(protocol proposal, work-in-progress, not authoritative) = Mutable Files = Mutable File Slots are places with a stable identifier that can hold data that changes over time. In contrast to CHK slots, for which the URI/identifier is derived from the contents themselves, the Mutable File Slot URI remains fixed for the life of the slot, regardless of what data is placed inside it. Each mutable slot is referenced by two different URIs. The "read-write" URI grants read-write access to its holder, allowing them to put whatever contents they like into the slot. The "read-only" URI is less powerful, only granting read access, and not enabling modification of the data. The read-write URI can be turned into the read-only URI, but not the other way around. The data in these slots is distributed over a number of servers, using the same erasure coding that CHK files use, with 3-of-10 being a typical choice of encoding parameters. The data is encrypted and signed in such a way that only the holders of the read-write URI will be able to set the contents of the slot, and only the holders of the read-only URI will be able to read those contents. Holders of either URI will be able to validate the contents as being written by someone with the read-write URI. The servers who hold the shares cannot read or modify them: the worst they can do is deny service (by deleting or corrupting the shares), or attempt a rollback attack (which can only succeed with the cooperation of at least k servers). == Consistency vs Availability == There is an age-old battle between consistency and availability. Epic papers have been written, elaborate proofs have been established, and generations of theorists have learned that you cannot simultaneously achieve guaranteed consistency with guaranteed reliability. In addition, the closer to 0 you get on either axis, the cost and complexity of the design goes up. Tahoe's design goals are to largely favor design simplicity, then slightly favor read availability, over the other criteria. As we develop more sophisticated mutable slots, the API may expose multiple read versions to the application layer. The tahoe philosophy is to defer most consistency recovery logic to the higher layers. Some applications have effective ways to merge multiple versions, so inconsistency is not necessarily a problem (i.e. directory nodes can usually merge multiple "add child" operations). == The Prime Coordination Directive: "Don't Do That" == The current rule for applications which run on top of Tahoe is "do not perform simultaneous uncoordinated writes". That means you need non-tahoe means to make sure that two parties are not trying to modify the same mutable slot at the same time. For example: * don't give the read-write URI to anyone else. Dirnodes in a private directory generally satisfy this case, as long as you don't use two clients on the same account at the same time * if you give a read-write URI to someone else, stop using it yourself. An inbox would be a good example of this. * if you give a read-write URI to someone else, call them on the phone before you write into it * build an automated mechanism to have your agents coordinate writes. For example, we expect a future release to include a FURL for a "coordination server" in the dirnodes. The rule can be that you must contact the coordination server and obtain a lock/lease on the file before you're allowed to modify it. If you do not follow this rule, Bad Things will happen. The worst-case Bad Thing is that the entire file will be lost. A less-bad Bad Thing is that one or more of the simultaneous writers will lose their changes. An observer of the file may not see monotonically-increasing changes to the file, i.e. they may see version 1, then version 2, then 3, then 2 again. Tahoe takes some amount of care to reduce the badness of these Bad Things. One way you can help nudge it from the "lose your file" case into the "lose some changes" case is to reduce the number of competing versions: multiple versions of the file that different parties are trying to establish as the one true current contents. Each simultaneous writer counts as a "competing version", as does the previous version of the file. If the count "S" of these competing versions is larger than N/k, then the file runs the risk of being lost completely. If at least one of the writers remains running after the collision is detected, it will attempt to recover, but if S>(N/k) and all writers crash after writing a few shares, the file will be lost. == Small Distributed Mutable Files == SDMF slots are suitable for small (<1MB) files that are editing by rewriting the entire file. The three operations are: * allocate (with initial contents) * set (with new contents) * get (old contents) The first use of SDMF slots will be to hold directories (dirnodes), which map encrypted child names to rw-URI/ro-URI pairs. === SDMF slots overview === Each SDMF slot is created with a DSA public/private key pair, using a system-wide common modulus and generator, in which