标签：dentry Pathname RCU component walk lookup path chatgpt

原文：https://www.kernel.org/doc/html/latest/filesystems/path-lookup.html

内核中文件系统相关的文档汇总：Filesystems in the Linux kernel

This write-up is based on three articles published at lwn.net:

• Pathname lookup in Linux
• RCU-walk: faster pathname lookup in Linux
• 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.

路径名查找最明显的方面是它的复杂性，几乎不需要太多探索就能发现。有许多规则、特殊情况和实现选择，它们共同使得读者困惑不已。计算机科学长期以来一直熟悉这种复杂性，并有工具来帮助管理它。我们将广泛使用的一种工具是"分而治之"的方法。在分析的早期阶段，我们将排除符号链接，直到最后一部分再处理它们。在我们接触到符号链接之前，根据虚拟文件系统（VFS）的锁定方法，我们还有另一个主要的划分，这将使我们能够单独审查"REF-walk"和"RCU-walk"。但我们不要过于急于求成。首先，我们需要澄清一些重要的低级区别。

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 "*at()" system calls such as openat().

路径名（有时称为“文件名”）用于标识文件系统中的对象，对大多数读者来说应该很熟悉。它们包含两种类型的元素：“斜杠”，即一个或多个“/”字符的序列，以及“组件”，即一个或多个非“/”字符的序列。这形成了两种类型的路径。以斜杠开头的路径是“绝对路径”，从文件系统根目录开始。其他路径是“相对路径”，从当前目录开始，或者从通过类似于openat()的“*at()”系统调用给定的文件描述符指定的其他位置开始。

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() passing the file descriptor, an empty path, and the AT_EMPTY_PATH flag.

诱人的是将第二种类型描述为以组件开头，但这并不总是准确的：路径名既可以缺少斜杠又可以缺少组件，换句话说，它可以为空。在POSIX中通常是禁止的，但是在Linux中的一些“*at()”系统调用中，当给定AT_EMPTY_PATH标志时，允许这种情况。例如，如果您在可执行文件上有一个打开的文件描述符，可以通过调用execveat()传递文件描述符、空路径和AT_EMPTY_PATH标志来执行它。

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.

这些路径可以分为两个部分：最后一个组件和其他所有内容。其他所有内容比较容易处理。在所有情况下，它必须标识一个已经存在的目录，否则将报告错误，例如ENOENT或ENOTDIR。

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.

最后一个组件就没有那么简单了。不仅不同的系统调用对其解释有很大差异（例如，有些创建它，有些不创建），而且它甚至可能不存在：空路径名和只有斜杠的路径名都没有最后一个组件。如果它存在，它可能是“.”或“..”，这两者与其他组件的处理方式有很大不同。

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:

如果路径名以斜杠结尾，例如“/tmp/foo/”，可能会诱使认为它具有一个空的最后一个组件。在许多情况下，这会得到正确的结果，但并非总是如此。特别是，mkdir()和rmdir()分别创建或删除由最后一个组件命名的目录，并且它们要求能够处理以“/”结尾的路径名。根据POSIX的规定：

A pathname that contains at least one non- character and that ends with one or more trailing characters shall not be resolved successfully unless the last pathname component before the trailing characters names an existing directory or a directory entry that is to be created for a directory immediately after the pathname is resolved.

只要路径名中至少包含一个非字符，并且以一个或多个尾部字符结尾，就不能成功解析该路径名，除非在尾部字符之前的最后一个路径名组件指定了一个现有目录或在路径名解析后立即创建一个目录的目录项。

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.

Linux的路径名解析代码（主要位于fs/namei.c中）处理了所有这些问题：将路径分解为组件，将“其他所有内容”与最后一个组件分开处理，并检查在不允许的情况下是否使用了尾部斜杠。它还解决了并发访问的重要问题。

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.

当一个进程查找路径名时，另一个进程可能正在进行会影响该查找的更改。一个相当极端的情况是，如果在另一个进程正在查找“a/b/..”时将“a/b”重命名为“a/c/b”，那个进程可能会成功解析为“a/c”。大多数竞争情况要复杂得多，路径名查找的任务的一个重要部分就是防止它们产生破坏性的影响。许多可能的竞争情况在“dcache”的上下文中最清楚地显示出来，理解这一点对于理解路径名查找至关重要。

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.

