Requirements For Solutions - Learning Module

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Correctness Proofs: Rigorous Verification

The Need for Mathematical Certainty

Throughout this module, we have evaluated synchronization algorithms by reasoning about their properties. But reasoning alone leaves room for error. Subtle bugs can hide in informal arguments. The history of computer science is filled with algorithms believed correct for years before subtle flaws were discovered.

This brings us to the gold standard of verification: formal correctness proofs. A formal proof provides mathematical certainty that an algorithm satisfies its requirements under stated assumptions. No testing can provide this level of assurance—only proof.

This page introduces the techniques of formal proof for synchronization algorithms. We will cover:

The mathematical foundations of correctness proofs
Invariant-based proof techniques
Structural induction for safety properties
Variant-based proofs for liveness properties
Complete proofs of classical algorithms

By the end, you will understand how to construct and verify rigorous proofs of synchronization correctness.

What You Will Learn

By the end of this page, you will understand the structure of formal correctness proofs, master invariant-based proof techniques for mutual exclusion, apply variant-based proofs for progress and bounded waiting, and be able to construct complete correctness proofs for synchronization algorithms.

Foundations of Formal Proofs

Formal proofs of concurrent algorithms differ from sequential program proofs. The non-determinism introduced by arbitrary scheduling requires techniques that account for all possible interleavings.

The Concurrent Program Model

We model a concurrent program as a transition system:

States (S): The set of all possible configurations of shared and local variables.

Initial States (S₀ ⊆ S): The starting configurations.

Transitions (→ ⊆ S × S): The set of atomic steps that take the system from one state to another.

A program execution is a sequence of states s₀ → s₁ → s₂ → ... where s₀ ∈ S₀ and each sᵢ → sᵢ₊₁ is a valid transition.

Safety vs. Liveness Properties

Properties of concurrent systems divide into two categories:

Safety Properties: "Nothing bad ever happens"

Mutual exclusion is a safety property
Formally: If the property is violated, it's violated at some finite point
Proved by showing an invariant that implies the safety condition

Liveness Properties: "Something good eventually happens"

Progress and bounded waiting are liveness properties
Formally: If the property is violated, we observe infinite waiting
Proved using fairness assumptions and variant functions

Safety vs. Liveness Properties
Property Type	Intuition	Examples	Proof Technique
Safety	Nothing bad happens	Mutual exclusion, deadlock-freedom	Invariants
Liveness	Something good happens	Progress, bounded waiting, termination	Variants + fairness

Fairness Assumptions

Liveness properties require fairness assumptions about the scheduler. Without them, a scheduler could indefinitely delay any process, violating any liveness property trivially.

Weak Fairness: If a process is continuously enabled (can take a step), it eventually takes a step.

Strong Fairness: If a process is infinitely often enabled (can take a step infinitely often), it eventually takes a step.

Most liveness proofs assume at least weak fairness. Without fairness, we can only prove safety properties.

Why Fairness Matters

Consider a spinlock: a process spins in 'while(locked){}'. With weak fairness, if the lock is eventually released, the process will eventually see it and proceed. Without fairness, the scheduler could always schedule other processes, and the spinner never gets a chance to check.

Invariant-Based Proofs for Safety

The primary technique for proving safety properties is invariant-based reasoning. An invariant is a predicate over states that, once established, remains true throughout execution.

Inductive Invariants

An invariant I is inductive if:

Initialization: I holds in all initial states
```
∀s ∈ S₀: I(s)
```
Preservation: I is preserved by all transitions
```
∀s, s': (I(s) ∧ s → s') → I(s')
```

If I is inductive, then I holds in all reachable states. This follows by induction on the length of executions.

