Transcript Fault Tolerance: Byzantine FT (TvS:Ch7)
System Reliability and Fault Tolerance
Reliable Communication
Byzantine Fault Tolerance
Recap: Replication
Write is handled only by the remote primary server, and the backups are updated accordingly; Read is performed locally.
Replicated services:
–
Sun Network Info Service (NIS, formerly Yellow Pages)
–
FAB: Building distributed enterprise disk arrays from commodity components, ASPLOS’04
Reliable Point-to-Point Comm.
Failure Models – Process failure : Sender vs Receiver » Fail-stop: a process crash which can be detected by other processes » How to detect such crash? Timeout can indicate only that a process is not responding – Comm. failure » send failure: A process completes a send, but the msg is not put in its outgoing msg buffer » receive failure: A msg is in incoming buf, but is not received by a process » Channel failure: fail while msg is transmitted from outgoing buf to incoming buf – Arbitrary failure (Byzantine failure) » Any type of error may occur. E.g. return wrong value. Reliable comm : – Validity : Any msg in the outgoing buf is eventually delivered to the incoming buf – Integrity : The msg received is identical to the one sent and no msgs are delivered twice.
© C. Xu, 1998-2007
RPC Failure Semantics
Five Possible Failures: – The client is unable to locate the server » Server is down or Stub mismatches with Skeleton » Throw UnknownHostException – The request message is lost » start with a timer » retransmission of the request message – The server crashes after receiving a request – The reply message is lost – The client crashes after sending a request In Java, all remote methods must be prepared to catch RemoteException
© C. Xu, 1998-2007
Server Crashes
a) A server in client-server communication Normal case b) c) Crash after execution Crash before execution • • •
At least once semantics At most once semantics Exactly once semantics java.rmi.ServerRuntimeException
© C. Xu, 1998-2007
Server Crash (Cont’)
Assume client requests server to print a msg – Send a completion msg (M) before print (P), or – Send a completion msg (M) after print (P) Combinations – M P C: crash after ack and print – M C( P): – P M C: – P C( M) – C( P M): crash before print and ack – C( M P)
© C. Xu, 1998-2007
When a crashed server recovers, the client can
never reissue a request (Never) always reissue a request (Always) reissue only if it received ack reissue only if it received no ack
Client Strategy M
P Server Strategy P
M Reissue strategy
Always Never Only when ACKed Only when not ACKed
MPC
DUP OK DUP OK
MC(P) C(MP)
OK OK ZERO ZERO OK ZERO ZERO OK
PMC PC(M) C(PM)
DUP DUP OK OK OK ZERO DUP OK OK DUP ZERO OK
ok: Text is printed once; dup: printed twice; zero: no printout © C. Xu, 1998-2007
Lost Reply Messages – Some requests can be re-executed with side-effects (e.g. Read 1024 bytes of a file); some not (idempotent).
– Solutions: » Structure requests in an idempotent way » Assign request a sequence number to be checked by server Client Crashes leading to orphan – computation extermination : client side logging of RPC about what to do; the log is checked after a reboot.
– reincarnation : client bcasts a new epoch when it reboots; server detects orphan computations based on epochs. » kill orphan remote computation or locate their owners – expiration : set a time quantum for each RPC request; if it cannot finish, more quanta are asked.
© C. Xu, 1998-2007
Reliable Multicast
Basic properties: – Validity : If a correct process multicasts message m, then it will eventually deliver m. – Integrity : a correct process delivers the msg at most once Atomic messages (aka agreement) – A message is delivered none to all members of a group, or to Message ordering guarantees – within group – across groups
© C. Xu, 1998-2007
Message Ordering
Different members may see messages in different orders Ordered group communication requires that all members agree about the order of messages Within each group, assign
global ordering
messages to Hold back messages that arrive out of order (delay their delivery)
© C. Xu, 1998-2007
(I) Unordered Multicasts
Process P1
mcast m1 mcast m2
Process P2
receives m1 receives m2
Process P3
receive m2 receives m1
(II) FIFO-ordered Multicasts Process P1
mcast m1 mcast m2
Process P2
receive m1 receives m3 receives m2 receives m4
Process P3
receives m3 receives m1 receives m2 receives m4
Process P4
mcast m3 mcast m4
If a process multicasts two msgs m and m’ in order, then every process in the group will deliver the msgs in the same order
(III) Causally-order Multicasts
C A (1) (1) (2) If mcast(g, m)
mcast(g, m’) Then any process in the group should deliver m before m’ B (1) Delayed (2) D (VI) Totally-ordered multicasts If a process delivers msg m before m’, then any other process that delivers m’ will deliver m before m’.
Centralized Impl of Total Ordering
Central ordering server (
sequencer
) assigns global sequence numbers Hosts apply to ordering server for numbers, or ordering server sends all messages itself Hold-back easy, since sequence numbers are sequential – Msgs will remain in hold-back queue until they can be delivered according to their sequence numbers. Sequencer: bottleneck and single point of failure – tricky protocol to deal with case where ordering server fails
© C. Xu, 1998-2007
Atomic Messages
Each recipient acks message, and sender retransmits if ack not received – Sender could crash before msg is delivered!!
