In order to support multiple filesystems, Linux contains a special kernel interface level called VFS (Virtual Filesystem Switch). This is similar to the vnode/vfs interface found in SVR4 derivatives (originally it came from BSD and Sun original implementations).
Linux inode cache is implemented in a single file, fs/inode.c, which consists of 977 lines of code. It is interesting to note that not many changes have been made to it for the last 5-7 years: one can still recognise some of the code comparing the latest version with, say, 1.3.42.
The structure of Linux inode cache is as follows:
inode_hashtable, where each inode is hashed by the value of the superblock pointer and 32bit inode number. Inodes without a superblock (inode->i_sb == NULL) are added to a doubly linked list headed by anon_hash_chain instead. Examples of anonymous inodes are sockets created by net/socket.c:sock_alloc(), by calling fs/inode.c:get_empty_inode().inode_in_use), which contains valid inodes with i_count>0 and i_nlink>0. Inodes newly allocated by get_empty_inode() and get_new_inode() are added to the inode_in_use list.inode_unused), which contains valid inodes with i_count = 0.sb->s_dirty) which contains valid inodes with i_count>0, i_nlink>0 and i_state & I_DIRTY. When inode is marked dirty, it is added to the sb->s_dirty list if it is also hashed. Maintaining a per-superblock dirty list of inodes allows to quickly sync inodes.inode_cachep. As inode objects are allocated and freed, they are taken from and returned to this SLAB cache.The type lists are anchored from inode->i_list, the hashtable from inode->i_hash. Each inode can be on a hashtable and one and only one type (in_use, unused or dirty) list.
All these lists are protected by a single spinlock: inode_lock.
The inode cache subsystem is initialised when inode_init() function is called from init/main.c:start_kernel(). The function is marked as __init, which means its code is thrown away later on. It is passed a single argument - the number of physical pages on the system. This is so that the inode cache can configure itself depending on how much memory is available, i.e. create a larger hashtable if there is enough memory.
The only stats information about inode cache is the number of unused inodes, stored in inodes_stat.nr_unused and accessible to user programs via files /proc/sys/fs/inode-nr and /proc/sys/fs/inode-state.
We can examine one of the lists from gdb running on a live kernel thus:
(gdb) printf "%d\n", (unsigned long)(&((struct inode *)0)->i_list) 8 (gdb) p inode_unused $34 = 0xdfa992a8 (gdb) p (struct list_head)inode_unused $35 = {next = 0xdfa992a8, prev = 0xdfcdd5a8} (gdb) p ((struct list_head)inode_unused).prev $36 = (struct list_head *) 0xdfcdd5a8 (gdb) p (((struct list_head)inode_unused).prev)->prev $37 = (struct list_head *) 0xdfb5a2e8 (gdb) set $i = (struct inode *)0xdfb5a2e0 (gdb) p $i->i_ino $38 = 0x3bec7 (gdb) p $i->i_count $39 = {counter = 0x0}
Note that we deducted 8 from the address 0xdfb5a2e8 to obtain the address of the struct inode (0xdfb5a2e0) according to the definition of list_entry() macro from include/linux/list.h.
To understand how inode cache works, let us trace a lifetime of an inode of a regular file on ext2 filesystem as it is opened and closed:
fd = open("file", O_RDONLY); close(fd);
The open(2) system call is implemented in fs/open.c:sys_open function and the real work is done by fs/open.c:filp_open() function, which is split into two parts:
open_namei(): fills in the nameidata structure containing the dentry and vfsmount structures.dentry_open(): given a dentry and vfsmount, this function allocates a new struct file and links them together; it also invokes the filesystem specific f_op->open() method which was set in inode->i_fop when inode was read in open_namei() (which provided inode via dentry->d_inode).The open_namei() function interacts with dentry cache via path_walk(), which in turn calls real_lookup(), which invokes the filesystem specific inode_operations->lookup() method. The role of this method is to find the entry in the parent directory with the matching name and then do iget(sb, ino) to get the corresponding inode - which brings us to the inode cache. When the inode is read in, the dentry is instantiated by means of d_add(dentry, inode). While we are at it, note that for UNIX-style filesystems which have the concept of on-disk inode number, it is the lookup method's job to map its endianness to current CPU format, e.g. if the inode number in raw (fs-specific) dir entry is in little-endian 32 bit format one could do:
unsigned long ino = le32_to_cpu(de->inode); inode = iget(sb, ino); d_add(dentry, inode);
So, when we open a file we hit iget(sb, ino) which is really iget4(sb, ino, NULL, NULL), which does:
inode_lock. If inode is found, its reference count (i_count) is incremented; if it was 0 prior to incrementation and the inode is not dirty, it is removed from whatever type list (inode->i_list) it is currently on (it has to be inode_unused list, of course) and inserted into inode_in_use type list; finally, inodes_stat.nr_unused is decremented.iget4() is guaranteed to return an unlocked inode.get_new_inode(), passing it the pointer to the place in the hashtable where it should be inserted to.get_new_inode() allocates a new inode from the inode_cachep SLAB cache but this operation can block (GFP_KERNEL allocation), so it must drop the inode_lock spinlock which guards the hashtable. Since it has dropped the spinlock, it must retry searching the inode in the hashtable afterwards; if it is found this time, it returns (after incrementing the reference by __iget) the one found in the hashtable and destroys the newly allocated one. If it is still not found in the hashtable, then the new inode we have just allocated is the one to be used; therefore it is initialised to the required values and the fs-specific sb->s_op->read_inode() method is invoked to populate the rest of the inode. This brings us from inode cache back to the filesystem code - remember that we came to the inode cache when filesystem-specific lookup() method invoked iget(). While the s_op->read_inode() method is reading the inode from disk, the inode is locked (i_state = I_LOCK); it is unlocked after the read_inode() method returns and all the waiters for it are woken up.Now, let's see what happens when we close this file descriptor. The close(2) system call is implemented in fs/open.c:sys_close() function, which calls do_close(fd, 1) which rips (replaces with NULL) the descriptor of the process' file descriptor table and invokes the filp_close() function which does most of the work. The interesting things happen in fput(), which checks if this was the last reference to the file, and if so calls fs/file_table.c:_fput() which calls __fput() which is where interaction with dcache (and therefore with inode cache - remember dcache is a Master of inode cache!) happens. The fs/dcache.c:dput() does dentry_iput() which brings us back to inode cache via iput(inode) so let us understand fs/inode.c:iput(inode):
sb->s_op->put_inode() method, it is invoked immediately with no spinlocks held (so it can block).inode_lock spinlock is taken and i_count is decremented. If this was NOT the last reference to this inode then we simply check if there are too many references to it and so i_count can wrap around the 32 bits allocated to it and if so we print a warning and return. Note that we call printk() while holding the inode_lock spinlock - this is fine because printk() can never block, therefore it may be called in absolutely any context (even from interrupt handlers!).The work performed by iput() on the last inode reference is rather complex so we separate it into a list of its own:
i_nlink == 0 (e.g. the file was unlinked while we held it open) then the inode is removed from hashtable and from its type list; if there are any data pages held in page cache for this inode, they are removed by means of truncate_all_inode_pages(&inode->i_data). Then the filesystem-specific s_op->delete_inode() method is invoked, which typically deletes the on-disk copy of the inode. If there is no s_op->delete_inode() method registered by the filesystem (e.g. ramfs) then we call clear_inode(inode), which invokes s_op->clear_inode() if registered and if inode corresponds to a block device, this device's reference count is dropped by bdput(inode->i_bdev).i_nlink != 0 then we check if there are other inodes in the same hash bucket and if there is none, then if inode is not dirty we delete it from its type list and add it to inode_unused list, incrementing inodes_stat.nr_unused. If there are inodes in the same hashbucket then we delete it from the type list and add to inode_unused list. If this was an anonymous inode (NetApp .snapshot) then we delete it from the type list and clear/destroy it completely.The Linux kernel provides a mechanism for new filesystems to be written with minimum effort. The historical reasons for this are:
Let us consider the steps required to implement a filesystem under Linux. The code to implement a filesystem can be either a dynamically loadable module or statically linked into the kernel, and the way it is done under Linux is very transparent. All that is needed is to fill in a struct file_system_type structure and register it with the VFS using the register_filesystem() function as in the following example from fs/bfs/inode.c:
#include <linux/module.h> #include <linux/init.h> static struct super_block *bfs_read_super(struct super_block *, void *, int); static DECLARE_FSTYPE_DEV(bfs_fs_type, "bfs", bfs_read_super); static int __init init_bfs_fs(void) { return register_filesystem(&bfs_fs_type); } static void __exit exit_bfs_fs(void) { unregister_filesystem(&bfs_fs_type); } module_init(init_bfs_fs) module_exit(exit_bfs_fs)
The module_init()/module_exit() macros ensure that, when BFS is compiled as a module, the functions init_bfs_fs() and exit_bfs_fs() turn into init_module() and cleanup_module() respectively; if BFS is statically linked into the kernel, the exit_bfs_fs() code vanishes as it is unnecessary.
