1099 lines
37 KiB
Plaintext
1099 lines
37 KiB
Plaintext
reftable
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--------
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Overview
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~~~~~~~~
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Problem statement
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^^^^^^^^^^^^^^^^^
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Some repositories contain a lot of references (e.g. android at 866k,
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rails at 31k). The existing packed-refs format takes up a lot of space
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(e.g. 62M), and does not scale with additional references. Lookup of a
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single reference requires linearly scanning the file.
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Atomic pushes modifying multiple references require copying the entire
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packed-refs file, which can be a considerable amount of data moved
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(e.g. 62M in, 62M out) for even small transactions (2 refs modified).
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Repositories with many loose references occupy a large number of disk
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blocks from the local file system, as each reference is its own file
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storing 41 bytes (and another file for the corresponding reflog). This
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negatively affects the number of inodes available when a large number of
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repositories are stored on the same filesystem. Readers can be penalized
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due to the larger number of syscalls required to traverse and read the
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`$GIT_DIR/refs` directory.
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Objectives
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^^^^^^^^^^
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* Near constant time lookup for any single reference, even when the
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repository is cold and not in process or kernel cache.
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* Near constant time verification if an object name is referred to by at least
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one reference (for allow-tip-sha1-in-want).
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* Efficient enumeration of an entire namespace, such as `refs/tags/`.
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* Support atomic push with `O(size_of_update)` operations.
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* Combine reflog storage with ref storage for small transactions.
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* Separate reflog storage for base refs and historical logs.
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Description
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^^^^^^^^^^^
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A reftable file is a portable binary file format customized for
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reference storage. References are sorted, enabling linear scans, binary
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search lookup, and range scans.
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Storage in the file is organized into variable sized blocks. Prefix
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compression is used within a single block to reduce disk space. Block
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size and alignment are tunable by the writer.
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Performance
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^^^^^^^^^^^
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Space used, packed-refs vs. reftable:
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[cols=",>,>,>,>,>",options="header",]
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|===============================================================
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|repository |packed-refs |reftable |% original |avg ref |avg obj
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|android |62.2 M |36.1 M |58.0% |33 bytes |5 bytes
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|rails |1.8 M |1.1 M |57.7% |29 bytes |4 bytes
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|git |78.7 K |48.1 K |61.0% |50 bytes |4 bytes
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|git (heads) |332 b |269 b |81.0% |33 bytes |0 bytes
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|===============================================================
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Scan (read 866k refs), by reference name lookup (single ref from 866k
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refs), and by SHA-1 lookup (refs with that SHA-1, from 866k refs):
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[cols=",>,>,>,>",options="header",]
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|=========================================================
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|format |cache |scan |by name |by SHA-1
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|packed-refs |cold |402 ms |409,660.1 usec |412,535.8 usec
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|packed-refs |hot | |6,844.6 usec |20,110.1 usec
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|reftable |cold |112 ms |33.9 usec |323.2 usec
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|reftable |hot | |20.2 usec |320.8 usec
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|=========================================================
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Space used for 149,932 log entries for 43,061 refs, reflog vs. reftable:
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[cols=",>,>",options="header",]
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|================================
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|format |size |avg entry
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|$GIT_DIR/logs |173 M |1209 bytes
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|reftable |5 M |37 bytes
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|================================
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Details
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~~~~~~~
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Peeling
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^^^^^^^
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References stored in a reftable are peeled, a record for an annotated
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(or signed) tag records both the tag object, and the object it refers
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to. This is analogous to storage in the packed-refs format.
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Reference name encoding
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^^^^^^^^^^^^^^^^^^^^^^^
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Reference names are an uninterpreted sequence of bytes that must pass
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linkgit:git-check-ref-format[1] as a valid reference name.
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Key unicity
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^^^^^^^^^^^
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Each entry must have a unique key; repeated keys are disallowed.
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Network byte order
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^^^^^^^^^^^^^^^^^^
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All multi-byte, fixed width fields are in network byte order.
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Varint encoding
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^^^^^^^^^^^^^^^
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Varint encoding is identical to the ofs-delta encoding method used
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within pack files.
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Decoder works as follows:
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....
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val = buf[ptr] & 0x7f
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while (buf[ptr] & 0x80) {
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ptr++
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val = ((val + 1) << 7) | (buf[ptr] & 0x7f)
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}
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....
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Ordering
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^^^^^^^^
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Blocks are lexicographically ordered by their first reference.
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Directory/file conflicts
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^^^^^^^^^^^^^^^^^^^^^^^^
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The reftable format accepts both `refs/heads/foo` and
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`refs/heads/foo/bar` as distinct references.
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This property is useful for retaining log records in reftable, but may
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confuse versions of Git using `$GIT_DIR/refs` directory tree to maintain
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references. Users of reftable may choose to continue to reject `foo` and
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`foo/bar` type conflicts to prevent problems for peers.
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File format
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~~~~~~~~~~~
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Structure
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^^^^^^^^^
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A reftable file has the following high-level structure:
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....