Using Invariants to Prove Mutual Exclusion

To prove mutual exclusion:

Define an invariant I that:
- Is inductive (holds initially, preserved by all transitions)
- Implies mutual exclusion: I(s) → at_most_one_in_cs(s)
Since I holds in all reachable states, mutual exclusion holds in all reachable states.

invariant_proof_structure.txt
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// Structure of an Invariant-Based Proof
 
// Goal: Prove mutual exclusion for algorithm A
 
// STEP 1: Define the invariant I
// I must be strong enough to imply mutual exclusion
// I must be preserved by every atomic action in the algorithm
 
Example invariant for a lock-based algorithm:
  I ≡ (in_cs[i] → lock == HELD_BY_i) for all i
 
// STEP 2: Prove initialization
// Show I holds in the initial state
Proof:
  Initially, no process is in CS.
  in_cs[i] = false for all i.
  Therefore, (in_cs[i] → lock == HELD_BY_i) is vacuously true.
  ∴ I holds initially. ✓
 
// STEP 3: Prove preservation
// For each atomic action a, show: I(s) ∧ s →ᵃ s' → I(s')
 
Action: Process i executes "acquire lock"
  Precondition: I(s), lock is free
  Effect: lock = HELD_BY_i
  Goal: Show I(s')
  Proof: No process was in CS (since lock was free and I(s)).
         After: if i enters CS, in_cs[i] = true, lock = HELD_BY_i.
         (in_cs[i] → lock == HELD_BY_i) holds. ✓
 
Action: Process i executes "release lock"
  Precondition: I(s), in_cs[i] = true
  Effect: lock = FREE, in_cs[i] = false
  Goal: Show I(s')
  Proof: After, in_cs[i] = false.
         (false → lock == HELD_BY_i) is true. ✓
 
[Continue for all actions...]
 
// STEP 4: Conclude mutual exclusion
// Show I → MUTEX where MUTEX ≡ ¬(in_cs[i] ∧ in_cs[j]) for i ≠ j
 
Proof:
  Suppose in_cs[i] ∧ in_cs[j] for some i ≠ j.
  By I: lock == HELD_BY_i (since in_cs[i])
  By I: lock == HELD_BY_j (since in_cs[j])
  But HELD_BY_i ≠ HELD_BY_j, contradiction.
  ∴ ¬(in_cs[i] ∧ in_cs[j]). MUTEX holds. ✓

Strengthening the Invariant

Sometimes the obvious invariant isn't strong enough to be inductive. We may need to strengthen it—add more conjuncts that help preservation even if they're not directly needed for the final property.

For example, if we need to prove P, and P alone isn't inductive, we find Q such that:

P ∧ Q is inductive
P ∧ Q → P (trivially true)

Finding the right strengthening is an art. Common techniques:

Include conditions linking control locations to variable values
Include conditions about what must be true when a process is at certain points
Include global properties that constrain the combined state

The Art of Invariant Discovery

When a simple invariant fails to be preserved by some transition, ask: 'What additional condition would make it preserved?' The answer often reveals the strengthening needed. Invariant discovery is iterative—propose, test preservation, refine.

Complete Proof: Peterson's Algorithm Mutual Exclusion

Let us construct a complete, rigorous proof that Peterson's algorithm satisfies mutual exclusion.

The Algorithm (Formalized)

peterson_formal.txt
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// Peterson's Algorithm - Formal Specification
 
// Shared Variables:
//   flag[0], flag[1] : Boolean (initially false)
//   turn : {0, 1}     (initially 0)
 
// Local Variables for process i:
//   pc[i] : {idle, want, wait, cs, exit}  (program counter)
 
// Process i (for i ∈ {0, 1}, let j = 1 - i):
 
// Location transitions:
// idle → want:  flag[i] := true
// want → wait:  turn := j
// wait → cs:    (flag[j] = false) ∨ (turn = i)  [guard]
// cs → exit:    <critical section complete>
// exit → idle:  flag[i] := false
 
// Initial State:
//   pc[0] = idle, pc[1] = idle
//   flag[0] = false, flag[1] = false
//   turn = 0

The Invariant

We propose the following invariant:

I ≡ ∀i ∈ {0,1}: (pc[i] ∈ {wait, cs, exit} → flag[i] = true)
    ∧
    ∀i ∈ {0,1}: (pc[i] = cs → (flag[j] = false ∨ turn = i))

The first conjunct says: if a process is waiting, in CS, or exiting, its flag is set.

The second conjunct says: if a process is in CS, either the other's flag is false or it's this process's turn.

