Kernel development

Kernel release status

The 4.2 merge window is open; see the separate article below for details.

Stable updates: 4.0.6, 3.14.45, and 3.10.81 were released on June 22.

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Quotes of the week

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4.2 Merge window part 1

Some of the more interesting user-visible changes include:

• The core file-name lookup code has been reworked to avoid using recursion. The visible effects are that stack pressure will be reduced (making things more robust when complex storage systems are in use) and the limit on the depth of nested symbolic links has been lifted.

• The performance-monitoring subsystem and associated perf tool have seen another long list of improvements; see this commit message for details.

• The jitter entropy random number generator has been merged.

• The KVM hypervisor has gained support for multiple address spaces and for system management mode, both of which are used to support secure boot in guests.

• New hardware support includes:
  • Processors and systems: ARM Ltd. Version 3 memory-management units.

  • GPIO: NXP LPC18XX/43XX GPIO controllers, Netlogic XLP GPIO controllers, Broadco BRCMSTB GPIO controllers, and Axis ETRAX FS GPIO controllers.

  • Input: Logitech M560 wireless mice and Sony motion controllers.

  • Miscellaneous: NXP LPC32xx/18xx/43xx timers, Marvell cryptographic engines, MediaTek SD/MMC card interfaces, Microchip TC74 single-input temperature sensor chips, Broadcom iProc PCIe BCMA buses, AppliedMicro X-Gene v1 PCIe MSI controllers, Cisco PCI-Express SCSI HBAs, Mikrotik RB4XX SPI controllers, Xilinx ZynqMP GQSPI controllers, Dialog Semiconductor DA9062 regulators, Qualcomm SPMI regulators, PowerNV flash controllers, STMicro LPC-based watchdogs and realtime clocks, and Broadcom STB NAND controllers.

  • Power supply: TI BQ24257 and BQ25890 battery chargers, X-Power AXP288 PMIC integrated chargers, and Richtek RT9455 battery chargers.

Changes visible to kernel developers include:

• The prototypes for the follow_link() and put_link() methods in struct inode_operations have changed to:

      const char *(*follow_link) (struct dentry *, void **);
      void (*put_link) (struct dentry *, void *);
  

  In-tree filesystems have been changed accordingly.

• The bulk of the read-copy-update (RCU) configuration options have been hidden behind an "expert" option; for the most part, RCU configuration is completely automatic now.

• Queued spinlocks are now used on the x86 architecture; they improve performance but should not bring any other visible changes. A queued reader/writer lock implementation has also been merged for x86.

• The x86 floating-point unit code has been massively rewritten. Ingo Molnar warns (in the commit message) that "these changes have a substantial regression risk", but none are known at the moment.

• Write-through mappings can now be created with ioremap_wt() or pgprot_writethrough(); this feature is currently supported on the x86 architecture.

The 4.2 merge window should remain open through July 5. If the usual schedule holds — and Linus doesn't take any ill-timed vacations this time — the final 4.2 release can be expected on August 23.

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Trying to head off kdbus

There has been a fair amount of opposition to kdbus—an effort to move much of the D-Bus interprocess communication (IPC) mechanism into the kernel—that has come up since it was first posted back in November 2014. Much of that criticism has been of the more or less unhinged variety, but there have also been a number of technical complaints about the IPC scheme. Two of the most persistent critics of the technical aspects of kdbus recently made preemptive attacks of a sort against kdbus—neither of which was particularly well-received by Greg Kroah-Hartman, who has been coordinating the kdbus effort.

