE. + In this case, the page will be relocated, relocked and if + that all succeeds, ->readpage will be called again. + + sync_page: called by the VM to notify the backing store to perform all + queued I/O operations for a page. I/O operations for other pages + associated with this address_space object may also be performed. + + This function is optional and is called only for pages with + PG_Writeback set while waiting for the writeback to complete. + + writepages: called by the VM to write out pages associated with the + address_space object. If wbc->sync_mode is WBC_SYNC_ALL, then + the writeback_control will specify a range of pages that must be + written out. If it is WBC_SYNC_NONE, then a nr_to_write is given + and that many pages should be written if possible. + If no ->writepages is given, then mpage_writepages is used + instead. This will choose pages from the address space that are + tagged as DIRTY and will pass them to ->writepage. + + set_page_dirty: called by the VM to set a page dirty. + This is particularly needed if an address space attaches + private data to a page, and that data needs to be updated when + a page is dirtied. This is called, for example, when a memory + mapped page gets modified. + If defined, it should set the PageDirty flag, and the + PAGECACHE_TAG_DIRTY tag in the radix tree. + + readpages: called by the VM to read pages associated with the address_space + object. This is essentially just a vector version of + readpage. Instead of just one page, several pages are + requested. + readpages is only used for read-ahead, so read errors are + ignored. If anything goes wrong, feel free to give up. + + write_begin: + Called by the generic buffered write code to ask the filesystem to + prepare to write len bytes at the given offset in the file. The + address_space should check that the write will be able to complete, + by allocating space if necessary and doing any other internal + housekeeping. If the write will update parts of any basic-blocks on + storage, then those blocks should be pre-read (if they haven't been + read already) so that the updated blocks can be written out properly. + + The filesystem must return the locked pagecache page for the specified + offset, in *pagep, for the caller to write into. + + It must be able to cope with short writes (where the length passed to + write_begin is greater than the number of bytes copied into the page). + + flags is a field for AOP_FLAG_xxx flags, described in + include/linux/fs.h. + + A void * may be returned in fsdata, which then gets passed into + write_end. + + Returns 0 on success; < 0 on failure (which is the error code), in + which case write_end is not called. + + write_end: After a successful write_begin, and data copy, write_end must + be called. len is the original len passed to write_begin, and copied + is the amount that was able to be copied (copied == len is always true + if write_begin was called with the AOP_FLAG_UNINTERRUPTIBLE flag). + + The filesystem must take care of unlocking the page and releasing it + refcount, and updating i_size. + + Returns < 0 on failure, otherwise the number of bytes (<= 'copied') + that were able to be copied into pagecache. + + bmap: called by the VFS to map a logical block offset within object to + physical block number. This method is used by the FIBMAP + ioctl and for working with swap-files. To be able to swap to + a file, the file must have a stable mapping to a block + device. The swap system does not go through the filesystem + but instead uses bmap to find out where the blocks in the file + are and uses those addresses directly. + + + invalidatepage: If a page has PagePrivate set, then invalidatepage + will be called when part or all of the page is to be removed + from the address space. This generally corresponds to either a + truncation or a complete invalidation of the address space + (in the latter case 'offset' will always be 0). + Any private data associated with the page should be updated + to reflect this truncation. If offset is 0, then + the private data should be released, because the page + must be able to be completely discarded. This may be done by + calling the ->releasepage function, but in this case the + release MUST succeed. + + releasepage: releasepage is called on PagePrivate pages to indicate + that the page should be freed if possible. ->releasepage + should remove any private data from the page and clear the + PagePrivate flag. If releasepage() fails for some reason, it must + indicate failure with a 0 return value. + releasepage() is used in two distinct though related cases. The + first is when the VM finds a clean page with no active users and + wants to make it a free page. If ->releasepage succeeds, the + page will be removed from the address_space and become free. + + The second case is when a request has been made to invalidate + some or all pages in an address_space. This can happen + through the fadvice(POSIX_FADV_DONTNEED) system call or by the + filesystem explicitly requesting it as nfs and 9fs do (when + they believe the cache may be out of date with storage) by + calling invalidate_inode_pages2(). + If the filesystem makes such a call, and needs to be certain + that all pages are invalidated, then its releasepage will + need to ensure this. Possibly it can clear the PageUptodate + bit if it cannot free private data yet. + + freepage: freepage is called once the page is no longer visible in + the page cache in order to allow the cleanup of any private + data. Since it may be called by the memory reclaimer, it + should not assume that the original address_space mapping still + exists, and it should not block. + + direct_IO: called by the generic read/write routines to perform + direct_IO - that is IO requests which bypass the page cache + and transfer data directly between the storage and the + application's address space. + + get_xip_page: called by the VM to translate a block number to a page. + The page is valid until the corresponding filesystem is unmounted. + Filesystems that want to use execute-in-place (XIP) need to implement + it. An example implementation can be found in fs/ext2/xip.c. + + migrate_page: This is used to compact the physical memory usage. + If the VM wants to relocate a page (maybe off a memory card + that is signalling imminent failure) it will pass a new page + and an old page to this function. migrate_page should + transfer any private data across and update any references + that it has to the page. + + launder_page: Called before freeing a page - it writes back the dirty page. To + prevent redirtying the page, it is kept locked during the whole + operation. + + error_remove_page: normally set to generic_error_remove_page if truncation + is ok for this address space. Used for memory failure handling. + Setting this implies you deal with pages going away under you, + unless you have them locked or reference counts increased. + + +The File Object +=============== + +A file object represents a file opened by a process. + + +struct file_operations +---------------------- + +This describes how the VFS can manipulate an open file. As of kernel +2.6.22, the following members are defined: + +struct file_operations { + struct module *owner; + loff_t (*llseek) (struct file *, loff_t, int); + ssize_t (*read) (struct file *, char __user *, size_t, loff_t *); + ssize_t (*write) (struct file *, const char __user *, size_t, loff_t *); + ssize_t (*aio_read) (struct kiocb *, const struct iovec *, unsigned long, loff_t); + ssize_t (*aio_write) (struct kiocb *, const struct iovec *, unsigned long, loff_t); + int (*readdir) (struct file *, void *, filldir_t); + unsigned int (*poll) (struct file *, struct poll_table_struct *); + long (*unlocked_ioctl) (struct file *, unsigned int, unsigned long); + long (*compat_ioctl) (struct file *, unsigned int, unsigned long); + int (*mmap) (struct file *, struct vm_area_struct *); + int (*open) (struct inode *, struct file *); + int (*flush) (struct file *); + int (*release) (struct inode *, struct file *); + int (*fsync) (struct file *, int datasync); + int (*aio_fsync) (struct kiocb *, int datasync); + int (*fasync) (int, struct file *, int); + int (*lock) (struct file *, int, struct file_lock *); + ssize_t (*readv) (struct file *, const struct iovec *, unsigned long, loff_t *); + ssize_t (*writev) (struct file *, const struct iovec *, unsigned long, loff_t *); + ssize_t (*sendfile) (struct file *, loff_t *, size_t, read_actor_t, void *); + ssize_t (*sendpage) (struct file *, struct page *, int, size_t, loff_t *, int); + unsigned long (*get_unmapped_area)(struct file *, unsigned long, unsigned long, unsigned long, unsigned long); + int (*check_flags)(int); + int (*flock) (struct file *, int, struct file_lock *); + ssize_t (*splice_write)(struct pipe_inode_info *, struct file *, size_t, unsigned int); + ssize_t (*splice_read)(struct file *, struct pipe_inode_info *, size_t, unsigned int); +}; + +Again, all methods are called without any locks being held, unless +otherwise noted. + + llseek: called when the VFS needs to move the file position index + + read: called by read(2) and related system calls + + aio_read: called by io_submit(2) and other asynchronous I/O operations + + write: called by write(2) and related system calls + + aio_write: called by io_submit(2) and other asynchronous I/O operations + + readdir: called when the VFS needs to read the directory contents + + poll: called by the VFS when a process wants to check if there is + activity on this file and (optionally) go to sleep until there + is activity. Called by the select(2) and poll(2) system calls + + unlocked_ioctl: called by the ioctl(2) system call. + + compat_ioctl: called by the ioctl(2) system call when 32 bit system calls + are used on 64 bit