Proof of Initialization

peterson_init_proof.txt
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// Proof: I holds in the initial state
 
Initial state:
  pc[0] = idle, pc[1] = idle
  flag[0] = false, flag[1] = false
  turn = 0
 
First conjunct for i = 0:
  pc[0] ∈ {wait, cs, exit}? No (pc[0] = idle)
  Therefore, (pc[0] ∈ {wait, cs, exit} → flag[0] = true) is vacuously true. ✓
 
First conjunct for i = 1:
  pc[1] ∈ {wait, cs, exit}? No (pc[1] = idle)
  Vacuously true. ✓
 
Second conjunct for i = 0:
  pc[0] = cs? No (pc[0] = idle)
  Vacuously true. ✓
 
Second conjunct for i = 1:
  pc[1] = cs? No (pc[1] = idle)
  Vacuously true. ✓
 
∴ I holds in the initial state. ✓

Proof of Preservation

We must show that each transition preserves I. There are 5 transitions per process; we analyze the critical ones:

peterson_preservation_proof.txt
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// Proof: each transition preserves I
 
// TRANSITION: idle → want (for process i)
// Effect: flag[i] := true, pc[i] := want
 
Before: Assume I holds.
After: 
  - First conjunct for i: pc[i] = want ∉ {wait, cs, exit}, so vacuously true. ✓
  - First conjunct for j: unchanged (we only modified process i's state). ✓ 
  - Second conjuncts: pc[i] ≠ cs, pc[j] unchanged. If pc[j] was not cs,
    still not cs. If pc[j] = cs, then by I before, (flag[i] = false ∨ turn = j).
    But we just set flag[i] = true! Does this break I?
    
    Hmm, potential issue. Let's check:
    If pc[j] = cs before, we had (flag[i] = false ∨ turn = j).
    After, flag[i] = true, so we need turn = j to maintain.
    
    But wait—we only set OUR flag. Turn was set during j's want→wait.
    Since j is in cs, j already passed the guard, so turn = j at that time.
    Turn only changes in want→wait transitions.
    Since we're in idle→want, we don't change turn.
    ∴ turn = j still holds. ✓
 
// TRANSITION: want → wait (for process i)
// Effect: turn := j, pc[i] := wait
 
Before: Assume I holds.
After:
  - First conjunct for i: pc[i] = wait, flag[i] was set in idle→want,
    and we haven't cleared it. But wait, prove that flag[i] = true.
    
    Issue: We need flag[i] = true when pc[i] = wait.
    Strengthen I: Add "pc[i] ∈ {want, wait, cs, exit} → flag[i] = true"
    (extend the set to include 'want')
    
    OK, assuming strengthened I:
    Before, pc[i] = want, so flag[i] = true.
    We don't modify flag[i] in this transition.
    After, pc[i] = wait, flag[i] = true. ✓
    
  - First conjunct for j: unchanged. ✓
  
  - Second conjunct for i: pc[i] = wait ≠ cs, vacuously true. ✓
  
  - Second conjunct for j: If pc[j] = cs, need (flag[i] = false ∨ turn = i).
    We just set turn = j, so turn ≠ i (since i ≠ j).
    We need flag[i] = false. But flag[i] = true!
    
    This seems like a problem... but wait:
    If pc[j] = cs, then j passed the guard: flag[i] = false ∨ turn = j.
    Since we're now setting turn = j, that's consistent.
    After this transition, turn = j, so (turn = j) is true.
    ∴ (flag[i] = false ∨ turn = j) holds.
    But the invariant for j says (flag[i] = false ∨ turn = j).
    After, turn = j, so this holds. ✓
 
// TRANSITION: wait → cs (for process i)
// Guard: (flag[j] = false) ∨ (turn = i)
// Effect: pc[i] := cs
 
Before: Assume I holds, and guard is true.
After:
  - First conjunct for i: pc[i] = cs, flag[i] was true (strengthened I). ✓
  
  - Second conjunct for i: pc[i] = cs, need (flag[j] = false ∨ turn = i).
    But this was exactly the guard! We only take this transition when
    the guard is true. ∴ The condition holds. ✓
 
// TRANSITION: exit → idle (for process i)  
// Effect: flag[i] := false, pc[i] := idle
 
Before: Assume I holds.
After:
  - First conjunct for i: pc[i] = idle ∉ {want, wait, cs, exit}. Vacuous. ✓
  - We set flag[i] = false. Does this affect second conjunct for j?
    If pc[j] = cs, we need (flag[i] = false ∨ turn = j).
    After, flag[i] = false, so this holds! ✓
 
∴ All transitions preserve I. ✓

Proof that I Implies Mutual Exclusion

peterson_implies_mutex.txt
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// Proof: I → MUTEX
 
// MUTEX ≡ ¬(pc[0] = cs ∧ pc[1] = cs)
 
Proof by contradiction:
  Assume I holds and pc[0] = cs ∧ pc[1] = cs.
  