the private key is a random 256 bit number, and the public key is a larger value (about 2048 bits) that can be derived with a bit of math from the private key. The public key is known as the "verification key", while the private key is called the "signature key". The 256-bit signature key is used verbatim as the "write capability". This can be converted into the 2048ish-bit verification key through a fairly cheap set of modular exponentiation operations; this is done any time the holder of the write-cap wants to read the data. (Note that the signature key can either be a newly-generated random value, or the hash of something else, if we found a need for a capability that's stronger than the write-cap). This results in a write-cap which is 256 bits long and can thus be expressed in an ASCII/transport-safe encoded form (base62 encoding, fits in 72 characters, including a local-node http: convenience prefix). The private key is hashed to form a 256-bit "salt". The public key is also hashed to form a 256-bit "pubkey hash". These two values are concatenated, hashed, and truncated to 192 bits to form the first 192 bits of the read-cap. The pubkey hash is hashed by itself and truncated to 64 bits to form the last 64 bits of the read-cap. The full read-cap is 256 bits long, just like the write-cap. The first 192 bits of the read-cap are hashed and truncated to form the first 64 bits of the storage index. The last 64 bits of the read-cap are hashed to form the last 64 bits of the storage index. This gives us a 128-bit storage index. The verification-cap is the first 64 bits of the storage index plus the pubkey hash, 320 bits total. The verification-cap doesn't need to be expressed in a printable transport-safe form, so it's ok that it's longer. The read-cap is hashed one way to form an AES encryption key that is used to encrypt the salt; this key is called the "salt key". The encrypted salt is stored in the share. The private key never changes, therefore the salt never changes, and the salt key is only used for a single purpose, so there is no need for an IV. The read-cap is hashed a different way to form the master data encryption key. A random "data salt" is generated each time the share's contents are replaced, and the master data encryption key is concatenated with the data salt, then hashed, to form the AES CTR-mode "read key" that will be used to encrypt the actual file data. This is to avoid key-reuse. An outstanding issue is how to avoid key reuse when files are modified in place instead of being replaced completely; this is not done in SDMF but might occur in MDMF. The private key is hashed one way to form the salt, and a different way to form the "write enabler master". For each storage server on which a share is kept, the write enabler master is concatenated with the server's nodeid and hashed, and the result is called the "write enabler" for that particular server. Note that multiple shares of the same slot stored on the same server will all get the same write enabler, i.e. the write enabler is associated with the "bucket", rather than the individual shares. The private key is hashed a third way to form the "data write key", which can be used by applications which wish to store some data in a form that is only available to those with a write-cap, and not to those with merely a read-cap. This is used to implement transitive read-onlyness of dirnodes. The public key is stored on the servers, as is the encrypted salt, the (non-encrypted) data salt, the encrypted data, and a signature. The container records the write-enabler, but of course this is not visible to readers. To make sure that every byte of the share can be verified by a holder of the verify-cap (and also by the storage server itself), the signature covers the version number, the sequence number, the root hash "R" of the share merkle tree, the encoding parameters, and the encrypted salt. "R" itself covers the hash trees and the share data. The read-write URI is just the private key. The read-only URI is the read-cap key. The verify-only URI contains the the pubkey hash and the first 64 bits of the storage index. FMW:b2a(privatekey) FMR:b2a(readcap) FMV:b2a(storageindex[:64])b2a(pubkey-hash) Note that this allows the read-only and verify-only URIs to be derived from the read-write URI without actually retrieving any data from the share, but instead by regenerating the public key from the private one. Uses of the read-only or verify-only caps must validate the public key against their pubkey hash (or its derivative) the first time they retrieve the pubkey, before trusting any signatures they see. The SDMF slot is allocated by sending a request to the storage server with a desired size, the storage index, and the write enabler for that server's nodeid. If granted, the write enabler is stashed inside the slot's backing store file. All further write requests must be accompanied by the write enabler or they will not be honored. The storage server does not share the write enabler with anyone