kernels. + + mmap: called by the mmap(2) system call + + open: called by the VFS when an inode should be opened. When the VFS + opens a file, it creates a new "struct file". It then calls the + open method for the newly allocated file structure. You might + think that the open method really belongs in + "struct inode_operations", and you may be right. I think it's + done the way it is because it makes filesystems simpler to + implement. The open() method is a good place to initialize the + "private_data" member in the file structure if you want to point + to a device structure + + flush: called by the close(2) system call to flush a file + + release: called when the last reference to an open file is closed + + fsync: called by the fsync(2) system call + + fasync: called by the fcntl(2) system call when asynchronous + (non-blocking) mode is enabled for a file + + lock: called by the fcntl(2) system call for F_GETLK, F_SETLK, and F_SETLKW + commands + + readv: called by the readv(2) system call + + writev: called by the writev(2) system call + + sendfile: called by the sendfile(2) system call + + get_unmapped_area: called by the mmap(2) system call + + check_flags: called by the fcntl(2) system call for F_SETFL command + + flock: called by the flock(2) system call + + splice_write: called by the VFS to splice data from a pipe to a file. This + method is used by the splice(2) system call + + splice_read: called by the VFS to splice data from file to a pipe. This + method is used by the splice(2) system call + +Note that the file operations are implemented by the specific +filesystem in which the inode resides. When opening a device node +(character or block special) most filesystems will call special +support routines in the VFS which will locate the required device +driver information. These support routines replace the filesystem file +operations with those for the device driver, and then proceed to call +the new open() method for the file. This is how opening a device file +in the filesystem eventually ends up calling the device driver open() +method. + + +Directory Entry Cache (dcache) +============================== + + +struct dentry_operations +------------------------ + +This describes how a filesystem can overload the standard dentry +operations. Dentries and the dcache are the domain of the VFS and the +individual filesystem implementations. Device drivers have no business +here. These methods may be set to NULL, as they are either optional or +the VFS uses a default. As of kernel 2.6.22, the following members are +defined: + +struct dentry_operations { + int (*d_revalidate)(struct dentry *, struct nameidata *); + int (*d_hash)(const struct dentry *, const struct inode *, + struct qstr *); + int (*d_compare)(const struct dentry *, const struct inode *, + const struct dentry *, const struct inode *, + unsigned int, const char *, const struct qstr *); + int (*d_delete)(const struct dentry *); + void (*d_release)(struct dentry *); + void (*d_iput)(struct dentry *, struct inode *); + char *(*d_dname)(struct dentry *, char *, int); + struct vfsmount *(*d_automount)(struct path *); + int (*d_manage)(struct dentry *, bool); +}; + + d_revalidate: called when the VFS needs to revalidate a dentry. This + is called whenever a name look-up finds a dentry in the + dcache. Most filesystems leave this as NULL, because all their + dentries in the dcache are valid + + d_revalidate may be called in rcu-walk mode (nd->flags & LOOKUP_RCU). + If in rcu-walk mode, the filesystem must revalidate the dentry without + blocking or storing to the dentry, d_parent and d_inode should not be + used without care (because they can go NULL), instead nd->inode should + be used. + + If a situation is encountered that rcu-walk cannot handle, return + -ECHILD and it will be called again in ref-walk mode. + + d_hash: called when the VFS adds a dentry to the hash table. The first + dentry passed to d_hash is the parent directory that the name is + to be hashed into. The inode is the dentry's inode. + + Same locking and synchronisation rules as d_compare regarding + what is safe to dereference etc. + + d_compare: called to compare a dentry name with a given name. The first + dentry is the parent of the dentry to be compared, the second is + the parent's inode, then the dentry and inode (may be NULL) of the + child dentry. len and name string are properties of the dentry to be + compared. qstr is the name to compare it with. + + Must be constant and idempotent, and should not take locks if + possible, and should not or store into the dentry or inodes. + Should not dereference pointers outside the dentry or inodes without + lots of care (eg. d_parent, d_inode, d_name should not be used). + + However, our vfsmount is pinned, and RCU held, so the dentries and + inodes won't disappear, neither will our sb or filesystem module. + ->i_sb and ->d_sb may be used. + + It is a tricky calling convention because it needs to be called under + "rcu-walk", ie. without any locks or references on things. + + d_delete: called when the last reference to a dentry is dropped and the + dcache is deciding whether or not to cache it. Return 1 to delete + immediately, or 0 to cache the dentry. Default is NULL which means to + always cache a reachable dentry. d_delete must be constant and + idempotent. + + d_release: called when a dentry is really deallocated + + d_iput: called when a dentry loses its inode (just prior to its + being deallocated). The default when this is NULL is that the + VFS calls iput(). If you define this method, you must call + iput() yourself + + d_dname: called when the pathname of a dentry should be generated. + Useful for some pseudo filesystems (sockfs, pipefs, ...) to delay + pathname generation. (Instead of doing it when dentry is created, + it's done only when the path is needed.). Real filesystems probably + dont want to use it, because their dentries are present in global + dcache hash, so their hash should be an invariant. As no lock is + held, d_dname() should not try to modify the dentry itself, unless + appropriate SMP safety is used. CAUTION : d_path() logic is quite + tricky. The correct way to return for example "Hello" is to put it + at the end of the buffer, and returns a pointer to the first char. + dynamic_dname() helper function is provided to take care of this. + + d_automount: called when an automount dentry is to be traversed (optional). + This should create a new VFS mount record and return the record to the + caller. The caller is supplied with a path parameter giving the + automount directory to describe the automount target and the parent + VFS mount record to provide inheritable mount parameters. NULL should + be returned if someone else managed to make the automount first. If + the vfsmount creation failed, then an error code should be returned. + If -EISDIR is returned, then the directory will be treated as an + ordinary directory and returned to pathwalk to continue walking. + + If a vfsmount is returned, the caller will attempt to mount it on the + mountpoint and will remove the vfsmount from its expiration list in + the case of failure. The vfsmount should be returned with 2 refs on + it to prevent automatic expiration - the caller will clean up the + additional ref. + + This function is only used if DCACHE_NEED_AUTOMOUNT is set on the + dentry. This is set by __d_instantiate() if S_AUTOMOUNT is set on the + inode being added. + + d_manage: called to allow the filesystem to manage the transition from a + dentry (optional). This allows autofs, for example, to hold up clients + waiting to explore behind a 'mountpoint' whilst letting the daemon go + past and construct the subtree there. 0 should be returned to let the + calling process continue. -EISDIR can be returned to tell pathwalk to + use this directory as an ordinary directory and to ignore anything + mounted on it and not to check the automount flag. Any other error + code will abort pathwalk completely. + + If the 'rcu_walk' parameter is true, then the caller is doing a + pathwalk in RCU-walk mode. Sleeping is not permitted in this mode, + and the caller can be asked to leave it and call again by returing + -ECHILD. + + This function is only used if DCACHE_MANAGE_TRANSIT is set on the + dentry being transited from. + +Example : + +static char *pipefs_dname(struct dentry *dent, char *buffer, int buflen) +{ + return dynamic_dname(dentry, buffer, buflen, "pipe:[%lu]", + dentry->d_inode->i_ino); +} + +Each dentry has a pointer to its parent dentry, as well as a hash list +of child dentries. Child dentries are basically like files in a +directory. + + +Directory Entry Cache API +-------------------------- + +There are a number of functions defined which permit a filesystem to +manipulate dentries: + + dget: open a new handle for an existing dentry (this just increments + the usage count) + + dput: close a handle for a dentry (decrements the usage count). If + the usage count drops to 0, and the dentry is still in its + parent's hash, the "d_delete" method is called to check whether + it should be cached. If it should not be cached, or if the dentry + is not hashed, it is deleted. Otherwise cached dentries are put + into an LRU list to be reclaimed on memory shortage. + + d_drop: this unhashes a dentry from its parents hash list. A + subsequent call to dput() will deallocate the dentry if its + usage count drops to 0 + + d_delete: delete a dentry. If there are no other open references to + the dentry then the dentry is turned into a negative dentry + (the d_iput() method is called). If there are other + references, then d_drop() is called instead + + d_add: add a dentry to its parents hash list and then calls + d_instantiate() + + d_instantiate: add a dentry to the alias hash list