  By second conjunct of I applied to process 0:
    pc[0] = cs → (flag[1] = false ∨ turn = 0)
    ∴ flag[1] = false ∨ turn = 0   ... (1)
    
  By second conjunct of I applied to process 1:
    pc[1] = cs → (flag[0] = false ∨ turn = 1)
    ∴ flag[0] = false ∨ turn = 1   ... (2)
    
  By first conjunct of I applied to process 0:
    pc[0] = cs → flag[0] = true
    ∴ flag[0] = true   ... (3)
    
  By first conjunct of I applied to process 1:
    pc[1] = cs → flag[1] = true
    ∴ flag[1] = true   ... (4)
    
  From (1) and (4): Since flag[1] = true, we must have turn = 0.
  From (2) and (3): Since flag[0] = true, we must have turn = 1.
  
  But turn cannot be both 0 and 1. Contradiction!
  
∴ ¬(pc[0] = cs ∧ pc[1] = cs). MUTEX holds. ∎

Mutual Exclusion Proven

We have rigorously proven that Peterson's algorithm satisfies mutual exclusion. The key insight is that the turn variable creates an asymmetry: both processes cannot have 'their turn' simultaneously, so both cannot satisfy the entry guard at the same time.

Variant-Based Proofs for Liveness

Liveness properties assert that something good eventually happens. We cannot prove such properties using invariants alone—invariants only constrain what CAN happen, not what WILL happen.

The Variant Method

To prove liveness, we use variant functions (also called ranking functions or progress measures):

A variant is a function V: States → W where:

W is a well-founded set (no infinite descending chains)
Progress: each step toward the goal decreases V
Bounded: V cannot decrease forever

Common choices for W:

Natural numbers (with usual ordering)
Tuples of naturals (lexicographic ordering)
Finite ordinals

Proving Progress

To prove progress (if CS is empty and processes contend, one enters):

Define a variant measuring 'distance' from entry
Show under fairness, the variant decreases
Since the variant can't decrease forever, entry occurs

variant_proof_structure.txt
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// Structure of a Variant-Based Liveness Proof
 
// Goal: Prove that waiting processes eventually enter CS
 
// STEP 1: Define the progress condition
// C(s) ≡ "processes are waiting but CS is empty"
 
// STEP 2: Define the variant V
// V: States → ℕ  (or some well-founded set)
// V measures "how far" we are from a process entering CS
 
Example for Peterson's:
  V(s) = (2 if both in wait, 1 if one in wait and one not, 0 otherwise)
  
  But this is too coarse. Better:
  
  V(s) = < number of processes ahead in "queue",
           number of steps until guard satisfied >
 
// STEP 3: Show V decreases while C holds
// Under fairness, show that eventually:
//   Either C becomes false (a process enters - goal achieved!)
//   Or V decreases
 
Proof sketch:
  Case: Both processes in wait state, turn = j.
  Then process i with turn = i will satisfy guard (flag[j] ∨ turn = i).
  Since turn = j (for j), process j is blocked.
  But process i's guard (flag[j] ∨ turn = i) evaluates to (true ∨ false) = true.
  Wait, let me reconsider...
  
  If turn = j, then:
  - Process i's guard: flag[j] = true (j is waiting), turn = i = (1-j) ≠ j.
    So flag[j] = true and turn ≠ i. Guard is false for i.
  - Process j's guard: flag[i] = true (i is waiting), turn = j. 
    So turn = j. Guard (flag[i]=false ∨ turn=j) becomes (false ∨ true) = true!
    
  So process j enters. V decreases (one fewer waiting). ✓
 
// STEP 4: Conclude
// Since W is well-founded and V decreases while C holds,
// eventually C becomes false (some process in CS).
// ∴ Progress holds.

