===============
  Pathname lookup
  ===============
  
  This write-up is based on three articles published at lwn.net:
  
  - <https://lwn.net/Articles/649115/> Pathname lookup in Linux
  - <https://lwn.net/Articles/649729/> RCU-walk: faster pathname lookup in Linux
  - <https://lwn.net/Articles/650786/> A walk among the symlinks
  
  Written by Neil Brown with help from Al Viro and Jon Corbet.
  It has subsequently been updated to reflect changes in the kernel
  including:
  
  - per-directory parallel name lookup.

  - ``openat2()`` resolution restriction flags.

  Introduction to pathname lookup
  ===============================

  
  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" 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 ...
  --------------------------

  .. _openat: http://man7.org/linux/man-pages/man2/openat.2.html

  
  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
  "``*at()``" system calls such as `openat() <openat_>`_.

  .. _execveat: http://man7.org/linux/man-pages/man2/execveat.2.html

  
  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 "``*at()``" 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() <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.

  .. _POSIX: https://pubs.opengroup.org/onlinepubs/9699919799/basedefs/V1_chap04.html#tag_04_12

  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_:

    A pathname that contains at least one non-<slash> character and
    that ends with one or more trailing <slash> characters shall not

    be resolved successfully unless the last pathname component before
    the trailing <slash> characters names an existing directory or a
    directory entry that is to be created for a directory immediately
    after the pathname is resolved.

  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.

  .. _Meet the Lockers: https://lwn.net/Articles/453685/

  
  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 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
  and 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.

  ``rename_lock`` is also used to detect and defend against potential attacks
  against ``LOOKUP_BENEATH`` and ``LOOKUP_IN_ROOT`` when resolving ".." (where
  the parent directory is moved outside the root, bypassing the ``path_equal()``
  check). If ``rename_lock`` is updated during the lookup and the path encounters
  a "..", a potential attack occurred and ``handle_dots()`` will bail out with
  ``-EAGAIN``.