for the inode and + updates the "d_inode" member. The "i_count" member in the + inode structure should be set/incremented. If the inode + pointer is NULL, the dentry is called a "negative + dentry". This function is commonly called when an inode is + created for an existing negative dentry + + d_lookup: look up a dentry given its parent and path name component + It looks up the child of that given name from the dcache + hash table. If it is found, the reference count is incremented + and the dentry is returned. The caller must use dput() + to free the dentry when it finishes using it. + +For further information on dentry locking, please refer to the document +Documentation/filesystems/dentry-locking.txt. + +Mount Options +============= + +Parsing options +--------------- + +On mount and remount the filesystem is passed a string containing a +comma separated list of mount options. The options can have either of +these forms: + + option + option=value + +The <linux/parser.h> header defines an API that helps parse these +options. There are plenty of examples on how to use it in existing +filesystems. + +Showing options +--------------- + +If a filesystem accepts mount options, it must define show_options() +to show all the currently active options. The rules are: + + - options MUST be shown which are not default or their values differ + from the default + + - options MAY be shown which are enabled by default or have their + default value + +Options used only internally between a mount helper and the kernel +(such as file descriptors), or which only have an effect during the +mounting (such as ones controlling the creation of a journal) are exempt +from the above rules. + +The underlying reason for the above rules is to make sure, that a +mount can be accurately replicated (e.g. umounting and mounting again) +based on the information found in /proc/mounts. + +A simple method of saving options at mount/remount time and showing +them is provided with the save_mount_options() and +generic_show_options() helper functions. Please note, that using +these may have drawbacks. For more info see header comments for these +functions in fs/namespace.c. + +Resources +========= + +(Note some of these resources are not up-to-date with the latest kernel + version.) + +Creating Linux virtual filesystems. 2002 + <http://lwn.net/Articles/13325/> + +The Linux Virtual File-system Layer by Neil Brown. 1999 + <http://www.cse.unsw.edu.au/~neilb/oss/linux-commentary/vfs.html> + +A tour of the Linux VFS by Michael K. Johnson. 1996 + <http://www.tldp.org/LDP/khg/HyperNews/get/fs/vfstour.html> + +A small trail through the Linux kernel by Andries Brouwer. 2001 + <http://www.win.tue.nl/~aeb/linux/vfs/trail.html> diff --git a/Documentation/filesystems/xfs-delayed-logging-design.txt b/Documentation/filesystems/xfs-delayed-logging-design.txt new file mode 100644 index 00000000..2ce36439 --- /dev/null +++ b/Documentation/filesystems/xfs-delayed-logging-design.txt @@ -0,0 +1,793 @@ +XFS Delayed Logging Design +-------------------------- + +Introduction to Re-logging in XFS +--------------------------------- + +XFS logging is a combination of logical and physical logging. Some objects, +such as inodes and dquots, are logged in logical format where the details +logged are made up of the changes to in-core structures rather than on-disk +structures. Other objects - typically buffers - have their physical changes +logged. The reason for these differences is to reduce the amount of log space +required for objects that are frequently logged. Some parts of inodes are more +frequently logged than others, and inodes are typically more frequently logged +than any other object (except maybe the superblock buffer) so keeping the +amount of metadata logged low is of prime importance. + +The reason that this is such a concern is that XFS allows multiple separate +modifications to a single object to be carried in the log at any given time. +This allows the log to avoid needing to flush each change to disk before +recording a new change to the object. XFS does this via a method called +"re-logging". Conceptually, this is quite simple - all it requires is that any +new change to the object is recorded with a *new copy* of all the existing +changes in the new transaction that is written to the log. + +That is, if we have a sequence of changes A through to F, and the object was +written to disk after change D, we would see in the log the following series +of transactions, their contents and the log sequence number (LSN) of the +transaction: + + Transaction Contents LSN + A A X + B A+B X+n + C A+B+C X+n+m + D A+B+C+D X+n+m+o + <object written to disk> + E E Y (> X+n+m+o) + F E+F Yٍ+p + +In other words, each time an object is relogged, the new transaction contains +the aggregation of all the previous changes currently held only in the log. + +This relogging technique also allows objects to be moved forward in the log so +that an object being relogged does not prevent the tail of the log from ever +moving forward. This can be seen in the table above by the changing +(increasing) LSN of each subsequent transaction - the LSN is effectively a +direct encoding of the location in the log of the transaction. + +This relogging is also used to implement long-running, multiple-commit +transactions. These transaction are known as rolling transactions, and require +a special log reservation known as a permanent transaction reservation. A +typical example of a rolling transaction is the removal of extents from an +inode which can only be done at a rate of two extents per transaction because +of reservation size limitations. Hence a rolling extent removal transaction +keeps relogging the inode and btree buffers as they get modified in each +removal operation. This keeps them moving forward in the log as the operation +progresses, ensuring that current operation never gets blocked by itself if the +log wraps around. + +Hence it can be seen that the relogging operation is fundamental to the correct +working of the XFS journalling subsystem. From the above description, most +people should be able to see why the XFS metadata operations writes so much to +the log - repeated operations to the same objects write the same changes to +the log over and over again. Worse is the fact that objects tend to get +dirtier as they get relogged, so each subsequent transaction is writing more +metadata into the log. + +Another feature of the XFS transaction subsystem is that most transactions are +asynchronous. That is, they don't commit to disk until either a log buffer is +filled (a log buffer can hold multiple transactions) or a synchronous operation +forces the log buffers holding the transactions to disk. This means that XFS is +doing aggregation of transactions in memory - batching them, if you like - to +minimise the impact of the log IO on transaction throughput. + +The limitation on asynchronous transaction throughput is the number and size of +log buffers made available by the log manager. By default there are 8 log +buffers available and the size of each is 32kB - the size can be increased up +to 256kB by use of a mount option. + +Effectively, this gives us the maximum bound of outstanding metadata changes +that can be made to the filesystem at any point in time - if all the log +buffers are full and under IO, then no more transactions can be committed until +the current batch completes. It is now common for a single current CPU core to +be to able to issue enough transactions to keep the log buffers full and under +IO permanently. Hence the XFS journalling subsystem can be considered to be IO +bound. + +Delayed Logging: Concepts +------------------------- + +The key thing to note about the asynchronous logging combined with the +relogging technique XFS uses is that we can be relogging changed objects +multiple times before they are committed to disk in the log buffers. If we +return to the previous relogging example, it is entirely possible that +transactions A through D are committed to disk in the same log buffer. + +That is, a single log buffer may contain multiple copies of the same object, +but only one of those copies needs to be there - the last one "D", as it +contains all the changes from the previous changes. In other words, we have one +necessary copy in the log buffer, and three stale copies that are simply +wasting space. When we are doing repeated operations on the same set of +objects, these "stale objects" can be over 90% of the space used in the log +buffers. It is clear that reducing the number of stale objects written to the +log would greatly reduce the amount of metadata we write to the log, and this +is the fundamental goal of delayed logging. + +From a conceptual point of view, XFS is already doing relogging in memory (where +memory == log buffer), only it is doing it extremely inefficiently. It is using +logical to physical formatting to do the relogging because there is no +infrastructure to keep track of logical changes in memory prior to physically +formatting the changes in a transaction to the log buffer. Hence we cannot avoid +accumulating stale objects in the log buffers. + +Delayed logging is the name we've given to keeping and tracking transactional +changes to objects in memory outside the log buffer infrastructure. Because of +the relogging concept fundamental to the XFS journalling subsystem, this is +actually relatively easy to do - all the changes to logged items are already +tracked in the current infrastructure. The big problem is how to accumulate +them and get them to the log in a consistent, recoverable manner. +Describing the problems and how they have been solved is the focus of this +document. + +One of the key changes that delayed logging makes to the operation of the +journalling subsystem is that it disassociates the amount of outstanding +metadata changes from the size and number of log buffers available. In other +words, instead of there only being a maximum of 2MB of transaction changes not +written to the log at any point in time, there may be a much greater amount +being accumulated in memory. Hence the potential for loss of metadata on a +crash is much greater than for the existing logging mechanism. + +It should be noted that this does not change the guarantee that log recovery +will result in a consistent filesystem. What it does mean is that as far as the +recovered filesystem is concerned, there may be many thousands of transactions +that simply did not occur as a result of the crash. This makes it even more +important that applications that care about their data use fsync() where they +need to ensure application level data integrity is maintained. + +It should be noted that delayed logging is not an innovative new concept that +warrants rigorous proofs to determine whether it is correct or not. The method +of accumulating changes in memory for some period before writing them to the +log is used effectively in many filesystems including ext3 and ext4. Hence +no time is spent in this document trying to convince the reader that the +concept is sound. Instead it is simply considered a "solved problem" and as +such implementing it in XFS is purely an exercise in software engineering. + +The fundamental requirements for delayed logging in XFS are simple: + + 1. Reduce the amount of metadata written to the log by at least + an order of magnitude. + 2. Supply sufficient statistics to validate Requirement #1. + 3. Supply sufficient new tracing infrastructure to be able to debug + problems with the new code. + 4. No on-disk format change (metadata or log format). + 5. Enable and disable with a mount option. + 6. No performance regressions for synchronous transaction workloads. + +Delayed Logging: Design +----------------------- + +Storing Changes + +The problem with accumulating changes at a logical level (i.e. just using the +existing log item dirty region tracking) is that when it comes to writing the +changes to the log buffers, we need to ensure that the object we are formatting +is not changing while we do this. This requires locking the object to prevent +concurrent modification. Hence flushing the logical changes to the log would +require us to lock every object, format them, and then unlock them again. + +This introduces lots of scope for deadlocks with transactions that are already +running. For example, a transaction has object A locked and modified, but needs +the delayed logging tracking lock to commit the transaction. However, the +flushing thread has the delayed logging tracking lock already held, and is +trying to get the lock on object A to flush it to the log buffer. This appears +to be an unsolvable deadlock condition, and it was solving this problem that +was the barrier to implementing delayed logging for so long. + +The solution is relatively simple - it just took a long time to recognise it. +Put simply, the current logging code formats the changes to each item into an +vector array that points to the changed regions in the item. The log write code +simply copies the memory these vectors point to into the log buffer during +transaction commit while the item is locked in the transaction. Instead of +using the log buffer as the destination of the formatting code, we can use an +allocated memory buffer big enough to fit the formatted vector. + +If we then copy the vector into the memory buffer and rewrite the vector to +point to the memory buffer rather than the object itself, we now have a copy of +the changes in a format that is compatible with the log buffer writing code. +that does not require us to lock the item to access. This formatting and +rewriting can all be done while the object is locked during transaction commit, +resulting in a vector that is transactionally consistent and can be accessed +without needing to lock the owning item. + +Hence we avoid the need to lock items when we need to flush outstanding +asynchronous transactions to the log. The differences between the existing +formatting method and the delayed logging formatting can be seen in the +diagram below. + +Current format log vector: + +Object +---------------------------------------------+ +Vector 1 +----+ +Vector 2 +----+ +Vector 3 +----------+ + +After formatting: + +Log Buffer +-V1-+-V2-+----V3----+ + +Delayed logging vector: + +Object +---------------------------------------------+ +Vector 1 +----+ +Vector 2 +----+ +Vector 3 +----------+ + +After formatting: + +Memory Buffer +-V1-+-V2-+----V3----+ +Vector 1 +----+ +Vector 2 +----+ +Vector 3 +----------+ + +The memory buffer and associated vector need to be passed as a single object, +but still need to be associated with the parent object so if the object is +relogged we can replace the current memory buffer with a new memory buffer that +contains the latest changes. + +The reason for keeping the vector around after we've formatted the memory +buffer is to support splitting vectors across log buffer boundaries correctly. +If we don't keep the vector around, we do not know where the region boundaries +are in the item, so we'd need a new encapsulation method for regions in the log +buffer writing (i.e. double encapsulation). This would be an on-disk format +change and as such is not desirable. It also means we'd have to write the log +region headers in the formatting stage, which is problematic as there is per +region state that needs to be placed into the headers during the log write. + +Hence we need to keep the vector, but by attaching the memory buffer to it and +rewriting the vector addresses to point at the memory buffer we end up with a +self-describing object that can be passed to the log buffer write code to be +handled in exactly the same manner as the existing log vectors are handled. +Hence we avoid needing a new on-disk format to handle items that have been +relogged in memory. + + +Tracking Changes + +Now that we can record transactional changes in memory in a form that allows +them to be used without limitations, we need to be able to track and accumulate +them so that they can be written to the log at some later point in time. The +log item is the natural place to store this vector and buffer, and also makes sense +to be the object that is used to track committed objects as it will always +exist once the object has been included in a transaction. + +The log item is already used to track the log items that have been written to +the log but not yet written to disk. Such log items are considered "active" +and as such are stored in the Active Item List (AIL) which is a LSN-ordered +double linked list. Items are inserted into this list during log buffer IO +completion, after which they are unpinned and can be written to disk. An object +that is in the AIL can be relogged, which causes the object to be pinned again +and then moved forward in the AIL when the log buffer IO completes for that +transaction. + +Essentially, this shows that an item that is in the AIL can still be modified +and relogged, so any tracking must be separate to the AIL infrastructure. As +such, we cannot reuse the AIL list pointers for tracking committed items, nor +can we store state in any field that is protected by the AIL lock. Hence the +committed item tracking needs it's own locks, lists and state fields in the log +item. + +Similar to the AIL, tracking of committed items is done through a new list +called the Committed Item List (CIL). The list tracks log items that have been +committed and have formatted memory buffers attached to them. It tracks objects +in transaction commit order, so when an object is relogged it is removed from +it's place in the list and re-inserted at the tail. This is entirely arbitrary +and done to make it easy for debugging - the last items in the list are the +ones that are most recently modified. Ordering of the CIL is not necessary for +transactional integrity (as discussed in the next section) so the ordering is +done for convenience/sanity of the developers. + + +Delayed Logging: Checkpoints + +When we have a log synchronisation event, commonly known as a "log force", +all the items in the CIL must be written into the log via the log buffers. +We need to write these items in the order that they exist in the CIL, and they +need to be written as an atomic transaction. The need for all the objects to be +written as an atomic transaction comes from the requirements of relogging and +log replay - all the changes in all the objects in a given transaction must +either be completely replayed during log recovery, or not replayed at all. If +a transaction is not replayed because it is not complete in the log, then +no later transactions should be replayed, either. + +To fulfill this requirement, we need to write the entire CIL in a single log +transaction. Fortunately, the XFS log code has no fixed limit on the size of a +transaction, nor does the log replay code. The only fundamental limit is that +the transaction cannot be larger than just under half the size of the log. The +reason for this limit is that to find the head and tail of the log, there must +be at least one complete transaction in the log at any given time. If a +transaction is larger than half the log, then there is the possibility that a +crash during the write of a such a transaction could partially overwrite the +only complete previous transaction in the log. This will result in a recovery +failure and an inconsistent filesystem and hence we must enforce the maximum +size of a checkpoint to be slightly less than a half the log. + +Apart from this size requirement, a