"dcache"缓存每个文件系统中有关名称的信息，以便快速进行查找。每个条目（称为"dentry"）包含三个重要字段：组件名称、指向父dentry的指针以及指向包含有关具有给定名称的父对象的进一步信息的"inode"指针。inode指针可以为NULL，表示该名称在父目录中不存在。虽然目录的dentry可以与子目录的dentry存在链接，但该链接不用于路径名查找，因此在此不予考虑。

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.

除了加速查找之外，"dcache"还有许多其他用途。其中一个特别相关的用途是它与记录哪个文件系统在哪里挂载的挂载表密切集成。挂载表实际上存储的是哪个dentry被挂载在哪个其他dentry之上。

When considering the dcache, we have another of our "two types" distinctions: there are two types of filesystems.

在考虑"dcache"时，我们有另一个"两种类型"的区分：有两种类型的文件系统。

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.

某些文件系统确保"dcache"中的信息始终完全准确（尽管不一定完整）。这可以使虚拟文件系统（VFS）在不与文件系统进行检查的情况下确定特定文件是否存在，并意味着VFS可以保护文件系统免受某些竞争和其他问题的影响。这些通常是"本地"文件系统，如ext3、XFS和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.

其他文件系统无法提供该保证，因为它们无法做到。这些通常是跨网络共享的文件系统，无论是像NFS和9P这样的远程文件系统，还是像ocfs2或cephfs这样的集群文件系统。这些文件系统允许VFS重新验证缓存的信息，并必须提供自己的保护措施来防止尴尬的竞争。VFS可以通过在dentry中设置DCACHE_OP_REVALIDATE标志来检测这些文件系统。

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.

在对所有这些部分进行仔细分类之后，我们现在可以开始研究沿着路径行走的实际过程了。特别是，我们将从处理路径名的“其他一切”部分开始，并重点关注“REF-walk”并发管理方法。如果忽略所有仅在设置了“LOOKUP_RCU”（表示使用RCU-walk）时运行的地方，可以在link_path_walk()函数中找到此代码。

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.

REF-walk在锁和引用计数方面相对较重。虽然不像旧的“大内核锁”时代那样重，但在需要时肯定不会害怕使用锁。它使用了各种不同的并发控制机制。假定您具备对各种原语的背景了解，或者可以从其他地方（例如《遇见锁匠》）获取相关信息。

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.

这里使用了lockref原语来提供自旋锁和引用计数。这个原语的特殊之处在于，概念上的序列"锁定；增加引用计数；解锁"通常可以通过单个原子内存操作来完成。

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.

持有dentry的引用可以确保dentry不会突然被释放并用于其他用途，因此各个字段中的值将按预期行为。它还在一定程度上保护了对inode的->d_inode引用。

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.

dentry与其inode之间的关联是相当持久的。例如，当文件被重命名时，dentry和inode一起移动到新位置。当创建文件时，dentry最初将为负（即d_inode为NULL），并作为创建操作的一部分分配给新的inode。

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.

当文件被删除时，可以通过将d_inode设置为NULL或从哈希表（稍后描述）中删除来在缓存中反映出来，该哈希表用于在父目录中查找名称。如果dentry仍在使用中，则使用第二个选项，因为在文件被删除后继续使用打开的文件是完全合法的，并且保留dentry有助于此。只有当dentry没有其他使用（即d_lockref中的引用计数为1）时，d_inode才会被设置为NULL。以这种方式进行操作对于非常常见的情况更加高效。

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

因此，只要对dentry持有计数引用，非NULL的->d_inode值将永远不会改变.

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.

d_lock是d_lockref中的自旋锁的同义词。对于我们的目的来说，持有此锁可以防止dentry被重命名或取消链接。特别是，它的父节点（d_parent）和名称（d_name）不能被更改，并且它不能从dentry哈希表中移除。

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.

在查找目录中的名称时，REF-walk会在哈希表中找到每个候选的dentry，并检查父节点和名称是否正确。因此，在搜索缓存时，它不会锁定父节点，只会锁定子节点。

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.