Proving Bounded Waiting

Bounded waiting requires a stronger argument: not just that SOME process enters, but that EACH waiting process enters within bounded bypasses.

We use a variant that tracks the specific waiting process:

Vᵢ(s) = "distance for process i to enter CS"

And show that:

Each bypass by other processes decreases Vᵢ (or maintains it)
The bound is computable from the initial Vᵢ

Fairness is Essential

Liveness proofs fundamentally require fairness assumptions. The variant decreases when the 'right' process takes a step. Fairness guarantees that the 'right' process eventually takes its step. Without fairness, we cannot conclude that the variant ever decreases.

Complete Proof: Peterson's Progress and Bounded Waiting

Let us complete the correctness proof for Peterson's algorithm by proving progress and bounded waiting.

Proof of Progress

peterson_progress_proof.txt
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// Proof: Peterson's algorithm satisfies Progress
 
// Goal: If CS is empty and at least one process is in entry (want or wait),
//       eventually some process enters CS.
 
// Assumption: Weak fairness (continuously enabled → eventually executes)
 
// Case 1: One process (say i) is in entry, the other (j) is in remainder.
 
  Process i is at want or wait.
  j is in remainder (idle state), so flag[j] = false.
  
  If i is at want:
    Under fairness, i eventually executes want → wait (sets turn = j).
    Now i is at wait.
    
  i is at wait:
    Guard: flag[j] = false ∨ turn = i
    Since flag[j] = false (j in remainder), guard is TRUE.
    Under fairness, i eventually executes wait → cs.
    
  ∴ Process i enters CS. Progress satisfied. ✓
 
// Case 2: Both processes are in entry.
 
  Both at want or wait states.
  
  Sub-case 2a: At least one is at want.
    Under fairness, it executes want → wait.
    Eventually, both are at wait.
    
  Sub-case 2b: Both at wait.
    flag[0] = true, flag[1] = true (by invariant).
    turn = 0 or turn = 1.
    
    If turn = 0:
      Process 0's guard: flag[1] = true, turn = 0. 
        Guard = (true ∨ true) = true... wait, that's wrong.
        Guard = (flag[1] = false) ∨ (turn = 0) = false ∨ true = true. ✓
      Process 0 can proceed.
      Under fairness, process 0 enters CS.
      
    If turn = 1:
      By symmetric argument, process 1 enters CS.
    
  ∴ Some process enters CS. Progress satisfied. ✓
 
// Conclusion: Under weak fairness, Peterson's algorithm has Progress. ∎

Proof of Bounded Waiting

peterson_bw_proof.txt
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// Proof: Peterson's algorithm satisfies Bounded Waiting with B = 1
 
// Goal: After process i enters its entry section, process j can enter
//       CS at most once before process i enters.
 
// Assume process i enters entry section at time t₀ (sets flag[i] = true).
 
// Case 1: Process j is not in entry/CS at time t₀.
  j is in remainder (idle), so flag[j] = false.
  
  As shown in progress proof, i will enter CS.
  j bypasses i 0 times. ✓
 
// Case 2: Process j is in CS at time t₀.
  j will eventually exit (finite CS execution assumption).
  When j exits:
    - j sets turn = i (in exit section)
    - j sets flag[j] = false
  
  Now, if j wants to enter again:
    j sets flag[j] = true
    j sets turn = i (in want → wait)
    j arrives at wait with turn = i
    
  In the meantime, i is in entry (want or wait).
  
  If i is at wait:
    turn = i (set by j's exit or j's entry!)
    i's guard: flag[j] = true, turn = i
    Guard = (true ∨ true) = true... no wait,
    Guard = (flag[j] = false) ∨ (turn = i) = false ∨ true = TRUE. ✓
    i can enter.
    
  If i is at want:
    Under fairness, i moves to wait.
    i sets turn = j.
    But then j's exit will set turn = i.
    Eventually i sees turn = i and enters.
    
  Key insight: After j exits CS and before i enters CS,
  if j sets turn again (in want→wait), it sets turn = i.
  This guarantees i's guard becomes true.
  
  j entered CS once (was already in at t₀), i then enters.
  j bypasses i at most 1 time. ✓
 
// Case 3: Process j is in entry at time t₀.
  j is at want or wait.
  turn is either 0 or 1.
  