In a post to the linux-kernel mailing list on June 22, Andy Lutomirski asked for Linus Torvalds's opinion on whether kdbus should be merged. He made it clear that he is not asking about timing (i.e. whether he will accept a—so far hypothetical—pull request from Kroah-Hartman in this merge window), but: "I mean whether you think that kdbus should be merged if it had appropriate review and people were okay with the implementation". He goes on to say that the uncertainty about the future of kdbus makes it harder for those who might be inclined to review the code (or to review it again, in his case) to decide whether to do that review:

He continued by listing three main reasons that he felt kdbus should not be merged. First off, he doesn't believe that it is needed to solve the problems that it claims to solve and that most of what needs to be done could be done in user space. He is also concerned that the buffering model used by kdbus does not take control groups using a memory controller into account. Lastly: "The sandbox model is, in my opinion, an experiment that isn't going to succeed."

As might be guessed, Kroah-Hartman was less than pleased by Lutomirski's message, calling it a "preemptive pull request denial". He was clearly irritated that Lutomirski had not had the courtesy to wait for his pull request—if he decided to make one—before attacking kdbus. But the technical complaints in the message went unread, as Kroah-Hartman seized on an incorrect, but offhand, comment early in Lutomirski's message to stop reading at that point.

Things took a different turn on June 23 with a request from Eric Biederman—another persistent kdbus critic—to remove the kdbus tree from linux-next. Biederman cited "significant work that was identified last merge window that has not yet been done, and who knows when it will be done" as the reason to remove it. Furthermore:

Once again, Kroah-Hartman was unamused: "No, that's not how this works. That's not how any of this works." While Biederman claimed that the project was ignoring feedback, Kroah-Hartman disagreed: "We are not ignoring _constructive_ feedback". There is "no valid reason" to remove kdbus from linux-next, he concluded.

But Biederman was not to be deterred. He maintained that linux-next is for code that is destined for the next merge window, which implies that kdbus will be offered up for merging soon (as the 4.2 merge window is open), but that the code is not even remotely ready for merging. He suggested that lots of constructive feedback has simply been ignored. In the final analysis, he is trying to stop another pull request for kdbus by having the code removed from linux-next—not a common tactic, or even one that has been used at all before.

Part of the trigger for both Lutomirski and Biederman's posts seem to have been the announcement of the sd-bus API that relies on the existing kdbus implementation (though it can fall back to using D-Bus), coupled with the most recent systemd release that depends on sd-bus. In addition, one of the kdbus maintainers, David Herrmann, recently put out a blog post that describes kdbus in terms of the AF_UNIX sockets that are familiar to many developers. That may have been seen as a prelude to another pull request for kdbus. Beyond that, the systemd release announcement encourages distributions to add kdbus to their kernels and to use it with systemd/sd-bus. That will simply lead to problems down the road, Lutomirski said:

But it is not uncommon for user-space code to be written to try out new kernel features, sometimes long before they get merged. It is a way to shake out problems with the new feature and to prove that it is useful. On the flip side, though, getting a feature into a distribution is a subtle (or not-so-subtle) way to apply pressure for the feature to be merged. Over the years, there have been a number of examples of features that were released in distributions and were merged into the kernel over fairly strenuous objections because of their widespread availability. Some Android features (e.g. binder) likely fall into that category, but things like AppArmor also probably qualify.

It is against that backdrop that Lutomirski and Biederman are making their moves. But, of course, Torvalds will have the final say and it turns out that he is inclined to merge kdbus:

Though it is clear that one of the arguments made by kdbus proponents doesn't hold any water with him: performance. The pull request Kroah-Hartman made for the 4.1 kernel ran into some serious questions about the performance gains claimed, with Torvalds himself expressing a fair amount of skepticism that speeding up D-Bus required a move into the kernel:

He hasn't changed his mind at all ("I am not AT ALL impressed by the performance argument"), but still is likely to merge kdbus at some point. Given that there have been no new postings of kdbus since the pull request from the last merge window, it would seem a little late to be asking for a new version to be pulled in for 4.2. And Kroah-Hartman has not done so; there may be work going on to address some of the outstanding comments and questions (perhaps including the performance question) so that kdbus won't be offered up for merging again until the next development cycle.