The struct file_system_type is declared in include/linux/fs.h:
struct file_system_type { const char *name; int fs_flags; struct super_block *(*read_super) (struct super_block *, void *, int); struct module *owner; struct vfsmount *kern_mnt; /* For kernel mount, if it's FS_SINGLE fs */ struct file_system_type * next; };
The fields thereof are explained thus:
/proc/filesystems file and is used as a key to find a filesystem by its name; this same name is used for the filesystem type in mount(2), and it should be unique: there can (obviously) be only one filesystem with a given name. For modules, name points to module's address spaces and not copied: this means cat /proc/filesystems can oops if the module was unloaded but filesystem is still registered.FS_REQUIRES_DEV for filesystems that can only be mounted on a block device, FS_SINGLE for filesystems that can have only one superblock, FS_NOMOUNT for filesystems that cannot be mounted from userspace by means of mount(2) system call: they can however be mounted internally using kern_mount() interface, e.g. pipefs.FS_SINGLE case where it will Oops in get_sb_single(), trying to dereference a NULL pointer in fs_type->kern_mnt->mnt_sb with (fs_type->kern_mnt = NULL).THIS_MODULE does the right thing automatically.FS_SINGLE filesystems only. This is set by kern_mount() (TODO: kern_mount() should refuse to mount filesystems if FS_SINGLE is not set).file_systems (see fs/super.c). The list is protected by the file_systems_lock read-write spinlock and functions register/unregister_filesystem() modify it by linking and unlinking the entry from the list.The job of the read_super() function is to fill in the fields of the superblock, allocate root inode and initialise any fs-private information associated with this mounted instance of the filesystem. So, typically the read_super() would do:
sb->s_dev argument, using buffer cache bread() function. If it anticipates to read a few more subsequent metadata blocks immediately then it makes sense to use breada() to schedule reading extra blocks asynchronously.sb->s_op to point to struct super_block_operations structure. This structure contains filesystem-specific functions implementing operations like "read inode", "delete inode", etc.d_alloc_root().sb->s_dirt to 1 and mark the buffer containing superblock dirty (TODO: why do we do this? I did it in BFS because MINIX did it...)Under Linux there are several levels of indirection between user file descriptor and the kernel inode structure. When a process makes open(2) system call, the kernel returns a small non-negative integer which can be used for subsequent I/O operations on this file. This integer is an index into an array of pointers to struct file. Each file structure points to a dentry via file->f_dentry. And each dentry points to an inode via dentry->d_inode.
Each task contains a field tsk->files which is a pointer to struct files_struct defined in include/linux/sched.h:
/* * Open file table structure */ struct files_struct { atomic_t count; rwlock_t file_lock; int max_fds; int max_fdset; int next_fd; struct file ** fd; /* current fd array */ fd_set *close_on_exec; fd_set *open_fds; fd_set close_on_exec_init; fd_set open_fds_init; struct file * fd_array[NR_OPEN_DEFAULT]; };
The file->count is a reference count, incremented by get_file() (usually called by fget()) and decremented by fput() and by put_filp(). The difference between fput() and put_filp() is that fput() does more work usually needed for regular files, such as releasing flock locks, releasing dentry, etc, while put_filp() is only manipulating file table structures, i.e. decrements the count, removes the file from the anon_list and adds it to the free_list, under protection of files_lock spinlock.
The tsk->files can be shared between parent and child if the child thread was created using clone() system call with CLONE_FILES set in the clone flags argument. This can be seen in kernel/fork.c:copy_files() (called by do_fork()) which only increments the file->count if CLONE_FILES is set instead of the usual copying file descriptor table in time-honoured tradition of classical UNIX fork(2).
When a file is opened, the file structure allocated for it is installed into current->files->fd[fd] slot and a fd bit is set in the bitmap current->files->open_fds . All this is done under the write protection of current->files->file_lock read-write spinlock. When the descriptor is closed a fd bit is cleared in current->files->open_fds and current->files->next_fd is set equal to fd as a hint for finding the first unused descriptor next time this process wants to open a file.
