Scalable Atomic Visibility with RAMP Transactions

Peter Bailis, Alan Fekete†, Ali Ghodsi, Joseph M. Hellerstein, Ion StoicaUC Berkeley and †University of Sydney

ABSTRACTDatabases can provide scalability by partitioning data across severalservers. However, multi-partition, multi-operation transactional ac-cess is often expensive, employing coordination-intensive locking,validation, or scheduling mechanisms. Accordingly, many real-world systems avoid mechanisms that provide useful semantics formulti-partition operations. This leads to incorrect behavior for alarge class of applications including secondary indexing, foreign keyenforcement, and materialized view maintenance. In this work, weidentify a new isolation model—Read Atomic (RA) isolation—thatmatches the requirements of these use cases by ensuring atomic vis-ibility: either all or none of each transaction’s updates are observedby other transactions. We present algorithms for Read Atomic Multi-Partition (RAMP) transactions that enforce atomic visibility whileoffering excellent scalability, guaranteed commit despite partialfailures (via synchronization independence), and minimized com-munication between servers (via partition independence). TheseRAMP transactions correctly mediate atomic visibility of updatesand provide readers with snapshot access to database state by usinglimited multi-versioning and by allowing clients to independentlyresolve non-atomic reads. We demonstrate that, in contrast with ex-isting algorithms, RAMP transactions incur limited overhead—evenunder high contention—and scale linearly to 100 servers.

1. INTRODUCTIONFaced with growing amounts of data and unprecedented query

volume, distributed databases increasingly split their data acrossmultiple servers, or partitions, such that no one partition containsan entire copy of the database [7,13,18,19,22,29,43]. This strategysucceeds in allowing near-unlimited scalability for operations thataccess single partitions. However, operations that access multiplepartitions must communicate across servers—often synchronously—in order to provide correct behavior. Designing systems and algo-rithms that tolerate these communication delays is a difficult taskbut is key to maintaining scalability [17, 28, 29, 35].

In this work, we address a largely underserved class of appli-cations requiring multi-partition, atomically visible1 transactionalaccess: cases where all or none of each transaction’s effects shouldbe visible. The status quo for these multi-partition atomic transac-tions provides an uncomfortable choice between algorithms that1Our use of “atomic” (specifically, Read Atomic isolation) concerns all-or-nothingvisibility of updates (i.e., the ACID isolation effects of ACID atomicity; Section 3).This differs from uses of “atomicity” to denote serializability [8] or linearizability [4].Permission to make digital or hard copies of all or part of this work for personal orclassroom use is granted without fee provided that copies are not made or distributedfor profit or commercial advantage and that copies bear this notice and the full citationon the first page. Copyrights for components of this work owned by others than theauthor(s) must be honored. Abstracting with credit is permitted. To copy otherwise, orrepublish, to post on servers or to redistribute to lists, requires prior specific permissionand/or a fee. Request permissions from permissions@acm.org.SIGMOD’14, June 22–27, 2014, Snowbird, UT, USA.Copyright is held by the owner/author(s). Publication rights licensed to ACM.ACM 978-1-4503-2376-5/14/06 ...$15.00.http://dx.doi.org/10.1145/2588555.2588562.

are fast but deliver inconsistent results and algorithms that de-liver consistent results but are often slow and unavailable underfailure. Many of the largest modern, real-world systems opt forprotocols that guarantee fast and scalable operation but providefew—if any—transactional semantics for operations on arbitrarysets of data items [11, 13, 15, 22, 26, 38, 44]. This results in incor-rect behavior for use cases that require atomic visibility, includ-ing secondary indexing, foreign key constraint enforcement, andmaterialized view maintenance (Section 2). In contrast, many tra-ditional transactional mechanisms correctly ensure atomicity ofupdates [8, 17, 43]. However, these algorithms—such as two-phaselocking and variants of optimistic concurrency control—are oftencoordination-intensive, slow, and, under failure, unavailable in adistributed environment [5, 18, 28, 35]. This dichotomy betweenscalability and atomic visibility has been described as “a fact oflife in the big cruel world of huge systems” [25]. The prolifera-tion of non-transactional multi-item operations is symptomatic of awidespread “fear of synchronization” at scale [9].

Our contribution in this paper is to demonstrate that atomicallyvisible transactions on partitioned databases are not at odds withscalability. Specifically, we provide high-performance implementa-tions of a new, non-serializable isolation model called Read Atomic(RA) isolation. RA ensures that all or none of each transaction’supdates are visible to others and that each transaction reads from anatomic snapshot of database state (Section 3)—this is useful in theapplications we target. We subsequently develop three new, scalablealgorithms for achieving RA isolation that we collectively title ReadAtomic Multi-Partition (RAMP) transactions (Section 4). RAMPtransactions guarantee scalability and outperform existing atomicalgorithms because they satisfy two key scalability constraints. First,RAMP transactions guarantee synchronization independence: oneclient’s transactions cannot cause another client’s transactions tostall or fail. Second, RAMP transactions guarantee partition inde-pendence: clients never need to contact partitions that their trans-actions do not directly reference. Together, these properties ensureguaranteed completion, limited coordination across partitions, andhorizontal scalability for multi-partition access.

RAMP transactions are scalable because they appropriately con-trol the visibility of updates without inhibiting concurrency. Ratherthan force concurrent reads and writes to stall, RAMP transactionsallow reads to “race” writes: RAMP transactions can autonomouslydetect the presence of non-atomic (partial) reads and, if necessary,repair them via a second round of communication with servers. Toaccomplish this, RAMP writers attach metadata to each write anduse limited multi-versioning to prevent readers from stalling. Thethree algorithms we present offer a trade-off between the size of thismetadata and performance. RAMP-Small transactions require con-stant space (a timestamp per write) and two round trip time delays(RTTs) for reads and writes. RAMP-Fast transactions require meta-data size that is linear in the number of writes in the transaction butonly require one RTT for reads in the common case and two in theworst case. RAMP-Hybrid transactions employ Bloom filters [10] toprovide an intermediate solution. Traditional techniques like locking

couple atomic visibility and mutual exclusion; RAMP transactionsprovide the benefits of the former without incurring the scalability,availability, or latency penalties of the latter.

In addition to providing a theoretical analysis and proofs of cor-rectness, we demonstrate that RAMP transactions deliver in practice.Our RAMP implementation achieves linear scalability to over 7million operations per second on a 100 server cluster (at overheadbelow 5% for a workload of 95% reads). Moreover, across a rangeof workload configurations, RAMP transactions incur limited over-head compared to other techniques and achieve higher performancethan existing approaches to atomic visibility (Section 5).

While the literature contains an abundance of isolation models [2,5], we believe that the large number of modern applications requiringRA isolation and the excellent scalability of RAMP transactionsjustify the addition of yet another model. RA isolation is too weakfor some applications, but, for the many that it can serve, RAMPtransactions offer substantial benefits.

2. OVERVIEW AND MOTIVATIONIn this paper, we consider the problem of making transactional

updates atomically visible to readers—a requirement that, as weoutline in this section, is found in several prominent use cases today.The basic property we provide is fairly simple: either all or none ofeach transaction’s updates should be visible to other transactions.For example, if a transaction T1 writes x = 1 and y = 1, then anothertransaction T2 should not read x = 1 and y = null. Instead, T2 shouldeither read x = 1 and y = 1 or, possibly, x = null and y = null.Informally, each transaction reads from an unchanging snapshot ofdatabase state that is aligned along transactional boundaries. We callthis property atomic visibility and formalize it via the Read Atomicisolation guarantee in Section 3.

The classic strategy for providing atomic visibility is to ensuremutual exclusion between readers and writers. For example, if atransaction like T1 above wants to update data items x and y, itcan acquire exclusive locks for each of x and y, update both items,then release the locks. No other transactions will observe partialupdates to x and y, ensuring atomic visibility. However, this solutionhas a drawback: while one transaction holds exclusive locks on xand y, no other transactions can access x and y for either reads orwrites. By using mutual exclusion to enforce the atomic visibilityof updates, we have also limited concurrency. In our example, if xand y are located on different servers, concurrent readers and writerswill be unable to perform useful work during communication delays.These communication delays form an upper bound on throughput:effectively, 1

message delay operations per second.To avoid this upper bound, we separate the problem of providing

atomic visibility from the problem of maintaining mutual exclusion.By achieving the former but avoiding the latter, the algorithms wedevelop in this paper are not subject to the scalability penaltiesof many prior approaches. To ensure that all servers successfullyexecute a transaction (or that none do), our algorithms employ anatomic commitment protocol (ACP). When coupled with a blockingconcurrency control mechanism like locking, ACPs are harmful toscalability and availability: arbitrary failures can (provably) causeany ACP implementation to stall [8]. (Optimistic concurrency con-trol mechanisms can similarly block during validation.) We insteaduse ACPs with non-blocking concurrency control mechanisms; thismeans that individual transactions can stall due to failures or com-munication delays without forcing other transactions to stall. In adeparture from traditional concurrency control, we allow multipleACP rounds to proceed in parallel over the same data.

The end result—our RAMP transactions—provide excellent scal-ability and performance under contention (e.g., in the event of write

hotspots) and are robust to partial failure. RAMP transactions’ non-blocking behavior means that they cannot provide certain guaranteeslike preventing concurrent updates. However, applications that canuse Read Atomic isolation will benefit from our algorithms. Theremainder of this section identifies several relevant use cases fromindustry that require atomic visibility for correctness.