else. The SDMF slot structure will be described in more detail below. The important pieces are: * a sequence number * a root hash "R" * the data salt * the encoding parameters (including k, N, file size, segment size) * a signed copy of [seqnum,R,data_salt,encoding_params] (using signature key) * the verification key (not encrypted) * the share hash chain (part of a Merkle tree over the share hashes) * the block hash tree (Merkle tree over blocks of share data) * the share data itself (erasure-coding of read-key-encrypted file data) * the salt, encrypted with the salt key The access pattern for read (assuming we hold the write-cap) is: * generate public key from the private one * hash private key to get the salt, hash public key, form read-cap * form storage-index * use storage-index to locate 'k' shares with identical 'R' values * either get one share, read 'k' from it, then read k-1 shares * or read, say, 5 shares, discover k, either get more or be finished * or copy k into the URIs * .. jump to "COMMON READ", below To read (assuming we only hold the read-cap), do: * hash read-cap pieces to generate storage index and salt key * use storage-index to locate 'k' shares with identical 'R' values * retrieve verification key and encrypted salt * decrypt salt * hash decrypted salt and pubkey to generate another copy of the read-cap, make sure they match (this validates the pubkey) * .. jump to "COMMON READ" * COMMON READ: * read seqnum, R, data salt, encoding parameters, signature * verify signature against verification key * hash data salt and read-cap to generate read-key * read share data, compute block-hash Merkle tree and root "r" * read share hash chain (leading from "r" to "R") * validate share hash chain up to the root "R" * submit share data to erasure decoding * decrypt decoded data with read-key * submit plaintext to application The access pattern for write is: * generate pubkey, salt, read-cap, storage-index as in read case * generate data salt for this update, generate read-key * encrypt plaintext from application with read-key * application can encrypt some data with the data-write-key to make it only available to writers (used for transitively-readonly dirnodes) * erasure-code crypttext to form shares * split shares into blocks * compute Merkle tree of blocks, giving root "r" for each share * compute Merkle tree of shares, find root "R" for the file as a whole * create share data structures, one per server: * use seqnum which is one higher than the old version * share hash chain has log(N) hashes, different for each server * signed data is the same for each server * include pubkey, encrypted salt, data salt * now we have N shares and need homes for them * walk through peers * if share is not already present, allocate-and-set * otherwise, try to modify existing share: * send testv_and_writev operation to each one * testv says to accept share if their(seqnum+R) <= our(seqnum+R) * count how many servers wind up with which versions (histogram over R) * keep going until N servers have the same version, or we run out of servers * if any servers wound up with a different version, report error to application * if we ran out of servers, initiate recovery process (described below) ==== Cryptographic Properties ==== This scheme protects the data's confidentiality with 192 bits of key material, since the read-cap contains 192 secret bits (derived from an encrypted salt, which is encrypted using those same 192 bits plus some additional public material). The integrity of the data (assuming that the signature is valid) is protected by the 256-bit hash which gets included in the signature. The privilege of modifying the data (equivalent to the ability to form a valid signature) is protected by a 256 bit random DSA private key, and the difficulty of computing a discrete logarithm in a 2048-bit field. There are a few weaker denial-of-service attacks possible. If N-k+1 of the shares are damaged or unavailable, the client will be unable to recover the file. Any coalition of more than N-k shareholders will be able to effect this attack by merely refusing to provide the desired share. The "write enabler" shared secret protects existing shares from being displaced by new ones, except by the holder of the write-cap. One server cannot affect the other shares of the same file, once those other shares are in place. The worst DoS attack is the "roadblock attack", which must be made before those shares get placed. Storage indexes are effectively random (being derived from the hash of a random value), so they are not guessable before the writer begins their upload, but there is a window of vulnerability during the beginning of the upload, when some servers have heard about the storage index but not all of them. The roadblock attack we want to prevent is when the first server that the uploader contacts quickly runs to all the other selected servers and places a bogus share under the same storage index, before the uploader can contact them. These shares will normally be