  inode->i_rwsem
  ~~~~~~~~~~~~~~

  ``i_rwsem`` is a read/write semaphore 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 or, optionally, when the list of entries in a

  directory is being retrieved with ``readdir()``.

  This has a complementary role to that of ``d_lock``: ``i_rwsem`` 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_rwsem`` 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_rwsem`` plays no role.

  The semaphore affects pathname lookup in two distinct ways.  Firstly it

  prevents changes during 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 a shared lock on ``i_rwsem``, 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 an exclusive lock on ``i_rwsem`` before performing the last lookup so

  that the required exclusion can be achieved.  How path lookup chooses

  to take, or not take, ``i_rwsem`` is one of the

  issues addressed in a subsequent section.

  If two threads attempt to look up the same name at the same time - a

  name that is not yet in the dcache - the shared lock on ``i_rwsem`` will

  not prevent them both adding new dentries with the same name.  As this
  would result in confusion an extra level of interlocking is used,

  based around a secondary hash table (``in_lookup_hashtable``) and a
  per-dentry flag bit (``DCACHE_PAR_LOOKUP``).

  
  To add a new dentry to the cache while only holding a shared lock on

  ``i_rwsem``, a thread must call ``d_alloc_parallel()``.  This allocates a

  dentry, stores the required name and parent in it, checks if there
  is already a matching dentry in the primary or secondary hash
  tables, and if not, stores the newly allocated dentry in the secondary

  hash table, with ``DCACHE_PAR_LOOKUP`` set.

  
  If a matching dentry was found in the primary hash table then that is
  returned and the caller can know that it lost a race with some other
  thread adding the entry.  If no matching dentry is found in either
  cache, the newly allocated dentry is returned and the caller can

  detect this from the presence of ``DCACHE_PAR_LOOKUP``.  In this case it

  knows that it has won any race and now is responsible for asking the
  filesystem to perform the lookup and find the matching inode.  When

  the lookup is complete, it must call ``d_lookup_done()`` which clears

  the flag and does some other house keeping, including removing the
  dentry from the secondary hash table - it will normally have been

  added to the primary hash table already.  Note that a ``struct
  waitqueue_head`` is passed to ``d_alloc_parallel()``, and
  ``d_lookup_done()`` must be called while this ``waitqueue_head`` is still

  in scope.
  
  If a matching dentry is found in the secondary hash table,

  ``d_alloc_parallel()`` has a little more work to do. It first waits for
  ``DCACHE_PAR_LOOKUP`` to be cleared, using a wait_queue that was passed
  to the instance of ``d_alloc_parallel()`` that won the race and that
  will be woken by the call to ``d_lookup_done()``.  It then checks to see

  if the dentry has now been added to the primary hash table.  If it
  has, the dentry is returned and the caller just sees that it lost any
  race.  If it hasn't been added to the primary hash table, the most
  likely explanation is that some other dentry was added instead using

  ``d_splice_alias()``.  In any case, ``d_alloc_parallel()`` repeats all the

  look ups from the start and will normally return something from the
  primary hash table.

  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.

  ``mount_lock`` is also used to detect and defend against potential attacks
  against ``LOOKUP_BENEATH`` and ``LOOKUP_IN_ROOT`` when resolving ".." (where
  the parent directory is moved outside the root, bypassing the ``path_equal()``
  check). If ``mount_lock`` is updated during the lookup and the path encounters
  a "..", a potential attack occurred and ``handle_dots()`` will bail out with
  ``-EAGAIN``.

  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``

  ----------------------------------------------

  .. _First edition Unix: https://minnie.tuhs.org/cgi-bin/utree.pl?file=V1/u2.s

  
  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`` or ``LAST_DOTDOT``.
  The ``last`` field is only valid if the type is ``LAST_NORM``.

  ``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.

  It should be noted that in the case of ``LOOKUP_IN_ROOT`` or
  ``LOOKUP_BENEATH``, the effective root becomes the directory file descriptor
  passed to ``openat2()`` (which exposes these ``LOOKUP_`` flags).

  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 ``i_rwsem``, 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 nameidata`` 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_rwsem`` 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_rwsem``, 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 autofs.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:
  
  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.
  
  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.
  
  RCU-walk - faster pathname lookup in Linux
  ==========================================
  
  RCU-walk is another algorithm for performing pathname lookup in Linux.
  It is in many ways similar to REF-walk and the two share quite a bit
  of code.  The significant difference in RCU-walk is how it allows for
  the possibility of concurrent access.
  
  We noted that REF-walk is complex because there are numerous details
  and special cases.  RCU-walk reduces this complexity by simply
  refusing to handle a number of cases -- it instead falls back to
  REF-walk.  The difficulty with RCU-walk comes from a different
  direction: unfamiliarity.  The locking rules when depending on RCU are
  quite different from traditional locking, so we will spend a little extra
  time when we come to those.
  
  Clear demarcation of roles
  --------------------------
  
  The easiest way to manage concurrency is to forcibly stop any other
  thread from changing the data structures that a given thread is
  looking at.  In cases where no other thread would even think of
  changing the data and lots of different threads want to read at the
  same time, this can be very costly.  Even when using locks that permit
  multiple concurrent readers, the simple act of updating the count of
  the number of current readers can impose an unwanted cost.  So the
  goal when reading a shared data structure that no other process is
  changing is to avoid writing anything to memory at all.  Take no
  locks, increment no counts, leave no footprints.
  
  The REF-walk mechanism already described certainly doesn't follow this
  principle, but then it is really designed to work when there may well
  be other threads modifying the data.  RCU-walk, in contrast, is
  designed for the common situation where there are lots of frequent
  readers and only occasional writers.  This may not be common in all
  parts of the filesystem tree, but in many parts it will be.  For the
  other parts it is important that RCU-walk can quickly fall back to
  using REF-walk.
  
  Pathname lookup always starts in RCU-walk mode but only remains there
  as long as what it is looking for is in the cache and is stable.  It
  dances lightly down the cached filesystem image, leaving no footprints
  and carefully watching where it is, to be sure it doesn't trip.  If it
  notices that something has changed or is changing, or if something