2.1 Read Atomic Isolation in the WildAs a simple example, consider a social networking application:

if two users, Sam and Mary, become “friends” (a bi-directionalrelationship), other users should never see that Sam is a friend ofMary but Mary is not a friend of Sam: either both relationshipsshould be visible, or neither should be. A transaction under ReadAtomic isolation would correctly enforce this behavior, and we canfurther classify three general use cases for Read Atomic isolation:1.) Foreign key constraints. Many database schemas contain in-formation about relationships between records in the form of foreignkey constraints. For example, Facebook’s TAO [11], LinkedIn’sEspresso [38], and Yahoo! PNUTS [15] store information aboutbusiness entities such as users, photos, and status updates as wellas relationships between them (e.g., the friend relationships above).Their data models often represent bi-directional edges as two dis-tinct uni-directional relationships. For example, in TAO, a userperforming a “like” action on a Facebook page produces updatesto both the LIKES and LIKED_BY associations [11]. PNUTS’s au-thors describe an identical scenario [15]. These applications requireforeign key maintenance and often, due to their unidirectional re-lationships, multi-entity update and access. Violations of atomicvisibility surface as broken bi-directional relationships (as with Samand Mary above) and dangling or incorrect references (e.g., Frank isan employee of department.id=5, but no such department existsin the department table).

With RAMP transactions, when inserting new entities, applica-tions can bundle relevant entities from each side of a foreign keyconstraint into a transaction. When deleting associations, userscan “tombstone” the opposite end of the association (i.e., deleteany entries with associations via a special record that signifies dele-tion) [45] to avoid dangling pointers.2.) Secondary indexing. Data is typically partitioned acrossservers according to a primary key (e.g., user ID). This allows fastlocation and retrieval of data via primary key lookups but makesaccess by secondary attributes (e.g., birth date) challenging. Thereare two dominant strategies for distributed secondary indexing. First,the local secondary index approach co-locates secondary indexesand primary data, so each server contains a secondary index that onlyreferences (and indexes) data stored on its server [7,38]. This allowseasy, single-server updates but requires contacting every partitionfor secondary attribute lookups (write-one, read-all), compromisingscalability for read-heavy workloads [11, 17, 38]. Alternatively, theglobal secondary index approach locates secondary indexes (whichmay be partitioned, but by a secondary attribute) separately fromprimary data [7, 15]. This alternative allows fast secondary lookups(read-one) but requires multi-partition update (at least write-two).

Real-world services employ either local secondary indexing (e.g.,Espresso [38], Cassandra, and Google Megastore’s local indexes [7])or non-atomic (incorrect) global secondary indexing (e.g., Espressoand Megastore’s global indexes, Yahoo! PNUTS’s proposed sec-ondary indexes [15]). The former is non-scalable but correct, whilethe latter is scalable but incorrect. For example, in a databasepartitioned by id with an incorrectly-maintained global secondaryindex on salary, the query ‘SELECT id, salary WHERE salary> 6�,���’ might return records with salary less than $60,000 andomit some records with salary greater than $60,000.

With RAMP transactions, the secondary index entry for a givenattribute can be updated atomically with base data. For example, ifa secondary index is stored as a mapping from secondary attributevalues to sets of item-versions matching the secondary attribute (e.g.,the secondary index entry for users with blue hair would containa list of user IDs and last-modified timestamps corresponding toall of the users with attribute hair-color=blue), then insertionsof new primary data require additions to the corresponding indexentry, deletions require removals, and updates require a “tombstone”deletion from one entry and an insertion into another.3.) Materialized view maintenance. Many applications precom-pute (i.e., materialize) queries over data, as in Twitter’s Rainbirdservice [44], Google’s Percolator [36], and LinkedIn’s Espressosystems [38]. As a simple example, Espresso stores a mailbox ofmessages for each user along with statistics about the mailbox mes-sages: for Espresso’s read-mostly workload, it is more efficient tomaintain (i.e., pre-materialize) a count of unread messages ratherthan scan all messages every time a user accesses her mailbox [38].In this case, any unread message indicators should remain in syncwith the messages in the mailbox. However, atomicity violationswill allow materialized views to diverge from the base data (e.g., Su-san’s mailbox displays a notification that she has unread messagesbut all 63,201 messages in her inbox are marked as read).

With RAMP transactions, base data and views can be updatedatomically. The physical maintenance of a view depends on itsspecification [14, 27], but RAMP transactions provide appropriateconcurrency control primitives for ensuring that changes are deliv-ered to the materialized view partition. For select-project views, asimple solution is to treat the view as a separate table and performmaintenance as needed: new rows can be inserted/deleted accordingto the specification, and, if necessary, the view can be (re-)computedon demand (i.e., lazy view maintenance [46]). For more complexviews, such as counters, users can execute RAMP transactions overspecialized data structures such as the CRDT G-Counter [40].In brief: Status Quo. Despite application requirements for ReadAtomic isolation, few large-scale production systems provide it. Forexample, the authors of Tao, Espresso, and PNUTS describe severalclasses of atomicity anomalies exposed by their systems, rangingfrom dangling pointers to the exposure of intermediate states andincorrect secondary index lookups, often highlighting these casesas areas for future research and design [11, 15, 38]. These systemsare not exceptions: data stores like Bigtable [13], Dynamo [22], andmany popular “NoSQL” [34] and even some “NewSQL” [5] storesdo not provide transactional guarantees for multi-item operations.

The designers of these Internet-scale, real-world systems havemade a conscious decision to provide scalability at the expense ofmulti-partition transactional semantics. Our goal with RAMP trans-actions is to preserve this scalability but deliver correct, atomicallyvisible behavior for the use cases we have described.

3. SEMANTICS AND SYSTEM MODELIn this section, we formalize Read Atomic isolation and, to cap-

ture scalability, formulate a pair of strict scalability criteria: syn-chronization and partition independence. Readers more interestedin RAMP algorithms may wish to proceed to Section 4.

3.1 RA Isolation: Formal SpecificationTo formalize RA isolation, as is standard [2], we consider ordered

sequences of reads and writes to arbitrary sets of items, or transac-tions. We call the set of items a transaction reads from and writesto its read set and write set. Each write creates a version of an itemand we identify versions of items by a unique timestamp taken from

a totally ordered set (e.g., rational numbers). Timestamps inducea total order on versions of each item (and a partial order acrossversions of different items). We denote version i of item x as xi.

A transaction Tj exhibits fractured reads if transaction Ti writesversions xm and yn (in any order, with x possibly but not necessarilyequal to y), Tj reads version xm and version yk, and k < n.

A system provides Read Atomic isolation (RA) if it prevents frac-tured reads anomalies and also prevents transactions from readinguncommitted, aborted, or intermediate data. Thus, RA providestransactions with a “snapshot” view of the database that respectstransaction boundaries (see the Appendix for more details, includinga discussion of transitivity). RA is simply a restriction on write visi-bility—if the ACID “Atomicity” property requires that all or none ofa transaction’s updates are performed, RA requires that all or noneof a transaction’s updates are made visible to other transactions.

3.2 RA Implications and LimitationsAs outlined in Section 2.1, RA isolation matches many of our

use cases. However, RA is not sufficient for all applications. RAdoes not prevent concurrent updates or provide serial access to dataitems. For example, RA is an incorrect choice for an applicationthat wishes to maintain positive bank account balances in the eventof withdrawals. RA is a better fit for our “friend” operation becausethe operation is write-only and correct execution (i.e., inserting bothrecords) is not conditional on concurrent updates.

From a programmer’s perspective, we have found RA isolationto be most easily understandable (at least initially) with read-onlyand write-only transactions; after all, because RA allows concur-rent writes, any values that are read might be changed at any time.However, read-write transactions are indeed well defined under RA.

3.3 System Model and ScalabilityWe consider databases that are partitioned, with the set of items

in the database spread over multiple servers. Each item has a singlelogical copy, stored on a server—called the item’s partition—whoseidentity can be calculated using the item. Clients forward opera-tions on each item to the item’s partition, where they are executed.Transaction execution terminates in commit, signaling success, orabort, signaling failure. In our examples, all data items have the nullvalue (?) at database initialization. We do not model replication ofdata items within a partition; this can happen at a lower level of thesystem than our discussion (see Section 4.6) as long as operationson each item are linearizable [4].Scalability criteria. As we hinted in Section 1, large-scale de-ployments often eschew transactional functionality on the premisethat it would be too expensive or unstable in the presence of failureand degraded operating modes [9, 11, 13, 15, 22, 25, 26, 38, 44]. Ourgoal in this paper is to provide robust and scalable transactionalfunctionality, and, so we first define criteria for “scalability”:

Synchronization independence ensures that one client’s transactionscannot cause another client’s to block and that, if a client can con-tact the partition responsible for each item in its transaction, thetransaction will eventually commit (or abort of its own volition).This prevents one transaction from causing another to abort—whichis particularly important in the presence of partial failures—andguarantees that each client is able to make useful progress. In theabsence of failures, this maximizes useful concurrency. In the dis-tributed systems literature, synchronization independence for repli-cated transactions is called transactional availability [5]. Note that“strong” isolation models like serializability and Snapshot Isolationviolate synchronization independence and limit scalability.

While many applications can limit their data accesses to a singlepartition via explicit data modeling [7,19,25,38] or planning [18,35],this is not always possible. In the case of secondary indexing, thereis a tangible cost associated with requiring single-partition updates(scatter-gather reads), while, in social networks like Facebook andlarge-scale hierarchical access patterns as in Rainbird, perfect parti-tioning of data accesses is near-impossible. Accordingly:

Partition independence ensures that, in order to execute a transac-tion, a client never has to contact partitions that its transaction doesnot access. Thus, a partition failure only affects transactions thataccess items contained on the partition. This also reduces load onservers not directly involved in a transaction’s execution. In the dis-tributed systems literature, partition independence for replicated datais called replica availability [5] or genuine partial replication [39].