The file structure is declared in include/linux/fs.h:
struct fown_struct { int pid; /* pid or -pgrp where SIGIO should be sent */ uid_t uid, euid; /* uid/euid of process setting the owner */ int signum; /* posix.1b rt signal to be delivered on IO */ }; struct file { struct list_head f_list; struct dentry *f_dentry; struct vfsmount *f_vfsmnt; struct file_operations *f_op; atomic_t f_count; unsigned int f_flags; mode_t f_mode; loff_t f_pos; unsigned long f_reada, f_ramax, f_raend, f_ralen, f_rawin; struct fown_struct f_owner; unsigned int f_uid, f_gid; int f_error; unsigned long f_version; /* needed for tty driver, and maybe others */ void *private_data; };
Let us look at the various fields of struct file:
sb->s_files list of all open files on this filesystem, if the corresponding inode is not anonymous, then dentry_open() (called by filp_open()) links the file into this list; b) fs/file_table.c:free_list, containing unused file structures; c) fs/file_table.c:anon_list, when a new file structure is created by get_empty_filp() it is placed on this list. All these lists are protected by the files_lock spinlock.open_namei() (or rather path_walk() which it calls) but the actual file->f_dentry field is set by dentry_open() to contain the dentry thus found.vfsmount structure of the filesystem containing the file. This is set by dentry_open() but is found as part of nameidata lookup by open_namei() (or rather path_init() which it calls).file_operations which contains various methods that can be invoked on the file. This is copied from inode->i_fop which is placed there by filesystem-specific s_op->read_inode() method during nameidata lookup. We will look at file_operations methods in detail later on in this section.get_file/put_filp/fput.O_XXX flags from open(2) system call copied there (with slight modifications by filp_open()) by dentry_open() and after clearing O_CREAT, O_EXCL, O_NOCTTY, O_TRUNC - there is no point in storing these flags permanently since they cannot be modified by F_SETFL (or queried by F_GETFL) fcntl(2) calls.dentry_open(). The point of the conversion is to store read and write access in separate bits so one could do easy checks like (f_mode & FMODE_WRITE) and (f_mode & FMODE_READ).long long, i.e. a 64bit value.SIGIO mechanism (see fs/fcntl.c:kill_fasync()).get_empty_filp(). If the file is a socket, used by ipv4 netfilter.fs/nfs/file.c and checked in mm/filemap.c:generic_file_write().event) whenever f_pos changes.file->f_dentry->d_inode->i_rdev.Now let us look at file_operations structure which contains the methods that can be invoked on files. Let us recall that it is copied from inode->i_fop where it is set by s_op->read_inode() method. It is declared in include/linux/fs.h:
struct file_operations { struct module *owner; loff_t (*llseek) (struct file *, loff_t, int); ssize_t (*read) (struct file *, char *, size_t, loff_t *); ssize_t (*write) (struct file *, const char *, size_t, loff_t *); int (*readdir) (struct file *, void *, filldir_t); unsigned int (*poll) (struct file *, struct poll_table_struct *); int (*ioctl) (struct inode *, 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 *, struct dentry *, 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 *); };
THIS_MODULE, filesystems can happily ignore it because their module counts are controlled at mount/umount time whilst the drivers need to control it at open/release time.fs/read_write.c:default_llseek() is used, which does the right thing (TODO: force all those who set it to NULL currently to use default_llseek - that way we save an if() in llseek())read(2) system call. Filesystems can use mm/filemap.c:generic_file_read() for regular files and fs/read_write.c:generic_read_dir() (which simply returns -EISDIR) for directories here.mm/filemap.c:generic_file_write() for regular files and ignore it for directories here.FIBMAP, FIGETBSZ, FIONREAD are implemented by higher levels so they never read f_op->ioctl() method.dentry_open(). Filesystems rarely use this, e.g. coda tries to cache the file locally at open time.release() method below). The only filesystem that uses this is NFS client to flush all dirty pages. Note that this can return an error which will be passed back to userspace which made the close(2) system call.file->f_count reaches 0. Although defined as returning int, the return value is ignored by VFS (see fs/file_table.c:__fput()).file = fget(fd)) and down/up inode->i_sem semaphore. Ext2 filesystem currently ignores the last argument and does exactly the same for fsync(2) and fdatasync(2).file->f_flags & FASYNC changes.posix_lock_file()), if it succeeds but the standard POSIX lock code fails then it will never be unlocked on fs-dependent level..Under Linux, information about mounted filesystems is kept in two separate structures - super_block and vfsmount. The reason for this is that Linux allows to mount the same filesystem (block device) under multiple mount points, which means that the same super_block can correspond to multiple vfsmount structures.