accepted, since storage servers create new shares on demand. The bogus shares would have randomly-generated write-enablers, which will of course be different than the real uploader's write-enabler, since the malicious server does not know the write-cap. If this attack were successful, the uploader would be unable to place any of their shares, because the slots have already been filled by the bogus shares. The uploader would probably try for peers further and further away from the desired location, but eventually they will hit a preconfigured distance limit and give up. In addition, the further the writer searches, the less likely it is that a reader will search as far. So a successful attack will either cause the file to be uploaded but not be reachable, or it will cause the upload to fail. If the uploader tries again (creating a new privkey), they may get lucky and the malicious servers will appear later in the query list, giving sufficient honest servers a chance to see their share before the malicious one manages to place bogus ones. The first line of defense against this attack is the timing challenges: the attacking server must be ready to act the moment a storage request arrives (which will only occur for a certain percentage of all new-file uploads), and only has a few seconds to act before the other servers will have allocated the shares (and recorded the write-enabler, terminating the window of vulnerability). The second line of defense is post-verification, and is possible because the storage index is partially derived from the public key hash. A storage server can, at any time, verify every public bit of the container as being signed by the verification key (this operation is recommended as a continual background process, when disk usage is minimal, to detect disk errors). The server can also hash the verification key to derive 64 bits of the storage index. If it detects that these 64 bits do not match (but the rest of the share validates correctly), then the implication is that this share was stored to the wrong storage index, either due to a bug or a roadblock attack. If an uploader finds that they are unable to place their shares because of "bad write enabler errors" (as reported by the prospective storage servers), it can "cry foul", and ask the storage server to perform this verification on the share in question. If the pubkey and storage index do not match, the storage server can delete the bogus share, thus allowing the real uploader to place their share. Of course the origin of the offending bogus share should be logged and reported to a central authority, so corrective measures can be taken. It may be necessary to have this "cry foul" protocol include the new write-enabler, to close the window during which the malicious server can re-submit the bogus share during the adjudication process. If the problem persists, the servers can be placed into pre-verification mode, in which this verification is performed on all potential shares before being committed to disk. This mode is more CPU-intensive (since normally the storage server ignores the contents of the container altogether), but would solve the problem completely. The mere existence of these potential defenses should be sufficient to deter any actual attacks. Note that the storage index only has 64 bits of pubkey-derived data in it, which is below the usual crypto guidelines for security factors. In this case it's a pre-image attack which would be needed, rather than a collision, and the actual attack would be to find a keypair for which the public key can be hashed three times to produce the desired portion of the storage index. We believe that 64 bits of material is sufficiently resistant to this form of pre-image attack to serve as a suitable deterrent. === Server Storage Protocol === The storage servers will provide a mutable slot container which is oblivious to the details of the data being contained inside it. Each storage index refers to a "bucket", and each bucket has one or more shares inside it. (In a well-provisioned network, each bucket will have only one share). The bucket is stored as a directory, using the base32-encoded storage index as the directory name. Each share is stored in a single file, using the share number as the filename. The container holds space for a container magic number (for versioning), the write enabler, the nodeid which accepted the write enabler (used for share migration, described below), a small number of lease structures, the embedded data itself, and expansion space for additional lease structures. # offset size name 1 0 32 magic verstr "tahoe mutable container v1" plus binary 2 32 20 write enabler's nodeid 3 52 32 write enabler 4 84 8 data size (actual share data present) (a) 5 92 8 offset of (8) count of extra leases (after data) 6 100 368 four leases, 92 bytes each 0 4 ownerid (0 means "no lease here") 4 4 expiration timestamp 8 32 renewal token 40 32 cancel token 72 20 nodeid which accepted the tokens 7 468 (a) data 8 ?? 