Download - Lecture XI: Distributed Transactions
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CMPT 401 Summer 2007
Dr. Alexandra Fedorova
Lecture XI: Distributed Transactions
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2CMPT 401 Summer 2007 © A. Fedorova
A Distributed Transaction
• A transaction is distributed across n processes.• Each process can decide to commit or abort the
transaction• A transaction must commit on all sites or abort on all sites
Application Program / Coordinating Server
Participating Site
Participating Site
commit
commit
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3CMPT 401 Summer 2007 © A. Fedorova
Comparison With Single-Site Transactions
• For single-site transactions we talked about:– Committing a transaction– Recovery (for atomicity, consistency and durability)– Concurrency control (for atomicity, consistency and
isolation)• For distributed transactions we will talk about:
– Distributed commitment– Distributed recovery– Distributed concurrency control
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4CMPT 401 Summer 2007 © A. Fedorova
A Distributed Transaction
# Some participant (the invoker) executes:send [T_START: transaction, Dc, participants] to all participants
# All participants (including the invoker) execute:upon (receipt of T_START: transaction, Dc, participants]Cknow = local_clock
# Perform operations requested by transactionif(willing and able to make updates permanent) then
vote = YESelse vote = NO#Decide commit or abort for the transactionatomic_commitment(transaction, participants)
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5CMPT 401 Summer 2007 © A. Fedorova
Atomic Commitment Problem
• All processes must decide to commit or abort– Is this possible in an asynchronous system?
• No. This is like consensus, and we saw that consensus is impossible
• But it is still used in asynchronous systems under certain assumptions but with no guarantees about termination:– Communication is reliable– Processes can crash– Processes recover from failure– Processes can log their state in stable storage– Stable storage survives crashes and is accessible upon restart
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6CMPT 401 Summer 2007 © A. Fedorova
Properties of Atomic Commitment
• Property 1: All participants that decide reach the same decision
• Property 2: If any participant decides commit, then all participants must have voted YES
• Property 3: If all participants vote YES and no failures occur, then all participants decide commit
• Property 4: Each participant decides at most once (a decision is irreversible)
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7CMPT 401 Summer 2007 © A. Fedorova
Components of Atomic Commitment• Normal execution
– The steps executed while no failures occur• Termination protocol
– When a site fails, the correct sites should still be able to decide on the outcome of pending transactions.
– They run a termination protocol to decide on all pending transactions.
• Recovery– When a site fails and then restarts it has to perform recovery for
all transactions that it has not yet committed appropriately– Single site recovery: abort all transactions that were active at the
time of the failure– Distributed system: might have to ask around; maybe an active
transaction was committed in the rest of the system
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8CMPT 401 Summer 2007 © A. Fedorova
Blocking vs. Non-Blocking Atomic Commitment
• Blocking Atomic Commitment: correct participants may be prevented from terminating the transaction due to failures of other part of the system
• Non-Blocking Atomic Commitment: transactions terminate consistently at all participating sites even in the presence of failures
• Non-blocking atomic commitment is not possible in systems with unreliable communication
• Non-blocking algorithms exist if one assumes reliable communication and a synchronous system
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9CMPT 401 Summer 2007 © A. Fedorova
Two-Phase Commit Protocol (2PC)
• An atomic commitment protocol– Phase 1: Decide commit or abort– Phase 2: Get the final decision from the coordinator, and execute
the final decision• We will study the protocol in a synchronous system
– Assume a message arrives within interval δ– Assume we can compute Dc – local time required to complete the
transaction– Assume we can compute Db – additional delay associated with
broadcast
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10CMPT 401 Summer 2007 © A. Fedorova
Two-Phase Commit Protocol (I)
# Executed by coordinatorprocedure atomic_commitment(transaction, participants)
send [VOTE_REQUEST] to all participantsset-timeout-to local_clock + 2δwait-for (receipt of [vote: vote] messages from all participants)
if (all votes are YES) thenbroadcast (commit, participants)
else broadcast (abort, participants)on-timeout
broadcast (abort, participants)
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11CMPT 401 Summer 2007 © A. Fedorova
Two-Phase Commit Protocol (II)# Executed by all participants (including the coordinator)
set-timeout-to Cknow + Dc + δwait-for (receipt of [VOTE_REQUEST] from coordinator)
send [vote: vote] to coordinatorif (vote = NO) then
decide abortelse
set-timeout-to Cknow + Dc + 2δ + Dbwait-for (delivery of decision message)
if(decision message is abort) thendecide abort
else decide commit on-timeout
decide according to termination_protocol()on-timeout
decide abort
May block the
participant
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12CMPT 401 Summer 2007 © A. Fedorova
The Need for Termination Protocol• If a participant
– Voted YES– Sent its decision to coordinator, but….– Received no final decision from coordinator
• A termination protocol must be run– Participants cannot simply decide to abort– If they already said they would commit, they cannot change their minds– The coordinator might have sent “commit” decisions to some
participants and then crashed– Since those participants might have committed, no other participant can