In addition to the above requirements, we limit the metadataoverhead of algorithms. There are many potential solutions forproviding atomic visibility that rely on storing prohibitive amountsof state. As a straw-man solution, each transaction could sendcopies of all of its writes to every partition it accesses so that readersobserve all of its writes by reading a single item. This provides RAisolation but requires considerable storage. Other solutions mayrequire extra data storage proportional to the number of servers inthe cluster or, worse, the database size (Section 6). We will attemptto minimize this metadata—that is, data that the transaction didnot itself write but which is required for correct execution. In ouralgorithms, we will specifically provide constant-factor metadataoverheads (RAMP-S, RAMP-H) or else overhead linear in transactionsize (but independent of data size; RAMP-F).

4. RAMP TRANSACTION ALGORITHMSGiven specifications for RA isolation and scalability, we present

algorithms for achieving both. For ease of understanding, we firstfocus on providing read-only and write-only transactions with a“last writer wins” overwrite policy, then subsequently discuss howto perform read/write transactions. Our focus in this section is onintuition and understanding; we defer all correctness and scalabilityproofs to the Appendix, providing salient details inline.

At a high level, RAMP transactions allow reads and writes toproceed concurrently. This provides excellent performance but, inturn, introduces a race condition: one transaction might only read asubset of another transaction’s writes, violating RA (i.e., fracturedreads might occur). Instead of preventing this race (hamperingscalability), RAMP readers autonomously detect the race (usingmetadata attached to each data item) and fetch any missing, in-flightwrites from their respective partitions. To make sure that readersnever have to block for writes to arrive at a partition, writers use atwo-phase (atomic commitment) protocol that ensures that once awrite is visible to readers on one partition, any other writes in thetransaction are present on and, if appropriately identified by version,readable from their respective partitions.

In this section, we present three algorithms that provide a trade-offbetween the amount of metadata required and the expected numberof extra reads to fetch missing writes. As discussed in Section 2,if techniques like distributed locking couple mutual exclusion withatomic visibility of writes, RAMP transactions correctly controlvisibility but allow concurrent and scalable execution.

4.1 RAMP-FastTo begin, we present a RAMP algorithm that, in the race-free

case, requires one RTT for reads and two RTTs for writes, calledRAMP-Fast (abbreviated RAMP-F; Algorithm 1). RAMP-F stores meta-data in the form of write sets (overhead linear in transaction size).

C1

PREPARE

PREPARE

lastCommit[x]= ∅versions={x1}

x1, md={y}

y1, md={x}lastCommit[y]= ∅

versions={y1}

COMMITtsc=1

COMMITtsc=1

GETi=x, tsreq= ∅ GET

lastCommit[x]=1versions={x1}

GET

x1, md={y}

y∅ , md={}

y1, md={x}

lastCommit[y]=1versions={y1}

vlatest ←{x:1, y:1}

prepared

RESPONSEprepared

i=y, tsreq= ∅

i=y, tsreq=1

lastCommit[x]= ∅versions={}

lastCommit[y]= ∅versions={}

BEGIN T1

committedcommitted

BEGIN T2

COMMIT T1

COMMIT T2resp={x1,y1}

RESPONSE

RESPONSE

RESPONSE

RESPONSE

RESPONSERESPONSE

C2 Px Py

[w(x1), w(y1)]

[r(x), r(y)]

Figure 1: Space-time diagram for RAMP-F execution for twotransactions T1 and T2 performed by clients C1 and C2 on par-titions Px and Py. Because T1 overlaps with T2, T2 must per-form a second round of reads to repair the fractured read be-tween x and y. T1’s writes are assigned timestamp 1. Lightly-shaded boxes represent current partition state (lastCommit andversions), while the single darkly-shaded box encapsulates allmessages exchanged during C2’s execution of transaction T2.

Overview. Each write in RAMP-F (lines 14–21) contains a times-tamp (line 15) that uniquely identifies the writing transaction aswell as a set of items written in the transaction (line 16). For now,combining a unique client ID and client-local sequence number issufficient for timestamp generation (see also Section 4.5).

RAMP-F write transactions proceed in two phases: a first roundof communication places each timestamped write on its respectivepartition. In this PREPARE phase, each partition adds the write to itslocal database (versions, lines 1, 17–19). A second round of com-munication marks versions as committed. In this COMMIT phase,each partition updates an index containing the highest-timestampedcommitted version of each item (lastCommit, lines 2, 20–21).

RAMP-F read transactions begin by first fetching the last (highest-timestamped) committed version for each item from its respectivepartition (lines 23–30). Using the results from this first round ofreads, each reader can calculate whether it is “missing” any versions(that is, versions that were prepared but not yet committed on theirpartitions). Combining the timestamp and set of items from eachversion read (i.e., its metadata) produces a mapping from items totimestamps that represent the highest-timestamped write for eachtransaction that appears in this first-round read set (lines 26–29). If

the reader has read a version of an item that has a lower timestampthan indicated in the mapping for that item, the reader issues asecond read to fetch the missing version (by timestamp) from itspartition (lines 30–32). Once all missing versions are fetched (whichcan be done in parallel), the client can return the resulting set ofversions—the first-round reads, with any missing versions replacedby the optional, second round of reads.By example. Consider the RAMP-F execution depicted in Figure 1.T1 writes to both x and y, performing the two-round write protocolon two partitions, Px and Py. However, T2 reads from x and y whileT1 is concurrently writing. Specifically, T2 reads from Px afterPx has committed T1’s write to x, but T2 reads from Py before Pyhas committed T1’s write to y. Therefore, T2’s first-round readsreturn x = x1 and y = ?, and returning this set of reads wouldviolate RA. Using the metadata attached to its first-round reads,T2 determines that it is missing y1 (since vlatest [y] = 1 and 1 >?)and so T2 subsequently issues a second read from Py to fetch y1 byversion. After completing its second-round read, T2 can safely returnits result set. T1’s progress is unaffected by T2, and T1 subsequentlycompletes by committing y1 on Py.Why it works. RAMP-F writers use metadata as a record of intent:a reader can detect if it has raced with an in-progress commit roundand use the metadata stored by the writer to fetch the missing data.Accordingly, RAMP-F readers only issue a second round of reads inthe event that they read from a partially-committed write transaction(where some but not all partitions have committed a write). In thisevent, readers will fetch the appropriate writes from the not-yet-committed partitions. Most importantly, RAMP-F readers never haveto stall waiting for a write that has not yet arrived at a partition:the two-round RAMP-F write protocol guarantees that, if a partitioncommits a write, all of the corresponding writes in the transaction arepresent on their respective partitions (though possibly not committedlocally). As long as a reader can identify the corresponding versionby timestamp, the reader can fetch the version from the respectivepartition’s set of pending writes without waiting. To enable this,RAMP-F writes contain metadata linear in the size of the writingtransaction’s write set (plus a timestamp per write).

RAMP-F requires 2 RTTs for writes: one for PREPARE and one forCOMMIT. For reads, RAMP-F requires one RTT in the absence ofconcurrent writes and two RTTs otherwise.

RAMP timestamps are only used to identify specific versions andin ordering concurrent writes to the same item; RAMP-F transactionsdo not require a “global” timestamp authority. For example, iflastCommit[k] = 2, there is no requirement that a transaction withtimestamp 1 has committed or even that such a transaction exists.

4.2 RAMP-Small: Trading Metadata for RTTsWhile RAMP-F requires linearly-sized metadata but provides best-

case one RTT for reads, RAMP-Small (RAMP-S) uses constant-sizemetadata but always requires two RTT for reads (Algorithm 2).RAMP-S and RAMP-F writes are identical, but, instead of attachingthe entire write set to each write, RAMP-S writers only store thetransaction timestamp (line 7). Unlike RAMP-F, RAMP-S readersissue a first round of reads to fetch the highest committed timestampfor each item from its respective partition (lines 3, 9–11). OnceRAMP-S readers have recieved the highest committed timestamp foreach item, the readers send the entire set of timestamps they receivedto the partitions in a second round of communication (lines 13–14).For each item in the read request, RAMP-S servers return the highest-timestamped version of the item that also appears in the supplied setof timestamps (lines 5–6). Readers subsequently return the resultsfrom the mandatory second round of requests.

Algorithm 1 RAMP-FastServer-side Data Structures

1: versions: set of versions hitem,value, timestamp tsv, metadata mdi2: latestCommit[i]: last committed timestamp for item i

Server-side Methods

3: procedure PREPARE(v : version)4: versions.add(v)5: return6: procedure COMMIT(tsc : timestamp)7: Its {w.item | w 2 versions^w.tsv = tsc}8: 8i 2 Its, latestCommit[i] max(latestCommit[i], tsc)

9: procedure GET(i : item, tsreq : timestamp)10: if tsreq = /0 then11: return v 2 versions : v.item = i^ v.tsv = latestCommit[item]12: else13: return v 2 versions : v.item = i^ v.tsv = tsreq

Client-side Methods

14: procedure PUT_ALL(W : set of hitem,valuei)15: tstx generate new timestamp16: Itx set of items in W17: parallel-for hi,vi 2W18: v hitem = i,value = v, tsv = tstx,md = (Itx�{i})i19: invoke PREPARE(v) on respective server (i.e., partition)20: parallel-for server s : s contains an item in W21: invoke COMMIT(tstx) on s

22: procedure GET_ALL(I : set of items)23: ret {}24: parallel-for i 2 I25: ret[i] GET(i, /0)26: vlatest {} (default value: �1)27: for response r 2 ret do28: for itx 2 r.md do29: vlatest [itx] max(vlatest [itx],r.tsv)

30: parallel-for item i 2 I31: if vlatest [i]> ret[i].tsv then32: ret[i] GET(i,vlatest [i])

33: return ret

By example. In Figure 1, under RAMP-S, Px and Py would respec-tively return the sets {1} and {?} in response to T2’s first round ofreads. T2 would subsequently send the set {1,?} to both Px and Py,which would return x1 and y1. (Including ? in the second-roundrequest is unnecessary, but we leave it in for ease of understanding.)Why it works. In RAMP-S, if a transaction has committed on somebut not all partitions, the transaction timestamp will be returnedin the first round of any concurrent read transaction accessing thecommitted partitions’ items. In the (required) second round of readrequests, any prepared-but-not-committed partitions will find thecommitted timestamp in the reader-provided set and return the ap-propriate version. In contrast with RAMP-F, where readers explicitlyprovide partitions with a specific version to return in the (optional)second round, RAMP-S readers defer the decision of which versionto return to the partition, which uses the reader-provided set to de-cide. This saves metadata but increases RTTs, and the size of theparameters of each second-round GET request is (worst-case) linearin the read set size. Unlike RAMP-F, there is no requirement to returnthe value of the last committed version in the first round (returningthe version, lastCommit[k], suffices in line 3).