» Simple approach: if sender crashes, a recipient volunteers to be “backup sender” for the message » re-sends message to everybody, waits for acks » use simple algorithm to choose volunteer » apply method again if backup sender crashes No single best solutions exist!
© C. Xu, 1998-2007
Reliability due to Replication
Blocking update, waiting till backups are updated – Blocking update of backup servers must be atomic so as to implement sequential consistency as the primary can serialize all incoming writes (in order ) and all processes see all writes in the same order from any backup servers. Total ordering due to the use of primary for centralized sequencer Atomic : – –
What happens if some W4 are Postive Ack and some are NAck?
Two-phase commit protocol : » W3: “prepare” msg from primary to other replicas » W3+: ack to “prepare” (
If in a prepared state, related objects be preserved in permanent storage; will eventually be able to commit it
.) » W4: “commit” or “abort” msg » W4+: ack to “commit/abort”
2PC Protocol in the Presence of Failures
If ack to “prepare” msg is timed out, primary can send “abort” to replicas and safely abort itself If replica waiting for “commit” or “abort” is timed out, – If its ack to “prepare” was negative, simply abort itself – If its ack to “prepare” was positive, it cannot “commit”, nor “abort”. Block, waiting for primary or network recovery How to handle crash/reboot, particularly primary failure?
– Cannot back out of a commit if already decided – Semantics of failure: store commit; cannot commit before store – Recovery protocol w/ non-volatile memory 2PC causes a long waiting time if primary fails after “prepare” msg is sent out – Three-phase commit protocol: Pre-Prepare, Prepare, Commit – Replica times out waiting for Commit msg will commit the trans – 2PC: execute transaction when everyone is willing to commit – 3PC: execute transaction when everyone knows it will commit
Recovery from Primary Failure
Need to pick up a new primary defining a new “view.” It could be set by human operator OR autonomic – Suppose the lowest-numbered live server is the primary – Replicas need to ping each other – Ping msg lose or delayed may lead to more than one primary Paxos protocol for fault-tolerant consensus – At most a single value is chosen – Agreement reached despite lost msgs and crashed nodes – Paxos protocol: eventually succeeds if a majority of replicas are reachable –
See Lamport’98 (submitted to TOCS in 90) and Chandra Toueg’96 for details
Handling Byzantine Failure
Byzantine Failure
– Failed replicas are not necessarily failure-stop – Failed replicas may generate arbitrary results!!
The Byzantine Generals Problem Leslie Lamport, Robert Shostak, and Marshall Pease in 1982
Byzantine Generals Problem
N divisions of Byzantine army surround city – Each division commanded by a general – Some of the N generals are traitors Generals communicate via messages – Traitors can send different values to different generals Requirements: – All loyal generals decide upon same plan of action – A “small” number of traitors cannot cause loyal generals to adopt a bad plan – NOT required to identify traitors
Restricted BGP
Restate problems as: – 1 commanding general – N-1 lieutenants Interactive consistency requirements – IC1: All loyal lieutenants obey the same order – IC2: If the commander is loyal, every loyal lieutenant obeys his/her order If we can solve this problem… – Original BGP problem reduces to N instances of this problem; one instance per general acting as commander
3-General Impossibility Result
Assume 2 loyal generals and 1 traitor (shaded) – Two messages: ATTACK or RETREAT If Lt.1 sees {Attack, “He said Retreat”} what to do?
–
If Lt2 is traitor (Fig1), L1 must attack to satisfy IC2
–
If Commander is traitor (Fig2), L1 and L2 must make same decision (always obeying commanders order over lieutenant’s violates IC1) Commander Attack Attack lieutenant He said Retreat lieutenant Attack lieutenant Commander Retreat He said Retreat lieutenant
General Impossibility Result
In general, no solutions with fewer than 3m+1 generals if there are m traitors Proof by contradiction: – Assume there is a solution for 3m Albanians, including m traitors – Let a Byzantine general simulate m Albanian generals – The problem is then reduced to 3-general problem
Solution Example
With one faulty process: f=1, N=4 1 st round: the cmd sends a value to each Lt 2 nd round: each Lt copies the value to all other Lts
p
1 (Commander)
p
1 (Commander)
p
2 1:v 2:1:v 1:v 3:1:u 1:v 2:1:v 4:1:v 4:1:v 3:1:w
p
3
p
2 1:u 2:1:u 1:v 3:1:w 1:w 2:1:u 4:1:v 4:1:v 3:1:w
p
3
p
4 Faulty processes are shown coloured
p
4
Practical Byzantine Fault Tolerance
Miguel Castro and Barbara Liskov OSDI’99
Assumptions
Asynchronous distributed systems Faulty nodes may behave arbitrarily – Due to malicious attacks or software errors Independent node failures Network may fail to deliver, delay, duplicate or deliver them out of order
An adversary may coordinate faulty nodes, delay comm, or delay correct nodes in order to cause the most damage to the service. BUT it cannot delay correct nodes indefinitely, nor subvert the cryptographic techniques
– Any network fault will be eventually repaired – E.g. cannot forge a valid signature of non-faulty node – E.g. Cannot find two msgs with the same digests
Objectives
To be used for the implementation of any deterministic replicated service with a
state
and some
operations
– Clients issue requests and block waiting for a reply
Safety
if no more than [(n-1)/3] faulty replicas (i.e. to tolerate f faulty nodes, at least n=3f+1 needed) – Safety: the replicated service satisfies linearizability – Behaves like a centralized implementation that executes ops atomically one at a time – Why 3f+1 the optimal resiliency?