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first_block {
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header
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first_ref_block
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}
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ref_block*
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ref_index*
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obj_block*
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obj_index*
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log_block*
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log_index*
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footer
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....
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A log-only file omits the `ref_block`, `ref_index`, `obj_block` and
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`obj_index` sections, containing only the file header and log block:
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....
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first_block {
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header
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}
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log_block*
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log_index*
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footer
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....
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In a log-only file, the first log block immediately follows the file
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header, without padding to block alignment.
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Block size
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^^^^^^^^^^
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The file's block size is arbitrarily determined by the writer, and does
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not have to be a power of 2. The block size must be larger than the
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longest reference name or log entry used in the repository, as
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references cannot span blocks.
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Powers of two that are friendly to the virtual memory system or
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filesystem (such as 4k or 8k) are recommended. Larger sizes (64k) can
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yield better compression, with a possible increased cost incurred by
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readers during access.
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The largest block size is `16777215` bytes (15.99 MiB).
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Block alignment
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^^^^^^^^^^^^^^^
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Writers may choose to align blocks at multiples of the block size by
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including `padding` filled with NUL bytes at the end of a block to round
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out to the chosen alignment. When alignment is used, writers must
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specify the alignment with the file header's `block_size` field.
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Block alignment is not required by the file format. Unaligned files must
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set `block_size = 0` in the file header, and omit `padding`. Unaligned
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files with more than one ref block must include the link:#Ref-index[ref
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index] to support fast lookup. Readers must be able to read both aligned
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and non-aligned files.
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Very small files (e.g. a single ref block) may omit `padding` and the ref
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index to reduce total file size.
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Header (version 1)
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^^^^^^^^^^^^^^^^^^
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A 24-byte header appears at the beginning of the file:
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....
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'REFT'
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uint8( version_number = 1 )
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uint24( block_size )
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uint64( min_update_index )
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uint64( max_update_index )
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....
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Aligned files must specify `block_size` to configure readers with the
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expected block alignment. Unaligned files must set `block_size = 0`.
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The `min_update_index` and `max_update_index` describe bounds for the
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`update_index` field of all log records in this file. When reftables are
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used in a stack for link:#Update-transactions[transactions], these
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fields can order the files such that the prior file's
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`max_update_index + 1` is the next file's `min_update_index`.
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Header (version 2)
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^^^^^^^^^^^^^^^^^^
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A 28-byte header appears at the beginning of the file:
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....
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'REFT'
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uint8( version_number = 2 )
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uint24( block_size )
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uint64( min_update_index )
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uint64( max_update_index )
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uint32( hash_id )
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....
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The header is identical to `version_number=1`, with the 4-byte hash ID
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("sha1" for SHA1 and "s256" for SHA-256) appended to the header.
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For maximum backward compatibility, it is recommended to use version 1 when
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writing SHA1 reftables.
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First ref block
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^^^^^^^^^^^^^^^
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The first ref block shares the same block as the file header, and is 24
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bytes smaller than all other blocks in the file. The first block
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immediately begins after the file header, at position 24.
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If the first block is a log block (a log-only file), its block header
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begins immediately at position 24.
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Ref block format
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^^^^^^^^^^^^^^^^
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A ref block is written as:
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....
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'r'
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uint24( block_len )
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ref_record+
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uint24( restart_offset )+
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uint16( restart_count )
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padding?
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....
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Blocks begin with `block_type = 'r'` and a 3-byte `block_len` which
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encodes the number of bytes in the block up to, but not including the
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optional `padding`. This is always less than or equal to the file's
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block size. In the first ref block, `block_len` includes 24 bytes for
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the file header.
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The 2-byte `restart_count` stores the number of entries in the
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`restart_offset` list, which must not be empty. Readers can use
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`restart_count` to binary search between restarts before starting a
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linear scan.
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Exactly `restart_count` 3-byte `restart_offset` values precede the
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`restart_count`. Offsets are relative to the start of the block and
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refer to the first byte of any `ref_record` whose name has not been
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prefix compressed. Entries in the `restart_offset` list must be sorted,
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ascending. Readers can start linear scans from any of these records.
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A variable number of `ref_record` fill the middle of the block,
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describing reference names and values. The format is described below.
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As the first ref block shares the first file block with the file header,
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all `restart_offset` in the first block are relative to the start of the
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file (position 0), and include the file header. This forces the first
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`restart_offset` to be `28`.
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ref record
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++++++++++
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A `ref_record` describes a single reference, storing both the name and
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its value(s). Records are formatted as:
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....
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varint( prefix_length )
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varint( (suffix_length << 3) | value_type )
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suffix
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varint( update_index_delta )
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value?
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....
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The `prefix_length` field specifies how many leading bytes of the prior
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reference record's name should be copied to obtain this reference's
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name. This must be 0 for the first reference in any block, and also must
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be 0 for any `ref_record` whose offset is listed in the `restart_offset`
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table at the end of the block.
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Recovering a reference name from any `ref_record` is a simple concat:
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....
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this_name = prior_name[0..prefix_length] + suffix
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....
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The `suffix_length` value provides the number of bytes available in
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`suffix` to copy from `suffix` to complete the reference name.