4.3 RAMP-Hybrid: An Intermediate SolutionRAMP-Hybrid (RAMP-H; Algorithm 3) strikes a compromise be-

tween RAMP-F and RAMP-S. RAMP-H and RAMP-S write protocols areidentical, but, instead of storing the entire write set (as in RAMP-F),

Algorithm 2 RAMP-SmallServer-side Data Structures

same as in RAMP-F (Algorithm 1)

Server-side Methods

PREPARE, COMMIT same as in RAMP-F

1: procedure GET(i : item, tsset : set of timestamps)2: if tsset = /0 then3: return v 2 versions : v.item = i^ v.tsv = latestCommit[k]4: else5: tsmatch = {t | t 2 tsset ^9v 2 versions : v.item = i^ v.tv = t}6: return v 2 versions : v.item = i^ v.tsv = max(tsmatch)

Client-side Methods

7: procedure PUT_ALL(W : set of hitem,valuei)same as RAMP-F PUT_ALL but do not instantiate md on line 18

8: procedure GET_ALL(I : set of items)9: tsset {}10: parallel-for i 2 I11: tsset .add(GET(i, /0).tsv)12: ret {}13: parallel-for item i 2 I14: ret[i] GET(i, tsset)

15: return ret

RAMP-H writers store a Bloom filter [10] representing the transactionwrite set (line 1). RAMP-H readers proceed as in RAMP-F, with afirst round of communication to fetch the last-committed version ofeach item from its partition (lines 3–5). Given this set of versions,RAMP-H readers subsequently compute a list of potentially higher-timestamped writes for each item (lines 7–10). Any potentiallymissing versions are fetched in a second round of reads (lines 12).By example. In Figure 1, under RAMP-H, x1 would contain aBloom filter with positives for x and y and y? would contain anempty Bloom filter. T2 would check for the presence of y in x1’sBloom filter (since x1’s version is 1 and 1 > ?) and, finding amatch, conclude that it is potentially missing a write (y1). T2 wouldsubsequently fetch y1 from Py.Why it works. RAMP-H is effectively a hybrid between RAMP-Fand RAMP-S. If the Bloom filter has no false positives, RAMP-H readsbehave like RAMP-F reads. If the Bloom filter has all false positives,RAMP-H reads behave like RAMP-S reads. Accordingly, the numberof (unnecessary) second-round reads (i.e., which would not be per-formed by RAMP-F) is controlled by the Bloom filter false positiverate, which is in turn (in expectation) proportional to the size of theBloom filter. Any second-round GET requests are accompanied by aset of timestamps that is also proportional in size to the false positiverate. Therefore, RAMP-H exposes a trade-off between metadata sizeand expected performance. To understand why RAMP-H is safe, wesimply have to show that any false positives (second-round reads)will not compromise the integrity of the result set; with uniquetimestamps, any reads due to false positives will return null.

4.4 Summary of Basic AlgorithmsThe RAMP algorithms allow readers to safely race writers with-

out requiring either to stall. The metadata attached to each writeallows readers in all three algorithms to safely handle concurrentand/or partial writes and in turn allows a trade-off between metadatasize and performance (Table 1): RAMP-F is optimized for fast reads,RAMP-S is optimized for small metadata, and RAMP-H is, as the namesuggests, a middle ground. RAMP-F requires metadata linear in trans-action size, while RAMP-S and RAMP-H require constant metadata.However, RAMP-S and RAMP-H require more RTTs for reads com-pared to RAMP-F when there is no race between readers and writers.

Algorithm 3 RAMP-HybridServer-side Data Structures

Same as in RAMP-F (Algorithm 1)

Server-side Methods

PREPARE, COMMIT same as in RAMP-FGET same as in RAMP-S

Client-side Methods

1: procedure PUT_ALL(W : set of hitem,valuei)same as RAMP-F PUT_ALL but instantiate md on line 18with Bloom filter containing Itx

2: procedure GET_ALL(I : set of items)3: ret {}4: parallel-for i 2 I5: ret[i] GET(i, /0)6: v f etch {}7: for version v 2 ret do8: for version v0 2 ret : v0 6= v do9: if v.tsv > v0.tsv ^ v.md.lookup(v0.item)! True then10: v f etch[v0.item].add(v.tsv)11: parallel-for item i 2 v f etch12: ret[i] GET(k,v f etch[i]) if GET(k,v f etch[i]) 6=?13: return ret

Algorithm RTTs/transaction Metadata (+stamp)W R (stable) R (O) Stored Per-Request

RAMP-F 2 1 2 txn items -RAMP-S 2 2 2 - stamp/itemRAMP-H 2 1+ e 2 Bloom filter stamp/item

Table 1: Comparison of basic algorithms: RTTs required forwrites (W), reads (R) without concurrent writes and in theworst case (O), stored metadata and metadata attached to readrequests (in addition to a timestamp for each).

When reads and writes race, in the worst case, all algorithms requiretwo RTTs for reads. Writes always require two RTTs to preventreaders from stalling due to missing, unprepared writes.

RAMP algorithms are scalable because clients only contact par-titions relative to their transactions (partition independence), andclients cannot stall one another (synchronization independence).More specifically, readers do not interfere with other readers, writ-ers do not interfere with other writers, and readers and writers canproceed concurrently. When a reader races a writer to the sameitems, the writer’s new versions will only become visible to thereader (i.e., be committed) once it is guaranteed that the reader willbe able to fetch all of them (possibly via a second round of com-munication). A reader will never have to stall waiting for writes toarrive at a partition (for details, see Invariant 1 in the Appendix).

4.5 Additional DetailsIn this section, we discuss relevant implementation details.

Multi-versioning and garbage collection. RAMP transactionsrely on multi-versioning to allow readers to access versions thathave not yet committed and/or have been overwritten. In our initialpresentation, we have used a completely multi-versioned storageengine; in practice, multi-versioning can be implemented by usinga single-versioned storage engine for retaining the last committedversion of each item and using a “look-aside” store for access toboth prepared-but-not-yet-committed writes and (temporarily) anyoverwritten versions. The look-aside store should make preparedversions durable but can—at the risk of aborting transactions in theevent of a server failure—simply store any overwritten versions in

memory. Thus, with some work, RAMP algorithms are portable tolegacy, non-multi-versioned storage systems.

In both architectures, each partition’s data will grow withoutbound if old versions are not removed. If a committed version ofan item is not the highest-timestamped committed version (i.e., acommitted version v of item k where v < lastCommit[k]), it can besafely discarded (i.e., garbage collected, or GCed) as long as noreaders will attempt to access it in the future (via second-roundGET requests). It is easiest to simply limit the running time of readtransactions and GC overwritten versions after a fixed amount of realtime has elapsed. Any read transactions that take longer than this GCwindow can be restarted [32, 33]. Therefore, the maximum numberof versions retained for each item is bounded by the item’s updaterate, and servers can reject any client GET requests for versions thathave been GCed (and the read transaction can be restarted). As amore principled solution, partitions can also gossip the timestampsof items that have been overwritten and have not been returned inthe first round of any ongoing read transactions.Read-write transactions. Until now, we have focused on read-only and write-only transactions. However, we can extend ouralgorithms to provide read-write transactions. If transactions pre-declare the data items they wish to read, then the client can executea GET_ALL transaction at the start of transaction execution to pre-fetch all items; subsequent accesses to those items can be servedfrom this pre-fetched set. Clients can buffer any writes and, upontransaction commit, send all new versions to servers (in parallel) viaa PUT_ALL request. As in Section 3, this may result in anomaliesdue to concurrent update but does not violate RA isolation. Giventhe benefits of pre-declared read/write sets [18, 35, 43] and writebuffering [17, 41], we believe this is a reasonable strategy. Forsecondary index lookups, clients can first look up secondary indexentries then subsequently (within the same transaction) read primarydata (specifying versions from index entries as appropriate).Timestamps. Timestamps should be unique across transactions,and, for “session” consistency (Appendix), increase on a per-clientbasis. Given unique client IDs, a client ID and sequence numberform unique transaction timestamps without coordination. Withoutunique client IDs, servers can assign unique timestamps with highprobability using UUIDs and by hashing transaction contents.Overwrites. In our algorithms, we have depicted a policy in whichversions are overwritten according to a highest-timestamp-winspolicy. In practice, and, for commutative updates, users may wishto employ a different policy upon COMMIT: for example, performset union. In this case, lastCommit[k] contains an abstract data type(e.g., set of versions) that can be updated with a merge operation [22,42] (instead of updateI f Greater) upon commit. This treats eachcommitted record as a set of versions, requiring additional metadata(that can be GCed as in Section 4.7).