decide “abort”• A termination protocol will try to contact other participants to find
out what they decided, and try to reach a decision
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13CMPT 401 Summer 2007 © A. Fedorova
Blocking Nature of Two-Phase Commit
• All protocols involving a termination protocol are blocking• A termination protocol may lead to a decision, but there
are scenarios when it will not:– The coordinator crashes during the broadcast of a decision– Several participants received the decision from coordinator,
applied it, and then crashed– All correct (not crashed) participants voted “YES”– Correct participants cannot decide until faulty participants
recover– If some faulty participants had decided “NO”, they would have
aborted– If some faulty participants had received a “commit” decision from
the coordinator, they would have committed
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14CMPT 401 Summer 2007 © A. Fedorova
Non-Blocking Atomic Commitment
• Property 4: Every correct participant that executes the atomic commitment protocol eventually decides
• Recall that 2PC was blocking because coordinator could broadcast decision to a part of participants and then crash
• Need a broadcast with stronger properties, such as Uniform Timed Reliable Broadcast (UTRB)
• Property of UTRB: If any participant (correct or not) receives a broadcast message m, then all correct participants eventually receive m
• With UTRB, if a participant times out waiting for final decision, it can unilaterally decide to abort
• The participant knows that if it has not received a message, then no one else has, so no one else could have committed
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15CMPT 401 Summer 2007 © A. Fedorova
A Simple UTRB Algorithm
procedure broadcast (m, G) // G is the set of participants
# Broadcaster executes:send [m] to all processes in Gdeliver m //local delivery of message m to the
process
# Process p ≠ broadcaster in G executes:upon (first receipt of [m])send [m] to all processes in Gdeliver m
Message will not be delivered to
local process before it is
forwarded to other participants
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16CMPT 401 Summer 2007 © A. Fedorova
2PC in Asynchronous Systems
• 2PC can be implemented in an asynchronous system with reliable communication channels
• This means that a message eventually gets delivered…• But we cannot set bounds on delivery time• So the process might have to wait forever…• Therefore, you cannot have non-blocking atomic commitment
in an asynchronous system• What if a participant whose message is waited on has crashed? • The expectation is that the participant will properly recover and
continue the protocol• So now let’s look at distributed recovery
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17CMPT 401 Summer 2007 © A. Fedorova
Distributed Recovery
• Remember single-site recovery: – transaction log records are kept on stable storage– upon reboot the system “undoes” updates from active on
uncommitted transactions– “replays” updates from committed transactions
• In a distributed system we cannot be sure whether a transaction that was active at the time of crash is:– Still active– Has committed– Has aborted– Maybe it has executed more updates while the recovering site
was crashing and rebooting
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18CMPT 401 Summer 2007 © A. Fedorova
Components of Distributed Recovery
• Actions performed by recovering site:– For each transaction that was active before the crash, try to
decide unilaterally based on log records– Ask others what they have decided
• Actions performed by other participants– Others may send the decision– Or a “don’t know” message
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19CMPT 401 Summer 2007 © A. Fedorova
Crash Before Local Decision
• Suppose a site crashes during the execution of transaction, before it reaches local decision (YES or NO)
• The transaction could have completed at other sites• What are the options? • Option 1: The crashed site restores its state with help of
other participants (restore the updates made while it was crashing and recovering)
• Option 2: The crashed site realizes that it crashed (by keeping the crash count), and sets local decision to NO
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20CMPT 401 Summer 2007 © A. Fedorova
Logging for Distributed Recovery
• Coordinator: forces “commit” decision to log before informing any participants – Like redo rule for single-site logging
• Participant: forces its vote (YES or NO) to disk before sending the vote to coordinator – – This way it knows that it must reach decision in agreement
with others• Participant: forces final decision (received from
coordinator) to the log, then responds to the coordinator– Once the coordinator receives responses from all
participants, it can remove its own decision log record
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21CMPT 401 Summer 2007 © A. Fedorova
Distributed Concurrency Control
• Multiple servers execute transactions, they share data distributed across sites
• A lock on data may be requested by many different servers• Distributed concurrency control methods:
– Centralized two-phase locking (C-2PL)– Distributed two-phase locking (D-2PL)– Optimistic concurrency control
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22CMPT 401 Summer 2007 © A. Fedorova
Distributed Concurrency Control: Notation
• SCH - global lock scheduler • Ti – a transaction• Oij(X) – an operation requiring a lock on X, executed at
server Sj as part of Ti• S1, S2, … - servers performing a distributed transaction• TM1, TM2, … - transaction managers – one for each server
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23CMPT 401 Summer 2007 © A. Fedorova
Centralized 2PL
• Let S1 be the server to which transaction Ti was submitted• Let S2 be the server maintaining data X• For each operation oi1(X), TM1 first requests the corresponding lock
from SCH (the central 2PL scheduler)• Once the lock is granted, the operation is forwarded to the server S2
maintaining X.