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The `update_index` that last modified the reference can be obtained by
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adding `update_index_delta` to the `min_update_index` from the file
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header: `min_update_index + update_index_delta`.
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The `value` follows. Its format is determined by `value_type`, one of
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the following:
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* `0x0`: deletion; no value data (see transactions, below)
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* `0x1`: one object name; value of the ref
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* `0x2`: two object names; value of the ref, peeled target
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* `0x3`: symbolic reference: `varint( target_len ) target`
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Symbolic references use `0x3`, followed by the complete name of the
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reference target. No compression is applied to the target name.
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Types `0x4..0x7` are reserved for future use.
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Ref index
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^^^^^^^^^
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The ref index stores the name of the last reference from every ref block
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in the file, enabling reduced disk seeks for lookups. Any reference can
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be found by searching the index, identifying the containing block, and
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searching within that block.
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The index may be organized into a multi-level index, where the 1st level
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index block points to additional ref index blocks (2nd level), which may
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in turn point to either additional index blocks (e.g. 3rd level) or ref
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blocks (leaf level). Disk reads required to access a ref go up with
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higher index levels. Multi-level indexes may be required to ensure no
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single index block exceeds the file format's max block size of
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`16777215` bytes (15.99 MiB). To achieve constant O(1) disk seeks for
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lookups the index must be a single level, which is permitted to exceed
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the file's configured block size, but not the format's max block size of
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15.99 MiB.
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If present, the ref index block(s) appears after the last ref block.
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If there are at least 4 ref blocks, a ref index block should be written
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to improve lookup times. Cold reads using the index require 2 disk reads
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(read index, read block), and binary searching < 4 blocks also requires
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<= 2 reads. Omitting the index block from smaller files saves space.
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If the file is unaligned and contains more than one ref block, the ref
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index must be written.
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Index block format:
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....
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'i'
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uint24( block_len )
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index_record+
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uint24( restart_offset )+
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uint16( restart_count )
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padding?
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....
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The index blocks begin with `block_type = 'i'` and a 3-byte `block_len`
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which encodes the number of bytes in the block, up to but not including
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the optional `padding`.
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The `restart_offset` and `restart_count` fields are identical in format,
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meaning and usage as in ref blocks.
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To reduce the number of reads required for random access in very large
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files the index block may be larger than other blocks. However, readers
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must hold the entire index in memory to benefit from this, so it's a
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time-space tradeoff in both file size and reader memory.
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Increasing the file's block size decreases the index size. Alternatively
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a multi-level index may be used, keeping index blocks within the file's
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block size, but increasing the number of blocks that need to be
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accessed.
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index record
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++++++++++++
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An index record describes the last entry in another block. Index records
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are written as:
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....
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varint( prefix_length )
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varint( (suffix_length << 3) | 0 )
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suffix
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varint( block_position )
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....
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Index records use prefix compression exactly like `ref_record`.
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Index records store `block_position` after the suffix, specifying the
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absolute position in bytes (from the start of the file) of the block
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that ends with this reference. Readers can seek to `block_position` to
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begin reading the block header.
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Readers must examine the block header at `block_position` to determine
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if the next block is another level index block, or the leaf-level ref
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block.
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Reading the index
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+++++++++++++++++
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Readers loading the ref index must first read the footer (below) to
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obtain `ref_index_position`. If not present, the position will be 0. The
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`ref_index_position` is for the 1st level root of the ref index.
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Obj block format
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^^^^^^^^^^^^^^^^
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Object blocks are optional. Writers may choose to omit object blocks,
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especially if readers will not use the object name to ref mapping.
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Object blocks use unique, abbreviated 2-31 byte object name keys, mapping to
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ref blocks containing references pointing to that object directly, or as
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the peeled value of an annotated tag. Like ref blocks, object blocks use
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the file's standard block size. The abbreviation length is available in
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the footer as `obj_id_len`.
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To save space in small files, object blocks may be omitted if the ref
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index is not present, as brute force search will only need to read a few
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ref blocks. When missing, readers should brute force a linear search of
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all references to lookup by object name.
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An object block is written as:
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....
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'o'
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uint24( block_len )
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obj_record+
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uint24( restart_offset )+
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uint16( restart_count )
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padding?
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....
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Fields are identical to ref block. Binary search using the restart table
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works the same as in reference blocks.
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Because object names are abbreviated by writers to the shortest unique
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abbreviation within the reftable, obj key lengths have a variable length. Their
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length must be at least 2 bytes. Readers must compare only for common prefix
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match within an obj block or obj index.
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obj record
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++++++++++
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An `obj_record` describes a single object abbreviation, and the blocks
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containing references using that unique abbreviation:
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....
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varint( prefix_length )
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varint( (suffix_length << 3) | cnt_3 )
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suffix
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varint( cnt_large )?
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varint( position_delta )*
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....
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Like in reference blocks, abbreviations are prefix compressed within an
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obj block. On large reftables with many unique objects, higher block
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sizes (64k), and higher restart interval (128), a `prefix_length` of 2
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or 3 and `suffix_length` of 3 may be common in obj records (unique
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abbreviation of 5-6 raw bytes, 10-12 hex digits).