4.6 Distribution and Fault ToleranceRAMP transactions operate in a distributed setting, which poses

challenges due to latency, partial failure, and network partitions.Synchronization independence ensures that failed clients do notcause other clients to fail, while partition independence ensures thatclients only have to contact partitions for items in their transactions.This provides fault tolerance and availability as long as clients canaccess relevant partitions, but here we further elucidate RAMPinteractions with replication and stalled operations.Replication. A variety of mechanisms including traditional databasemaster-slave replication with failover, quorum-based protocols, andstate machine replication and can ensure availability of individualpartitions in the event of individual server failure [8]. To control

durability, clients can wait until the effects of their operations (e.g.,modifications to versions and lastCommit) are persisted locally ontheir respective partitions and/or to multiple physical servers beforereturning from PUT_ALL calls (either via master-to-slave replicationor via quorum replication and by performing two-phase commitacross multiple active servers). Notably, because RAMP trans-actions can safely overlap in time, replicas can process differenttransactions’ PREPARE and COMMIT requests in parallel.Stalled Operations. RAMP writes use a two-phase atomic com-mitment protocol that ensures readers never block waiting for writesto arrive. As discussed in Section 2, every ACP may block dur-ing failures [8]. However, due to synchronization independence, ablocked transaction (due to failed clients, failed servers, or networkpartitions) cannot cause other transactions to block. Blocked writesinstead act as “resource leaks” on partitions: partitions will retainprepared versions indefinitely unless action is taken.

To “free” these leaks, RAMP servers can use the CooperativeTermination Protocol (CTP) described in [8]. CTP can alwayscomplete the transaction except when every partition has performedPREPARE but no partition has performed COMMIT. In CTP, if aserver Sp has performed PREPARE for transaction T but times outwaiting for a COMMIT, Sp can check the status of T on any otherpartitions for items in T ’s write set. If another server Sc has receivedCOMMIT for T , then Sp can COMMIT T . If Sa, a server responsiblefor an item in T , has not received PREPARE for T , Sa and Sp canpromise never to PREPARE or COMMIT T in the future and Sp cansafely discard its versions. A client recovering from a failure canread from the servers to determine if they unblocked its write. Writesthat block mid-COMMIT will also become visible on all partitions.

CTP (evaluated in Section 5) only runs when writes block (ortime-outs fire) and runs asynchronously with respect to other op-erations. CTP requires that PREPARE messages contain a list ofservers involved in the transaction (a subset of RAMP-F metadata buta superset of RAMP-H and RAMP-S) and that servers remember whenthey COMMIT and “abort” writes (e.g., in a log file). Compared toalternatives (e.g., replicating clients [24]), we have found CTP to beboth lightweight and effective.

4.7 Further OptimizationsRAMP algorithms also allow several possible optimizations:Faster commit detection. If a server returns a version in responseto a GET request and the version’s timestamp is greater than thehighest committed version of that item (i.e., lastCommit), then trans-action writing the version has committed on at least one partition.In this case, the server can mark the version as committed. Thisscenario will occur when all partitions have performed PREPAREand at least one server but not all partitions have performed COMMIT(as in CTP). This allows faster updates to lastCommit (and thereforefewer expected RAMP-F and RAMP-H RTTs).Metadata garbage collection. Once all of transaction T ’s writesare committed on each respective partition (i.e., are reflected inlastCommit), readers are guaranteed to read T ’s writes (or laterwrites). Therefore, non-timestamp metadata for T ’s writes stored inRAMP-F and RAMP-H (write sets and Bloom filters) can therefore bediscarded. Detecting that all servers have performed COMMIT canbe performed asynchronously via a third round of communicationperformed by either clients or servers.One-phase writes. We have considered two-phase writes, but, ifa user does not wish to read her writes (thereby sacrificing sessionguarantees outlined in the Appendix), the client can return afterissuing its PREPARE round (without sacrificing durability). Theclient can subsequently execute the COMMIT phase asynchronously,

or, similar to optimizations presented in Paxos Commit [24], theservers can exchange PREPARE acknowledgements with one anotherand decide to COMMIT autonomously. This optimization is safebecause multiple PREPARE phases can safely overlap.

5. EXPERIMENTAL EVALUATIONWe proceed to experimentally demonstrate RAMP transaction

scalability as compared to existing transactional and non-transactionalmechanisms. RAMP-F, RAMP-H, and often RAMP-S outperform exist-ing solutions across a range of workload conditions while exhibit-ing overheads typically within 8% and no more than 48% of peakthroughput. As expected from our theoretical analysis, the per-formance of our RAMP algorithms does not degrade substantiallyunder contention and scales linearly to over 7.1 million operationsper second on 100 servers. These outcomes validate our choice topursue synchronization- and partition-independent algorithms.

5.1 Experimental SetupTo demonstrate the effect of concurrency control on performance

and scalability, we implemented several concurrency control algo-rithms in a partitioned, multi-versioned, main-memory databaseprototype. Our prototype is in Java and employs a custom RPCsystem with Kryo 2.20 for serialization. Servers are arranged asa distributed hash table with partition placement determined byrandom hashing. As in stores like Dynamo [22], clients can con-nect to any server to execute operations, which the server will per-form on their behalf (i.e., each server acts as a client in our RAMPpseudocode). We implemented RAMP-F, RAMP-S, and RAMP-H andconfigure a wall-clock GC window of 5 seconds as described inSection 4.5. RAMP-H uses a 256-bit Bloom filter based on an im-plementation of MurmurHash2.0, with four hashes per entry; todemonstrate the effects of filter saturation, we do not modify theseparameters in our experiments. Our prototype utilizes the “Fastercommit detection” optimization from Section 4.5 but we chose notto employ the latter two optimizations in order to preserve sessionguarantees and because metadata overheads were generally minor.Algorithms for comparison. As a baseline, we do not employ anyconcurrency control (denoted NWNR, for no write and no read locks);reads and writes take one RTT and are executed in parallel.

We also consider three lock-based mechanisms: long write locksand long read locks, providing Repeatable Read isolation (PL-2.99;denoted LWLR), long write locks with short read locks, providingRead Committed isolation (PL-2L; denoted LWSR; does not provideRA), and long write locks with no read locks, providing Read Un-committed isolation [2] (LWNR; also does not provide RA). Whileonly LWLR provides RA, LWSR and LWNR provide a useful basis forcomparison, particularly in measuring concurrency-related lockingoverheads. To avoid deadlocks, the system lexicographically orderslock requests by item and performs them sequentially. When locksare not used (as for reads in LWNR and reads and writes for NWNR),the system parallelizes operations.

We also consider an algorithm where, for each transaction, des-ignated “coordinator” servers enforce RA isolation—effectively,the Eiger system’s 2PC-PCI mechanism [33] (denoted E-PCI; Sec-tion 6). Writes proceed via prepare and commit rounds, but any readsthat arrive at a partition and overlap with a concurrent write to thesame item must contact a (randomly chosen, per-write-transaction)“coordinator” partition to determine whether the coordinator’s pre-pared writes have been committed. Writes require two RTTs, whilereads require one RTT during quiescence and two RTTs in the pres-ence of concurrent updates (to a variable number of coordinatorpartitions—linear in the number of concurrent writes to the item).

Using a coordinator violates partition independence but not syn-chronization independence. We optimize 2PC-PCI reads by havingclients determine a read timestamp for each transaction (eliminatingan RTT) and do not include happens-before metadata.

This range of lock-based strategies (LWNR, LWSR, LWNR), recentcomparable approach (E-PCI), and best-case (NWNR; no concurrencycontrol) baseline provides a spectrum of strategies for comparison.Environment and benchmark. We evaluate each algorithm us-ing the YCSB benchmark [16] and deploy variably-sized sets ofservers on public cloud infrastructure. We employ cr1.8xlargeinstances on Amazon EC2 and, by default, deploy five partitions onfive servers. We group sets of reads and sets of writes into read-onlyand write-only transactions (default size: 4 operations), and use thedefault YCSB workload (workloada, with Zipfian distributed itemaccesses) but with a 95% read and 5% write proportion, reflectingread-heavy applications (Section 2, [11, 33, 44]; e.g., Tao’s 500 to1 reads-to-writes [11, 33], Espresso’s 1000 to 1 Mailbox applica-tion [38], and Spanner’s 3396 to 1 advertising application [17]).

By default, we use 5000 concurrent clients split across 5 separateEC2 instances and, to fully expose our metadata overheads, use avalue size of 1 byte per write. We found that lock-based algorithmswere highly inefficient for YCSB’s default 1K item database, so weincreased the database size to 1M items by default. Each versioncontains a timestamp (64 bits), and, with YCSB keys (i.e., item IDs)of size 11 bytes and a transaction length L, RAMP-F requires 11Lbytes of metadata per version, while RAMP-H requires 32 bytes. Wesuccessively vary several parameters, including number of clients,read proportion, transaction length, value size, database size, andnumber of servers and report the average of three sixty-second trials.

5.2 Experimental Results: ComparisonOur first set of experiments focuses on two metrics: performance

compared to baseline and performance compared to existing tech-niques. The overhead of RAMP algorithms is typically less than8% compared to baseline (NWNR) throughput, is sometimes zero, andis never greater than 50%. RAMP-F and RAMP-H always outperformthe lock-based and E-PCI techniques, while RAMP-S outperformslock-based techniques and often outperforms E-PCI. We proceed todemonstrate this behavior over a variety of conditions:Number of clients. RAMP performance scales well with increasedload and incurs little overhead (Figure 2). With few concurrentclients, there are few concurrent updates and therefore few second-round reads; performance for RAMP-F and RAMP-H is close to oreven matches that of NWNR. At peak throughput (at 10,000 clients),RAMP-F and RAMP-H pay a throughput overhead of 4.2% comparedto NWNR. RAMP-F and RAMP-H exhibit near-identical performance; theRAMP-H Bloom filter triggers few false positives (and therefore fewextra RTTs compared to RAMP-F). RAMP-S incurs greater overheadand peaks at almost 60% of the throughput of NWNR. Its guaranteedtwo-round trip reads are expensive and it acts as an effective lowerbound on RAMP-F and RAMP-H performance. In all configurations,the algorithms achieve low latency (RAMP-F, RAMP-H, NWNR less than35ms on average and less than 10 ms at 5,000 clients; RAMP-S lessthan 53ms, 14.3 ms at 5,000 clients).