  If turn = j (i.e., it's j's turn):
    j can enter (guard satisfied).
    j enters CS (bypass 1).
    j exits, sets turn = i.
    Now i's guard: (flag[j] = false) ∨ (turn = i).
    After j's exit, flag[j] = false, so guard is true.
    i enters.
    j bypassed i once. ✓
    
  If turn = i:
    j's guard: flag[i] = true, turn = i ≠ j. Guard false.
    j cannot enter.
    i's guard: flag[j] = true, turn = i. Guard true.
    i enters.
    j bypassed i zero times. ✓
 
// Conclusion: In all cases, j enters at most once before i.
// Bounded Waiting holds with B = 1. ∎

Complete Proof Achieved

We have now rigorously proven that Peterson's algorithm satisfies all three requirements: Mutual Exclusion (via invariant proof), Progress (via case analysis with fairness), and Bounded Waiting with B=1 (via bypass counting). Peterson's algorithm is correct!

Summary of Proof Techniques

Let us consolidate the proof techniques we have developed:

Technique Taxonomy

Proof Techniques by Property
Property	Type	Technique	Key Elements
Mutual Exclusion	Safety	Invariant	Find I such that I is inductive and I → MUTEX
Absence of Deadlock	Safety	Invariant	Show ¬(all waiting ∧ none can proceed)
Progress	Liveness	Variant + Fairness	Show variant decreases until entry
Bounded Waiting	Liveness	Bypass Counting	Count max bypasses per waiting process
Termination	Liveness	Variant	Show execution eventually completes

Common Invariant Patterns

Certain invariant patterns recur across synchronization algorithms:

Reusable Invariant Patterns

•Location → Variable: at(L) → var = val (e.g., in CS implies lock held)
•Gate condition: in_cs[i] → gate_condition[i] where gate prevents simultaneous entry
•Turn exclusivity: turn has one value at a time, so two different turn values cannot hold
•Flag semantics: flag[i] = true ↔ process i is interested
•Conservation: Σ in_cs[i] + Σ waiting[i] + Σ idle[i] = n (all processes accounted for)

Common Proof Pitfalls

Common Errors in Correctness Proofs
Pitfall	Description	Remedy
Weak invariant	Invariant doesn't imply property	Strengthen with additional conjuncts
Non-inductive	Invariant not preserved by all transitions	Strengthen or relax invariant
Missed transition	Forgot to check some program action	Systematically enumerate all actions
Conflated atomicity	Treated non-atomic as atomic	Decompose into true atomic steps
Omitted fairness	Proved liveness without fairness	State fairness assumption explicitly
Case explosion	Didn't use structure to simplify	Use symmetry and case merging

Proof Development Process

Start with the property to prove. 2. Guess an invariant (for safety) or variant (for liveness). 3. Try to prove initialization and preservation. 4. When a step fails, analyze why and strengthen. 5. Iterate until complete. 6. Verify the invariant implies the desired property.

Mechanized Proofs and Verification Tools

While hand-written proofs are valuable for understanding, they are error-prone. Mechanized proofs use computer assistance to verify correctness with mathematical certainty.

Interactive Theorem Provers

Provers like Coq, Isabelle/HOL, and Lean allow humans to construct proofs interactively, with the computer verifying each step:

Human provides proof strategy and key insights
Computer checks all logical steps rigorously
Final proof is machine-verified

Model Checkers (State-Space Exploration)

Tools like SPIN, TLA+, and UPPAAL exhaustively explore the state space:

Automatically check all reachable states
Report counterexamples if property fails
Limited by state explosion, but highly automated

tlaplus_peterson.tla
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--------------------------- MODULE Peterson ---------------------------
EXTENDS Integers
 
VARIABLES flag, turn, pc
 
vars == <<flag, turn, pc>>
 
Init == /\ flag = [i \in {0, 1} |-> FALSE]
        /\ turn \in {0, 1}
        /\ pc = [i \in {0, 1} |-> "idle"]
 