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Each record contains `position_count` number of positions for matching
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ref blocks. For 1-7 positions the count is stored in `cnt_3`. When
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`cnt_3 = 0` the actual count follows in a varint, `cnt_large`.
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The use of `cnt_3` bets most objects are pointed to by only a single
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reference, some may be pointed to by a couple of references, and very
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few (if any) are pointed to by more than 7 references.
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A special case exists when `cnt_3 = 0` and `cnt_large = 0`: there are no
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`position_delta`, but at least one reference starts with this
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abbreviation. A reader that needs exact reference names must scan all
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references to find which specific references have the desired object.
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Writers should use this format when the `position_delta` list would have
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overflowed the file's block size due to a high number of references
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pointing to the same object.
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The first `position_delta` is the position from the start of the file.
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Additional `position_delta` entries are sorted ascending and relative to
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the prior entry, e.g. a reader would perform:
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....
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pos = position_delta[0]
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prior = pos
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for (j = 1; j < position_count; j++) {
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pos = prior + position_delta[j]
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prior = pos
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}
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....
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With a position in hand, a reader must linearly scan the ref block,
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starting from the first `ref_record`, testing each reference's object names
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(for `value_type = 0x1` or `0x2`) for full equality. Faster searching by
|
|
object name within a single ref block is not supported by the reftable format.
|
|
Smaller block sizes reduce the number of candidates this step must
|
|
consider.
|
|
|
|
Obj index
|
|
^^^^^^^^^
|
|
|
|
The obj index stores the abbreviation from the last entry for every obj
|
|
block in the file, enabling reduced disk seeks for all lookups. It is
|
|
formatted exactly the same as the ref index, but refers to obj blocks.
|
|
|
|
The obj index should be present if obj blocks are present, as obj blocks
|
|
should only be written in larger files.
|
|
|
|
Readers loading the obj index must first read the footer (below) to
|
|
obtain `obj_index_position`. If not present, the position will be 0.
|
|
|
|
Log block format
|
|
^^^^^^^^^^^^^^^^
|
|
|
|
Unlike ref and obj blocks, log blocks are always unaligned.
|
|
|
|
Log blocks are variable in size, and do not match the `block_size`
|
|
specified in the file header or footer. Writers should choose an
|
|
appropriate buffer size to prepare a log block for deflation, such as
|
|
`2 * block_size`.
|
|
|
|
A log block is written as:
|
|
|
|
....
|
|
'g'
|
|
uint24( block_len )
|
|
zlib_deflate {
|
|
log_record+
|
|
uint24( restart_offset )+
|
|
uint16( restart_count )
|
|
}
|
|
....
|
|
|
|
Log blocks look similar to ref blocks, except `block_type = 'g'`.
|
|
|
|
The 4-byte block header is followed by the deflated block contents using
|
|
zlib deflate. The `block_len` in the header is the inflated size
|
|
(including 4-byte block header), and should be used by readers to
|
|
preallocate the inflation output buffer. A log block's `block_len` may
|
|
exceed the file's block size.
|
|
|
|
Offsets within the log block (e.g. `restart_offset`) still include the
|
|
4-byte header. Readers may prefer prefixing the inflation output buffer
|
|
with the 4-byte header.
|
|
|
|
Within the deflate container, a variable number of `log_record` describe
|
|
reference changes. The log record format is described below. See ref
|
|
block format (above) for a description of `restart_offset` and
|
|
`restart_count`.
|
|
|
|
Because log blocks have no alignment or padding between blocks, readers
|
|
must keep track of the bytes consumed by the inflater to know where the
|
|
next log block begins.
|
|
|
|
log record
|
|
++++++++++
|
|
|
|
Log record keys are structured as:
|
|
|
|
....
|
|
ref_name '\0' reverse_int64( update_index )
|
|
....
|
|
|
|
where `update_index` is the unique transaction identifier. The
|
|
`update_index` field must be unique within the scope of a `ref_name`.
|
|
See the update transactions section below for further details.
|
|
|
|
The `reverse_int64` function inverses the value so lexicographical
|
|
ordering the network byte order encoding sorts the more recent records
|
|
with higher `update_index` values first:
|
|
|
|
....
|
|
reverse_int64(int64 t) {
|
|
return 0xffffffffffffffff - t;
|
|
}
|
|
....
|
|
|
|
Log records have a similar starting structure to ref and index records,
|
|
utilizing the same prefix compression scheme applied to the log record
|
|
key described above.
|
|
|
|
....
|
|
varint( prefix_length )
|
|
varint( (suffix_length << 3) | log_type )
|
|
suffix
|
|
log_data {
|
|
old_id
|
|
new_id
|
|
varint( name_length ) name
|
|
varint( email_length ) email
|
|
varint( time_seconds )
|
|
sint16( tz_offset )
|
|
varint( message_length ) message
|
|
}?
|
|
....
|
|
|
|
Log record entries use `log_type` to indicate what follows:
|
|
|
|
* `0x0`: deletion; no log data.