In comparison, the remaining algorithms perform less favorably.In contrast with the RAMP algorithms, E-PCI servers must check acoordinator server for each in-flight write transaction to determinewhether to reveal writes to clients. For modest load, the overheadof these commit checks places E-PCI performance between thatof RAMP-S and RAMP-H. However, the number of in-flight writesincreases with load (and is worsened due to YCSB’s Zipfian dis-tributed accesses), increasing the number of E-PCI commit checks.

RAMP-F RAMP-S RAMP-H NWNR LWNR LWSR LWLR E-PCI

0 2000 4000 6000 8000 100000

30K

60K

90K

120K

150K

180K

Thro

ughp

ut(tx

n/s)

2000 4000 6000 8000 10000Concurrent Clients

10�1

100

101

102

Avg.

Late

ncy

(ms)

Figure 2: Throughput and latency under varying client load.We omit latencies for LWLR, which peaked at over 1.5s.

This in turn decreases throughput, and, with 10,000 concurrentclients, E-PCI performs so many commit checks per read (over 20%of reads trigger a commit check, and, on servers with hot items,each commit check requires indirected coordinator checks for anaverage of 9.84 transactions) that it underperforms the LWNR lock-based scheme. Meanwhile, multi-partition locking is expensive [35]:with 10,000 clients, the most efficient algorithm, LWNR, attains only28.6% of the throughput of NWNR, while the least efficient, LWLR,attains only 1.6% (peaking at 3,412 transactions per second).

We subsequently varied several other workload parameters, whichwe briefly discuss below and plot in Figure 3:Read proportion. Increased write activity leads to a greaternumber of races between reads and writes and therefore additionalsecond-round RTTs for RAMP-F and RAMP-H reads. With all writetransactions, all RAMP algorithms are equivalent (two RTT) andachieve approximately 65% of the throughput of NWNR. With allreads, RAMP-F, RAMP-S, NWNR, and E-PCI are identical, with a sin-gle RTT. Between these extremes, RAMP-F and RAMP-S scale near-linearly with the write proportion. In contrast, lock-based protocolsfare poorly as contention increases, while E-PCI again incurs penal-ties due to commit checks.Transaction length. Increased transaction lengths have variableimpact on the relative performance of RAMP algorithms. Synchro-nization independence does not penalize long-running transactions,but, with longer transactions, metadata overheads increase. RAMP-Frelative throughput decreases due to additional metadata (linear intransaction length) and RAMP-H relative performance also decreasesas its Bloom filters saturate. (However, YCSB’s Zipfian-distributedaccess patterns result in a non-linear relationship between lengthand throughput.) As discussed above, we explicitly decided not totune RAMP-H Bloom filter size but believe a logarithmic increase infilter size could improve RAMP-H performance for large transactionlengths (e.g., 1024 bit filters should lower the false positive rate fortransactions of length 256 from over 92% to slightly over 2%).Value size. Value size similarly does not seriously impact relativethroughput. At a value size of 1B, RAMP-F is within 2.3% of NWNR.However, at a value size of 100KB, RAMP-F performance nearlymatches that of NWNR: the overhead due to metadata decreases, andwrite request rates slow, decreasing concurrent writes (and subse-

0 25 50 75 100Percentage Reads)

0

30K

60K

90K

120K

150K

180K

Thro

ughp

ut(tx

n/s)

2 4 8 16 32 64 128Transaction Size (operations)

0

30K

60K

90K

120K

150K

180K

Thro

ughp

ut(tx

n/s)

1 10 100 1000 10000 100000Value Size (bytes)

0

30K

60K

90K

120K

150K

180K

Thro

ughp

ut(tx

n/s)

10 100 1000 10K 100K 1M 10MDatabase Size (items)

0

30K

60K

90K

120K

150K

180K

Thro

ughp

ut(tx

n/s)

Figure 3: Algorithm performance across varying workload con-ditions. RAMP-F and RAMP-H exhibit similar performance to NWNRbaseline, while RAMP-S’s 2 RTT reads incur a greater perfor-mance penalty across almost all configurations. RAMP trans-actions consistently outperform RA isolated alternatives.

quently second-round RTTs). Nonetheless, absolute throughputdrops by a factor of 24 as value sizes moves from 1B to 100KB.Database size. RAMP algorithms are robust to high contention fora small set of items: with only 1000 items in the database, RAMP-Fachieves throughput within 3.1% of NWNR. RAMP algorithms arelargely agnostic to read/write contention, although, with fewer itemsin the database, the probability of races between readers and in-progress writers increases, resulting in additional second-roundreads for RAMP-F and RAMP-H. In contrast, lock-based algorithmsfare poorly under high contention, while E-PCI indirected commitchecks again incurred additional overhead. By relying on clients(rather than additional partitions) to repair fractured writes, RAMP-F,RAMP-H, and RAMP-S performance is less affected by hot items.

Overall, RAMP-F and RAMP-H exhibit performance close to that ofno concurrency control due to their independence properties andguaranteed worst-case performance. As the proportion of writes

NWNR RAMP-F RAMP-H RAMP-S

0 25 50 75 1000

2M

4M

6M

8MTh

roug

hput

(ops

/s)

0 25 50 75 100Number of Servers

040K80K

120K160K200K

oper

atio

ns/s

/ser

ver

Figure 4: RAMP transactions scale linearly to over 7 millionoperations/s with comparable performance to NWNR baseline.

increases, an increasing proportion of RAMP-F and RAMP-H opera-tions take two RTTs and performance trends towards that of RAMP-S,which provides a constant two RTT overhead. In contrast, lock-based protocols perform poorly under contention while E-PCI trig-gers more commit checks than RAMP-F and RAMP-H trigger secondround reads (but still performs well without contention and for par-ticularly read-heavy workloads). The ability to allow clients toindependently verify read sets enables good performance despite arange of (sometimes adverse) conditions (e.g., high contention).

5.3 Experimental Results: CTP OverheadWe also evaluated the overhead of blocked writes in our imple-

mentation of the Cooperative Termination Protocol discussed inSection 4.6. To simulate blocked writes, we artificially droppeda percentage of COMMIT commands in PUT_ALL calls such thatclients returned from writes early and partitions were forced to com-plete the commit via CTP. This behavior is worse than expectedbecause “blocked” clients continue to issue new operations. Thetable below reports the throughput reduction as the proportion ofblocked writes increases (compared to no blocked writes) for aworkload of 100% RAMP-F write transactions:

Blocked % 0.01% 0.1% 25% 50%Throughput No change 99.86% 77.53% 67.92%

As these results demonstrate, CTP can reduce throughput becauseeach commit check consumes resources (here, network and CPUcapacity). However, CTP only performs commit checks in the eventof blocked writes (or time-outs; set to 5s in our experiments), soa modest failure rate of 1 in 1000 writes has a limited effect. Thehigher failure rates produce a near-linear throughput reduction but,in practice, a blocking rate of even a few percent is likely indicativeof larger systemic failures. As Figure 3 hints, the effect of additionalmetadata for the participant list in RAMP-H and RAMP-S is limited,and, for our default workload of 5% writes, we observe similartrends but with throughput degradation of 10% or less across theabove configurations. This validates our initial motivation behindthe choice of CTP: average-case overheads are small.

5.4 Experimental Results: ScalabilityWe finally validate our chosen scalability criteria by demonstrat-

ing linear scalability of RAMP transactions to 100 servers. Wedeployed an increasing number of servers within the us-west-2EC2 region and, to mitigate the effects of hot items during scal-ing, configured uniform random access to items. We were unable

to include more than 20 instances in an EC2 “placement group,”which guarantees 10 GbE connections between instances, so, past 20servers, servers communicated over a degraded network. Around 40servers, we exhausted the us-west-2b “availability zone” (datacen-ter) capacity and had to allocate our instances across the remainingzones, further degrading network performance. However, as shownin Figure 4, each RAMP algorithm scales linearly, even though inexpectation, at 100 servers, all but one in 100M transactions is amulti-partition operation. In particular, RAMP-F achieves slightlyunder 7.1 million operations per second, or 1.79 million transactionsper second on a set of 100 servers (71,635 operations per partitionper second). At all scales, RAMP-F throughput was always within10% of NWNR. With 100 servers, RAMP-F was within 2.6%, RAMP-Swithin 3.4%, and RAMP-S was within 45% of NWNR. In light of ourscalability criteria, this behavior is unsurprising.

6. RELATED WORKReplicated databases offer a broad spectrum of isolation guaran-

tees at varying costs to performance and availability [8]:Serializability. At the strong end of the isolation spectrum isserializability, which provides transactions with the equivalent ofa serial execution (and therefore also provides RA). A range oftechniques can enforce serializability in distributed databases [3, 8],multi-version concurrency control (e.g. [37]) locking (e.g. [31]), andoptimistic concurrency control [41]. These useful semantics comewith costs in the form of decreased concurrency (e.g., contentionand/or failed optimistic operations) and limited availability duringpartial failure [5, 21]. Many designs [19, 29] exploit cheap serializ-ability within a single partition but face scalability challenges fordistributed operations. Recent industrial efforts like F1 [41] andSpanner [17] have improved performance via aggressive hardwareadvances but, their reported throughput is still limited to 20 and 250writes per item per second. Multi-partition serializable transactionsare expensive and, especially under adverse conditions, are likely toremain expensive [18, 28, 35].Weak isolation. The remainder of the isolation spectrum is morevaried. Most real-world databases offer (and often default to) non-serializable isolation models [5, 34]. These “weak isolation” levelsallow greater concurrency and fewer system-induced aborts com-pared to serializable execution but provide weaker semantic guaran-tees. For example, the popular choice of Snapshot Isolation preventsLost Update anomalies but not Write Skew anomalies [2]; by pre-venting Lost Update, concurrency control mechanisms providingSnapshot Isolation violate synchronization independence [5]. Inrecent years, many “NoSQL” designs have avoided cross-partitiontransactions entirely, effectively providing Read Uncommitted isola-tion in many industrial databases such PNUTS [15], Dynamo [22],TAO [11], Espresso [38], Rainbird [44], and BigTable [13]. Thesesystems avoid penalties associated with stronger isolation but in turnsacrifice transactional guarantees (and therefore do not offer RA).Related mechanisms. There are several algorithms that are closelyrelated to our choice of RA and RAMP algorithm design.