在查找给定名称的父节点（处理“..”）时，REF-walk可以使用d_lock来获取对d_parent的稳定引用，但它首先尝试一种更轻量级的方法。如dget_parent()中所示，如果可以在父节点上获得引用，并且随后可以看到d_parent没有发生变化，那么实际上不需要对子节点进行锁定。

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.

在给定的目录中查找给定名称涉及从两个值（名称和目录的dentry）计算哈希，访问哈希表中的该槽，并搜索找到的链表。

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.

当dentry被重命名时，名称和父dentry都可能发生变化，因此哈希几乎肯定也会发生变化。这将使dentry移动到哈希表中的不同链。如果文件名搜索恰好正在查看以这种方式移动的dentry，它可能会继续沿着错误的链进行搜索，从而错过正确链的一部分。

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.

名称查找过程（d_lookup()）不会尝试阻止这种情况发生，而只是检测它何时发生。rename_lock是一个序列锁，每当重命名任何dentry时都会更新它。如果d_lookup发现在无法成功扫描哈希表中的链时发生了重命名，它只会再次尝试。

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.

rename_lock还用于检测和防御解析“..”时针对LOOKUP_BENEATH和LOOKUP_IN_ROOT的潜在攻击（其中父目录移动到根目录之外，绕过path_equal()检查）。如果在查找过程中更新了rename_lock并且路径遇到“..”，则发生了潜在攻击，handle_dots()将以-EAGAIN退出。

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().

i_rwsem是一个读/写信号量，用于串行化对特定目录的所有更改。这确保了例如unlink()和rename()不能同时发生。它还在文件系统被要求查找当前不在dcache中的名称，或者在使用readdir()检索目录中的条目列表时，保持目录的稳定性。

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.

这与d_lock具有互补的作用：目录上的i_rwsem保护该目录中的所有名称，而名称上的d_lock仅保护目录中的一个名称。对dcache的大多数更改都会在相关目录索引节点上持有i_rwsem，并在更改发生时短暂地获取一个或多个dentry上的d_lock。一个例外是当空闲的dentry由于内存压力而从dcache中删除时。这使用d_lock，但i_rwsem不起作用。

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.

该信号量以两种不同的方式影响路径名查找。首先，它防止在查找目录中的名称期间进行更改。walk_component()首先使用lookup_fast()，它再次使用只有d_lock锁定的缓存来检查名称是否存在。如果找不到该名称，那么walk_component()将回退到lookup_slow()，它在i_rwsem上获取一个共享锁，再次检查名称是否存在于缓存中，然后调用文件系统以获得确定的答案。无论结果如何，都将向缓存中添加一个新的dentry。

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.

其次，当路径名查找到达最后一个组件时，有时需要在执行最后一次查找之前对i_rwsem进行独占锁定，以实现所需的排他性。路径查找选择是否获取i_rwsem的方式是后续部分中解决的问题之一。

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).

如果两个线程同时尝试查找相同的尚未在dcache中的名称，则对i_rwsem的共享锁不会阻止它们都添加具有相同名称的新dentry。由于这会导致混乱，使用了额外的互锁级别，基于一个辅助哈希表（in_lookup_hashtable）和一个每个dentry的标志位（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.

要在仅持有i_rwsem的共享锁的情况下向缓存中添加一个新的dentry，线程必须调用d_alloc_parallel()。这将分配一个dentry，将所需的名称和父目录存储在其中，检查主哈希表或辅助哈希表中是否已经存在匹配的dentry，如果没有，则将新分配的dentry存储在辅助哈希表中，并设置DCACHE_PAR_LOOKUP。

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.

如果在主哈希表中找到了匹配的dentry，则返回该dentry，并且调用者可以知道它在与其他线程添加该条目的竞争中失败了。如果在任何缓存中都找不到匹配的dentry，则返回新分配的dentry，并且调用者可以通过DCACHE_PAR_LOOKUP的存在来检测到这一点。在这种情况下，它知道它赢得了任何竞争，现在负责要求文件系统执行查找并找到匹配的索引节点。完成查找后，必须调用d_lookup_done()，该函数清除标志并进行其他一些清理工作，包括从辅助哈希表中删除dentry - 它通常已经添加到主哈希表中。请注意，在调用d_alloc_parallel()时传递了一个struct waitqueue_head，并且必须在此waitqueue_head仍然在作用域内时调用d_lookup_done()。

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.