  isn't in the cache, then it tries to stop gracefully and switch to
  REF-walk.
  
  This stopping requires getting a counted reference on the current

  ``vfsmount`` and ``dentry``, and ensuring that these are still valid -

  that a path walk with REF-walk would have found the same entries.
  This is an invariant that RCU-walk must guarantee.  It can only make
  decisions, such as selecting the next step, that are decisions which
  REF-walk could also have made if it were walking down the tree at the
  same time.  If the graceful stop succeeds, the rest of the path is
  processed with the reliable, if slightly sluggish, REF-walk.  If
  RCU-walk finds it cannot stop gracefully, it simply gives up and
  restarts from the top with REF-walk.
  
  This pattern of "try RCU-walk, if that fails try REF-walk" can be

  clearly seen in functions like ``filename_lookup()``,
  ``filename_parentat()``, ``filename_mountpoint()``,
  ``do_filp_open()``, and ``do_file_open_root()``.  These five

  correspond roughly to the four ``path_*()`` functions we met earlier,
  each of which calls ``link_path_walk()``.  The ``path_*()`` functions are

  called using different mode flags until a mode is found which works.

  They are first called with ``LOOKUP_RCU`` set to request "RCU-walk".  If
  that fails with the error ``ECHILD`` they are called again with no

  special flag to request "REF-walk".  If either of those report the

  error ``ESTALE`` a final attempt is made with ``LOOKUP_REVAL`` set (and no
  ``LOOKUP_RCU``) to ensure that entries found in the cache are forcibly

  revalidated - normally entries are only revalidated if the filesystem
  determines that they are too old to trust.

  The ``LOOKUP_RCU`` attempt may drop that flag internally and switch to

  REF-walk, but will never then try to switch back to RCU-walk.  Places
  that trip up RCU-walk are much more likely to be near the leaves and
  so it is very unlikely that there will be much, if any, benefit from
  switching back.
  
  RCU and seqlocks: fast and light
  --------------------------------
  
  RCU is, unsurprisingly, critical to RCU-walk mode.  The

  ``rcu_read_lock()`` is held for the entire time that RCU-walk is walking

  down a path.  The particular guarantee it provides is that the key
  data structures - dentries, inodes, super_blocks, and mounts - will
  not be freed while the lock is held.  They might be unlinked or
  invalidated in one way or another, but the memory will not be
  repurposed so values in various fields will still be meaningful.  This
  is the only guarantee that RCU provides; everything else is done using
  seqlocks.
  
  As we saw above, REF-walk holds a counted reference to the current
  dentry and the current vfsmount, and does not release those references
  before taking references to the "next" dentry or vfsmount.  It also

  sometimes takes the ``d_lock`` spinlock.  These references and locks are

  taken to prevent certain changes from happening.  RCU-walk must not
  take those references or locks and so cannot prevent such changes.
  Instead, it checks to see if a change has been made, and aborts or
  retries if it has.
  
  To preserve the invariant mentioned above (that RCU-walk may only make
  decisions that REF-walk could have made), it must make the checks at
  or near the same places that REF-walk holds the references.  So, when
  REF-walk increments a reference count or takes a spinlock, RCU-walk

  samples the status of a seqlock using ``read_seqcount_begin()`` or a

  similar function.  When REF-walk decrements the count or drops the
  lock, RCU-walk checks if the sampled status is still valid using

  ``read_seqcount_retry()`` or similar.

  
  However, there is a little bit more to seqlocks than that.  If
  RCU-walk accesses two different fields in a seqlock-protected
  structure, or accesses the same field twice, there is no a priori
  guarantee of any consistency between those accesses.  When consistency
  is needed - which it usually is - RCU-walk must take a copy and then

  use ``read_seqcount_retry()`` to validate that copy.

  ``read_seqcount_retry()`` not only checks the sequence number, but also

  imposes a memory barrier so that no memory-read instruction from
  *before* the call can be delayed until *after* the call, either by the
  CPU or by the compiler.  A simple example of this can be seen in

  ``slow_dentry_cmp()`` which, for filesystems which do not use simple

  byte-wise name equality, calls into the filesystem to compare a name
  against a dentry.  The length and name pointer are copied into local

  variables, then ``read_seqcount_retry()`` is called to confirm the two
  are consistent, and only then is ``->d_compare()`` called.  When
  standard filename comparison is used, ``dentry_cmp()`` is called

  instead.  Notably it does *not* use ``read_seqcount_retry()``, but

  instead has a large comment explaining why the consistency guarantee

  isn't necessary.  A subsequent ``read_seqcount_retry()`` will be

  sufficient to catch any problem that could occur at this point.
  
  With that little refresher on seqlocks out of the way we can look at
  the bigger picture of how RCU-walk uses seqlocks.

  ``mount_lock`` and ``nd->m_seq``

  ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~

  We already met the ``mount_lock`` seqlock when REF-walk used it to

  ensure that crossing a mount point is performed safely.  RCU-walk uses
  it for that too, but for quite a bit more.

  Instead of taking a counted reference to each ``vfsmount`` as it
  descends the tree, RCU-walk samples the state of ``mount_lock`` at the

  start of the walk and stores this initial sequence number in the

  ``struct nameidata`` in the ``m_seq`` field.  This one lock and one
  sequence number are used to validate all accesses to all ``vfsmounts``,

  and all mount point crossings.  As changes to the mount table are
  relatively rare, it is reasonable to fall back on REF-walk any time
  that any "mount" or "unmount" happens.

  ``m_seq`` is checked (using ``read_seqretry()``) at the end of an RCU-walk

  sequence, whether switching to REF-walk for the rest of the path or
  when the end of the path is reached.  It is also checked when stepping

  down over a mount point (in ``__follow_mount_rcu()``) or up (in
  ``follow_dotdot_rcu()``).  If it is ever found to have changed, the

  whole RCU-walk sequence is aborted and the path is processed again by
  REF-walk.

  If RCU-walk finds that ``mount_lock`` hasn't changed then it can be sure

  that, had REF-walk taken counted references on each vfsmount, the
  results would have been the same.  This ensures the invariant holds,
  at least for vfsmount structures.

  ``dentry->d_seq`` and ``nd->seq``

  ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~

  In place of taking a count or lock on ``d_reflock``, RCU-walk samples
  the per-dentry ``d_seq`` seqlock, and stores the sequence number in the
  ``seq`` field of the nameidata structure, so ``nd->seq`` should always be
  the current sequence number of ``nd->dentry``.  This number needs to be

  revalidated after copying, and before using, the name, parent, or
  inode of the dentry.
  
  The handling of the name we have already looked at, and the parent is

  only accessed in ``follow_dotdot_rcu()`` which fairly trivially follows

  the required pattern, though it does so for three different cases.