COPS-GT’s two-round read-only transaction protocol [32] issimilar to RAMP-F reads—client read transactions identify causallyinconsistent versions by timestamp and fetch them from servers.While COPS-GT provides causal consistency (requiring additionalmetadata), it does not support RA isolation for multi-item writes.

Eiger provides its write-only transactions [33] by electing a coor-dinator server for each write. As discussed in Section 5 (E-PCI), thenumber of “commit checks” performed during its read-only trans-actions is proportional to the number of concurrent writes. Usinga coordinator violates partition independence but in turn provides

causal consistency. This coordinator election is analogous to G-Store’s dynamic key grouping [19] but with weaker isolation guaran-tees; each coordinator effectively contains a partitioned completedtransaction list from [12]. Instead of relying on indirection, RAMPtransaction clients autonomously assemble reads and only requireconstant factor (or, for RAMP-F, linear in transaction size) metadatasize compared to Eiger’s PL-2L (worst-case linear in database size).

RAMP transactions are inspired by our earlier proposal for Mono-tonic Atomic View (MAV) isolation: transactions read from a mono-tonically advancing view of database state [5]. MAV is strictlyweaker than RA and does not prevent fractured reads, as requiredfor our applications (i.e., reads are not guaranteed to be transac-tionally aligned). The prior MAV algorithm we briefly sketchedin [5] is similar to RAMP-F but, as a consequence of its weakersemantics, allows one-round read transactions. The RAMP algo-rithms described here are portable to the highly available (i.e., non-linearizable, “AP/EL” [1, 23]) replicated setting of [5], albeit withnecessary penalties to latency between updates and their visibility.

Overall, we are not aware of a concurrency control mechanism forpartitioned databases that provides synchronization independence,partition independence, and at least RA isolation.

7. CONCLUSIONThis paper described how to achieve atomically visible multi-

partition transactions without incurring the performance and avail-ability penalties of traditional algorithms. We first identified a newisolation level—Read Atomic isolation—that provides atomic visi-bility and matches the requirements of a large class of real-worldapplications. We subsequently achieved RA isolation via scalable,contention-agnostic RAMP transactions. In contrast with techniquesthat use inconsistent but fast updates, RAMP transactions providecorrect semantics for applications requiring secondary indexing, for-eign key constraints, and materialized view maintenance while main-taining scalability and performance. By leveraging multi-versioningwith a variable but small (and, in two of three algorithms, constant)amount of metadata per write, RAMP transactions allow clients todetect and assemble atomic sets of versions in one to two rounds ofcommunication with servers (depending on the RAMP implemen-tation). The choice of synchronization and partition independentalgorithms allowed us to achieve near-baseline performance across avariety of workload configurations and scale linearly to 100 servers.While RAMP transactions are not appropriate for all applications,the many for which they are well suited will benefit measurably.

Acknowledgments The authors would like to thank Peter Alvaro,Giselle Cheung, Neil Conway, Aaron Davidson, Mike Franklin, Au-rojit Panda, Nuno Preguiça, Edward Ribeiro, Shivaram Venkatara-man, and the SIGMOD reviewers for their insightful feedback.This research is supported by NSF CISE Expeditions award CCF-1139158 and DARPA XData Award FA8750-12-2-0331, the Na-tional Science Foundation Graduate Research Fellowship (grantDGE-1106400), and gifts from Amazon Web Services, Google, SAP,Apple, Inc., Cisco, Clearstory Data, Cloudera, EMC, Ericsson, Face-book, GameOnTalis, General Electric, Hortonworks, Huawei, Intel,Microsoft, NetApp, NTT Multimedia Communications Laborato-ries, Oracle, Samsung, Splunk, VMware, WANdisco and Yahoo!.

8. REFERENCES[1] D. J. Abadi. Consistency tradeoffs in modern distributed database system design:

CAP is only part of the story. IEEE Computer, 45(2):37–42, 2012.[2] A. Adya. Weak consistency: a generalized theory and optimistic implementations

for distributed transactions. PhD thesis, MIT, 1999.[3] D. Agrawal and V. Krishnaswamy. Using multiversion data for non-interfering

execution of write-only transactions. In SIGMOD 1991.

[4] H. Attiya and J. Welch. Distributed Computing: Fundamentals, Simulations andAdvanced Topics (2nd edition). John Wiley Interscience, March 2004.

[5] P. Bailis, A. Davidson, A. Fekete, A. Ghodsi, J. M. Hellerstein, and I. Stoica.Highly Available Transactions: Virtues and Limitations. In VLDB 2014.

[6] P. Bailis, A. Fekete, A. Ghodsi, J. M. Hellerstein, and I. Stoica. The potentialdangers of causal consistency and an explicit solution. In SOCC 2012.

[7] J. Baker, C. Bond, J. Corbett, J. Furman, et al. Megastore: Providing scalable,highly available storage for interactive services. In CIDR 2011.

[8] P. Bernstein, V. Hadzilacos, and N. Goodman. Concurrency control and recoveryin database systems. Addison-wesley New York, 1987.

[9] K. Birman, G. Chockler, and R. van Renesse. Toward a cloud computing researchagenda. SIGACT News, 40(2):68–80, June 2009.

[10] B. H. Bloom. Space/time trade-offs in hash coding with allowable errors. CACM,13(7):422–426, 1970.

[11] N. Bronson, Z. Amsden, G. Cabrera, P. Chukka, P. Dimov, et al. TAO:Facebook’s distributed data store for the social graph. In USENIX ATC 2013.

[12] A. Chan and R. Gray. Implementing distributed read-only transactions. IEEETransactions on Software Engineering, (2):205–212, 1985.

[13] F. Chang, J. Dean, S. Ghemawat, W. C. Hsieh, D. A. Wallach, M. Burrows, et al.Bigtable: A distributed storage system for structured data. In OSDI 2006.

[14] R. Chirkova and J. Yang. Materialized views. Foundations and Trends inDatabases, 4(4):295–405, 2012.

[15] B. F. Cooper, R. Ramakrishnan, U. Srivastava, A. Silberstein, P. Bohannon, et al.PNUTS: Yahoo!’s hosted data serving platform. In VLDB 2008.

[16] B. F. Cooper, A. Silberstein, E. Tam, R. Ramakrishnan, and R. Sears.Benchmarking cloud serving systems with YCSB. In ACM SOCC 2010.

[17] J. C. Corbett, J. Dean, M. Epstein, A. Fikes, C. Frost, J. J. Furman, et al. Spanner:Google’s globally-distributed database. In OSDI 2012.

[18] C. Curino, E. Jones, Y. Zhang, and S. Madden. Schism: a workload-drivenapproach to database replication and partitioning. In VLDB 2010.

[19] S. Das, D. Agrawal, and A. El Abbadi. G-store: a scalable data store fortransactional multi key access in the cloud. In ACM SOCC 2010.

[20] K. Daudjee and K. Salem. Lazy database replication with ordering guarantees. InICDE 2004, pages 424–435.

[21] S. Davidson, H. Garcia-Molina, and D. Skeen. Consistency in partitionednetworks. ACM Computing Surveys, 17(3):341–370, 1985.

[22] G. DeCandia, D. Hastorun, M. Jampani, G. Kakulapati, A. Lakshman, et al.Dynamo: Amazon’s highly available key-value store. In SOSP 2007.

[23] S. Gilbert and N. Lynch. Brewer’s conjecture and the feasibility of consistent,available, partition-tolerant web services. SIGACT News, 33(2):51–59, 2002.

[24] J. Gray and L. Lamport. Consensus on transaction commit. ACM TODS,31(1):133–160, Mar. 2006.

[25] P. Helland. Life beyond distributed transactions: an apostate’s opinion. In CIDR2007.

[26] S. Hull. 20 obstacles to scalability. Commun. ACM, 56(9):54–59, 2013.[27] N. Huyn. Maintaining global integrity constraints in distributed databases.

Constraints, 2(3/4):377–399, Jan. 1998.[28] E. P. Jones, D. J. Abadi, and S. Madden. Low overhead concurrency control for

partitioned main memory databases. In SIGMOD 2010.[29] R. Kallman, H. Kimura, J. Natkins, A. Pavlo, et al. H-Store: a high-performance,

distributed main memory transaction processing system. In VLDB 2008.[30] R. J. Lipton and J. S. Sandberg. PRAM: a scalable shared memory. Technical

Report TR-180-88, Princeton University, September 1988.[31] F. Llirbat, E. Simon, D. Tombroff, et al. Using versions in update transactions:

Application to integrity checking. In VLDB 1997.[32] W. Lloyd, M. J. Freedman, et al. Don’t settle for eventual: scalable causal

consistency for wide-area storage with COPS. In SOSP 2011.[33] W. Lloyd, M. J. Freedman, M. Kaminsky, and D. G. Andersen. Stronger

semantics for low-latency geo-replicated storage. In NSDI 2013.[34] C. Mohan. History repeats itself: Sensible and NonsenSQL aspects of the

NoSQL hoopla. In EDBT 2013.[35] A. Pavlo, C. Curino, and S. Zdonik. Skew-aware automatic database partitioning

in shared-nothing, parallel OLTP systems. In SIGMOD 2012.[36] D. Peng and F. Dabek. Large-scale incremental processing using distributed

transactions and notifications. In OSDI 2010.[37] S. H. Phatak and B. Badrinath. Multiversion reconciliation for mobile databases.