One of the key aspects of the dispute concerns the addressing of review comments. Lutomirski, Biederman, and others believe that there have been substantive constructive comments that have simply been ignored—though, with no new code posted, that may well be a premature conclusion. But "constructive", it seems, is in the eye of the beholder, so the kdbus developers may also feel that they have addressed those comments by disagreeing with them—and explaining why. It would seem to be a classic standoff situation that requires a final arbiter to make a determination. Torvalds is that arbiter, of course, and he seems to have decided—now it appears to just be a question of when, not if, kdbus is merged.

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Pathname lookup in Linux

One of the changes that was recently merged for Linux 4.2 is a substantial rewrite of parts of the pathname-lookup code in the Linux virtual filesystem (VFS) layer. This rewrite primarily affects the handling of symbolic links — though, like many such rewrites, opportunities were found to rationalize and improve other parts of the code, too. With all this change, it seems like a good opportunity to document just how pathname lookup works. While such documentation cannot stay current indefinitely, writing it immediately after a big change might increase the time until it becomes inaccurate.

The most obvious aspect of pathname lookup, which very little exploration is needed to discover, is that it is complex. There are many rules, special cases, and implementation alternatives that all combine to confuse the unwary reader. Computer science has long been acquainted with such complexity and has tools to help manage it. One tool that we will make extensive use of is "divide and conquer". For the early parts of the analysis we will divide off symlinks, leaving them until the final part. Well before we get to symlinks we have another major division based on the VFS's approach to locking which will allow us to review "REF-walk" and "RCU-walk" (a pair of algorithms that have been described previously) separately. But we are getting ahead of ourselves. There are some important low-level distinctions we need to clarify first.

There are two sorts of ...

Pathnames (sometimes "file names"), used to identify objects in the filesystem, will be familiar to most readers. They contain two sorts of elements: "slashes" that are sequences of one or more "/" characters, and "components" that are sequences of one or more non-"/" characters. These form two kinds of paths. Those that start with slashes are "absolute" and start from the filesystem root. The others are "relative" and start from the current directory, or from some other location specified by a file descriptor given to a "xxxat()" system call such as openat().

It is tempting to describe the second kind as starting with a component, but that isn't always accurate: a pathname can lack both slashes and components; it can be empty, in other words. This is generally forbidden in POSIX, but some of those "xxxat()" system calls in Linux permit it when the AT_EMPTY_PATH flag is given. For example, if you have an open file descriptor on an executable file, you can execute it by calling execveat(), passing the file descriptor, an empty path, and the AT_EMPTY_PATH flag.

These paths can be divided into two sections: the final component and everything else. The "everything else" is the easy bit. In all cases it must identify a directory that already exists, otherwise an error such as ENOENT or ENOTDIR will be reported. The final component is not so simple. Not only do different system calls interpret it quite differently (e.g. some create it, some do not), but it might not even exist: neither the empty pathname nor the pathname that is just slashes have a final component. If it does exist, it could be "." or ".." which are handled quite differently from other components.

If a pathname ends with a slash, such as "/tmp/foo/" it might be tempting to consider that to have an empty final component. In many ways that would lead to correct results, but not always. In particular, mkdir() and rmdir() each create or remove a directory named by the final component, and they are required to work with pathnames ending in "/". According to POSIX:

The Linux pathname-walking code (mostly in fs/namei.c) deals with all of these issues: breaking the path into components, handling the "everything else" quite separately from the final component, and checking that the trailing slash is not used where it isn't permitted. It also addresses the important issue of concurrent access.

While one process is looking up a pathname, another might be making changes that affect that lookup. One fairly extreme case is that if a/b were renamed to a/c/b while another process were looking up a/b/.., that process might successfully resolve on a/c. Most races are much more subtle, and a big part of the task of pathname lookup is to prevent them from having damaging effects. Many of the possible races are seen most clearly in the context of the "dcache" and an understanding of that is central to understanding pathname lookup.