(* Process i transitions *)
Enter(i) == /\ pc[i] = "idle"
            /\ pc' = [pc EXCEPT ![i] = "want"]
            /\ flag' = [flag EXCEPT ![i] = TRUE]
            /\ UNCHANGED turn
 
SetTurn(i) == LET j == 1 - i IN
              /\ pc[i] = "want"
              /\ pc' = [pc EXCEPT ![i] = "wait"]
              /\ turn' = j
              /\ UNCHANGED flag
 
EnterCS(i) == LET j == 1 - i IN
              /\ pc[i] = "wait"
              /\ (flag[j] = FALSE \/ turn = i)
              /\ pc' = [pc EXCEPT ![i] = "cs"]
              /\ UNCHANGED <<flag, turn>>
 
ExitCS(i) == LET j == 1 - i IN
             /\ pc[i] = "cs"
             /\ pc' = [pc EXCEPT ![i] = "idle"]
             /\ flag' = [flag EXCEPT ![i] = FALSE]
             /\ turn' = j
 
Next == \E i \in {0, 1}: Enter(i) \/ SetTurn(i) \/ EnterCS(i) \/ ExitCS(i)
 
(* Mutual exclusion property *)
MutualExclusion == ~(pc[0] = "cs" /\ pc[1] = "cs")
 
(* Specification *)
Spec == Init /\ [][Next]_vars /\ WF_vars(Next)
 
THEOREM Spec => []MutualExclusion
 
=========================================================================

Industrial Use of Verification

Formal verification is used in high-stakes domains:

Hardware verification: CPU designs are now routinely model-checked
Aerospace: Flight control systems require formal proofs
Security protocols: TLS, authentication protocols are verified
Distributed systems: Amazon uses TLA+ for DynamoDB, S3
Compilers: CompCert is a formally verified C compiler

For synchronization specifically, libraries like liblfds (lock-free data structures) have undergone formal verification.

When to Use Formal Verification

Formal verification is most valuable when: (1) Correctness is critical (safety, security, financial), (2) Testing is insufficient (concurrency, rare conditions), (3) The system will run for long periods without restart, (4) Bugs have severe consequences (data loss, safety hazards). For synchronization algorithms—especially in OS kernels—formal verification is increasingly standard.

Summary: Correctness Proofs

We have explored the rigorous mathematical foundations of correctness proofs for synchronization algorithms. Let us consolidate the key concepts.

Key Takeaways

•Formal proofs provide mathematical certainty that algorithms meet their requirements, transcending what testing can achieve.
•Safety properties (like mutual exclusion) are proved using inductive invariants—predicates true initially and preserved by all transitions.
•Liveness properties (like progress) require fairness assumptions and variant functions that decrease toward the goal.
•Proving mutual exclusion involves finding an invariant that implies no two processes are in CS simultaneously, then proving the invariant is inductive.
•Proving progress involves showing that from any waiting state, under fairness, some process eventually enters CS.
•Proving bounded waiting involves counting the maximum bypasses a process can experience and showing it's finite.
•Mechanized verification tools (TLA+, Coq, SPIN) provide machine-checked certainty and are used in industry for critical systems.

Module Complete

With this page, we have completed Module 3: Requirements for Solutions. You now possess:

Deep understanding of the three fundamental requirements: mutual exclusion, progress, and bounded waiting
Analytical skills to evaluate any synchronization solution systematically
Proof techniques to rigorously verify that algorithms are correct

These foundations prepare you for studying specific algorithms (Peterson's, Dekker's) in depth, and for understanding the hardware and OS mechanisms (atomic operations, locks, semaphores) that implement synchronization in practice.

Next chapters will explore:

Peterson's and Dekker's algorithms in detail
Hardware atomic operations (test-and-set, compare-and-swap)
OS synchronization primitives (mutexes, semaphores, monitors)
Classic synchronization problems (producer-consumer, readers-writers, dining philosophers)

Module Complete

Congratulations! You have mastered the requirements for solutions to the critical section problem—mutual exclusion, progress, and bounded waiting—along with systematic evaluation techniques and formal proof methods. This foundational knowledge will serve you throughout your study of operating systems and concurrent programming.

