Putting consistency back into eventual consistency Putting consistency back into eventual consistency
Paper summary Many distributed databases geo-replicate data for (i) lower read latency and (ii) higher fault tolerance in the face of an extreme failure (e.g. lightning striking a data center). Implementing strong consistency over a geo-replicated system can incur tremendous write latency, as updates have to coordinate between geographically distant data centers. On the other hand, weak consistency is a real brain buster. This paper introduces a new consistency model between weak and strong consistency, explicit consistency, which takes into account user specified invariants. It also presents an explicitly consistent system which 1. performs static analysis to determine which operations can be executed without coordination, 2. uses invariant-repair or violation-avoidance to resolve or avoid conflicts, and 3. instruments user code with calls to middleware. The system is called Indigo and is built on an existing causally consistent key-value store with various properties. Explicit Consistency. A database is a collection of objects replicated across data centers. Users issue reads and writes as part of transactions, and these transactions are asynchronously replicated between data centers. We denote by t(S) the database state achieved by applying transaction t to database state S. S_n is the database state achieved after the nth transaction. That is, $S_n = t_n(...(t_1(t_init))...)$. $T(S_n) = \{t_1, ..., t_n\}$ is the set of transactions used to create S_n. We say a transaction t_a happens before a transaction $t_b$, denoted $t_a \rightarrow t_b$, if t_a is in $T(S_b)$. $O = (T,\rightarrow)$ is a partial order. $O' = (T, <)$ is a serialization of $O$ if $<$ is total and respects $\rightarrow$. Given an invariant $I$, we say $S$ is $I$-valid if $I(S) = true$. $(T, <)$ is an I-valid serialization if I holds on all prefixes of the serialization. If a system ensures that all serializations are I-valid, it provides explicit consistency. In other words, an explicitly consistent database ensures invariants always hold. This builds off of Bailis et al.'s notion of invariant-confluence. Determining $I$-offender Sets. An invariant is a universally quantified first order logic formula in prenex normal form. The invariant can include uninterpreted functions like Player($P$) and enrolled($P, T$). A postcondition states how operations affect the truth values of the uninterpreted functions in invariants. Every operation is annotated with postconditions. A predicate clause directly alters the truth assignments of a predicate (e.g. not Player(P)). A function clause relates old and new database states (e.g. nrPlayers(T) = nrPlayers(T) + 1). This language is rather expressive, as evidenced by multiple examples in the paper. A set of transactions is an I-offender if it is not invariant-confluent. First, pairs of operations are checked to see if a contradictory truth assignment is formed (e.g. Player(P) and not Player(P)). Then, every pair of transactions is considered. Given the weakest liberal precondition of the transactions, we substitute the effects of the transactions into the invariant to get a formula. We then check for the validity of the formula using Z3. If the formula is valid, the transactions are invariant-confluent. Handling I-offender Sets. There are two ways to handle I-offenders: invariant-repair and violation-avoidance. Invariant-repair involves CRDTs; the bulk of this paper focuses on violation-avoidance which leverages existing reservation and escrow transaction techniques. - UID generation. Unique identifiers can easily be generated without coordination. For example, a node can concatenate an incrementing counter with its MAC address. - Multi-level lock reservation. Locking is the most general form of reservation. Locks come in three flavors: (i) shared forbid, (ii) shared allow, and (iii) exclusive allow. Transactions acquire locks to avoid invariant violation. For example, an enrollTournament could acquire a sharedForbid lock on removing players, and removePlayer could acquire a sharedAllow lock. Exclusive allow are used for self-conflicting operations. - Multi-level mask reservation. If our invariant is a disjunction P1 or ... or Pn, then to preserve the invariant, we only need to guarantee that at least one of the disjuncts remains true. A mask reservation is a vector of locks where an operation can falsify one of the disjuncts only after acquiring a lock on another true disjunct preventing it from being falsified. - Escrow reservation. Imagine our invariant is x >= k and x has initial value x0. Escrow transactions allocate x0 - k rights. A transaction can decrement x only after acquiring and spending a right. When x is incremented, a right is generated. This gets tricker for invariants like |A| >= k where concurrent additions could generate too many rights leading to an invariant violation. Here, we use escrow transactions for conditions, where a primary is allocated for each reservation. Rights are not immediately generated; instead, the primary is responsible for generating rights. - Partition lock reservation. Partition locks allow operations to lock a small part of an object. For example, an operation could lock part of a timeline to ensure there are no overlapping timespans. There are many ways to use reservations to avoid invariant violations. Indigo uses heuristics and estimated operation frequencies to try and minimize reservation acquisitions. Implementation. Indigo can run over any key-value store that offers (i) causally consistency, (ii) snapshot transactions with CRDTs, and (iii) linearizability within a data center. Currently, it uses Swiftcloud. Its fault tolerance leverages the underlying fault tolerance of the key-value store. Each reservation is stored as an object in the key-value store, where operations are structured as transfers to avoid some concurrency oddities.

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