More than just a cache

The dcache caches information about names in each filesystem to make them quickly available for lookup. Each entry (known as a "dentry") contains three significant fields: a component name, a pointer to a parent dentry, and a pointer to the "inode" which contains further information about the object in that parent with the given name. The inode pointer can be NULL indicating that the name doesn't exist in the parent. While there can be linkage in the dentry of a directory to the dentries of the children, that linkage is not used for pathname lookup, and so will not be considered here.

The dcache has a number of uses apart from accelerating lookup. One that will be particularly relevant is that it is closely integrated with the mount table that records which filesystem is mounted where. What the mount table actually stores is which dentry is mounted on top of which other dentry.

When considering the dcache, we have another of our "two types" distinctions: there are two types of filesystems. Some filesystems ensure that the information in the dcache is always completely accurate (though not necessarily complete). This can allow the VFS to determine if a particular file does or doesn't exist without checking with the filesystem, and means that the VFS can protect the filesystem against certain races and other problems. These are typically "local" filesystems such as ext3, XFS, and Btrfs.

Other filesystems don't provide that guarantee because they cannot. These are typically filesystems that are shared across a network, whether remote filesystems like NFS and 9P, or cluster filesystems like ocfs2 or cephfs. These filesystems allow the VFS to revalidate cached information and must provide their own protection against awkward races. The VFS can detect these filesystems by the DCACHE_OP_REVALIDATE flag being set in the dentry.

REF-walk: simple concurrency management with refcounts and spinlocks

With all of those divisions carefully classified, we can now start looking at the actual process of walking along a path. In particular we will start with the handling of the "everything else" part of a pathname, and focus on the "REF-walk" approach to concurrency management. This code is found in the link_path_walk() function, if you ignore all the places that only run when "LOOKUP_RCU" (indicating the use of RCU-walk) is set.

REF-walk is fairly heavy-handed with locks and reference counts. Not as heavy-handed as in the old "big kernel lock" days, but certainly not afraid of taking a lock when one is needed. It uses a variety of different concurrency controls. A background understanding of the various primitives is assumed, or can be gleaned from elsewhere such as in Meet the Lockers.

The locking mechanisms used by REF-walk include:

dentry->d_lockref

This uses the relatively recently introduced lockref primitive to provide both a spinlock and a reference count. The special sauce of this primitive is that the conceptual sequence "lock; inc_ref; unlock;" can often be performed with a single, atomic memory operation.

Holding a reference on a dentry ensures that the dentry won't suddenly be freed and used for something else, so the values in various fields will behave as expected. It also protects the ->d_inode reference to the inode to some extent.

The association between a dentry and its inode is fairly permanent. For example, when a file is renamed, the dentry and inode move together to the new location. When a file is created the dentry will initially be negative (i.e. d_inode is NULL), and will be assigned to the new inode as part of the act of creation.

When a file is deleted, this can be reflected in the cache either by setting d_inode to NULL, or by removing it from the hash table (described shortly) used to look up the name in the parent directory. If the dentry is still in use, the second option is used, as it is perfectly legal to keep using an open file after it has been deleted; having the dentry around helps. If the dentry is not otherwise in use (i.e. if the refcount in d_lockref is one), only then will d_inode be set to NULL. Doing it this way is more efficient for a very common case.

So as long as a counted reference is held to a dentry, a non-NULL ->d_inode value will never be changed.

dentry->d_lock

d_lock is a synonym for the spinlock that is part of d_lockref above. For our purposes, holding this lock protects against the dentry being renamed or unlinked. In particular, its parent (d_parent), and its name (d_name) cannot be changed, and it cannot be removed from the dentry hash table.

When looking for a name in a directory, REF-walk takes d_lock on each candidate dentry that it finds in the hash table and then checks that the parent and name are correct. So it doesn't lock the parent while searching in the cache; it only locks children.