checkpoint transaction looks no different +to any other transaction - it contains a transaction header, a series of +formatted log items and a commit record at the tail. From a recovery +perspective, the checkpoint transaction is also no different - just a lot +bigger with a lot more items in it. The worst case effect of this is that we +might need to tune the recovery transaction object hash size. + +Because the checkpoint is just another transaction and all the changes to log +items are stored as log vectors, we can use the existing log buffer writing +code to write the changes into the log. To do this efficiently, we need to +minimise the time we hold the CIL locked while writing the checkpoint +transaction. The current log write code enables us to do this easily with the +way it separates the writing of the transaction contents (the log vectors) from +the transaction commit record, but tracking this requires us to have a +per-checkpoint context that travels through the log write process through to +checkpoint completion. + +Hence a checkpoint has a context that tracks the state of the current +checkpoint from initiation to checkpoint completion. A new context is initiated +at the same time a checkpoint transaction is started. That is, when we remove +all the current items from the CIL during a checkpoint operation, we move all +those changes into the current checkpoint context. We then initialise a new +context and attach that to the CIL for aggregation of new transactions. + +This allows us to unlock the CIL immediately after transfer of all the +committed items and effectively allow new transactions to be issued while we +are formatting the checkpoint into the log. It also allows concurrent +checkpoints to be written into the log buffers in the case of log force heavy +workloads, just like the existing transaction commit code does. This, however, +requires that we strictly order the commit records in the log so that +checkpoint sequence order is maintained during log replay. + +To ensure that we can be writing an item into a checkpoint transaction at +the same time another transaction modifies the item and inserts the log item +into the new CIL, then checkpoint transaction commit code cannot use log items +to store the list of log vectors that need to be written into the transaction. +Hence log vectors need to be able to be chained together to allow them to be +detached from the log items. That is, when the CIL is flushed the memory +buffer and log vector attached to each log item needs to be attached to the +checkpoint context so that the log item can be released. In diagrammatic form, +the CIL would look like this before the flush: + + CIL Head + | + V + Log Item <-> log vector 1 -> memory buffer + | -> vector array + V + Log Item <-> log vector 2 -> memory buffer + | -> vector array + V + ...... + | + V + Log Item <-> log vector N-1 -> memory buffer + | -> vector array + V + Log Item <-> log vector N -> memory buffer + -> vector array + +And after the flush the CIL head is empty, and the checkpoint context log +vector list would look like: + + Checkpoint Context + | + V + log vector 1 -> memory buffer + | -> vector array + | -> Log Item + V + log vector 2 -> memory buffer + | -> vector array + | -> Log Item + V + ...... + | + V + log vector N-1 -> memory buffer + | -> vector array + | -> Log Item + V + log vector N -> memory buffer + -> vector array + -> Log Item + +Once this transfer is done, the CIL can be unlocked and new transactions can +start, while the checkpoint flush code works over the log vector chain to +commit the checkpoint. + +Once the checkpoint is written into the log buffers, the checkpoint context is +attached to the log buffer that the commit record was written to along with a +completion callback. Log IO completion will call that callback, which can then +run transaction committed processing for the log items (i.e. insert into AIL +and unpin) in the log vector chain and then free the log vector chain and +checkpoint context. + +Discussion Point: I am uncertain as to whether the log item is the most +efficient way to track vectors, even though it seems like the natural way to do +it. The fact that we walk the log items (in the CIL) just to chain the log +vectors and break the link between the log item and the log vector means that +we take a cache line hit for the log item list modification, then another for +the log vector chaining. If we track by the log vectors, then we only need to +break the link between the log item and the log vector, which means we should +dirty only the log item cachelines. Normally I wouldn't be concerned about one +vs two dirty cachelines except for the fact I've seen upwards of 80,000 log +vectors in one checkpoint transaction. I'd guess this is a "measure and +compare" situation that can be done after a working and reviewed implementation +is in the dev tree.... + +Delayed Logging: Checkpoint Sequencing + +One of the key aspects of the XFS transaction subsystem is that it tags +committed transactions with the log sequence number of the transaction commit. +This allows transactions to be issued asynchronously even though there may be +future operations that cannot be completed until that transaction is fully +committed to the log. In the rare case that a dependent operation occurs (e.g. +re-using a freed metadata extent for a data extent), a special, optimised log +force can be issued to force the dependent transaction to disk immediately. + +To do this, transactions need to record the LSN of the commit record of the +transaction. This LSN comes directly from the log buffer the transaction is +written into. While this works just fine for the existing transaction +mechanism, it does not work for delayed logging because transactions are not +written directly into the log buffers. Hence some other method of sequencing +transactions is required. + +As discussed in the checkpoint section, delayed logging uses per-checkpoint +contexts, and as such it is simple to assign a sequence number to each +checkpoint. Because the switching of checkpoint contexts must be done +atomically, it is simple to ensure that each new context has a monotonically +increasing sequence number assigned to it without the need for an external +atomic counter - we can just take the current context sequence number and add +one to it for the new context. + +Then, instead of assigning a log buffer LSN to the transaction commit LSN +during the commit, we can assign the current checkpoint sequence. This allows +operations that track transactions that have not yet completed know what +checkpoint sequence needs to be committed before they can continue. As a +result, the code that forces the log to a specific LSN now needs to ensure that +the log forces to a specific checkpoint. + +To ensure that we can do this, we need to track all the checkpoint contexts +that are currently committing to the log. When we flush a checkpoint, the +context gets added to a "committing" list which can be searched. When a +checkpoint commit completes, it is removed from the committing list. Because +the checkpoint context records the LSN of the commit record for the checkpoint, +we can also wait on the log buffer that contains the commit record, thereby +using the existing log force mechanisms to execute synchronous forces. + +It should be noted that the synchronous forces may need to be extended with +mitigation algorithms similar to the current log buffer code to allow +aggregation of multiple synchronous transactions if there are already +synchronous transactions being flushed. Investigation of the performance of the +current design is needed before making any decisions here. + +The main concern with log forces is to ensure that all the previous checkpoints +are also committed to disk before the one we need to wait for. Therefore we +need to check that all the prior contexts in the committing list are also +complete before waiting on the one we need to complete. We do this +synchronisation in the log force code so that we don't need to wait anywhere +else for such serialisation - it only matters when we do a log force. + +The only remaining complexity is that a log force now also has to handle the +case where the forcing sequence number is the same as the current context. That +is, we need to flush the CIL and potentially wait for it to complete. This is a +simple addition to the existing log forcing code to check the sequence numbers +and push if required. Indeed, placing the current sequence checkpoint flush in +the log force code enables the current mechanism for issuing synchronous +transactions to remain untouched (i.e. commit an asynchronous transaction, then +force the log at the LSN of that transaction) and so the higher level code +behaves the same regardless of whether delayed logging is being used or not. + +Delayed Logging: Checkpoint Log Space Accounting + +The big issue for a checkpoint transaction is the log space reservation for the +transaction. We don't know how big a checkpoint transaction is going to be +ahead of time, nor how many log buffers it will take to write out, nor the +number of split log vector regions are going to be used. We can track the +amount of log space required as we add items to the commit item list, but we +still need to reserve the space in the log for the checkpoint. + +A typical transaction reserves enough space in the log for the worst case space +usage of the transaction. The reservation accounts for log record headers, +transaction and region headers, headers for split regions, buffer tail padding, +etc. as well as the actual space for all the changed metadata in the +transaction. While some of this is fixed overhead, much of it is dependent on +the size of the transaction and the number