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A memory cache of device blocks indexed by device identity and block number. A hit avoids slow device access; a miss obtains a buffer and reads the block. Dirty buffers carry modified data that must be written before reuse, and pinned buffers cannot be evicted while an operation uses them. Deferred writes improve throughput but need explicit durability handling.

A buffer cache can hold both file data blocks and storage containing file metadata, such as inodes and block allocation bitmaps. Reusing a buffered device block saves repeated input-output; retaining several modifications before buffer cache writeback can combine them into one transfer. Section 40.7 of Operating Systems: Three Easy Pieces by Remzi H. Arpaci-Dusseau and others explains these benefits for file systems.^[ref]

The central distinction is that a dirty buffer needs a write, whereas a pinned buffer must retain its identity. A clean buffer can still be pinned, and a dirty buffer can have no client references. Neither releasing a reference nor obtaining a cache hit establishes file-system transaction durability.

An association between a name in a file-system directory and a file. In the toy file system, the pairs are (a, 42) and (b, 42): each name resolves to the same inode. The name belongs to the directory entry, rather than being the single authoritative name stored in that inode.

A storage unit holding part of a file's contents. The toy inode references two four-byte file data blocks for six bytes of contents: allocated capacity is $2\cdot 4=8$ bytes, while logical size is $4+2=6$ bytes. This counts only the two file data blocks, excluding storage for the inode and directory entries.

A nonnegative integer identifying an open access in one operating-system process. The program passes its value to operations such as reading and closing; the operating-system kernel interprets it through that operating-system process's descriptor table. In the toy example, table entry 3 refers to an open file description for inode 42. Another operating-system process's entry 3 can reach a different file, and a closed file descriptor value can be reused. ^[ref]

Information describing a file, such as type, owner, access permissions, size, timestamps and hard link count. In the toy layout, the inode carries these attributes; the directory entries carry the names. The distinction matters because renaming a file need not copy its contents.

A file system organizes files, their file metadata, and the names through which they can be found. In the Unix model used here, a file-system directory contains directory entries associating names with inodes; an inode describes the underlying file and how to locate its contents. The file-system namespace answers "what does this name reach?", while the inode answers "which file is it?". These are distinct responsibilities, even though they cooperate on every lookup.

For a computed toy example, let one file-system directory contain names a and b, both referring to inode 42 in the same file system. That inode has size 6 bytes and two hard links. Its direct references lead to four-byte file data blocks 7 and 9: block 7 holds hell, while block 9 holds o followed by a newline, with two unused slots. Reading either name therefore obtains the same six-byte contents hello followed by a newline. There is one file with two names, not two copies. This direct-reference layout and tiny block size are teaching choices, not requirements on every file system.

The separation is useful because a rename within this file system, from a to a previously unused name c, can change a directory entry without copying the six bytes of contents. The remaining name b still reaches inode 42. A successful write changing the first byte through either name changes what both names subsequently read, to jello followed by a newline. This sharing is the attraction of hard links, and also why a hard link is not an independent backup. A program that replaces a named file with a different file can instead make the names diverge. ^[ref]

A file containing directory entries. Each entry supplies a name in that file-system directory; two entries in the same file-system directory cannot have the same name. In an inode-based file system, a file-system directory has its own inode, and its contents describe names of other objects. ^[ref]

File-system journaling makes a group of related file system updates recoverable after a crash by recording them in a journal before changing their final storage locations. The useful promise is file-system crash consistency: after journal replay, a protected file-system transaction has taken effect completely or has not taken effect. File-system transaction durability adds a different promise: an acknowledged update survives the crash. The protected group can contain only file metadata, or both file metadata and file data blocks, depending on the file-system journaling mode.

Consider a tiny file containing ABCD in four-byte file data block 7. Its inode records size 4 and a reference to block 7. Appending EFGH in previously free block 9 requires three new stored records: $I$, the inode with size 8 and references to blocks 7 and 9; $B$, the block allocation bitmap marking block 9 allocated; and $D$, file data block 9 containing EFGH. An inode reference to a block still marked free lets a later allocation overwrite live contents; allocating a block without recording its owner can leak space. Even mutually consistent file metadata can expose stale contents if $D$ has not reached storage. Chapter 42 of Crash Consistency: FSCK and Journaling (section 42.1, pp. 2–4) develops this three-record append problem; the names, block numbers and tiny contents here are original teaching choices.

The first diagram protects all three records with write-ahead logging. Each journal payload record identifies a final destination and its complete new contents. A storage flush makes the complete journal payload durable before a separate journal commit record is written and made durable. Only then may a journal checkpoint copy the records to their final locations. The journal remains available until those copies have themselves passed a storage flush; it can then be retired and its space reused. Journal replay repairs an interrupted journal checkpoint using the retained new contents.^[ref]

These are persistence dependencies, not a claim that issuing writes in this order is sufficient. A device can report completion while its volatile memory cache still holds the only copy. The Linux kernel documents both flushing earlier completed requests and making completion of a particular request depend on persistence. The timeline uses separate storage flushes to make the dependencies explicit; an implementation can use equivalent device-supported operations. A commit write can persist before its storage flush returns: recovery uses the actual durable marker, while acknowledgment waits for successful completion.^[ref]

The appeal is that recovery knows exactly which updates to finish. It does not have to infer a user's intended contents from contradictory inodes and block allocation bitmaps. The cost includes writing the journal payload as well as its final copies and waiting for persistence boundaries. The journal records a recent transition, not a historical collection of recoverable files; file-system journaling is not a file-system backup.