Let us look at struct super_block first, declared in include/linux/fs.h:
struct super_block { struct list_head s_list; /* Keep this first */ kdev_t s_dev; unsigned long s_blocksize; unsigned char s_blocksize_bits; unsigned char s_lock; unsigned char s_dirt; struct file_system_type *s_type; struct super_operations *s_op; struct dquot_operations *dq_op; unsigned long s_flags; unsigned long s_magic; struct dentry *s_root; wait_queue_head_t s_wait; struct list_head s_dirty; /* dirty inodes */ struct list_head s_files; struct block_device *s_bdev; struct list_head s_mounts; /* vfsmount(s) of this one */ struct quota_mount_options s_dquot; /* Diskquota specific options */ union { struct minix_sb_info minix_sb; struct ext2_sb_info ext2_sb; ..... all filesystems that need sb-private info ... void *generic_sbp; } u; /* * The next field is for VFS *only*. No filesystems have any business * even looking at it. You had been warned. */ struct semaphore s_vfs_rename_sem; /* Kludge */ /* The next field is used by knfsd when converting a (inode number based) * file handle into a dentry. As it builds a path in the dcache tree from * the bottom up, there may for a time be a subpath of dentrys which is not * connected to the main tree. This semaphore ensure that there is only ever * one such free path per filesystem. Note that unconnected files (or other * non-directories) are allowed, but not unconnected diretories. */ struct semaphore s_nfsd_free_path_sem; };
The various fields in the super_block structure are:
FS_REQUIRES_DEV filesystems, this is the i_dev of the block device. For others (called anonymous filesystems) this is an integer MKDEV(UNNAMED_MAJOR, i) where i is the first unset bit in unnamed_dev_in_use array, between 1 and 255 inclusive. See fs/super.c:get_unnamed_dev()/put_unnamed_dev(). It has been suggested many times that anonymous filesystems should not use s_dev field.lock_super()/unlock_super().struct file_system_type of the corresponding filesystem. Filesystem's read_super() method doesn't need to set it as VFS fs/super.c:read_super() sets it for you if fs-specific read_super() succeeds and resets to NULL if it fails.super_operations structure which contains fs-specific methods to read/write inodes etc. It is the job of filesystem's read_super() method to initialise s_op correctly.read_super() to read the root inode from the disk and pass it to d_alloc_root() to allocate the dentry and instantiate it. Some filesystems spell "root" other than "/" and so use more generic d_alloc() function to bind the dentry to a name, e.g. pipefs mounts itself on "pipe:" as its own root instead of "/".inode->i_state & I_DIRTY) then it is on superblock-specific dirty list linked via inode->i_list.fs/file_table.c:fs_may_remount_ro() which goes through sb->s_files list and denies remounting if there are files opened for write (file->f_mode & FMODE_WRITE) or files with pending unlink (inode->i_nlink == 0).FS_REQUIRES_DEV, this points to the block_device structure describing the device the filesystem is mounted on.vfsmount structures, one for each mounted instance of this superblock.The superblock operations are described in the super_operations structure declared in include/linux/fs.h:
struct super_operations { void (*read_inode) (struct inode *); void (*write_inode) (struct inode *, int); void (*put_inode) (struct inode *); void (*delete_inode) (struct inode *); void (*put_super) (struct super_block *); void (*write_super) (struct super_block *); int (*statfs) (struct super_block *, struct statfs *); int (*remount_fs) (struct super_block *, int *, char *); void (*clear_inode) (struct inode *); void (*umount_begin) (struct super_block *); };
fs/inode.c:get_new_inode() from iget4() (and therefore iget()). If a filesystem wants to use iget() then read_inode() must be implemented - otherwise get_new_inode() will panic. While inode is being read it is locked (inode->i_state = I_LOCK). When the function returns, all waiters on inode->i_wait are woken up. The job of the filesystem's read_inode() method is to locate the disk block which contains the inode to be read and use buffer cache bread() function to read it in and initialise the various fields of inode structure, for example the inode->i_op and inode->i_fop so that VFS level knows what operations can be performed on the inode or corresponding file. Filesystems that don't implement read_inode() are ramfs and pipefs. For example, ramfs has its own inode-generating function ramfs_get_inode() with all the inode operations calling it as needed.read_inode() in that it needs to locate the relevant block on disk and interact with buffer cache by calling mark_buffer_dirty(bh). This method is called on dirty inodes (those marked dirty with mark_inode_dirty()) when the inode needs to be sync'd either individually or as part of syncing the entire filesystem.inode->i_count and inode->i_nlink reach 0. Filesystem deletes the on-disk copy of the inode and calls clear_inode() on VFS inode to "terminate it with extreme prejudice".brelse() the block containing the superblock and kfree() any bitmaps allocated for free blocks, inodes, etc.sb-private area) and mark_buffer_dirty(bh) . It should also clear sb->s_dirt flag.struct statfs passed as argument is a kernel pointer, not a user pointer so we don't need to do any I/O to userspace. If not implemented then statfs(2) will fail with ENOSYS.clear_inode(). Filesystems that attach private data to inode structure (via generic_ip field) must free it here.So, let us look at what happens when we mount a on-disk (FS_REQUIRES_DEV) filesystem. The implementation of the mount(2) system call is in fs/super.c:sys_mount() which is the just a wrapper that copies the options, filesystem type and device name for the do_mount() function which does the real work:
do_mount() calling get_fs_type() and once by get_sb_dev() calling get_filesystem() if read_super() was successful. The first increment is to prevent module unloading while we are inside read_super() method and the second increment is to indicate that the module is in use by this mounted instance. Obviously, do_mount() decrements the count before returning, so overall the count only grows by 1 after each mount.fs_type->fs_flags & FS_REQUIRES_DEV is true, the superblock is initialised by a call to get_sb_bdev() which obtains the reference to the block device and interacts with the filesystem's read_super() method to fill in the superblock. If all goes well, the super_block structure is initialised and we have an extra reference to the filesystem's module and a reference to the underlying block device.vfsmount structure is allocated and linked to sb->s_mounts list and to the global vfsmntlist list. The vfsmount field mnt_instances allows to find all instances mounted on the same superblock as this one. The mnt_list field allows to find all instances for all superblocks system-wide. The mnt_sb field points to this superblock and mnt_root has a new reference to the sb->s_root dentry.As a simple example of Linux filesystem that does not require a block device for mounting, let us consider pipefs from fs/pipe.c. The filesystem's preamble is rather straightforward and requires little explanation:
static DECLARE_FSTYPE(pipe_fs_type, "pipefs", pipefs_read_super, FS_NOMOUNT|FS_SINGLE); static int __init init_pipe_fs(void) { int err = register_filesystem(&pipe_fs_type); if (!err) { pipe_mnt = kern_mount(&pipe_fs_type); err = PTR_ERR(pipe_mnt); if (!IS_ERR(pipe_mnt)) err = 0; } return err; } static void __exit exit_pipe_fs(void) { unregister_filesystem(&pipe_fs_type); kern_umount(pipe_mnt); } module_init(init_pipe_fs) module_exit(exit_pipe_fs)
The filesystem is of type FS_NOMOUNT|FS_SINGLE, which means it cannot be mounted from userspace and can only have one superblock system-wide. The FS_SINGLE file also means that it must be mounted via kern_mount() after it is successfully registered via register_filesystem(), which is exactly what happens in init_pipe_fs(). The only bug in this function is that if kern_mount() fails (e.g. because kmalloc() failed in add_vfsmnt()) then the filesystem is left as registered but module initialisation fails. This will cause cat /proc/filesystems to Oops. (have just sent a patch to Linus mentioning that although this is not a real bug today as pipefs can't be compiled as a module, it should be written with the view that in the future it may become modularised).
The result of register_filesystem() is that pipe_fs_type is linked into the file_systems list so one can read /proc/filesystems and find "pipefs" entry in there with "nodev" flag indicating that FS_REQUIRES_DEV was not set. The /proc/filesystems file should really be enhanced to support all the new FS_ flags (and I made a patch to do so) but it cannot be done because it will break all the user applications that use it. Despite Linux kernel interfaces changing every minute (only for the better) when it comes to the userspace compatibility, Linux is a very conservative operating system which allows many applications to be used for a long time without being recompiled.
The result of kern_mount() is that:
unnamed_dev_in_use bitmap; if there are no more bits then kern_mount() fails with EMFILE.get_empty_super(). The get_empty_super() function walks the list of superblocks headed by super_block and looks for empty entry, i.e. s->s_dev == 0. If no such empty superblock is found then a new one is allocated using kmalloc() at GFP_USER priority. The maximum system-wide number of superblocks is checked in get_empty_super() so if it starts failing, one can adjust the tunable /proc/sys/fs/super-max.pipe_fs_type->read_super() method, i.e. pipefs_read_super(), is invoked which allocates root inode and root dentry sb->s_root, and sets sb->s_op to be &pipefs_ops.kern_mount() calls add_vfsmnt(NULL, sb->s_root, "none") which allocates a new vfsmount structure and links it into vfsmntlist and sb->s_mounts.pipe_fs_type->kern_mnt is set to this new vfsmount structure and it is returned. The reason why the return value of kern_mount() is a vfsmount structure is because even FS_SINGLE filesystems can be mounted multiple times and so their mnt->mnt_sb will point to the same thing which would be silly to return from multiple calls to kern_mount().Now that the filesystem is registered and inkernel-mounted we can use it. The entry point into the pipefs filesystem is the pipe(2) system call, implemented in arch-dependent function sys_pipe() but the real work is done by a portable fs/pipe.c:do_pipe() function. Let us look at do_pipe() then. The interaction with pipefs happens when do_pipe() calls get_pipe_inode() to allocate a new pipefs inode. For this inode, inode->i_sb is set to pipefs' superblock pipe_mnt->mnt_sb, the file operations i_fop is set to rdwr_pipe_fops and the number of readers and writers (held in inode->i_pipe) is set to 1. The reason why there is a separate inode field i_pipe instead of keeping it in the fs-private union is that pipes and FIFOs share the same code and FIFOs can exist on other filesystems which use the other access paths within the same union which is very bad C and can work only by pure luck. So, yes, 2.2.x kernels work only by pure luck and will stop working as soon as you slightly rearrange the fields in the inode.