of regions being logged (the number +of log vectors in the transaction). + +An example of the differences would be logging directory changes versus logging +inode changes. If you modify lots of inode cores (e.g. chmod -R g+w *), then +there are lots of transactions that only contain an inode core and an inode log +format structure. That is, two vectors totaling roughly 150 bytes. If we modify +10,000 inodes, we have about 1.5MB of metadata to write in 20,000 vectors. Each +vector is 12 bytes, so the total to be logged is approximately 1.75MB. In +comparison, if we are logging full directory buffers, they are typically 4KB +each, so we in 1.5MB of directory buffers we'd have roughly 400 buffers and a +buffer format structure for each buffer - roughly 800 vectors or 1.51MB total +space. From this, it should be obvious that a static log space reservation is +not particularly flexible and is difficult to select the "optimal value" for +all workloads. + +Further, if we are going to use a static reservation, which bit of the entire +reservation does it cover? We account for space used by the transaction +reservation by tracking the space currently used by the object in the CIL and +then calculating the increase or decrease in space used as the object is +relogged. This allows for a checkpoint reservation to only have to account for +log buffer metadata used such as log header records. + +However, even using a static reservation for just the log metadata is +problematic. Typically log record headers use at least 16KB of log space per +1MB of log space consumed (512 bytes per 32k) and the reservation needs to be +large enough to handle arbitrary sized checkpoint transactions. This +reservation needs to be made before the checkpoint is started, and we need to +be able to reserve the space without sleeping. For a 8MB checkpoint, we need a +reservation of around 150KB, which is a non-trivial amount of space. + +A static reservation needs to manipulate the log grant counters - we can take a +permanent reservation on the space, but we still need to make sure we refresh +the write reservation (the actual space available to the transaction) after +every checkpoint transaction completion. Unfortunately, if this space is not +available when required, then the regrant code will sleep waiting for it. + +The problem with this is that it can lead to deadlocks as we may need to commit +checkpoints to be able to free up log space (refer back to the description of +rolling transactions for an example of this). Hence we *must* always have +space available in the log if we are to use static reservations, and that is +very difficult and complex to arrange. It is possible to do, but there is a +simpler way. + +The simpler way of doing this is tracking the entire log space used by the +items in the CIL and using this to dynamically calculate the amount of log +space required by the log metadata. If this log metadata space changes as a +result of a transaction commit inserting a new memory buffer into the CIL, then +the difference in space required is removed from the transaction that causes +the change. Transactions at this level will *always* have enough space +available in their reservation for this as they have already reserved the +maximal amount of log metadata space they require, and such a delta reservation +will always be less than or equal to the maximal amount in the reservation. + +Hence we can grow the checkpoint transaction reservation dynamically as items +are added to the CIL and avoid the need for reserving and regranting log space +up front. This avoids deadlocks and removes a blocking point from the +checkpoint flush code. + +As mentioned early, transactions can't grow to more than half the size of the +log. Hence as part of the reservation growing, we need to also check the size +of the reservation against the maximum allowed transaction size. If we reach +the maximum threshold, we need to push the CIL to the log. This is effectively +a "background flush" and is done on demand. This is identical to +a CIL push triggered by a log force, only that there is no waiting for the +checkpoint commit to complete. This background push is checked and executed by +transaction commit code. + +If the transaction subsystem goes idle while we still have items in the CIL, +they will be flushed by the periodic log force issued by the xfssyncd. This log +force will push the CIL to disk, and if the transaction subsystem stays idle, +allow the idle log to be covered (effectively marked clean) in exactly the same +manner that is done for the existing logging method. A discussion point is +whether this log force needs to be done more frequently than the current rate +which is once every 30s. + + +Delayed Logging: Log Item Pinning + +Currently log items are pinned during transaction commit while the items are +still locked. This happens just after the items are formatted, though it could +be done any time before the items are unlocked. The result of this mechanism is +that items get pinned once for every transaction that is committed to the log +buffers. Hence items that are relogged in the log buffers will have a pin count +for every outstanding transaction they were dirtied in. When each of these +transactions is completed, they will unpin the item once. As a result, the item +only becomes unpinned when all the transactions complete and there are no +pending transactions. Thus the pinning and unpinning of a log item is symmetric +as there is a 1:1 relationship with transaction commit and log item completion. + +For delayed logging, however, we have an asymmetric transaction commit to +completion relationship. Every time an object is relogged in the CIL it goes +through the commit process without a corresponding completion being registered. +That is, we now have a many-to-one relationship between transaction commit and +log item completion. The result of this is that pinning and unpinning of the +log items becomes unbalanced if we retain the "pin on transaction commit, unpin +on transaction completion" model. + +To keep pin/unpin symmetry, the algorithm needs to change to a "pin on +insertion into the CIL, unpin on checkpoint completion". In other words, the +pinning and unpinning becomes symmetric around a checkpoint context. We have to +pin the object the first time it is inserted into the CIL - if it is already in +the CIL during a transaction commit, then we do not pin it again. Because there +can be multiple outstanding checkpoint contexts, we can still see elevated pin +counts, but as each checkpoint completes the pin count will retain the correct +value according to it's context. + +Just to make matters more slightly more complex, this checkpoint level context +for the pin count means that the pinning of an item must take place under the +CIL commit/flush lock. If we pin the object outside this lock, we cannot +guarantee which context the pin count is associated with. This is because of +the fact pinning the item is dependent on whether the item is present in the +current CIL or not. If we don't pin the CIL first before we check and pin the +object, we have a race with CIL being flushed between the check and the pin +(or not pinning, as the case may be). Hence we must hold the CIL flush/commit +lock to guarantee that we pin the items correctly. + +Delayed Logging: Concurrent Scalability + +A fundamental requirement for the CIL is that accesses through transaction +commits must scale to many concurrent commits. The current transaction commit +code does not break down even when there are transactions coming from 2048 +processors at once. The current transaction code does not go any faster than if +there was only one CPU using it, but it does not slow down either. + +As a result, the delayed logging transaction commit code needs to be designed +for concurrency from the ground up. It is obvious that there are serialisation +points in the design - the three important ones are: + + 1. Locking out new transaction commits while flushing the CIL + 2. Adding items to the CIL and updating item space accounting + 3. Checkpoint commit ordering + +Looking at the transaction commit and CIL flushing interactions, it is clear +that we have a many-to-one interaction here. That is, the only restriction on +the number of concurrent transactions that can be trying to commit at once is +the amount of space available in the log for their reservations. The practical +limit here is in the order of several hundred concurrent transactions for a +128MB log, which means that it is generally one per CPU in a machine. + +The amount of time a transaction commit needs to hold out a flush is a +relatively long period of time - the pinning of log items needs to be done +while we are holding out a CIL flush, so at the moment that means it is held +across the formatting of the objects into memory buffers (i.e. while memcpy()s +are in progress). Ultimately a two pass algorithm where the formatting is done +separately to the pinning of objects could be used to reduce the hold time of +the transaction commit side. + +Because of the number of potential transaction commit side holders, the lock +really needs to be a sleeping lock - if the CIL flush takes the lock, we do not +want every other CPU in the machine spinning on the CIL lock. Given that +flushing the CIL could involve walking a list of tens of thousands of log +items, it will get held for a significant time and so spin contention is a +significant concern. Preventing lots of CPUs spinning doing nothing is the +main reason for choosing a sleeping lock even though nothing in either the +transaction commit or CIL flush side sleeps with the lock held. + +It should also be noted that CIL flushing is also a relatively rare operation +compared to