  When not at a mount point, ``d_parent`` is followed and its ``d_seq`` is

  collected.  When we are at a mount point, we instead follow the

  ``mnt->mnt_mountpoint`` link to get a new dentry and collect its
  ``d_seq``.  Then, after finally finding a ``d_parent`` to follow, we must

  check if we have landed on a mount point and, if so, must find that

  mount point and follow the ``mnt->mnt_root`` link.  This would imply a

  somewhat unusual, but certainly possible, circumstance where the
  starting point of the path lookup was in part of the filesystem that
  was mounted on, and so not visible from the root.

  The inode pointer, stored in ``->d_inode``, is a little more

  interesting.  The inode will always need to be accessed at least
  twice, once to determine if it is NULL and once to verify access
  permissions.  Symlink handling requires a validated inode pointer too.
  Rather than revalidating on each access, a copy is made on the first

  access and it is stored in the ``inode`` field of ``nameidata`` from where

  it can be safely accessed without further validation.

  ``lookup_fast()`` is the only lookup routine that is used in RCU-mode,
  ``lookup_slow()`` being too slow and requiring locks.  It is in
  ``lookup_fast()`` that we find the important "hand over hand" tracking

  of the current dentry.

  The current ``dentry`` and current ``seq`` number are passed to
  ``__d_lookup_rcu()`` which, on success, returns a new ``dentry`` and a
  new ``seq`` number.  ``lookup_fast()`` then copies the inode pointer and
  revalidates the new ``seq`` number.  It then validates the old ``dentry``
  with the old ``seq`` number one last time and only then continues.  This
  process of getting the ``seq`` number of the new dentry and then
  checking the ``seq`` number of the old exactly mirrors the process of

  getting a counted reference to the new dentry before dropping that for
  the old dentry which we saw in REF-walk.

  No ``inode->i_rwsem`` or even ``rename_lock``

  ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~

  A semaphore is a fairly heavyweight lock that can only be taken when it is

  permissible to sleep.  As ``rcu_read_lock()`` forbids sleeping,
  ``inode->i_rwsem`` plays no role in RCU-walk.  If some other thread does
  take ``i_rwsem`` and modifies the directory in a way that RCU-walk needs

  to notice, the result will be either that RCU-walk fails to find the
  dentry that it is looking for, or it will find a dentry which

  ``read_seqretry()`` won't validate.  In either case it will drop down to

  REF-walk mode which can take whatever locks are needed.

  Though ``rename_lock`` could be used by RCU-walk as it doesn't require
  any sleeping, RCU-walk doesn't bother.  REF-walk uses ``rename_lock`` to

  protect against the possibility of hash chains in the dcache changing
  while they are being searched.  This can result in failing to find
  something that actually is there.  When RCU-walk fails to find
  something in the dentry cache, whether it is really there or not, it
  already drops down to REF-walk and tries again with appropriate
  locking.  This neatly handles all cases, so adding extra checks on
  rename_lock would bring no significant value.

  ``unlazy walk()`` and ``complete_walk()``

  -----------------------------------------

  
  That "dropping down to REF-walk" typically involves a call to

  ``unlazy_walk()``, so named because "RCU-walk" is also sometimes
  referred to as "lazy walk".  ``unlazy_walk()`` is called when

  following the path down to the current vfsmount/dentry pair seems to
  have proceeded successfully, but the next step is problematic.  This
  can happen if the next name cannot be found in the dcache, if
  permission checking or name revalidation couldn't be achieved while

  the ``rcu_read_lock()`` is held (which forbids sleeping), if an

  automount point is found, or in a couple of cases involving symlinks.

  It is also called from ``complete_walk()`` when the lookup has reached

  the final component, or the very end of the path, depending on which
  particular flavor of lookup is used.
  
  Other reasons for dropping out of RCU-walk that do not trigger a call

  to ``unlazy_walk()`` are when some inconsistency is found that cannot be
  handled immediately, such as ``mount_lock`` or one of the ``d_seq``

  seqlocks reporting a change.  In these cases the relevant function

  will return ``-ECHILD`` which will percolate up until it triggers a new

  attempt from the top using REF-walk.

  For those cases where ``unlazy_walk()`` is an option, it essentially

  takes a reference on each of the pointers that it holds (vfsmount,
  dentry, and possibly some symbolic links) and then verifies that the
  relevant seqlocks have not been changed.  If there have been changes,

  it, too, aborts with ``-ECHILD``, otherwise the transition to REF-walk

  has been a success and the lookup process continues.
  
  Taking a reference on those pointers is not quite as simple as just
  incrementing a counter.  That works to take a second reference if you
  already have one (often indirectly through another object), but it
  isn't sufficient if you don't actually have a counted reference at

  all.  For ``dentry->d_lockref``, it is safe to increment the reference

  counter to get a reference unless it has been explicitly marked as

  "dead" which involves setting the counter to ``-128``.
  ``lockref_get_not_dead()`` achieves this.

  For ``mnt->mnt_count`` it is safe to take a reference as long as
  ``mount_lock`` is then used to validate the reference.  If that

  validation fails, it may *not* be safe to just drop that reference in

  the standard way of calling ``mnt_put()`` - an unmount may have
  progressed too far.  So the code in ``legitimize_mnt()``, when it

  finds that the reference it got might not be safe, checks the

  ``MNT_SYNC_UMOUNT`` flag to determine if a simple ``mnt_put()`` is

  correct, or if it should just decrement the count and pretend none of
  this ever happened.
  
  Taking care in filesystems

  --------------------------

  
  RCU-walk depends almost entirely on cached information and often will
  not call into the filesystem at all.  However there are two places,
  besides the already-mentioned component-name comparison, where the
  file system might be included in RCU-walk, and it must know to be
  careful.
  
  If the filesystem has non-standard permission-checking requirements -
  such as a networked filesystem which may need to check with the server

  - the ``i_op->permission`` interface might be called during RCU-walk.
  In this case an extra "``MAY_NOT_BLOCK``" flag is passed so that it
  knows not to sleep, but to return ``-ECHILD`` if it cannot complete
  promptly.  ``i_op->permission`` is given the inode pointer, not the

  dentry, so it doesn't need to worry about further consistency checks.
  However if it accesses any other filesystem data structures, it must

  ensure they are safe to be accessed with only the ``rcu_read_lock()``
  held.  This typically means they must be freed using ``kfree_rcu()`` or

  similar.

  .. _READ_ONCE: https://lwn.net/Articles/624126/

  
  If the filesystem may need to revalidate dcache entries, then

  ``d_op->d_revalidate`` may be called in RCU-walk too.  This interface
  *is* passed the dentry but does not have access to the ``inode`` or the
  ``seq`` number from the ``nameidata``, so it needs to be extra careful