如果在辅助哈希表中找到了匹配的dentry，d_alloc_parallel()需要做更多的工作。它首先等待DCACHE_PAR_LOOKUP被清除，使用传递给赢得竞争的d_alloc_parallel()实例的等待队列，并在调用d_lookup_done()时唤醒该队列。然后，它检查是否已将dentry添加到主哈希表中。如果已经添加到主哈希表中，则返回dentry，并且调用者只会看到它在任何竞争中失败。如果尚未将其添加到主哈希表中，则最有可能的解释是使用d_splice_alias()添加了其他某个dentry。无论如何，d_alloc_parallel()会从头开始重复所有查找，并且通常会从主哈希表中返回某个结果。

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.

mnt_count是“mount”结构上的每个CPU的引用计数器。这里的每个CPU表示增加计数是廉价的，因为它只使用CPU本地内存，但是检查计数是否为零是昂贵的，因为它需要与每个CPU进行检查。通过mnt_count引用可以防止挂载结构在常规卸载操作的结果中消失，但不能防止“惰性”卸载。因此，持有mnt_count并不能确保挂载点保留在命名空间中，特别是不能稳定挂载点的链接。然而，它确保挂载数据结构保持一致，并提供对挂载文件系统的根dentry的引用。因此，通过->mnt_count进行引用可以提供对已挂载dentry的稳定引用，但不能提供对挂载点dentry的引用。

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.

mount_lock是一个全局的序列锁，有点像rename_lock。它可以用来检查是否对任何挂载点进行了更改。

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.

在向下遍历树（远离根目录）时，当穿越一个挂载点时会使用该锁来检查穿越是否安全。也就是说，首先读取序列锁中的值，然后找到当前目录上挂载的挂载点（如果有的话），并增加mnt_count。最后，将mount_lock中的值与旧值进行比较。如果没有更改，那么穿越是安全的。如果有更改，mnt_count将递减，并重新尝试整个过程。

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.

当通过跟随“..”链接向上遍历树（朝向根目录）时，需要更加小心。在这种情况下，序列锁（包含计数器和自旋锁）被完全锁定，以防止在向上步进时对任何挂载点进行更改。这种锁定是为了稳定与挂载的dentry之间的链接，而挂载本身的引用计数并不能确保这一点。

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.

mount_lock还用于检测和防御解析“..”时针对LOOKUP_BENEATH和LOOKUP_IN_ROOT的潜在攻击（其中父目录移动到根目录之外，绕过了path_equal()检查）。如果在查找过程中更新了mount_lock，并且路径遇到“..”，则发生了潜在攻击，handle_dots()将以-EAGAIN退出。

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.

最后，全局（但非常轻量级）的RCU读锁会不时地被持有，以确保某些数据结构不会意外释放。

In particular it is held while scanning chains in the dcache hash table, and the mount point hash table.

特别是在扫描dcache哈希表和挂载点哈希表中的链表时会持有该锁。

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):

在走过路径的整个过程中，当前状态存储在一个名为nameidata的结构体中，"namei"是该函数的传统名称，它可以将一个"名称"转换为一个"inode"，这个名称可以追溯到第一版的Unix。struct nameidata包含以下字段（以及其他字段）：

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.

路径包含一个struct vfsmount（嵌入在struct mount中）和一个struct dentry。它们一起记录了路径的当前状态。它们最初指向起始点（当前工作目录、根目录或由文件描述符标识的其他目录），并在每一步更新。始终保持通过d_lockref和mnt_count的引用。

struct qstr last

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

这是一个字符串，加上一个长度（即非空终止符），它是路径名中的"下一个"组成部分。

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.

它可以是LAST_NORM、LAST_ROOT、LAST_DOT或LAST_DOTDOT中的一个。如果类型是LAST_NORM，则最后一个字段才有效。

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.

这是用来保存文件系统有效根目录的引用。通常情况下，这个引用是不需要的，所以这个字段只在第一次使用时被赋值，或者在请求非标准根目录时被赋值。在nameidata中保持一个引用可以确保整个路径遍历过程中只有一个根目录生效，即使它与chroot()系统调用同时进行。

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).