When looking for the parent for a given name (to handle ".."), REF-walk can take d_lock to get a stable reference to d_parent, but it first tries a more lightweight approach. As seen in dget_parent(), if a reference can be claimed on the parent, and if subsequently d_parent can be seen to have not changed, then there is no need to actually take the lock on the child.

rename_lock

Looking up a given name in a given directory involves computing a hash from the two values (the name and the dentry of the directory), accessing that slot in a hash table, and searching the linked list that is found there. When a dentry is renamed, the name and the parent dentry can both change so the hash will almost certainly change too. This would move the dentry to a different chain in the hash table. If a filename search happened to be looking at a dentry that was moved in this way, it might end up continuing the search down the wrong chain, and so miss out on part of the correct chain.

The name-lookup process (d_lookup()) does not try to prevent this from happening, but only to detect when it happens. rename_lock is a seqlock that is updated whenever any dentry is renamed. If d_lookup finds that a rename happened while it unsuccessfully scanned a chain in the hash table, it simply tries again.

inode->i_mutex

i_mutex is a mutex that serializes all changes to a particular directory. This ensures that, for example, an unlink() and a rename() cannot both happen at the same time. It also keeps the directory stable while the filesystem is asked to look up a name that is not currently in the dcache.

This has a complementary role to that of d_lock: i_mutex on a directory protects all of the names in that directory, while d_lock on a name protects just one name in a directory. Most changes to the dcache hold i_mutex on the relevant directory inode and briefly take d_lock on one or more the dentries while the change happens. One exception is when idle dentries are removed from the dcache due to memory pressure. This uses d_lock, but i_mutex plays no role.

The mutex affects pathname lookup in two distinct ways. Firstly, it serializes lookup of a name in a directory. walk_component() uses lookup_fast() first which, in turn, checks to see if the name is in the cache, using only d_lock locking. If the name isn't found, then walk_component() falls back to lookup_slow(), which takes i_mutex, checks again that the name isn't in the cache, and then calls in to the filesystem to get a definitive answer. A new dentry will be added to the cache regardless of the result.

Secondly, when pathname lookup reaches the final component, it will sometimes need to take i_mutex before performing the last lookup so that the required exclusion can be achieved. How path lookup chooses to take, or not take, i_mutex is one of the issues addressed in a subsequent section.

mnt->mnt_count

mnt_count is a per-CPU reference counter on "mount" structures. Per-CPU here means that incrementing the count is cheap as it only uses CPU-local memory, but checking if the count is zero is expensive as it needs to check with every CPU. Taking a mnt_count reference prevents the mount structure from disappearing as the result of regular unmount operations, but does not prevent a "lazy" unmount. So holding mnt_count doesn't ensure that the mount remains in the namespace and, in particular, doesn't stabilize the link to the mounted-on dentry. It does, however, ensure that the mount data structure remains coherent, and it provides a reference to the root dentry of the mounted filesystem. So a reference through ->mnt_count provides a stable reference to the mounted dentry, but not the mounted-on dentry.

mount_lock

mount_lock is a global seqlock, a bit like rename_lock. It can be used to check if any change has been made to any mount points.

While walking down the tree (away from the root), this lock is used when crossing a mount point to check that the crossing was safe. That is, the value in the seqlock is read, then the code finds the mount that is mounted on the current directory, if there is one, and increments the mnt_count. Finally, the value in mount_lock is checked against the old value. If there is no change, then the crossing was safe. If there was a change, the mnt_count is decremented and the whole process is retried.

When walking up the tree (towards the root) by following a ".." link, a little more care is needed. In this case the seqlock (which contains both a counter and a spinlock) is fully locked to prevent any changes to any mount points while stepping up. This locking is needed to stabilize the link to the mounted-on dentry, which the refcount on the mount itself doesn't ensure.

RCU

Finally the global (but extremely lightweight) RCU read lock is held from time to time to ensure certain data structures don't get freed unexpectedly. In particular, it is held while scanning chains in the dcache hash table, and the mount point hash table.