  when accessing fields in the dentry.  This "extra care" typically

  involves using  `READ_ONCE() <READ_ONCE_>`_ to access fields, and verifying the

  result is not NULL before using it.  This pattern can be seen in

  ``nfs_lookup_revalidate()``.

  
  A pair of patterns
  ------------------
  
  In various places in the details of REF-walk and RCU-walk, and also in
  the big picture, there are a couple of related patterns that are worth
  being aware of.
  
  The first is "try quickly and check, if that fails try slowly".  We
  can see that in the high-level approach of first trying RCU-walk and

  then trying REF-walk, and in places where ``unlazy_walk()`` is used to

  switch to REF-walk for the rest of the path.  We also saw it earlier

  in ``dget_parent()`` when following a "``..``" link.  It tries a quick way

  to get a reference, then falls back to taking locks if needed.
  
  The second pattern is "try quickly and check, if that fails try

  again - repeatedly".  This is seen with the use of ``rename_lock`` and
  ``mount_lock`` in REF-walk.  RCU-walk doesn't make use of this pattern -

  if anything goes wrong it is much safer to just abort and try a more
  sedate approach.
  
  The emphasis here is "try quickly and check".  It should probably be

  "try quickly *and carefully*, then check".  The fact that checking is

  needed is a reminder that the system is dynamic and only a limited
  number of things are safe at all.  The most likely cause of errors in
  this whole process is assuming something is safe when in reality it
  isn't.  Careful consideration of what exactly guarantees the safety of
  each access is sometimes necessary.
  
  A walk among the symlinks
  =========================
  
  There are several basic issues that we will examine to understand the
  handling of symbolic links:  the symlink stack, together with cache
  lifetimes, will help us understand the overall recursive handling of
  symlinks and lead to the special care needed for the final component.
  Then a consideration of access-time updates and summary of the various
  flags controlling lookup will finish the story.
  
  The symlink stack
  -----------------
  
  There are only two sorts of filesystem objects that can usefully
  appear in a path prior to the final component: directories and symlinks.
  Handling directories is quite straightforward: the new directory
  simply becomes the starting point at which to interpret the next
  component on the path.  Handling symbolic links requires a bit more
  work.
  
  Conceptually, symbolic links could be handled by editing the path.  If
  a component name refers to a symbolic link, then that component is
  replaced by the body of the link and, if that body starts with a '/',
  then all preceding parts of the path are discarded.  This is what the

  "``readlink -f``" command does, though it also edits out "``.``" and
  "``..``" components.

  
  Directly editing the path string is not really necessary when looking
  up a path, and discarding early components is pointless as they aren't
  looked at anyway.  Keeping track of all remaining components is
  important, but they can of course be kept separately; there is no need
  to concatenate them.  As one symlink may easily refer to another,
  which in turn can refer to a third, we may need to keep the remaining
  components of several paths, each to be processed when the preceding
  ones are completed.  These path remnants are kept on a stack of
  limited size.
  
  There are two reasons for placing limits on how many symlinks can
  occur in a single path lookup.  The most obvious is to avoid loops.
  If a symlink referred to itself either directly or through
  intermediaries, then following the symlink can never complete

  successfully - the error ``ELOOP`` must be returned.  Loops can be

  detected without imposing limits, but limits are the simplest solution
  and, given the second reason for restriction, quite sufficient.

  .. _outlined recently: http://thread.gmane.org/gmane.linux.kernel/1934390/focus=1934550

  The second reason was `outlined recently`_ by Linus:

     Because it's a latency and DoS issue too. We need to react well to
     true loops, but also to "very deep" non-loops. It's not about memory
     use, it's about users triggering unreasonable CPU resources.

  Linux imposes a limit on the length of any pathname: ``PATH_MAX``, which

  is 4096.  There are a number of reasons for this limit; not letting the
  kernel spend too much time on just one path is one of them.  With
  symbolic links you can effectively generate much longer paths so some
  sort of limit is needed for the same reason.  Linux imposes a limit of
  at most 40 symlinks in any one path lookup.  It previously imposed a
  further limit of eight on the maximum depth of recursion, but that was
  raised to 40 when a separate stack was implemented, so there is now
  just the one limit.

  The ``nameidata`` structure that we met in an earlier article contains a

  small stack that can be used to store the remaining part of up to two
  symlinks.  In many cases this will be sufficient.  If it isn't, a
  separate stack is allocated with room for 40 symlinks.  Pathname
  lookup will never exceed that stack as, once the 40th symlink is
  detected, an error is returned.
  
  It might seem that the name remnants are all that needs to be stored on
  this stack, but we need a bit more.  To see that, we need to move on to
  cache lifetimes.
  
  Storage and lifetime of cached symlinks
  ---------------------------------------
  
  Like other filesystem resources, such as inodes and directory
  entries, symlinks are cached by Linux to avoid repeated costly access
  to external storage.  It is particularly important for RCU-walk to be
  able to find and temporarily hold onto these cached entries, so that
  it doesn't need to drop down into REF-walk.

  .. _object-oriented design pattern: https://lwn.net/Articles/446317/

  
  While each filesystem is free to make its own choice, symlinks are
  typically stored in one of two places.  Short symlinks are often

  stored directly in the inode.  When a filesystem allocates a ``struct
  inode`` it typically allocates extra space to store private data (a
  common `object-oriented design pattern`_ in the kernel).  This will

  sometimes include space for a symlink.  The other common location is
  in the page cache, which normally stores the content of files.  The
  pathname in a symlink can be seen as the content of that symlink and
  can easily be stored in the page cache just like file content.
  
  When neither of these is suitable, the next most likely scenario is
  that the filesystem will allocate some temporary memory and copy or
  construct the symlink content into that memory whenever it is needed.
  
  When the symlink is stored in the inode, it has the same lifetime as
  the inode which, itself, is protected by RCU or by a counted reference
  on the dentry.  This means that the mechanisms that pathname lookup
  uses to access the dcache and icache (inode cache) safely are quite
  sufficient for accessing some cached symlinks safely.  In these cases,

  the ``i_link`` pointer in the inode is set to point to wherever the

  symlink is stored and it can be accessed directly whenever needed.
  
  When the symlink is stored in the page cache or elsewhere, the