|
|
* `0x1`: standard git reflog data using `log_data` above.
|
|
|
|
The `log_type = 0x0` is mostly useful for `git stash drop`, removing an
|
|
entry from the reflog of `refs/stash` in a transaction file (below),
|
|
without needing to rewrite larger files. Readers reading a stack of
|
|
reflogs must treat this as a deletion.
|
|
|
|
For `log_type = 0x1`, the `log_data` section follows
|
|
linkgit:git-update-ref[1] logging and includes:
|
|
|
|
* two object names (old id, new id)
|
|
* varint string of committer's name
|
|
* varint string of committer's email
|
|
* varint time in seconds since epoch (Jan 1, 1970)
|
|
* 2-byte timezone offset in minutes (signed)
|
|
* varint string of message
|
|
|
|
`tz_offset` is the absolute number of minutes from GMT the committer was
|
|
at the time of the update. For example `GMT-0800` is encoded in reftable
|
|
as `sint16(-480)` and `GMT+0230` is `sint16(150)`.
|
|
|
|
The committer email does not contain `<` or `>`, it's the value normally
|
|
found between the `<>` in a git commit object header.
|
|
|
|
The `message_length` may be 0, in which case there was no message
|
|
supplied for the update.
|
|
|
|
Contrary to traditional reflog (which is a file), renames are encoded as
|
|
a combination of ref deletion and ref creation. A deletion is a log
|
|
record with a zero new_id, and a creation is a log record with a zero old_id.
|
|
|
|
Reading the log
|
|
+++++++++++++++
|
|
|
|
Readers accessing the log must first read the footer (below) to
|
|
determine the `log_position`. The first block of the log begins at
|
|
`log_position` bytes since the start of the file. The `log_position` is
|
|
not block aligned.
|
|
|
|
Importing logs
|
|
++++++++++++++
|
|
|
|
When importing from `$GIT_DIR/logs` writers should globally order all
|
|
log records roughly by timestamp while preserving file order, and assign
|
|
unique, increasing `update_index` values for each log line. Newer log
|
|
records get higher `update_index` values.
|
|
|
|
Although an import may write only a single reftable file, the reftable
|
|
file must span many unique `update_index`, as each log line requires its
|
|
own `update_index` to preserve semantics.
|
|
|
|
Log index
|
|
^^^^^^^^^
|
|
|
|
The log index stores the log key
|
|
(`refname \0 reverse_int64(update_index)`) for the last log record of
|
|
every log block in the file, supporting bounded-time lookup.
|
|
|
|
A log index block must be written if 2 or more log blocks are written to
|
|
the file. If present, the log index appears after the last log block.
|
|
There is no padding used to align the log index to block alignment.
|
|
|
|
Log index format is identical to ref index, except the keys are 9 bytes
|
|
longer to include `'\0'` and the 8-byte `reverse_int64(update_index)`.
|
|
Records use `block_position` to refer to the start of a log block.
|
|
|
|
Reading the index
|
|
+++++++++++++++++
|
|
|
|
Readers loading the log index must first read the footer (below) to
|
|
obtain `log_index_position`. If not present, the position will be 0.
|
|
|
|
Footer
|
|
^^^^^^
|
|
|
|
After the last block of the file, a file footer is written. It begins
|
|
like the file header, but is extended with additional data.
|
|
|
|
....
|
|
HEADER
|
|
|
|
uint64( ref_index_position )
|
|
uint64( (obj_position << 5) | obj_id_len )
|
|
uint64( obj_index_position )
|
|
|
|
uint64( log_position )
|
|
uint64( log_index_position )
|
|
|
|
uint32( CRC-32 of above )
|
|
....
|
|
|
|
If a section is missing (e.g. ref index) the corresponding position
|
|
field (e.g. `ref_index_position`) will be 0.
|
|
|
|
* `obj_position`: byte position for the first obj block.
|
|
* `obj_id_len`: number of bytes used to abbreviate object names in
|
|
obj blocks.
|
|
* `log_position`: byte position for the first log block.
|
|
* `ref_index_position`: byte position for the start of the ref index.
|
|
* `obj_index_position`: byte position for the start of the obj index.
|
|
* `log_index_position`: byte position for the start of the log index.
|
|
|
|
The size of the footer is 68 bytes for version 1, and 72 bytes for
|
|
version 2.
|
|
|
|
Reading the footer
|
|
++++++++++++++++++
|
|
|
|
Readers must first read the file start to determine the version
|
|
number. Then they seek to `file_length - FOOTER_LENGTH` to access the
|
|
footer. A trusted external source (such as `stat(2)`) is necessary to
|
|
obtain `file_length`. When reading the footer, readers must verify:
|
|
|
|
* 4-byte magic is correct
|
|
* 1-byte version number is recognized
|
|
* 4-byte CRC-32 matches the other 64 bytes (including magic, and
|
|
version)
|
|
|
|
Once verified, the other fields of the footer can be accessed.
|
|
|
|
Empty tables
|
|
++++++++++++
|
|
|
|
A reftable may be empty. In this case, the file starts with a header
|
|
and is immediately followed by a footer.