Each pipe(2) system call increments a reference count on the pipe_mnt mount instance.
Under Linux, pipes are not symmetric (bidirection or STREAM pipes), i.e. two sides of the file have different file->f_op operations - the read_pipe_fops and write_pipe_fops respectively. The write on read side returns EBADF and so does read on write side.
As a simple example of ondisk Linux filesystem, let us consider BFS. The preamble of the BFS module is in fs/bfs/inode.c:
static DECLARE_FSTYPE_DEV(bfs_fs_type, "bfs", bfs_read_super); static int __init init_bfs_fs(void) { return register_filesystem(&bfs_fs_type); } static void __exit exit_bfs_fs(void) { unregister_filesystem(&bfs_fs_type); } module_init(init_bfs_fs) module_exit(exit_bfs_fs)
A special fstype declaration macro DECLARE_FSTYPE_DEV() is used which sets the fs_type->flags to FS_REQUIRES_DEV to signify that BFS requires a real block device to be mounted on.
The module's initialisation function registers the filesystem with VFS and the cleanup function (only present when BFS is configured to be a module) unregisters it.
With the filesystem registered, we can proceed to mount it, which would invoke out fs_type->read_super() method which is implemented in fs/bfs/inode.c:bfs_read_super(). It does the following:
set_blocksize(s->s_dev, BFS_BSIZE): since we are about to interact with the block device layer via the buffer cache, we must initialise a few things, namely set the block size and also inform VFS via fields s->s_blocksize and s->s_blocksize_bits.bh = bread(dev, 0, BFS_BSIZE): we read block 0 of the device passed via s->s_dev. This block is the filesystem's superblock.BFS_MAGIC number and, if valid, stored in the sb-private field s->su_sbh (which is really s->u.bfs_sb.si_sbh).kmalloc(GFP_KERNEL) and clear all bits to 0 except the first two which we set to 1 to indicate that we should never allocate inodes 0 and 1. Inode 2 is root and the corresponding bit will be set to 1 a few lines later anyway - the filesystem should have a valid root inode at mounting time!s->s_op, which means that we can from this point invoke inode cache via iget() which results in s_op->read_inode() to be invoked. This finds the block that contains the specified (by inode->i_ino and inode->i_dev) inode and reads it in. If we fail to get root inode then we free the inode bitmap and release superblock buffer back to buffer cache and return NULL. If root inode was read OK, then we allocate a dentry with name / (as becometh root) and instantiate it with this inode.iput() - we don't hold a reference to it longer than needed.s->s_dirt flag (TODO: why do I do this? Originally, I did it because minix_read_super() did but neither minix nor BFS seem to modify superblock in the read_super()).fs/super.c:read_super().After the read_super() function returns successfully, VFS obtains the reference to the filesystem module via call to get_filesystem(fs_type) in fs/super.c:get_sb_bdev() and a reference to the block device.
Now, let us examine what happens when we do I/O on the filesystem. We already examined how inodes are read when iget() is called and how they are released on iput(). Reading inodes sets up, among other things, inode->i_op and inode->i_fop; opening a file will propagate inode->i_fop into file->f_op.
Let us examine the code path of the link(2) system call. The implementation of the system call is in fs/namei.c:sys_link():
getname() function which does the error checking.path_init()/path_walk() interaction with dcache. The result is stored in old_nd and nd structures.old_nd.mnt != nd.mnt then "cross-device link" EXDEV is returned - one cannot link between filesystems, in Linux this translates into - one cannot link between mounted instances of a filesystem (or, in particular between filesystems).nd by lookup_create() .vfs_link() function is called which checks if we can create a new entry in the directory and invokes the dir->i_op->link() method which brings us back to filesystem-specific fs/bfs/dir.c:bfs_link() function.bfs_link(), we check if we are trying to link a directory and if so, refuse with EPERM error. This is the same behaviour as standard (ext2).bfs_add_entry() which goes through all entries looking for unused slot (de->ino == 0) and, when found, writes out the name/inode pair into the corresponding block and marks it dirty (at non-superblock priority).inode->i_nlink, update inode->i_ctime and mark this inode dirty as well as instantiating the new dentry with the inode.Other related inode operations like unlink()/rename() etc work in a similar way, so not much is gained by examining them all in details.