  situation is not so straightforward.  A reference on a dentry or even
  on an inode does not imply any reference on cached pages of that

  inode, and even an ``rcu_read_lock()`` is not sufficient to ensure that

  a page will not disappear.  So for these symlinks the pathname lookup
  code needs to ask the filesystem to provide a stable reference and,
  significantly, needs to release that reference when it is finished
  with it.
  
  Taking a reference to a cache page is often possible even in RCU-walk
  mode.  It does require making changes to memory, which is best avoided,
  but that isn't necessarily a big cost and it is better than dropping
  out of RCU-walk mode completely.  Even filesystems that allocate

  space to copy the symlink into can use ``GFP_ATOMIC`` to often successfully

  allocate memory without the need to drop out of RCU-walk.  If a
  filesystem cannot successfully get a reference in RCU-walk mode, it

  must return ``-ECHILD`` and ``unlazy_walk()`` will be called to return to

  REF-walk mode in which the filesystem is allowed to sleep.

  The place for all this to happen is the ``i_op->follow_link()`` inode

  method.  In the present mainline code this is never actually called in
  RCU-walk mode as the rewrite is not quite complete.  It is likely that

  in a future release this method will be passed an ``inode`` pointer when

  called in RCU-walk mode so it both (1) knows to be careful, and (2) has the

  validated pointer.  Much like the ``i_op->permission()`` method we
  looked at previously, ``->follow_link()`` would need to be careful that

  all the data structures it references are safe to be accessed while
  holding no counted reference, only the RCU lock.  Though getting a

  reference with ``->follow_link()`` is not yet done in RCU-walk mode, the

  code is ready to release the reference when that does happen.
  
  This need to drop the reference to a symlink adds significant
  complexity.  It requires a reference to the inode so that the

  ``i_op->put_link()`` inode operation can be called.  In REF-walk, that

  reference is kept implicitly through a reference to the dentry, so

  keeping the ``struct path`` of the symlink is easiest.  For RCU-walk,

  the pointer to the inode is kept separately.  To allow switching from
  RCU-walk back to REF-walk in the middle of processing nested symlinks
  we also need the seq number for the dentry so we can confirm that
  switching back was safe.
  
  Finally, when providing a reference to a symlink, the filesystem also

  provides an opaque "cookie" that must be passed to ``->put_link()`` so that it

  knows what to free.  This might be the allocated memory area, or a

  pointer to the ``struct page`` in the page cache, or something else

  completely.  Only the filesystem knows what it is.
  
  In order for the reference to each symlink to be dropped when the walk completes,
  whether in RCU-walk or REF-walk, the symlink stack needs to contain,
  along with the path remnants:

  - the ``struct path`` to provide a reference to the inode in REF-walk
  - the ``struct inode *`` to provide a reference to the inode in RCU-walk
  - the ``seq`` to allow the path to be safely switched from RCU-walk to REF-walk
  - the ``cookie`` that tells ``->put_path()`` what to put.

  
  This means that each entry in the symlink stack needs to hold five
  pointers and an integer instead of just one pointer (the path
  remnant).  On a 64-bit system, this is about 40 bytes per entry;
  with 40 entries it adds up to 1600 bytes total, which is less than
  half a page.  So it might seem like a lot, but is by no means
  excessive.

  Note that, in a given stack frame, the path remnant (``name``) is not

  part of the symlink that the other fields refer to.  It is the remnant
  to be followed once that symlink has been fully parsed.
  
  Following the symlink
  ---------------------

  The main loop in ``link_path_walk()`` iterates seamlessly over all

  components in the path and all of the non-final symlinks.  As symlinks

  are processed, the ``name`` pointer is adjusted to point to a new

  symlink, or is restored from the stack, so that much of the loop

  doesn't need to notice.  Getting this ``name`` variable on and off the

  stack is very straightforward; pushing and popping the references is
  a little more complex.

  When a symlink is found, ``walk_component()`` returns the value ``1``
  (``0`` is returned for any other sort of success, and a negative number
  is, as usual, an error indicator).  This causes ``get_link()`` to be

  called; it then gets the link from the filesystem.  Providing that

  operation is successful, the old path ``name`` is placed on the stack,
  and the new value is used as the ``name`` for a while.  When the end of
  the path is found (i.e. ``*name`` is ``'\0'``) the old ``name`` is restored

  off the stack and path walking continues.
  
  Pushing and popping the reference pointers (inode, cookie, etc.) is more
  complex in part because of the desire to handle tail recursion.  When
  the last component of a symlink itself points to a symlink, we
  want to pop the symlink-just-completed off the stack before pushing
  the symlink-just-found to avoid leaving empty path remnants that would
  just get in the way.
  
  It is most convenient to push the new symlink references onto the

  stack in ``walk_component()`` immediately when the symlink is found;
  ``walk_component()`` is also the last piece of code that needs to look at the

  old symlink as it walks that last component.  So it is quite

  convenient for ``walk_component()`` to release the old symlink and pop

  the references just before pushing the reference information for the

  new symlink.  It is guided in this by two flags; ``WALK_GET``, which

  gives it permission to follow a symlink if it finds one, and

  ``WALK_PUT``, which tells it to release the current symlink after it has been
  followed.  ``WALK_PUT`` is tested first, leading to a call to
  ``put_link()``.  ``WALK_GET`` is tested subsequently (by
  ``should_follow_link()``) leading to a call to ``pick_link()`` which sets

  up the stack frame.

  Symlinks with no final component
  ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~

  
  A pair of special-case symlinks deserve a little further explanation.

  Both result in a new ``struct path`` (with mount and dentry) being set
  up in the ``nameidata``, and result in ``get_link()`` returning ``NULL``.

  The more obvious case is a symlink to "``/``".  All symlinks starting
  with "``/``" are detected in ``get_link()`` which resets the ``nameidata``

  to point to the effective filesystem root.  If the symlink only

  contains "``/``" then there is nothing more to do, no components at all,
  so ``NULL`` is returned to indicate that the symlink can be released and

  the stack frame discarded.

  The other case involves things in ``/proc`` that look like symlinks but

  aren't really (and are therefore commonly referred to as "magic-links")::

       $ ls -l /proc/self/fd/1
       lrwx------ 1 neilb neilb 64 Jun 13 10:19 /proc/self/fd/1 -> /dev/pts/4

  Every open file descriptor in any process is represented in ``/proc`` by

  something that looks like a symlink.  It is really a reference to the

  target file, not just the name of it.  When you ``readlink`` these

  objects you get a name that might refer to the same file - unless it