|
|
|
|
Binary search
|
|
^^^^^^^^^^^^^
|
|
|
|
Binary search within a block is supported by the `restart_offset` fields
|
|
at the end of the block. Readers can binary search through the restart
|
|
table to locate between which two restart points the sought reference or
|
|
key should appear.
|
|
|
|
Each record identified by a `restart_offset` stores the complete key in
|
|
the `suffix` field of the record, making the compare operation during
|
|
binary search straightforward.
|
|
|
|
Once a restart point lexicographically before the sought reference has
|
|
been identified, readers can linearly scan through the following record
|
|
entries to locate the sought record, terminating if the current record
|
|
sorts after (and therefore the sought key is not present).
|
|
|
|
Restart point selection
|
|
+++++++++++++++++++++++
|
|
|
|
Writers determine the restart points at file creation. The process is
|
|
arbitrary, but every 16 or 64 records is recommended. Every 16 may be
|
|
more suitable for smaller block sizes (4k or 8k), every 64 for larger
|
|
block sizes (64k).
|
|
|
|
More frequent restart points reduces prefix compression and increases
|
|
space consumed by the restart table, both of which increase file size.
|
|
|
|
Less frequent restart points makes prefix compression more effective,
|
|
decreasing overall file size, with increased penalties for readers
|
|
walking through more records after the binary search step.
|
|
|
|
A maximum of `65535` restart points per block is supported.
|
|
|
|
Considerations
|
|
~~~~~~~~~~~~~~
|
|
|
|
Lightweight refs dominate
|
|
^^^^^^^^^^^^^^^^^^^^^^^^^
|
|
|
|
The reftable format assumes the vast majority of references are single
|
|
object names valued with common prefixes, such as Gerrit Code Review's
|
|
`refs/changes/` namespace, GitHub's `refs/pulls/` namespace, or many
|
|
lightweight tags in the `refs/tags/` namespace.
|
|
|
|
Annotated tags storing the peeled object cost an additional object name per
|
|
reference.
|
|
|
|
Low overhead
|
|
^^^^^^^^^^^^
|
|
|
|
A reftable with very few references (e.g. git.git with 5 heads) is 269
|
|
bytes for reftable, vs. 332 bytes for packed-refs. This supports
|
|
reftable scaling down for transaction logs (below).
|
|
|
|
Block size
|
|
^^^^^^^^^^
|
|
|
|
For a Gerrit Code Review type repository with many change refs, larger
|
|
block sizes (64 KiB) and less frequent restart points (every 64) yield
|
|
better compression due to more references within the block compressing
|
|
against the prior reference.
|
|
|
|
Larger block sizes reduce the index size, as the reftable will require
|
|
fewer blocks to store the same number of references.
|
|
|
|
Minimal disk seeks
|
|
^^^^^^^^^^^^^^^^^^
|
|
|
|
Assuming the index block has been loaded into memory, binary searching
|
|
for any single reference requires exactly 1 disk seek to load the
|
|
containing block.
|
|
|
|
Scans and lookups dominate
|
|
^^^^^^^^^^^^^^^^^^^^^^^^^^
|
|
|
|
Scanning all references and lookup by name (or namespace such as
|
|
`refs/heads/`) are the most common activities performed on repositories.
|
|
Object names are stored directly with references to optimize this use case.
|
|
|
|
Logs are infrequently read
|
|
^^^^^^^^^^^^^^^^^^^^^^^^^^
|
|
|
|
Logs are infrequently accessed, but can be large. Deflating log blocks
|
|
saves disk space, with some increased penalty at read time.
|
|
|
|
Logs are stored in an isolated section from refs, reducing the burden on
|
|
reference readers that want to ignore logs. Further, historical logs can
|
|
be isolated into log-only files.
|
|
|
|
Logs are read backwards
|
|
^^^^^^^^^^^^^^^^^^^^^^^
|
|
|
|
Logs are frequently accessed backwards (most recent N records for master
|
|
to answer `master@{4}`), so log records are grouped by reference, and
|
|
sorted descending by update index.
|
|
|
|
Repository format
|
|
~~~~~~~~~~~~~~~~~
|
|
|
|
Version 1
|
|
^^^^^^^^^
|
|
|
|
A repository must set its `$GIT_DIR/config` to configure reftable:
|
|
|
|
....
|
|
[core]
|
|
repositoryformatversion = 1
|
|
[extensions]
|
|
refStorage = reftable
|
|
....
|
|
|
|
Layout
|
|
^^^^^^
|
|
|
|
A collection of reftable files are stored in the `$GIT_DIR/reftable/` directory.
|
|
Their names should have a random element, such that each filename is globally
|
|
unique; this helps avoid spurious failures on Windows, where open files cannot
|
|
be removed or overwritten. It suggested to use
|
|
`${min_update_index}-${max_update_index}-${random}.ref` as a naming convention.
|
|
|
|
Log-only files use the `.log` extension, while ref-only and mixed ref
|
|
and log files use `.ref`. extension.