transaction commit for asynchronous transaction workloads - only +time will tell if using a read-write semaphore for exclusion will limit +transaction commit concurrency due to cache line bouncing of the lock on the +read side. + +The second serialisation point is on the transaction commit side where items +are inserted into the CIL. Because transactions can enter this code +concurrently, the CIL needs to be protected separately from the above +commit/flush exclusion. It also needs to be an exclusive lock but it is only +held for a very short time and so a spin lock is appropriate here. It is +possible that this lock will become a contention point, but given the short +hold time once per transaction I think that contention is unlikely. + +The final serialisation point is the checkpoint commit record ordering code +that is run as part of the checkpoint commit and log force sequencing. The code +path that triggers a CIL flush (i.e. whatever triggers the log force) will enter +an ordering loop after writing all the log vectors into the log buffers but +before writing the commit record. This loop walks the list of committing +checkpoints and needs to block waiting for checkpoints to complete their commit +record write. As a result it needs a lock and a wait variable. Log force +sequencing also requires the same lock, list walk, and blocking mechanism to +ensure completion of checkpoints. + +These two sequencing operations can use the mechanism even though the +events they are waiting for are different. The checkpoint commit record +sequencing needs to wait until checkpoint contexts contain a commit LSN +(obtained through completion of a commit record write) while log force +sequencing needs to wait until previous checkpoint contexts are removed from +the committing list (i.e. they've completed). A simple wait variable and +broadcast wakeups (thundering herds) has been used to implement these two +serialisation queues. They use the same lock as the CIL, too. If we see too +much contention on the CIL lock, or too many context switches as a result of +the broadcast wakeups these operations can be put under a new spinlock and +given separate wait lists to reduce lock contention and the number of processes +woken by the wrong event. + + +Lifecycle Changes + +The existing log item life cycle is as follows: + + 1. Transaction allocate + 2. Transaction reserve + 3. Lock item + 4. Join item to transaction + If not already attached, + Allocate log item + Attach log item to owner item + Attach log item to transaction + 5. Modify item + Record modifications in log item + 6. Transaction commit + Pin item in memory + Format item into log buffer + Write commit LSN into transaction + Unlock item + Attach transaction to log buffer + + <log buffer IO dispatched> + <log buffer IO completes> + + 7. Transaction completion + Mark log item committed + Insert log item into AIL + Write commit LSN into log item + Unpin log item + 8. AIL traversal + Lock item + Mark log item clean + Flush item to disk + + <item IO completion> + + 9. Log item removed from AIL + Moves log tail + Item unlocked + +Essentially, steps 1-6 operate independently from step 7, which is also +independent of steps 8-9. An item can be locked in steps 1-6 or steps 8-9 +at the same time step 7 is occurring, but only steps 1-6 or 8-9 can occur +at the same time. If the log item is in the AIL or between steps 6 and 7 +and steps 1-6 are re-entered, then the item is relogged. Only when steps 8-9 +are entered and completed is the object considered clean. + +With delayed logging, there are new steps inserted into the life cycle: + + 1. Transaction allocate + 2. Transaction reserve + 3. Lock item + 4. Join item to transaction + If not already attached, + Allocate log item + Attach log item to owner item + Attach log item to transaction + 5. Modify item + Record modifications in log item + 6. Transaction commit + Pin item in memory if not pinned in CIL + Format item into log vector + buffer + Attach log vector and buffer to log item + Insert log item into CIL + Write CIL context sequence into transaction + Unlock item + + <next log force> + + 7. CIL push + lock CIL flush + Chain log vectors and buffers together + Remove items from CIL + unlock CIL flush + write log vectors into log + sequence commit records + attach checkpoint context to log buffer + + <log buffer IO dispatched> + <log buffer IO completes> + + 8. Checkpoint completion + Mark log item committed + Insert item into AIL + Write commit LSN into log item + Unpin log item + 9. AIL traversal + Lock item + Mark log item clean + Flush item to disk + <item IO completion> + 10. Log item removed from AIL + Moves log tail + Item unlocked + +From this, it can be seen that the only life cycle differences between the two +logging methods are in the middle of the life cycle - they still have the same +beginning and end and execution constraints. The only differences are in the +committing of the log items to the log itself and the completion processing. +Hence delayed logging should not introduce any constraints on log item +behaviour, allocation or freeing that don't already exist. + +As a result of this zero-impact "insertion" of delayed logging infrastructure +and the design of the internal structures to avoid on disk format changes, we +can basically switch between delayed logging and the existing mechanism with a +mount option. Fundamentally, there is no reason why the log manager would not +be able to swap methods automatically and transparently depending on load +characteristics, but this should not be necessary if delayed logging works as +designed. diff --git a/Documentation/filesystems/xfs.txt b/Documentation/filesystems/xfs.txt new file mode 100644 index 00000000..3fc0c31a --- /dev/null +++ b/Documentation/filesystems/xfs.txt @@ -0,0 +1,252 @@ + +The SGI XFS Filesystem +====================== + +XFS is a high performance journaling filesystem which originated +on the SGI IRIX platform. It is completely multi-threaded, can +support large files and large filesystems, extended attributes, +variable block sizes, is extent based, and makes extensive use of +Btrees (directories, extents, free space) to aid both performance +and scalability. + +Refer to the documentation at http://oss.sgi.com/projects/xfs/ +for further details. This implementation is on-disk compatible +with the IRIX version of XFS. + + +Mount Options +============= + +When mounting an XFS filesystem, the following options are accepted. + + allocsize=size + Sets the buffered I/O end-of-file preallocation size when + doing delayed allocation writeout (default size is 64KiB). + Valid values for this option are page size (typically 4KiB) + through to 1GiB, inclusive, in power-of-2 increments. + + attr2/noattr2 + The options enable/disable (default is disabled for backward + compatibility on-disk) an "opportunistic" improvement to be + made in the way inline extended attributes are stored on-disk. + When the new form is used for the first time (by setting or + removing extended attributes) the on-disk superblock feature + bit field will be updated to reflect this format being in use. + + barrier + Enables the use of block layer write barriers for writes into + the journal and unwritten extent conversion. This allows for + drive level write caching to be enabled, for devices that + support write barriers. + + discard + Issue command to let the block device reclaim space freed by the + filesystem. This is useful for SSD devices, thinly provisioned + LUNs and virtual machine images, but may have a performance + impact. This option is incompatible with the nodelaylog option. + + dmapi + Enable the DMAPI (Data Management API) event callouts. + Use with the "mtpt" option. + + grpid/bsdgroups and nogrpid/sysvgroups + These options define what group ID a newly created file gets. + When grpid is set, it takes the group ID of the directory in + which it is created; otherwise (the default) it takes the fsgid + of the current process, unless the directory has the setgid bit + set, in which case it takes the gid from the parent directory, + and also gets the setgid bit set if it is a directory itself. + + ihashsize=value + In memory inode hashes have been removed, so this option has + no function as of August 2007. Option is deprecated. + + ikeep/noikeep + When ikeep is specified, XFS does not delete empty inode clusters + and keeps them around on disk. ikeep is the traditional XFS + behaviour. When noikeep is specified, empty inode clusters + are returned to the free space pool. The default is noikeep for + non-DMAPI mounts, while ikeep is the default when DMAPI is in use. + + inode64 + Indicates that XFS is allowed to create inodes at any location + in the filesystem, including those which will result in inode + numbers occupying more than 32 bits of significance. This is + provided for backwards compatibility, but causes problems for + backup applications that cannot handle large inode numbers. + + largeio/nolargeio + If "nolargeio" is specified, the optimal I/O reported in + st_blksize by stat(2) will be as small as possible to allow user + applications to avoid inefficient read/modify/write I/O. + If "largeio" specified, a filesystem that has a "swidth" specified + will return the "swidth" value (in bytes) in st_blksize. If the + filesystem does not have a "swidth" specified but does specify + an "allocsize" then "allocsize" (in bytes) will be returned + instead. + If neither of these two options are specified, then filesystem + will behave as if "nolargeio" was specified. + + logbufs=value + Set the number of in-memory log buffers. Valid numbers range + from 2-8 inclusive. + The default value is 8 buffers for filesystems with a + blocksize of 64KiB, 4 buffers for filesystems with a blocksize + of 32KiB, 3 buffers for filesystems with a blocksize of 16KiB + and 2 buffers for all other configurations. Increasing the + number of buffers may increase performance on some workloads + at the cost of the memory used for the additional log buffers + and their associated control structures. + + logbsize=value + Set the size of each in-memory log buffer. + Size may be specified in bytes, or in kilobytes with a "k" suffix. + Valid sizes for version 1 and version 2 logs are 16384 (16k) and + 32768 (32k). Valid sizes for version 2 logs also include + 65536 (64k), 131072 (128k) and 262144 (256k). + The default value for machines with more than 32MiB of memory + is 32768, machines with less memory use 16384 by default. + + logdev=device and rtdev=device + Use an external log (metadata journal) and/or real-time device. + An XFS filesystem has up to three parts: a data section, a log + section, and a real-time section. The real-time section is + optional, and the log section can be separate from the data + section or contained within it. + + mtpt=mountpoint + Use with the "dmapi" option. The value specified here will be + included in the DMAPI mount event, and should be the path of + the actual mountpoint that is used. + + noalign + Data allocations will not be aligned at stripe unit boundaries. + + noatime + Access timestamps are not updated when a file is read. + + norecovery + The filesystem will be mounted without running log recovery. + If the filesystem was not cleanly unmounted, it is likely to + be inconsistent when mounted in "norecovery" mode. + Some files or directories may not be accessible because of this. + Filesystems mounted "norecovery" must be mounted read-only or + the mount will fail. + + nouuid + Don't check for double mounted file systems using the file system uuid. + This is useful to mount LVM snapshot volumes. + + uquota/usrquota/uqnoenforce/quota + User disk quota accounting enabled, and limits (optionally) + enforced. Refer to xfs_quota(8) for further details. + + gquota/grpquota/gqnoenforce + Group disk quota accounting enabled and limits (optionally) + enforced. Refer to xfs_quota(8) for further details. + + pquota/prjquota/pqnoenforce + Project disk quota accounting enabled and limits (optionally) + enforced. Refer to xfs_quota(8) for further details. + + sunit=value and swidth=value + Used to specify the stripe unit and width for a RAID device or + a stripe volume. "value" must be specified in 512-byte block + units. + If this option is not specified and the filesystem was made on + a stripe volume or the stripe width or unit were specified for + the RAID device at mkfs time, then the mount system call will + restore the value from the superblock. For filesystems that + are made directly on RAID devices, these options can be used + to override the information in the superblock if the underlying + disk layout changes after the filesystem has been created. + The "swidth" option is required if the "sunit" option has been + specified, and must be a multiple of the "sunit" value. + + swalloc + Data allocations will be rounded up to stripe width boundaries + when the current end of file is being extended and the file + size is larger than the stripe width size. + + +sysctls +======= + +The following sysctls are available for the XFS filesystem: + + fs.xfs.stats_clear (Min: 0 Default: 0 Max: 1) + Setting this to "1" clears accumulated XFS statistics + in /proc/fs/xfs/stat. It then immediately resets to "0". + + fs.xfs.xfssyncd_centisecs (Min: 100 Default: 3000 Max: 720000) + The interval at which the xfssyncd thread flushes metadata + out to disk. This thread will flush log activity out, and + do some processing on unlinked inodes. + + fs.xfs.xfsbufd_centisecs (Min: 50 Default: 100 Max: 3000) + The interval at which xfsbufd scans the dirty metadata buffers list. + + fs.xfs.age_buffer_centisecs (Min: 100 Default: 1500 Max: 720000) + The age at which xfsbufd flushes dirty metadata buffers to disk. + + fs.xfs.error_level (Min: 0 Default: 3 Max: 11) + A volume knob for error reporting when internal errors occur. + This will generate detailed messages & backtraces for filesystem + shutdowns, for example. Current threshold values are: + + XFS_ERRLEVEL_OFF: 0 + XFS_ERRLEVEL_LOW: 1 + XFS_ERRLEVEL_HIGH: 5 + + fs.xfs.panic_mask (Min: 0 Default: 0 Max: 127) + Causes certain error conditions to call BUG(). Value is a bitmask; + AND together the tags which represent errors which should cause panics: + + XFS_NO_PTAG 0 + XFS_PTAG_IFLUSH 0x00000001 + XFS_PTAG_LOGRES 0x00000002 + XFS_PTAG_AILDELETE 0x00000004 + XFS_PTAG_ERROR_REPORT 0x00000008 + XFS_PTAG_SHUTDOWN_CORRUPT 0x00000010 + XFS_PTAG_SHUTDOWN_IOERROR 0x00000020 + XFS_PTAG_SHUTDOWN_LOGERROR 0x00000040 + + This option is intended for debugging only. + + fs.xfs.irix_symlink_mode (Min: 0 Default: 0 Max: 1) + Controls whether symlinks are created with mode 0777 (default) + or whether their mode is affected by the umask (irix mode). + + fs.xfs.irix_sgid_inherit (Min: 0 Default: 0 Max: 1) + Controls files created in SGID directories. + If the group ID of the new file does not match the effective group + ID or one of the supplementary group IDs of the parent dir, the + ISGID bit is cleared if the irix_sgid_inherit compatibility sysctl + is set. + + fs.xfs.inherit_sync (Min: 0 Default: 1 Max: 1) + Setting this to "1" will cause the "sync" flag set + by the xfs_io(8) chattr command on a directory to be + inherited by files in that directory. + + fs.xfs.inherit_nodump (Min: 0 Default: 1 Max: 1) + Setting this to "1" will cause the "nodump" flag set + by the xfs_io(8) chattr command on a directory to be + inherited by files in that directory. + + fs.xfs.inherit_noatime (Min: 0 Default: 1 Max: 1) + Setting this to "1" will cause the "noatime" flag set + by the xfs_io(8) chattr command on a directory to be + inherited by files in that directory. + + fs.xfs.inherit_nosymlinks (Min: 0 Default: 1 Max: 1) + Setting this to "1" will cause the "nosymlinks" flag set + by the xfs_io(8) chattr command on a directory to be + inherited by files in that directory. + + fs.xfs.rotorstep (Min: 1 Default: 1 Max: 256) + In "inode32" allocation mode, this option determines how many + files the allocator attempts to allocate in the same allocation + group before moving to the next allocation group. The intent + is to control the rate at which the allocator moves between + allocation groups when allocating extents for new files. diff --git a/Documentation/filesystems/xip.txt b/Documentation/filesystems/xip.txt new file mode 100644 index 00000000..0466ee56 --- /dev/null +++ b/Documentation/filesystems/xip.txt @@ -0,0 +1,68 @@ +Execute-in-place for file mappings +---------------------------------- + +Motivation +---------- +File mappings are performed by mapping page cache pages to userspace. In +addition, read&write type file operations also transfer data from/to the page +cache. + +For memory backed storage devices that use the block device interface, the page +cache pages are in fact copies of the original storage. Various approaches +exist to work around the need for an extra copy. The ramdisk driver for example +does read the data into the page cache, keeps a reference, and discards the +original data behind later on. + +Execute-in-place solves this issue the other way around: instead of keeping +data in the page cache, the need to have a page cache copy is eliminated +completely. With execute-in-place, read&write type operations are performed +directly from/to the memory backed storage device. For file mappings, the +storage device itself is mapped directly into userspace. + +This implementation was initially written for shared memory segments between +different virtual machines on s390 hardware to allow multiple machines to +share the same binaries and libraries. + +Implementation +-------------- +Execute-in-place is implemented in three steps: block device operation, +address space operation, and file operations. + +A block device operation named direct_access is used to retrieve a +reference (pointer) to a block on-disk. The reference is supposed to be +cpu-addressable, physical address and remain valid until the release operation +is performed. A struct block_device reference is used to address the device, +and a sector_t argument is used to identify the individual block. As an +alternative, memory technology devices can be used for this. + +The block device operation is optional, these block devices support it as of +today: +- dcssblk: s390 dcss block device driver + +An address space operation named get_xip_mem is used to retrieve references +to a page frame number and a kernel address. To obtain these values a reference +to an address_space is provided. This function assigns values to the kmem and +pfn parameters. The third argument indicates whether the function should allocate +blocks if needed. + +This address space operation is mutually exclusive with readpage&writepage that +do page cache read/write operations. +The following filesystems support it as of today: +- ext2: the second extended filesystem, see Documentation/filesystems/ext2.txt + +A set of file operations that do utilize get_xip_page can be found in +mm/filemap_xip.c . The following file operation implementations are provided: +- aio_read/aio_write +- readv/writev +- sendfile + +The generic file operations do_sync_read/do_sync_write can be used to implement +classic synchronous IO calls. + +Shortcomings +------------ +This implementation is limited to storage devices that are cpu addressable at +all times (no highmem or such). It works well on rom/ram, but enhancements are +needed to make it work with flash in read+write mode. +Putting the Linux kernel and/or its modules on a xip filesystem does not mean +they are not copied. |