需要注意的是，在LOOKUP_IN_ROOT或LOOKUP_BENEATH的情况下，有效根目录变为传递给openat2()的目录文件描述符（这暴露了这些LOOKUP_标志）。

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().

当满足以下两个条件之一时，需要使用根目录：（1）路径名或符号链接以“'/'”开头，或者（2）正在处理一个“..”组件，因为“..”从根目录开始必须始终保持在根目录。通常情况下，使用的值是调用进程的当前根目录。可以提供替代根目录，例如当sysctl()调用file_open_root()时，以及当NFSv4或Btrfs调用mount_subtree()时。在每种情况下，都在文件系统的一个非常特定的部分查找路径名，并且查找不能允许逃离该子树。它的工作方式有点像本地的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:

忽略符号链接的处理，我们现在可以描述一下“link_path_walk()”函数，它处理除最后一个组件之外的所有查找操作：

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.

给定一个路径（名称）和一个nameidata结构（nd），检查当前目录是否具有执行权限，然后将名称移动到下一个组件，同时更新last_type和last。如果这是最后一个组件，则返回，否则调用walk_component()并从头开始重复。

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.

walk_component()更简单。如果组件是LAST_DOTS，它调用handle_dots()来执行先前描述的必要锁定操作。如果它找到一个LAST_NORM组件，首先调用“lookup_fast()”只在dcache中查找，但如果是这种类型的文件系统，它将要求文件系统重新验证结果。如果没有得到好的结果，它调用“lookup_slow()”，它获取i_rwsem，重新检查缓存，然后要求文件系统找到一个确定的答案。

As the last step of walk_component(), step_into() will be called either directly from walk_component() or from handle_dots(). It calls handle_mounts(), to check and handle mount points, in which a new struct path is created 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. Then if there is a symbolic link, step_into() calls pick_link() to deal with it, otherwise it installs the new struct path in the struct nameidata, and drops the unneeded references.

在walk_component()的最后一步，将直接从walk_component()或handle_dots()调用step_into()。它调用handle_mounts()来检查和处理挂载点，在此过程中创建一个新的struct path，其中包含对新的dentry的计数引用和对新的vfsmount的引用，只有在它与先前的vfsmount不同时才计数。然后，如果存在符号链接，step_into()调用pick_link()来处理它，否则将新的struct path安装在struct nameidata中，并丢弃不需要的引用。

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.

在放弃对先前dentry的引用之前获取对新dentry的引用的这种“逐个传递”的顺序可能显而易见，但值得指出的是，这样我们将能够在“RCU-walk”版本中识别出它的类似之处。

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() and path_openat() each of which handles the differing requirements of different system calls.

link_path_walk()函数只会遍历到设置nd->last和nd->last_type以指向路径的最后一个组件。它不会在最后一次调用时调用walk_component()函数。处理最后一个组件的工作留给调用者来解决。这些调用者分别是path_lookupat()、path_parentat()和path_openat()，它们处理不同系统调用的不同要求。

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_parentat()显然是最简单的 - 它只是在link_path_walk()周围包装了一点点的管理工作，并将父目录和最后一个组件返回给调用者。调用者可能是为了创建一个名称（通过filename_create()）或者删除或重命名一个名称（在这种情况下使用user_path_parent()）。他们将使用i_rwsem来排除其他更改，同时验证并执行他们的操作。

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. It is worth noting that when flag LOOKUP_MOUNTPOINT is set, path_lookupat() will unset LOOKUP_JUMPED in nameidata so that in the subsequent path traversal d_weak_revalidate() won't be called. This is important when unmounting a filesystem that is inaccessible, such as one provided by a dead NFS server.

path_lookupat()几乎同样简单 - 当需要一个现有对象（例如通过stat()或chmod()）时使用它。它基本上只是通过调用lookup_last()在最后一个组件上调用walk_component()。path_lookupat()只返回最后的dentry。值得注意的是，当设置了LOOKUP_MOUNTPOINT标志时，path_lookupat()会在nameidata中取消LOOKUP_JUMPED，以便在后续的路径遍历中不会调用d_weak_revalidate()。这在卸载一个无法访问的文件系统时非常重要，比如由一个已经停止运行的NFS服务器提供的文件系统。

Finally path_openat() is used for the open() system call; it contains, in support functions starting with "open_last_lookups()", 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. "open_last_lookups()" will sometimes, but not always, take i_rwsem, depending on what it finds.