Bringing it together with struct nameidata

Throughout the process of walking a path, the current status is stored in a struct nameidata, "namei" being the traditional name — dating all the way back to First Edition Unix — of the function that converts a "name" to an "inode". struct nameidata contains (among other fields):

struct path path

A path contains a struct vfsmount (which is embedded in a struct mount) and a struct dentry. Together these record the current status of the walk. They start out referring to the starting point (the current working directory, the root directory, or some other directory identified by a file descriptor) and are updated on each step. A reference through d_lockref and mnt_count is always held.

struct qstr last

This is a string together with a length (i.e. not nul terminated) that is the "next" component in the pathname.

int last_type

This is one of LAST_NORM, LAST_ROOT, LAST_DOT, LAST_DOTDOT, or LAST_BIND. The last field is only valid if the type is LAST_NORM. LAST_BIND is used when following a symlink and no components of the symlink have been processed yet. Others should be fairly self-explanatory.

struct path root

This is used to hold a reference to the effective root of the filesystem. Often that reference won't be needed, so this field is only assigned the first time it is used, or when a non-standard root is requested. Keeping a reference in the nameidata ensures that only one root is in effect for the entire path walk, even if it races with a chroot() system call.

The root is needed when either of two conditions holds: (1) either the pathname or a symbolic link starts with a "/", or (2) a ".." component is being handled, since ".." from the root must always stay at the root. The value used is usually the current root directory of the calling process. An alternate root can be provided as when sysctl() calls file_open_root(), and when NFSv4 or Btrfs call mount_subtree(). In each case, a pathname is being looked up in a very specific part of the filesystem, and the lookup must not be allowed to escape that subtree. It works a bit like a local chroot().

Ignoring the handling of symbolic links, we can now describe the link_path_walk() function, which handles the lookup of everything except the final component as:

Given a path (name) and a nameidata structure (nd), check that the current directory has execute permission and then advance name over one component while updating last_type and last. If that was the final component, then return, otherwise call walk_component() and repeat from the top.

walk_component() is even easier. If the component is LAST_DOTS, it calls handle_dots(), which does the necessary locking as already described. If it finds a LAST_NORM component, it first calls lookup_fast(), which only looks in the dcache, but will ask the filesystem to revalidate the result if it is that sort of filesystem. If that doesn't get a good result, it calls lookup_slow(), which takes the i_mutex, rechecks the cache, and then asks the filesystem to find a definitive answer. Each of these will call follow_managed() (as described below) to handle any mount points.

In the absence of symbolic links, walk_component() creates a new struct path containing a counted reference to the new dentry and a reference to the new vfsmount, which is only counted if it is different from the previous vfsmount. It then calls path_to_nameidata() to install the new struct path in the struct nameidate and drop the unneeded references.

This "hand-over-hand" sequencing of getting a reference to the new dentry before dropping the reference to the previous dentry may seem obvious, but is worth pointing out so that we will recognize its analogue in the "RCU-walk" version.

Handling the final component

link_path_walk() only walks as far as setting nd->last and nd->last_type to refer to the final component of the path. It does not call walk_component() that last time. Handling that final component remains for the caller to sort out. Those callers are path_lookupat(), path_parentat(), path_mountpoint(), and path_openat(), each of which handles the differing requirements of different system calls.

path_parentat() is clearly the simplest; it just wraps a little bit of housekeeping around link_path_walk() and returns the parent directory and final component to the caller. The caller will be either aiming to create a name (via filename_create()) or remove or rename a name (in which case user_path_parent() is used). They will use i_mutex to exclude other changes while they validate and then perform their operation.

path_lookupat() is nearly as simple; it is used when an existing object is wanted such as by stat() or chmod(). It essentially just calls walk_component() on the final component through a call to lookup_last(). path_lookupat() returns just the final dentry.

path_mountpoint() handles the special case of unmounting, which must not try to revalidate the mounted filesystem. It effectively contains, through a call to mountpoint_last(), an alternate implementation of lookup_slow(), which skips that step. This is important when unmounting a filesystem that is inaccessible, such as one provided by a dead NFS server.