In ICDE 1999.[38] L. Qiao, K. Surlaker, S. Das, T. Quiggle, et al. On brewing fresh Espresso:

LinkedIn’s distributed data serving platform. In SIGMOD 2013.[39] N. Schiper, P. Sutra, and F. Pedone. P-store: Genuine partial replication in wide

area networks. In IEEE SRDS 2010.[40] M. Shapiro et al. A comprehensive study of convergent and commutative

replicated data types. Technical Report 7506, INRIA, 2011.[41] J. Shute et al. F1: A distributed SQL database that scales. In VLDB 2013.[42] D. B. Terry, A. J. Demers, K. Petersen, M. J. Spreitzer, M. M. Theimer, and B. B.

Welch. Session guarantees for weakly consistent replicated data. In PDIS 1994.[43] A. Thomson, T. Diamond, S. Weng, K. Ren, P. Shao, and D. Abadi. Calvin: Fast

distributed transactions for partitioned database systems. In SIGMOD 2012.[44] K. Weil. Rainbird: Real-time analytics at Twitter. Strata 2011

http://slidesha.re/hjMOui.[45] S. B. Zdonik. Object-oriented type evolution. In DBPL, pages 277–288, 1987.[46] J. Zhou et al. Lazy maintenance of materialized views. In VLDB 2007.

APPENDIX: Proofs and Isolation DetailsRAMP-F Correctness. To prove RAMP-F provides RA isolation, we showthat the two-round read protocol returns a transactionally atomic set ofversions. To do so, we formalize criteria for atomic (read) sets of versionsin the form of companion sets. We will call the set of versions producedby a transaction sibling versions and call two items from the same write setsibling items.

Given two versions xi and y j , we say that xi is a companion to y j if xiis a transactional sibling of y j or x is a sibling item of y j and i > j. Wesay that a set of versions V is a companion set if, for every pair (xi,y j) ofversions in V where x is a sibling item of y j , xi is a companion to y j . InFigure 1, the versions returned by T2’s first round of reads ({x1,y?}) donot comprise a companion set because y? has a lower timestamp than x1’ssibling version of y (that is, x1 has sibling version y1 and but ?< 1 so y?has too low of a timestamp). Subsets of companion sets are also companionsets and companion sets also have a useful property for RA isolation:

Claim 1 (Companion sets are atomic). Companion sets do not containfractured reads.Proof. Claim 1 follows from the definitions of companion sets and fracturedreads. If V is a companion set, then every version xi 2V is also a companionto every other version y j 2V where v j contains x in its sibling items. If Vcontained fractured reads, V would contain two versions xi,y j such that thetransaction that wrote y j also wrote a version xk , i < k. However, in this case,xi would not be a companion to y j , a contradiction. Therefore, V cannotcontain fractured reads.

To provide RA, RAMP-F clients assemble a companion set for the requesteditems (in vlatest ), which we prove below:

Claim 2. RAMP-F provides Read Atomic isolation.Proof. Each write in RAMP-F contains information regarding its siblings,which can be identified by item and timestamp. Given a set of RAMP-Fversions, recording the highest timestamped version of each item (as recordedeither in the version itself or via sibling metadata) yields a companion setof item-timestamp pairs: if a client reads two versions xi and y j such thatx is in y j’s sibling items but i < j, then vlatest [x] will contain j and noti. Accordingly, given the versions returned by the first round of RAMP-Freads, clients calculate a companion set containing versions of the requesteditems. Given this companion set, clients check the first-round versionsagainst this set by timestamp and issue a second round of reads to fetchany companions that were not returned in the first round. The resultingset of versions will be a subset of the computed companion set and willtherefore also be a companion set. This ensures that the returned results donot contain fractured reads. RAMP-F first-round reads access lastCommit, soeach transaction corresponding to a first-round version is committed, and,therefore, any siblings requested in the (optional) second round of readsare also committed. Accordingly, RAMP-F never reads aborted or non-final(intermediate) writes. This establishes that RAMP-F provides RA.

RAMP-F Scalability and Independence. RAMP-F also provides the in-dependence guarantees from Section 3.3. The following invariant overlastCommit is core to RAMP-F GET request completion:

Invariant 1 (Companions present). If a version xi is referenced by lastCommit(that is, lastCommit[x] = i), then each of xi’s sibling versions are present inversions on their respective partitions.

Invariant 1 is maintained by RAMP-F’s two-phase write protocol. lastCommitis only updated once a transaction’s writes have been placed into versionsby a first round of PREPARE messages. Siblings will be present in versions(but not necessarily lastCommit).

Claim 3. RAMP-F provides synchronization independence.Proof. Clients in RAMP-F do not communicate or coordinate with one an-other and only contact servers. Accordingly, to show that RAMP-F providessynchronization independence, it suffices to show that server-side operationsalways terminate. PREPARE and COMMIT methods only access data storedon the local partition and do not block due to external coordination or othermethod invocations; therefore, they complete. GET requests issued in thefirst round of reads have tsreq =? and therefore will return the version cor-responding to lastCommit[k], which was placed into versions in a previouslycompleted PREPARE round. GET requests issued in the second round ofclient reads have tsreq set to the client’s calculated vlatest [k]. vlatest [k] is asibling of a version returned from lastCommit in the first round, so, due to

Invariant 1, the requested version will be present in versions. Therefore, GETinvocations are guaranteed access to their requested version and can returnwithout waiting. The success of RAMP-F operations do not depend on thesuccess or failure of other clients’ RAMP-F operations.

Claim 4. RAMP-F provides partition independence.Proof. RAMP-F transactions do not access partitions that are unrelated to eachtransaction’s specified data items and servers do not contact other servers inorder to provide a safe response for operations.

RAMP-S Correctness. RAMP-S writes and first-round reads proceed iden-tically to RAMP-F writes, but the metadata written and returned is different.Therefore, the proof is similar to RAMP-F, with a slight modification for thesecond round of reads.

Claim 5. RAMP-S provides Read Atomic isolation.Proof. To show that RAMP-S provides RA, it suffices to show that RAMP-Ssecond-round reads (resp) are a companion set. Given two versions xi,y j 2resp such that x 6= y, if x is a sibling item of y j , then xi must be a companionto y j . If xi were not a companion to y j , then it would imply that x is nota sibling item of y j (so we are done) or that j > i. If j > i, then, dueto Invariant 1 (which also holds for RAMP-S writes due to identical writeprotocols), y j’s sibling is present in versions on the partition for x and wouldhave been returned by the server (line 6), a contradiction. Each second-roundGET request returns only one version, so we are done.

RAMP-S Scalability and Independence. RAMP-S provides synchroniza-tion independence and partition independence. For brevity, we again omitfull proofs, which closely resemble those of RAMP-F.RAMP-H Correctness. The probabilistic behavior of the RAMP-H Bloomfilter admits false positives. However, given unique transaction timestamps(Section 4.5), requesting false siblings by timestamp and item does not affectcorrectness:

Claim 6. RAMP-H provides Read Atomic isolation.Proof. To show that RAMP-H provides Read Atomic isolation, it suffices toshow that any versions requested by RAMP-H second-round reads that wouldnot have been requested by RAMP-F second-round reads (call this set v f alse)do not compromise the validity of RAMP-H’s returned companion set. Anyversions in v f alse do not exist: timestamps are unique, so, for each versionxi, there are no versions x j of non-sibling items with the same timestamp asxi (i.e., where i = j). Therefore, requesting versions in v f alse do not changethe set of results collected in the second round.

RAMP-H Scalability and Independence. RAMP-H provides synchroniza-tion independence and partition independence. We omit full proofs, whichclosely resemble those of RAMP-F. The only significant difference fromRAMP-F is that second-round GET requests may return ?, but, as we showedabove, these empty responses correspond to false positives in the Bloomfilter and therefore do not affect correctness.Comparison to other isolation levels. The fractured reads anomaly issimilar to Adya’s “Missing Transaction Updates” definition, only applied toimmediate read dependencies (rather than all transitive dependencies). RAis stronger than PL-2 (Read Committed), but weaker than PL-SI, PL-CS, andPL-2.99 (notably, RA does not prevent anti-dependency cycles, or Adya’sG2 or G-SIa—informally, it allows concurrent updates) [2].

RA does not (by itself) provide ordering guarantees across transactions.Our RAMP implementations provide a variant of PRAM consistency, where,for each item, each user’s writes are serialized [30] (i.e., “session” order-ing [20]), and, once a user’s operation completes, all other users will observeits effects (regular register semantics, applied at the transaction level). Thisprovides transitivity with respect to each user’s operations. For example,if a user updates her privacy settings and subsequently posts a new photo,the photo cannot be read without the privacy setting change [15]. However,PRAM does not respect the happens-before relation [4] across users. IfSam reads Mary’s comment and replies to it, other users may read Sam’scomment without Mary’s comment. In this case, RAMP transactions canleverage explicit causality [6] via foreign key dependencies, but happens-before is not provided by default. If required, we believe it is possible toenforce happens-before but, due to scalability concerns regarding metadataand partition independence (e.g., [6] and Section 5), do not further explorethis possibility. An “active-active” replicated implementation can provideavailable [5, 23] operation at the cost of these recency guarantees.