4 count of extra leases 9 ?? n*92 extra leases The "extra leases" field must be copied and rewritten each time the size of the enclosed data changes. The hope is that most buckets will have four or fewer leases and this extra copying will not usually be necessary. The (4) "data size" field contains the actual number of bytes of data present in field (7), such that a client request to read beyond 504+(a) will result in an error. This allows the client to (one day) read relative to the end of the file. The container size (that is, (8)-(7)) might be larger, especially if extra size was pre-allocated in anticipation of filling the container with a lot of data. The offset in (5) points at the *count* of extra leases, at (8). The actual leases (at (9)) begin 4 bytes later. If the container size changes, both (8) and (9) must be relocated by copying. The server will honor any write commands that provide the write token and do not exceed the server-wide storage size limitations. Read and write commands MUST be restricted to the 'data' portion of the container: the implementation of those commands MUST perform correct bounds-checking to make sure other portions of the container are inaccessible to the clients. The two methods provided by the storage server on these "MutableSlot" share objects are: * readv(ListOf(offset=int, length=int)) * returns a list of bytestrings, of the various requested lengths * offset < 0 is interpreted relative to the end of the data * spans which hit the end of the data will return truncated data * testv_and_writev(write_enabler, test_vector, write_vector) * this is a test-and-set operation which performs the given tests and only applies the desired writes if all tests succeed. This is used to detect simultaneous writers, and to reduce the chance that an update will lose data recently written by some other party (written after the last time this slot was read). * test_vector=ListOf(TupleOf(offset, length, opcode, specimen)) * the opcode is a string, from the set [gt, ge, eq, le, lt, ne] * each element of the test vector is read from the slot's data and compared against the specimen using the desired (in)equality. If all tests evaluate True, the write is performed * write_vector=ListOf(TupleOf(offset, newdata)) * offset < 0 is not yet defined, it probably means relative to the end of the data, which probably means append, but we haven't nailed it down quite yet * write vectors are executed in order, which specifies the results of overlapping writes * return value: * error: OutOfSpace * error: something else (io error, out of memory, whatever) * (True, old_test_data): the write was accepted (test_vector passed) * (False, old_test_data): the write was rejected (test_vector failed) * both 'accepted' and 'rejected' return the old data that was used for the test_vector comparison. This can be used by the client to detect write collisions, including collisions for which the desired behavior was to overwrite the old version. In addition, the storage server provides several methods to access these share objects: * allocate_mutable_slot(storage_index, sharenums=SetOf(int)) * returns DictOf(int, MutableSlot) * get_mutable_slot(storage_index) * returns DictOf(int, MutableSlot) * or raises KeyError We intend to add an interface which allows small slots to allocate-and-write in a single call, as well as do update or read in a single call. The goal is to allow a reasonably-sized dirnode to be created (or updated, or read) in just one round trip (to all N shareholders in parallel). ==== migrating shares ==== If a share must be migrated from one server to another, two values become invalid: the write enabler (since it was computed for the old server), and the lease renew/cancel tokens. Suppose that a slot was first created on nodeA, and was thus initialized with WE(nodeA) (= H(WEM+nodeA)). Later, for provisioning reasons, the share is moved from nodeA to nodeB. Readers may still be able to find the share in its new home, depending upon how many servers are present in the grid, where the new nodeid lands in the permuted index for this particular storage index, and how many servers the reading client is willing to contact. When a client attempts to write to this migrated share, it will get a "bad write enabler" error, since the WE it computes for nodeB will not match the WE(nodeA) that was embedded in the share. When this occurs, the "bad write enabler" message must include the old nodeid (e.g. nodeA) that was in the share. The client then computes H(nodeB+H(WEM+nodeA)), which is the same as H(nodeB+WE(nodeA)). The client sends this along with the new WE(nodeB), which is H(WEM+nodeB). Note that the client only sends WE(nodeB) to nodeB, never to anyone else. Also note that the client does not send a value to nodeB that would allow the node to impersonate the client to a third node: everything sent to nodeB will include something specific to nodeB in it. The server locally computes H(nodeB+WE(nodeA)), using its own node id and the old write enabler from the share. It