  has been unlinked or mounted over.  When ``walk_component()`` follows
  one of these, the ``->follow_link()`` method in "procfs" doesn't return
  a string name, but instead calls ``nd_jump_link()`` which updates the
  ``nameidata`` in place to point to that target.  ``->follow_link()`` then
  returns ``NULL``.  Again there is no final component and ``get_link()``
  reports this by leaving the ``last_type`` field of ``nameidata`` as
  ``LAST_BIND``.

  
  Following the symlink in the final component
  --------------------------------------------

  All this leads to ``link_path_walk()`` walking down every component, and

  following all symbolic links it finds, until it reaches the final

  component.  This is just returned in the ``last`` field of ``nameidata``.

  For some callers, this is all they need; they want to create that

  ``last`` name if it doesn't exist or give an error if it does.  Other

  callers will want to follow a symlink if one is found, and possibly
  apply special handling to the last component of that symlink, rather
  than just the last component of the original file name.  These callers

  potentially need to call ``link_path_walk()`` again and again on

  successive symlinks until one is found that doesn't point to another
  symlink.

  This case is handled by the relevant caller of ``link_path_walk()``, such as
  ``path_lookupat()`` using a loop that calls ``link_path_walk()``, and then

  handles the final component.  If the final component is a symlink

  that needs to be followed, then ``trailing_symlink()`` is called to set
  things up properly and the loop repeats, calling ``link_path_walk()``

  again.  This could loop as many as 40 times if the last component of
  each symlink is another symlink.
  
  The various functions that examine the final component and possibly

  report that it is a symlink are ``lookup_last()``, ``mountpoint_last()``
  and ``do_last()``, each of which use the same convention as
  ``walk_component()`` of returning ``1`` if a symlink was found that needs

  to be followed.

  Of these, ``do_last()`` is the most interesting as it is used for
  opening a file.  Part of ``do_last()`` runs with ``i_rwsem`` held and this
  part is in a separate function: ``lookup_open()``.

  Explaining ``do_last()`` completely is beyond the scope of this article,

  but a few highlights should help those interested in exploring the
  code.

  1. Rather than just finding the target file, ``do_last()`` needs to open
     it.  If the file was found in the dcache, then ``vfs_open()`` is used for
     this.  If not, then ``lookup_open()`` will either call ``atomic_open()`` (if
     the filesystem provides it) to combine the final lookup with the open, or
     will perform the separate ``lookup_real()`` and ``vfs_create()`` steps
     directly.  In the later case the actual "open" of this newly found or
     created file will be performed by ``vfs_open()``, just as if the name
     were found in the dcache.
  
  2. ``vfs_open()`` can fail with ``-EOPENSTALE`` if the cached information
     wasn't quite current enough.  Rather than restarting the lookup from
     the top with ``LOOKUP_REVAL`` set, ``lookup_open()`` is called instead,
     giving the filesystem a chance to resolve small inconsistencies.
     If that doesn't work, only then is the lookup restarted from the top.

  
  3. An open with O_CREAT **does** follow a symlink in the final component,

     unlike other creation system calls (like ``mkdir``).  So the sequence::

            ln -s bar /tmp/foo
            echo hello > /tmp/foo

     will create a file called ``/tmp/bar``.  This is not permitted if
     ``O_EXCL`` is set but otherwise is handled for an O_CREAT open much
     like for a non-creating open: ``should_follow_link()`` returns ``1``, and
     so does ``do_last()`` so that ``trailing_symlink()`` gets called and the
     open process continues on the symlink that was found.

  
  Updating the access time
  ------------------------
  
  We previously said of RCU-walk that it would "take no locks, increment
  no counts, leave no footprints."  We have since seen that some
  "footprints" can be needed when handling symlinks as a counted
  reference (or even a memory allocation) may be needed.  But these
  footprints are best kept to a minimum.
  
  One other place where walking down a symlink can involve leaving
  footprints in a way that doesn't affect directories is in updating access times.
  In Unix (and Linux) every filesystem object has a "last accessed

  time", or "``atime``".  Passing through a directory to access a file

  within is not considered to be an access for the purposes of

  ``atime``; only listing the contents of a directory can update its ``atime``.
  Symlinks are different it seems.  Both reading a symlink (with ``readlink()``)

  and looking up a symlink on the way to some other destination can
  update the atime on that symlink.

  .. _clearest statement: https://pubs.opengroup.org/onlinepubs/9699919799/basedefs/V1_chap04.html#tag_04_08

  
  It is not clear why this is the case; POSIX has little to say on the

  subject.  The `clearest statement`_ is that, if a particular implementation

  updates a timestamp in a place not specified by POSIX, this must be
  documented "except that any changes caused by pathname resolution need
  not be documented".  This seems to imply that POSIX doesn't really
  care about access-time updates during pathname lookup.

  .. _Linux 1.3.87: https://git.kernel.org/cgit/linux/kernel/git/history/history.git/diff/fs/ext2/symlink.c?id=f806c6db77b8eaa6e00dcfb6b567706feae8dbb8

  An examination of history shows that prior to `Linux 1.3.87`_, the ext2

  filesystem, at least, didn't update atime when following a link.
  Unfortunately we have no record of why that behavior was changed.
  
  In any case, access time must now be updated and that operation can be
  quite complex.  Trying to stay in RCU-walk while doing it is best

  avoided.  Fortunately it is often permitted to skip the ``atime``
  update.  Because ``atime`` updates cause performance problems in various
  areas, Linux supports the ``relatime`` mount option, which generally
  limits the updates of ``atime`` to once per day on files that aren't

  being changed (and symlinks never change once created).  Even without

  ``relatime``, many filesystems record ``atime`` with a one-second

  granularity, so only one update per second is required.

  It is easy to test if an ``atime`` update is needed while in RCU-walk

  mode and, if it isn't, the update can be skipped and RCU-walk mode

  continues.  Only when an ``atime`` update is actually required does the

  path walk drop down to REF-walk.  All of this is handled in the

  ``get_link()`` function.

  
  A few flags
  -----------
  
  A suitable way to wrap up this tour of pathname walking is to list

  the various flags that can be stored in the ``nameidata`` to guide the

  lookup process.  Many of these are only meaningful on the final

  component, others reflect the current state of the pathname lookup, and some
  apply restrictions to all path components encountered in the path lookup.

  And then there is ``LOOKUP_EMPTY``, which doesn't fit conceptually with

  the others.  If this is not set, an empty pathname causes an error
  very early on.  If it is set, empty pathnames are not considered to be
  an error.

  Global state flags
  ~~~~~~~~~~~~~~~~~~