|
|
|
|
The stack ordering file is `$GIT_DIR/reftable/tables.list` and lists the
|
|
current files, one per line, in order, from oldest (base) to newest
|
|
(most recent):
|
|
|
|
....
|
|
$ cat .git/reftable/tables.list
|
|
00000001-00000001-RANDOM1.log
|
|
00000002-00000002-RANDOM2.ref
|
|
00000003-00000003-RANDOM3.ref
|
|
....
|
|
|
|
Readers must read `$GIT_DIR/reftable/tables.list` to determine which
|
|
files are relevant right now, and search through the stack in reverse
|
|
order (last reftable is examined first).
|
|
|
|
Reftable files not listed in `tables.list` may be new (and about to be
|
|
added to the stack by the active writer), or ancient and ready to be
|
|
pruned.
|
|
|
|
Backward compatibility
|
|
^^^^^^^^^^^^^^^^^^^^^^
|
|
|
|
Older clients should continue to recognize the directory as a git
|
|
repository so they don't look for an enclosing repository in parent
|
|
directories. To this end, a reftable-enabled repository must contain the
|
|
following dummy files
|
|
|
|
* `.git/HEAD`, a regular file containing `ref: refs/heads/.invalid`.
|
|
* `.git/refs/`, a directory
|
|
* `.git/refs/heads`, a regular file
|
|
|
|
Readers
|
|
^^^^^^^
|
|
|
|
Readers can obtain a consistent snapshot of the reference space by
|
|
following:
|
|
|
|
1. Open and read the `tables.list` file.
|
|
2. Open each of the reftable files that it mentions.
|
|
3. If any of the files is missing, goto 1.
|
|
4. Read from the now-open files as long as necessary.
|
|
|
|
Update transactions
|
|
^^^^^^^^^^^^^^^^^^^
|
|
|
|
Although reftables are immutable, mutations are supported by writing a
|
|
new reftable and atomically appending it to the stack:
|
|
|
|
1. Acquire `tables.list.lock`.
|
|
2. Read `tables.list` to determine current reftables.
|
|
3. Select `update_index` to be most recent file's
|
|
`max_update_index + 1`.
|
|
4. Prepare temp reftable `tmp_XXXXXX`, including log entries.
|
|
5. Rename `tmp_XXXXXX` to `${update_index}-${update_index}-${random}.ref`.
|
|
6. Copy `tables.list` to `tables.list.lock`, appending file from (5).
|
|
7. Rename `tables.list.lock` to `tables.list`.
|
|
|
|
During step 4 the new file's `min_update_index` and `max_update_index`
|
|
are both set to the `update_index` selected by step 3. All log records
|
|
for the transaction use the same `update_index` in their keys. This
|
|
enables later correlation of which references were updated by the same
|
|
transaction.
|
|
|
|
Because a single `tables.list.lock` file is used to manage locking, the
|
|
repository is single-threaded for writers. Writers may have to busy-spin
|
|
(with backoff) around creating `tables.list.lock`, for up to an
|
|
acceptable wait period, aborting if the repository is too busy to
|
|
mutate. Application servers wrapped around repositories (e.g. Gerrit
|
|
Code Review) can layer their own lock/wait queue to improve fairness to
|
|
writers.
|
|
|
|
Reference deletions
|
|
^^^^^^^^^^^^^^^^^^^
|
|
|
|
Deletion of any reference can be explicitly stored by setting the `type`
|
|
to `0x0` and omitting the `value` field of the `ref_record`. This serves
|
|
as a tombstone, overriding any assertions about the existence of the
|
|
reference from earlier files in the stack.
|
|
|
|
Compaction
|
|
^^^^^^^^^^
|
|
|
|
A partial stack of reftables can be compacted by merging references
|
|
using a straightforward merge join across reftables, selecting the most
|
|
recent value for output, and omitting deleted references that do not
|
|
appear in remaining, lower reftables.
|
|
|
|
A compacted reftable should set its `min_update_index` to the smallest
|
|
of the input files' `min_update_index`, and its `max_update_index`
|
|
likewise to the largest input `max_update_index`.
|
|
|
|
For sake of illustration, assume the stack currently consists of
|
|
reftable files (from oldest to newest): A, B, C, and D. The compactor is
|
|
going to compact B and C, leaving A and D alone.
|
|
|
|
1. Obtain lock `tables.list.lock` and read the `tables.list` file.
|
|
2. Obtain locks `B.lock` and `C.lock`. Ownership of these locks
|
|
prevents other processes from trying to compact these files.
|
|
3. Release `tables.list.lock`.
|
|
4. Compact `B` and `C` into a temp file
|
|
`${min_update_index}-${max_update_index}_XXXXXX`.
|
|
5. Reacquire lock `tables.list.lock`.
|
|
6. Verify that `B` and `C` are still in the stack, in that order. This
|
|
should always be the case, assuming that other processes are adhering to
|
|
the locking protocol.
|
|
7. Rename `${min_update_index}-${max_update_index}_XXXXXX` to
|
|
`${min_update_index}-${max_update_index}-${random}.ref`.