compares this against the value supplied by the client. If they match, this serves as proof that the client was able to compute the old write enabler. The server then accepts the client's new WE(nodeB) and writes it into the container. This WE-fixup process requires an extra round trip, and requires the error message to include the old nodeid, but does not require any public key operations on either client or server. Migrating the leases will require a similar protocol. This protocol will be defined concretely at a later date. === Code Details === The current FileNode class will be renamed ImmutableFileNode, and a new MutableFileNode class will be created. Instances of this class will contain a URI and a reference to the client (for peer selection and connection). The methods of MutableFileNode are: * replace(newdata) -> OK, ConsistencyError, NotEnoughPeersError * get() -> [deferred] newdata, NotEnoughPeersError * if there are multiple retrieveable versions in the grid, get() returns the first version it can reconstruct, and silently ignores the others. In the future, a more advanced API will signal and provide access to the multiple heads. The peer-selection and data-structure manipulation (and signing/verification) steps will be implemented in a separate class in allmydata/mutable.py . === SMDF Slot Format === This SMDF data lives inside a server-side MutableSlot container. The server is generally oblivious to this format, but it may look inside the container when verification is desired. This data is tightly packed. There are no gaps left between the different fields, and the offset table is mainly present to allow future flexibility of key sizes. # offset size name 1 0 1 version byte, \x01 for this format 2 1 8 sequence number. 2^64-1 must be handled specially, TBD 3 9 32 "R" (root of share hash Merkle tree) 4 41 32 data salt (readkey is H(readcap+data_salt)) 5 73 32 encrypted salt (AESenc(key=H(readcap), salt) 6 105 18 encoding parameters: 105 1 k 106 1 N 107 8 segment size 115 8 data length (of original plaintext) 7 123 36 offset table: 127 4 (9) signature 131 4 (10) share hash chain 135 4 (11) block hash tree 139 4 (12) share data 143 8 (13) EOF 8 151 256 verification key (2048bit DSA key) 9 407 40 signature=DSAsig(H([1,2,3,4,5,6])) 10 447 (a) share hash chain, encoded as: "".join([pack(">H32s", shnum, hash) for (shnum,hash) in needed_hashes]) 11 ?? (b) block hash tree, encoded as: "".join([pack(">32s",hash) for hash in block_hash_tree]) 12 ?? LEN share data 13 ?? -- EOF (a) The share hash chain contains ceil(log(N)) hashes, each 32 bytes long. This is the set of hashes necessary to validate this share's leaf in the share Merkle tree. For N=10, this is 4 hashes, i.e. 128 bytes. (b) The block hash tree contains ceil(length/segsize) hashes, each 32 bytes long. This is the set of hashes necessary to validate any given block of share data up to the per-share root "r". Each "r" is a leaf of the share has tree (with root "R"), from which a minimal subset of hashes is put in the share hash chain in (8). === Recovery === The first line of defense against damage caused by colliding writes is the Prime Coordination Directive: "Don't Do That". The second line of defense is to keep "S" (the number of competing versions) lower than N/k. If this holds true, at least one competing version will have k shares and thus be recoverable. Note that server unavailability counts against us here: the old version stored on the unavailable server must be included in the value of S. The third line of defense is our use of testv_and_writev() (described below), which increases the convergence of simultaneous writes: one of the writers will be favored (the one with the highest "R"), and that version is more likely to be accepted than the others. This defense is least effective in the pathological situation where S simultaneous writers are active, the one with the lowest "R" writes to N-k+1 of the shares and then dies, then the one with the next-lowest "R" writes to N-2k+1 of the shares and dies, etc, until the one with the highest "R" writes to k-1 shares and dies. Any other sequencing will allow the highest "R" to write to at least k shares and establish a new revision. The fourth line of defense is the fact that each client keeps writing until at least one version has N shares. This uses additional servers, if necessary, to make sure that either the client's version or some newer/overriding version is highly available. The fifth line of defense is the recovery algorithm, which seeks to make sure that at least *one* version is highly available, even if that version is somebody else's. The write-shares-to-peers algorithm is as follows: * permute peers according to storage index * walk through peers, trying to assign one share per peer * for each peer: * send testv_and_writev, using "old(seqnum+R) <= our(seqnum+R)" as the test * this means that we will overwrite any old versions, and we will overwrite simultaenous writers of the same version if our R is higher. We will