Finally path_openat() is used for the open() system call; it contains, in support functions starting with do_last(), all the complexity needed to handle the different subtleties of O_CREAT (with or without O_EXCL), final "/" characters, and trailing symbolic links. We will revisit this in the final part of this series, which focuses on those symbolic links. do_last() will sometimes, but not always, take i_mutex, depending on what it finds.

Each of these, or the functions which call them, need to be alert to the possibility that the final component is not LAST_NORM. If the goal of the lookup is to create something, then any value for last_type other than LAST_NORM will result in an error. For example if path_parentat() reports LAST_DOTDOT, then the caller won't try to create that name. They also check for trailing slashes by testing last.name[last.len]. If there is any character beyond the final component, it must be a trailing slash.

Revalidation and automounts

Apart from symbolic links, there are only two parts of the "REF-walk" process not yet covered. One is the handling of stale cache entries and the other is automounts.

On filesystems that require it, the lookup routines will call the ->d_revalidate() dentry method to ensure that the cached information is current. This will often confirm validity or update a few details from a server. In some cases it may find that there has been change further up the path and that something that was thought to be valid previously isn't really. When this happens the lookup of the whole path is aborted and retried with the LOOKUP_REVAL flag set. This forces revalidation to be more thorough. We will see more details of this retry process in the next article.

Automount points are locations in the filesystem where an attempt to lookup a name can trigger changes to how that lookup should be handled, in particular by mounting a filesystem there. These are covered in greater detail in autofs4.txt in the Linux documentation tree, but a few notes specifically related to path lookup are in order here.

The Linux VFS has a concept of "managed" dentries which is reflected in function names such as follow_managed(). There are three potentially interesting things about these dentries corresponding to three different flags that might be set in dentry->d_flags:

DCACHE_MANAGE_TRANSIT

If this flag has been set, then the filesystem has requested that the d_manage() dentry operation be called before handling any possible mount point. This can perform two particular services:

1. It can block to avoid races. If an automount point is being unmounted, the d_manage() function will usually wait for that process to complete before letting the new lookup proceed and possibly trigger a new automount.

2. It can selectively allow only some processes to transit through a mount point. When a server process is managing automounts, it may need to access a directory without triggering normal automount processing. That server process can identify itself to the autofs filesystem, which will then give it a special pass through d_manage() by returning -EISDIR.

DCACHE_MOUNTED

This flag is set on every dentry that is mounted on. As Linux supports multiple filesystem namespaces, it is possible that the dentry may not be mounted on in this namespace, just in some other. So this flag is seen as a hint, not a promise. If this flag is set, and d_manage() didn't return -EISDIR, lookup_mnt() is called to examine the mount hash table (honoring the mount_lock described earlier) and possibly return a new vfsmount and a new dentry (both with counted references).

DCACHE_NEED_AUTOMOUNT

If d_manage() allowed us to get this far, and lookup_mnt() didn't find a mount point, then this flag causes the d_automount() dentry operation to be called.

The d_automount() operation can be arbitrarily complex and may communicate with server processes etc., but it should ultimately either report that there was an error, that there was nothing to mount, or should provide an updated struct path with new dentry and vfsmount. In the latter case, finish_automount() will be called to safely install the new mount point into the mount table.

There is no new locking of import here and it is important that no locks (only counted references) are held over this processing due to the very real possibility of extended delays. This will become more important next time when we examine RCU-walk which is particularly sensitive to delays.

To be continued

With a dcache and a mount table, slashes at the start and the end, final components handled multiple ways and with a variety of locking primitives; we have already examined, and hopefully understood, a fair degree of complexity. There is still more to come though. Next week we will build on this to explore the very different challenges that arise when we try to perform pathname lookup without taking any locks at all.

Finally, I'd like to make a note of thanks to Al Viro for reviewing an early draft of this article and highlighting a number of errors.

Comments (7 posted)