  We have already met two global state flags: ``LOOKUP_RCU`` and
  ``LOOKUP_REVAL``.  These select between one of three overall approaches

  to lookup: RCU-walk, REF-walk, and REF-walk with forced revalidation.

  ``LOOKUP_PARENT`` indicates that the final component hasn't been reached

  yet.  This is primarily used to tell the audit subsystem the full
  context of a particular access being audited.

  ``LOOKUP_ROOT`` indicates that the ``root`` field in the ``nameidata`` was

  provided by the caller, so it shouldn't be released when it is no
  longer needed.

  ``LOOKUP_JUMPED`` means that the current dentry was chosen not because

  it had the right name but for some other reason.  This happens when

  following "``..``", following a symlink to ``/``, crossing a mount point

  or accessing a "``/proc/$PID/fd/$FD``" symlink (also known as a "magic
  link"). In this case the filesystem has not been asked to revalidate the
  name (with ``d_revalidate()``).  In such cases the inode may still need
  to be revalidated, so ``d_op->d_weak_revalidate()`` is called if

  ``LOOKUP_JUMPED`` is set when the look completes - which may be at the

  final component or, when creating, unlinking, or renaming, at the penultimate component.

  Resolution-restriction flags
  ~~~~~~~~~~~~~~~~~~~~~~~~~~~~
  
  In order to allow userspace to protect itself against certain race conditions
  and attack scenarios involving changing path components, a series of flags are
  available which apply restrictions to all path components encountered during
  path lookup. These flags are exposed through ``openat2()``'s ``resolve`` field.
  
  ``LOOKUP_NO_SYMLINKS`` blocks all symlink traversals (including magic-links).
  This is distinctly different from ``LOOKUP_FOLLOW``, because the latter only
  relates to restricting the following of trailing symlinks.
  
  ``LOOKUP_NO_MAGICLINKS`` blocks all magic-link traversals. Filesystems must
  ensure that they return errors from ``nd_jump_link()``, because that is how
  ``LOOKUP_NO_MAGICLINKS`` and other magic-link restrictions are implemented.
  
  ``LOOKUP_NO_XDEV`` blocks all ``vfsmount`` traversals (this includes both
  bind-mounts and ordinary mounts). Note that the ``vfsmount`` which contains the
  lookup is determined by the first mountpoint the path lookup reaches --
  absolute paths start with the ``vfsmount`` of ``/``, and relative paths start
  with the ``dfd``'s ``vfsmount``. Magic-links are only permitted if the
  ``vfsmount`` of the path is unchanged.
  
  ``LOOKUP_BENEATH`` blocks any path components which resolve outside the
  starting point of the resolution. This is done by blocking ``nd_jump_root()``
  as well as blocking ".." if it would jump outside the starting point.
  ``rename_lock`` and ``mount_lock`` are used to detect attacks against the
  resolution of "..". Magic-links are also blocked.
  
  ``LOOKUP_IN_ROOT`` resolves all path components as though the starting point

  were the filesystem root. ``nd_jump_root()`` brings the resolution back to

  the starting point, and ".." at the starting point will act as a no-op. As with
  ``LOOKUP_BENEATH``, ``rename_lock`` and ``mount_lock`` are used to detect
  attacks against ".." resolution. Magic-links are also blocked.

  Final-component flags
  ~~~~~~~~~~~~~~~~~~~~~

  
  Some of these flags are only set when the final component is being
  considered.  Others are only checked for when considering that final
  component.

  ``LOOKUP_AUTOMOUNT`` ensures that, if the final component is an automount

  point, then the mount is triggered.  Some operations would trigger it

  anyway, but operations like ``stat()`` deliberately don't.  ``statfs()``
  needs to trigger the mount but otherwise behaves a lot like ``stat()``, so
  it sets ``LOOKUP_AUTOMOUNT``, as does "``quotactl()``" and the handling of
  "``mount --bind``".

  ``LOOKUP_FOLLOW`` has a similar function to ``LOOKUP_AUTOMOUNT`` but for

  symlinks.  Some system calls set or clear it implicitly, while

  others have API flags such as ``AT_SYMLINK_FOLLOW`` and
  ``UMOUNT_NOFOLLOW`` to control it.  Its effect is similar to
  ``WALK_GET`` that we already met, but it is used in a different way.

  ``LOOKUP_DIRECTORY`` insists that the final component is a directory.

  Various callers set this and it is also set when the final component
  is found to be followed by a slash.

  Finally ``LOOKUP_OPEN``, ``LOOKUP_CREATE``, ``LOOKUP_EXCL``, and
  ``LOOKUP_RENAME_TARGET`` are not used directly by the VFS but are made
  available to the filesystem and particularly the ``->d_revalidate()``

  method.  A filesystem can choose not to bother revalidating too hard
  if it knows that it will be asked to open or create the file soon.

  These flags were previously useful for ``->lookup()`` too but with the
  introduction of ``->atomic_open()`` they are less relevant there.

  
  End of the road
  ---------------
  
  Despite its complexity, all this pathname lookup code appears to be
  in good shape - various parts are certainly easier to understand now
  than even a couple of releases ago.  But that doesn't mean it is
  "finished".   As already mentioned, RCU-walk currently only follows
  symlinks that are stored in the inode so, while it handles many ext4
  symlinks, it doesn't help with NFS, XFS, or Btrfs.  That support
  is not likely to be long delayed.