Liveness
: clients eventually receive replies to their requests, – At most [(n-1)/3] faulty replicas – Comm delay is bounded with unknown bounds; delay is the latency from the time of first sending to the time of receipt by the destination
Algorithm in a nutshell
Backup f + 1 Match (OK) Client Primary Backup Backup
Replicas and Views
Set of replicas (R): |R| ≥ 3f + 1 R2 ……… R|R-1| For view v primary p assigned such that p= v mod |R|
Normal Case Operation
{REQUEST, o, t, c} Client Primary o – Operation t – Timestamp c - Client
Timestamps are totally ordered such that later requests have higher timestamps than earlier ones
Normal Case Operation
state of each replica is stored in a
message log
Primary
p
receives a client request phase protocol
m
, it starts a three to atomically multicast the request to the replicas –
Pre-prepare, Prepare, Commit Pre-Prepare
and
Prepare
phases are for total ordering of requests sent in the same view, even when the primary is faulty Prepare and Commit phases are to ensure committed requests are totally ordered across views
Pre-Prepare Phase
Backup Primary <
v – view number n – sequence number m – message d – digest of the message
Backup
Prepare Phase
If replica i accepts the PRE-PREPARE message it enters
prepare
phase by multicasting < PREPARE , v, n, d, i> to all other replicas and adds both messages to its log
Otherwise does nothing
A replica accepts the PRE-PREPARE message provided, – The signatures are valid and the digest matches m – It is in view v – It has not accepted a PRE-PREPARE for the same v and n – Sequence number is within accepted bounds
Commit Phase
When replica i receives
2f matched
PREPARE msg, the replica gets into Commit Phase by multicasting
i
executes required operation after it has accepted 2f+1
matched
commit msgs from different replicas. Replica i’s state reflects the seq execution of all requests with lower sequence numbers. This ensures all non-faulty replicas execute requests in same order.
To guarantee exactly-once semantics , replicas discard requests whose timestamp is lower than the timestamp in the last reply they sent to the client.
Normal Operation Reply
All replicas sends the reply
Normal Case Operation: Summery
Pre-prepare Prepare Commit C Request Reply Primary: 0 1 2 Faulty: 3 X
Safeguards
If the client does not receive replies soon enough , it broadcasts the request to all replicas If the request has already been processed, the replicas simply re-send the reply If the replica is not the primary, it relays the request to the primary If the primary does not multicast the request to the group, it will eventually be suspected to be faulty by enough replicas to cause a view change
View Changes
Timer
is set when a request is received, recording the waiting time for the request to execute.
If the timer of replica expires in view
v
, the replica starts a view change to move to view
v+1 by,
– Stop accepting
pre-prepare/prepare/commit
messages – Multicasting a
VIEW-CHANGE
message
s
known to i C = 2f + 1 checkpoint msgs proving correctness of s P = {Pm: for each
m
prepared, its #seq >n} Pm = pre-prepare and 2f matching prepare msgs
New Primary
When primary p’ of view v+1 receives 2f valid VIEW-CHANGE messages
– It multicasts a
For a replica that accepts NEW-VIEW
– Sends PREPARE messages for every pre-prepare in set O – Moves to view v+1
References
See OSDI’99 for
– Optimization, implementation, and evaluation – 3% overhead in NFS daemon
See tech report for
– Formal presentation of the algorithm in I/O automation model – Proof of safety and liveness
See OSDI’00 for
– Proactive recovery
Further Readings
In synchronous systems, assume msg exchanges take place in rounds timeout and processes can detect the absence of a msg through a – At least f+1 rounds of msgs are needed
(Fisher-Lynch, 82)
In async systems with unbounded delay, a crashed process becomes indistinguishable from a slow one. – [Impossibility] No algorithm can guarantee to reach consensus in such systems, even with one process crash failure.
(Fisher-Lynch-Paterson, J. ACM’85
) Approaches to working around the impossibility – In partially async systems with bounded but unknown delay » Practical Byzantine Fault Tolerant Alg (Castro-Liskov’99) – Using failure detectors: unresponsive process be treated as failure and discard their subsequent msgs. – Consensus can be reached, even with an unreliable failure detector, if fewer than N/2 processes crashes and comm is reliable. (Chandra Toueg’96) – Statistical consensus: “no guarantee” doesn’t mean “cannot”. » Introduce an element of chance in processes’ behaviors so that the adversary cannot exercise its thwarting strategy effectively.