Linux supports loading user application binaries from disk. More interestingly, the binaries can be stored in different formats and the operating system's response to programs via system calls can deviate from norm (norm being the Linux behaviour) as required, in order to emulate formats found in other flavours of UNIX (COFF, etc) and also to emulate system calls behaviour of other flavours (Solaris, UnixWare, etc). This is what execution domains and binary formats are for.
Each Linux task has a personality stored in its task_struct (p->personality). The currently existing (either in the official kernel or as addon patch) personalities include support for FreeBSD, Solaris, UnixWare, OpenServer and many other popular operating systems. The value of current->personality is split into two parts:
STICKY_TIMEOUTS, WHOLE_SECONDS, etc.By changing the personality, we can change the way the operating system treats certain system calls, for example adding a STICKY_TIMEOUT to current->personality makes select(2) system call preserve the value of last argument (timeout) instead of storing the unslept time. Some buggy programs rely on buggy operating systems (non-Linux) and so Linux provides a way to emulate bugs in cases where the source code is not available and so bugs cannot be fixed.
Execution domain is a contiguous range of personalities implemented by a single module. Usually a single execution domain implements a single personality but sometimes it is possible to implement "close" personalities in a single module without too many conditionals.
Execution domains are implemented in kernel/exec_domain.c and were completely rewritten for 2.4 kernel, compared with 2.2.x. The list of execution domains currently supported by the kernel, along with the range of personalities they support, is available by reading the /proc/execdomains file. Execution domains, except the PER_LINUX one, can be implemented as dynamically loadable modules.
The user interface is via personality(2) system call, which sets the current process' personality or returns the value of current->personality if the argument is set to impossible personality 0xffffffff. Obviously, the behaviour of this system call itself does not depend on personality..
The kernel interface to execution domains registration consists of two functions:
int register_exec_domain(struct exec_domain *): registers the execution domain by linking it into single-linked list exec_domains under the write protection of the read-write spinlock exec_domains_lock. Returns 0 on success, non-zero on failure.int unregister_exec_domain(struct exec_domain *): unregisters the execution domain by unlinking it from the exec_domains list, again using exec_domains_lock spinlock in write mode. Returns 0 on success.The reason why exec_domains_lock is a read-write is that only registration and unregistration requests modify the list, whilst doing cat /proc/filesystems calls fs/exec_domain.c:get_exec_domain_list(), which needs only read access to the list. Registering a new execution domain defines a "lcall7 handler" and a signal number conversion map. Actually, ABI patch extends this concept of exec domain to include extra information (like socket options, socket types, address family and errno maps).
The binary formats are implemented in a similar manner, i.e. a single-linked list formats is defined in fs/exec.c and is protected by a read-write lock binfmt_lock. As with exec_domains_lock, the binfmt_lock is taken read on most occasions except for registration/unregistration of binary formats. Registering a new binary format enhances the execve(2) system call with new load_binary()/load_shlib() functions as well as ability to core_dump() . The load_shlib() method is used only by the old uselib(2) system call while the load_binary() method is called by the search_binary_handler() from do_execve() which implements execve(2) system call.
The personality of the process is determined at binary format loading by the corresponding format's load_binary() method using some heuristics. For example to determine UnixWare7 binaries one first marks the binary using the elfmark(1) utility, which sets the ELF header's e_flags to the magic value 0x314B4455 which is detected at ELF loading time and current->personality is set to PER_UW7. If this heuristic fails, then a more generic one, such as treat ELF interpreter paths like /usr/lib/ld.so.1 or /usr/lib/libc.so.1 to indicate a SVR4 binary, is used and personality is set to PER_SVR4. One could write a little utility program that uses Linux's ptrace(2) capabilities to single-step the code and force a running program into any personality.
Once personality (and therefore current->exec_domain) is known, the system calls are handled as follows. Let us assume that a process makes a system call by means of lcall7 gate instruction. This transfers control to ENTRY(lcall7) of arch/i386/kernel/entry.S because it was prepared in arch/i386/kernel/traps.c:trap_init(). After appropriate stack layout conversion, entry.S:lcall7 obtains the pointer to exec_domain from current and then an offset of lcall7 handler within the exec_domain (which is hardcoded as 4 in asm code so you can't shift the handler field around in C declaration of struct exec_domain) and jumps to it. So, in C, it would look like this:
static void UW7_lcall7(int segment, struct pt_regs * regs) { abi_dispatch(regs, &uw7_funcs[regs->eax & 0xff], 1); }
where abi_dispatch() is a wrapper around the table of function pointers that implement this personality's system calls uw7_funcs.