最后，path_openat()用于open()系统调用；它包含了以"open_last_lookups()"开头的支持函数，用于处理O_CREAT（带或不带O_EXCL）、最后的"/"字符和尾部符号链接的不同细节。我们将在本系列的最后一部分重新讨论这个问题，重点关注这些符号链接。"open_last_lookups()"有时会获取i_rwsem，但并不总是，这取决于它找到了什么。

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.

这些函数中的每一个，或者调用它们的函数，都需要注意最后一个组件可能不是LAST_NORM的可能性。如果查找的目标是创建某个东西，那么除了LAST_NORM之外的任何last_type值都会导致错误。例如，如果path_parentat()报告LAST_DOTDOT，那么调用者将不会尝试创建该名称。它们还通过测试last.name[last.len]来检查是否有尾部斜杠。如果最后一个组件之后还有任何字符，那么它必须是一个尾部斜杠。

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.

除了符号链接外，“REF-walk”过程还有两个部分尚未涵盖。一个是处理过期缓存条目，另一个是自动挂载。

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.

在需要的文件系统上，查找例程将调用->d_revalidate() dentry方法以确保缓存的信息是最新的。这通常会确认有效性或从服务器更新一些细节。在某些情况下，它可能会发现路径上方发生了变化，之前认为有效的某些内容实际上并非如此。当发生这种情况时，整个路径的查找将被中止，并使用“LOOKUP_REVAL”标志重试。这会强制重新验证更加彻底。我们将在下一篇文章中详细了解此重试过程的更多细节。

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 - how it works 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. 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(), do_filp_open(), and do_file_open_root(). These four correspond roughly to the three 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.

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() 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.

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 (MAXSYMLINKS) 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.

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->get_link() inode method. This is called both in RCU-walk and REF-walk. In RCU-walk the dentry* argument is NULL, ->get_link() can return -ECHILD to drop out of RCU-walk. Much like the i_op->permission() method we looked at previously, ->get_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. A callback struct delayed_called will be passed to ->get_link(): file systems can set their own put_link function and argument through set_delayed_call(). Later on, when VFS wants to put link, it will call do_delayed_call() to invoke that callback function with the argument.

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 previous path

• the const char * to provide a reference to the to previous name

• the seq to allow the path to be safely switched from RCU-walk to REF-walk

• the struct delayed_call for later invocation.

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() calls pick_link() via step_into() which returns 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 three flags: WALK_NOFOLLOW which forbids it from following a symlink if it finds one, WALK_MORE which indicates that it is yet too early to release the current symlink, and WALK_TRAILING which indicates that it is on the final component of the lookup, so we will check userspace flag LOOKUP_FOLLOW to decide whether follow it when it is a symlink and call may_follow_link() to check if we have privilege to follow it.

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 pick_link() returning NULL.

The more obvious case is a symlink to "/". All symlinks starting with "/" are detected in pick_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 ->get_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. ->get_link() then returns NULL. Again there is no final component and pick_link() returns NULL.

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 relevant callers of link_path_walk(), such as path_lookupat(), path_openat() using a loop that calls link_path_walk(), and then handles the final component by calling open_last_lookups() or lookup_last(). If it is a symlink that needs to be followed, open_last_lookups() or lookup_last() will set things up properly and return the path so that 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.

Of the various functions that examine the final component, open_last_lookups() is the most interesting as it works in tandem with do_open() for opening a file. Part of open_last_lookups() runs with i_rwsem held and this part is in a separate function: lookup_open().

Explaining open_last_lookups() and do_open() 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_open() is used after open_last_lookup() 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 i_op->lookup() and i_op->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. If it's in RCU-walk -ECHILD will be returned otherwise -ESTALE is returned. When -ESTALE is returned, the caller may retry with LOOKUP_REVAL flag set.

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: lookup_last() or open_last_lookup() returns a non NULL value, and link_path_walk() 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.

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.

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.

ND_ROOT_PRESET 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.

ND_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 ND_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.

标签：dentry,Pathname,RCU,component,walk,lookup,path,chatgpt
From： https://www.cnblogs.com/pengdonglin137/p/17517919.html