not overwrite writers using a higher seqnum. * record the version that each share winds up with. If the write was accepted, this is our own version. If it was rejected, read the old_test_data to find out what version was retained. * if old_test_data indicates the seqnum was equal or greater than our own, mark the "Simultanous Writes Detected" flag, which will eventually result in an error being reported to the writer (in their close() call). * build a histogram of "R" values * repeat until the histogram indicate that some version (possibly ours) has N shares. Use new servers if necessary. * If we run out of servers: * if there are at least shares-of-happiness of any one version, we're happy, so return. (the close() might still get an error) * not happy, need to reinforce something, goto RECOVERY RECOVERY: * read all shares, count the versions, identify the recoverable ones, discard the unrecoverable ones. * sort versions: locate max(seqnums), put all versions with that seqnum in the list, sort by number of outstanding shares. Then put our own version. (TODO: put versions with seqnum us ahead of us?). * for each version: * attempt to recover that version * if not possible, remove it from the list, go to next one * if recovered, start at beginning of peer list, push that version, continue until N shares are placed * if pushing our own version, bump up the seqnum to one higher than the max seqnum we saw * if we run out of servers: * schedule retry and exponential backoff to repeat RECOVERY * admit defeat after some period? presumeably the client will be shut down eventually, maybe keep trying (once per hour?) until then. == Medium Distributed Mutable Files == These are just like the SDMF case, but: * we actually take advantage of the Merkle hash tree over the blocks, by reading a single segment of data at a time (and its necessary hashes), to reduce the read-time alacrity * we allow arbitrary writes to the file (i.e. seek() is provided, and O_TRUNC is no longer required) * we write more code on the client side (in the MutableFileNode class), to first read each segment that a write must modify. This looks exactly like the way a normal filesystem uses a block device, or how a CPU must perform a cache-line fill before modifying a single word. * we might implement some sort of copy-based atomic update server call, to allow multiple writev() calls to appear atomic to any readers. MDMF slots provide fairly efficient in-place edits of very large files (a few GB). Appending data is also fairly efficient, although each time a power of 2 boundary is crossed, the entire file must effectively be re-uploaded (because the size of the block hash tree changes), so if the filesize is known in advance, that space ought to be pre-allocated (by leaving extra space between the block hash tree and the actual data). MDMF1 uses the Merkle tree to enable low-alacrity random-access reads. MDMF2 adds cache-line reads to allow random-access writes. == Large Distributed Mutable Files == LDMF slots use a fundamentally different way to store the file, inspired by Mercurial's "revlog" format. They enable very efficient insert/remove/replace editing of arbitrary spans. Multiple versions of the file can be retained, in a revision graph that can have multiple heads. Each revision can be referenced by a cryptographic identifier. There are two forms of the URI, one that means "most recent version", and a longer one that points to a specific revision. Metadata can be attached to the revisions, like timestamps, to enable rolling back an entire tree to a specific point in history. LDMF1 provides deltas but tries to avoid dealing with multiple heads. LDMF2 provides explicit support for revision identifiers and branching. == TODO == improve allocate-and-write or get-writer-buckets API to allow one-call (or maybe two-call) updates. The challenge is in figuring out which shares are on which machines. First cut will have lots of round trips. (eventually) define behavior when seqnum wraps. At the very least make sure it can't cause a security problem. "the slot is worn out" is acceptable. (eventually) define share-migration lease update protocol. Including the nodeid who accepted the lease is useful, we can use the same protocol as we do for updating the write enabler. However we need to know which lease to update.. maybe send back a list of all old nodeids that we find, then try all of them when we accept the update? We now do this in a specially-formatted IndexError exception: "UNABLE to renew non-existent lease. I have leases accepted by " + "nodeids: '12345','abcde','44221' ." Every node in a given tahoe grid must have the same common DSA moduli and exponent, but different grids could use different parameters. We haven't figured out how to define a "grid id" yet, but I think the DSA parameters should be part of that identifier. In practical terms, this might mean that the Introducer tells each node what parameters to use, or perhaps the node could have a config file which specifies them instead.