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24CMPT 401 Summer 2007 © A. Fedorova
Centralized 2PL
Read(x)
Release locks
Read(x)
TM1 SCH S2
Grant lock
Operation executed
...commit
Request lock
S1
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25CMPT 401 Summer 2007 © A. Fedorova
Distributed 2PL
• Each server has its own local 2PL scheduler SCH• Let S1 be the server to which transaction Ti was
submitted:– For each operation Oi1(X), TM1 forwards the operation to the TM2
of server S2 maintaining X– The remote site first acquires a lock on X and then submits the
execution of the operation.
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26CMPT 401 Summer 2007 © A. Fedorova
Distributed 2PL
Read(X)
Release locks
TM1 TM2 SCH of S2
Operation executed
...commit
Read(X)
S1 S2
Request lock Submit op
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27CMPT 401 Summer 2007 © A. Fedorova
Optimistic Concurrency Control
• Locking is conservative– Locking overhead even if no conflicts– Deadlock detection/resolution (especially problematic in
distributed environment)– Lock manager can fail independently
• Optimistic concurrency control– Perform operation first – Check for conflicts only later (e.g., at commit time)
• Centralized systems never used optimistic CC• Interesting alternative for distributed environment
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28CMPT 401 Summer 2007 © A. Fedorova
Optimistic Concurrency Control
• Working Phase:– If first operation on X, then load last committed version from DB and cache– Otherwise read/write cached version– Keep WriteSet containing objects written – Keep ReadSet containing objects read
• Validation Phase– Check whether transaction conflicts with other transactions
• Update Phase– Upon successful validation, cached version of updated objects are written back to
DB (= changes are made public)• Assumption (for simplicity):
– Only one transaction may be in validation phase and update phase at a time (e.g., implemented as critical section)
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29CMPT 401 Summer 2007 © A. Fedorova
Backward Validation: Working Phase
Maintain transaction counter Tn, initialize Tn := 0 Upon begin of transaction Ti (could be first read/write request)
StartTn(Ti) := Tn WriteSet(Ti) := ReadSet(Ti) := { }
Upon Read(X), Write(X) request of Ti If first operation on X, load last committed version from DB and
cache, Read/write cached version of X If Read(X), then ReadSet(Ti) := ReadSet(Ti) { X } If Write(X), then WriteSet(Ti) := WriteSet(Ti) { X }
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30CMPT 401 Summer 2007 © A. Fedorova
Backward Validation: Validation and Update Phases
At time of validation of transaction Ti For all j where StartTn(Ti) < j ≤ Tn
Let WriteSet(Tk) be the write set of transaction Tk with finishTn(Tk) = j
If WriteSet(Tk) ReadSet(Ti) { }, then abort Ti After validation (update)
Tn = Tn + 1 finishTn(Ti) = Tn Write WriteSet(Ti) to DB
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31CMPT 401 Summer 2007 © A. Fedorova
Optimistic Concurrency Control
Working Validation Update
T1
T2
T3
T5(active)
StartTN(T5)
finishTN(T1)
finishTN(T2)
finishTN(T3)
Tn = finishTN(T3)
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32CMPT 401 Summer 2007 © A. Fedorova
Optimistic Concurrency Control: Discussion
• Validation and Update Phase in Critical section– Reduced concurrency– Optimizations possible
• Space overhead:– To validate Ti, must have WriteSets for all Tj where
finishTn(Tj) > StartTn(Ti) and Tj was active when Ti began.– Each transaction must record read/write activity
• No deadlock but potentially more aborts
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33CMPT 401 Summer 2007 © A. Fedorova
Distributed Deadlock Resolution
• Similar remedies as for single-site deadlock resolution:– Prevention (lock ordering)– Avoidance (abort transaction that waits for too long)– Detection (maintain a wait-for graph, abort transactions involved
in a cycle)• Deadlock avoidance and detection require keeping
dependency graphs, or wait-for-graphs (WFGs)• WFGs are more difficult to construct in a distributed
system (takes more time, must use vector clocks or distributed snapshots)
• Deadlock managers can fail independently
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34CMPT 401 Summer 2007 © A. Fedorova
Summary• Atomic commitment
– Two-phase commit– Non-blocking implementation possible in a synchronous system
with reliable communication channel– Possible in an asynchronous system, but not guaranteed to
terminate (blocking)• Distributed recovery
– Keep state on stable storage– When reboot, ask around to recover the most current state
• Distributed concurrency control – Centralized lock manager– Distributed lock manager– Optimistic concurrency control