|
|
8. Write the new stack to `tables.list.lock`, replacing `B` and `C`
|
|
with the file from (4).
|
|
9. Rename `tables.list.lock` to `tables.list`.
|
|
10. Delete `B` and `C`, perhaps after a short sleep to avoid forcing
|
|
readers to backtrack.
|
|
|
|
This strategy permits compactions to proceed independently of updates.
|
|
|
|
Each reftable (compacted or not) is uniquely identified by its name, so
|
|
open reftables can be cached by their name.
|
|
|
|
Windows
|
|
^^^^^^^
|
|
|
|
On windows, and other systems that do not allow deleting or renaming to open
|
|
files, compaction may succeed, but other readers may prevent obsolete tables
|
|
from being deleted.
|
|
|
|
On these platforms, the following strategy can be followed: on closing a
|
|
reftable stack, reload `tables.list`, and delete any tables no longer mentioned
|
|
in `tables.list`.
|
|
|
|
Irregular program exit may still leave about unused files. In this case, a
|
|
cleanup operation should proceed as follows:
|
|
|
|
* take a lock `tables.list.lock` to prevent concurrent modifications
|
|
* refresh the reftable stack, by reading `tables.list`
|
|
* for each `*.ref` file, remove it if
|
|
** it is not mentioned in `tables.list`, and
|
|
** its max update_index is not beyond the max update_index of the stack
|
|
|
|
|
|
Alternatives considered
|
|
~~~~~~~~~~~~~~~~~~~~~~~
|
|
|
|
bzip packed-refs
|
|
^^^^^^^^^^^^^^^^
|
|
|
|
`bzip2` can significantly shrink a large packed-refs file (e.g. 62 MiB
|
|
compresses to 23 MiB, 37%). However the bzip format does not support
|
|
random access to a single reference. Readers must inflate and discard
|
|
while performing a linear scan.
|
|
|
|
Breaking packed-refs into chunks (individually compressing each chunk)
|
|
would reduce the amount of data a reader must inflate, but still leaves
|
|
the problem of indexing chunks to support readers efficiently locating
|
|
the correct chunk.
|
|
|
|
Given the compression achieved by reftable's encoding, it does not seem
|
|
necessary to add the complexity of bzip/gzip/zlib.
|
|
|
|
Michael Haggerty's alternate format
|
|
^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
|
|
|
|
Michael Haggerty proposed
|
|
link:https://lore.kernel.org/git/CAMy9T_HCnyc1g8XWOOWhe7nN0aEFyyBskV2aOMb_fe%2BwGvEJ7A%40mail.gmail.com/[an
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alternate] format to reftable on the Git mailing list. This format uses
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smaller chunks, without the restart table, and avoids block alignment
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with padding. Reflog entries immediately follow each ref, and are thus
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interleaved between refs.
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Performance testing indicates reftable is faster for lookups (51%
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faster, 11.2 usec vs. 5.4 usec), although reftable produces a slightly
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larger file (+ ~3.2%, 28.3M vs 29.2M):
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|
[cols=">,>,>,>",options="header",]
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|=====================================
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|
|format |size |seek cold |seek hot
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|mh-alt |28.3 M |23.4 usec |11.2 usec
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|reftable |29.2 M |19.9 usec |5.4 usec
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|
|=====================================
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|
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|
JGit Ketch RefTree
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|
^^^^^^^^^^^^^^^^^^
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|
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https://dev.eclipse.org/mhonarc/lists/jgit-dev/msg03073.html[JGit Ketch]
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proposed
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link:https://lore.kernel.org/git/CAJo%3DhJvnAPNAdDcAAwAvU9C4RVeQdoS3Ev9WTguHx4fD0V_nOg%40mail.gmail.com/[RefTree],
|
|
an encoding of references inside Git tree objects stored as part of the
|
|
repository's object database.
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|
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|
The RefTree format adds additional load on the object database storage
|
|
layer (more loose objects, more objects in packs), and relies heavily on
|
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the packer's delta compression to save space. Namespaces which are flat
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|
(e.g. thousands of tags in refs/tags) initially create very large loose
|
|
objects, and so RefTree does not address the problem of copying many
|
|
references to modify a handful.
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|
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Flat namespaces are not efficiently searchable in RefTree, as tree
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objects in canonical formatting cannot be binary searched. This fails
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the need to handle a large number of references in a single namespace,
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such as GitHub's `refs/pulls`, or a project with many tags.
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LMDB
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|
^^^^
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|
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|
David Turner proposed
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|
https://lore.kernel.org/git/1455772670-21142-26-git-send-email-dturner@twopensource.com/[using
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|
LMDB], as LMDB is lightweight (64k of runtime code) and GPL-compatible
|
|
license.
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|
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|
A downside of LMDB is its reliance on a single C implementation. This
|
|
makes embedding inside JGit (a popular reimplementation of Git)
|
|
difficult, and hoisting onto virtual storage (for JGit DFS) virtually
|
|
impossible.
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|
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|
A common format that can be supported by all major Git implementations
|
|
(git-core, JGit, libgit2) is strongly preferred.
|