Correctness Proofs: Rigorous Verification

The Need for Mathematical Certainty

This page introduces the techniques of formal proof for synchronization algorithms. We will cover:

The mathematical foundations of correctness proofs
Invariant-based proof techniques
Structural induction for safety properties
Variant-based proofs for liveness properties
Complete proofs of classical algorithms

By the end, you will understand how to construct and verify rigorous proofs of synchronization correctness.

What You Will Learn

Foundations of Formal Proofs

Formal proofs of concurrent algorithms differ from sequential program proofs. The non-determinism introduced by arbitrary scheduling requires techniques that account for all possible interleavings.

The Concurrent Program Model

We model a concurrent program as a transition system:

States (S): The set of all possible configurations of shared and local variables.

Initial States (S₀ ⊆ S): The starting configurations.

Transitions (→ ⊆ S × S): The set of atomic steps that take the system from one state to another.

A program execution is a sequence of states s₀ → s₁ → s₂ → ... where s₀ ∈ S₀ and each sᵢ → sᵢ₊₁ is a valid transition.

Safety vs. Liveness Properties

Properties of concurrent systems divide into two categories:

Safety Properties: "Nothing bad ever happens"

Mutual exclusion is a safety property
Formally: If the property is violated, it's violated at some finite point
Proved by showing an invariant that implies the safety condition

Liveness Properties: "Something good eventually happens"

Progress and bounded waiting are liveness properties
Formally: If the property is violated, we observe infinite waiting
Proved using fairness assumptions and variant functions

Safety vs. Liveness Properties
Property Type	Intuition	Examples	Proof Technique
Safety	Nothing bad happens	Mutual exclusion, deadlock-freedom	Invariants
Liveness	Something good happens	Progress, bounded waiting, termination	Variants + fairness

Fairness Assumptions

Liveness properties require fairness assumptions about the scheduler. Without them, a scheduler could indefinitely delay any process, violating any liveness property trivially.

Weak Fairness: If a process is continuously enabled (can take a step), it eventually takes a step.

Strong Fairness: If a process is infinitely often enabled (can take a step infinitely often), it eventually takes a step.

Most liveness proofs assume at least weak fairness. Without fairness, we can only prove safety properties.

Why Fairness Matters

Invariant-Based Proofs for Safety

The primary technique for proving safety properties is invariant-based reasoning. An invariant is a predicate over states that, once established, remains true throughout execution.

Inductive Invariants

An invariant I is inductive if:

Initialization: I holds in all initial states
```
∀s ∈ S₀: I(s)
```
Preservation: I is preserved by all transitions
```
∀s, s': (I(s) ∧ s → s') → I(s')
```

If I is inductive, then I holds in all reachable states. This follows by induction on the length of executions.

Using Invariants to Prove Mutual Exclusion

To prove mutual exclusion:

Define an invariant I that:
- Is inductive (holds initially, preserved by all transitions)
- Implies mutual exclusion: I(s) → at_most_one_in_cs(s)
Since I holds in all reachable states, mutual exclusion holds in all reachable states.

invariant_proof_structure.txt
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// Structure of an Invariant-Based Proof
 
// Goal: Prove mutual exclusion for algorithm A
 
// STEP 1: Define the invariant I
// I must be strong enough to imply mutual exclusion
// I must be preserved by every atomic action in the algorithm
 
Example invariant for a lock-based algorithm:
  I ≡ (in_cs[i] → lock == HELD_BY_i) for all i
 
// STEP 2: Prove initialization
// Show I holds in the initial state
Proof:
  Initially, no process is in CS.
  in_cs[i] = false for all i.
  Therefore, (in_cs[i] → lock == HELD_BY_i) is vacuously true.
  ∴ I holds initially. ✓
 
// STEP 3: Prove preservation
// For each atomic action a, show: I(s) ∧ s →ᵃ s' → I(s')
 
Action: Process i executes "acquire lock"
  Precondition: I(s), lock is free
  Effect: lock = HELD_BY_i
  Goal: Show I(s')
  Proof: No process was in CS (since lock was free and I(s)).
         After: if i enters CS, in_cs[i] = true, lock = HELD_BY_i.
         (in_cs[i] → lock == HELD_BY_i) holds. ✓
 
Action: Process i executes "release lock"
  Precondition: I(s), in_cs[i] = true
  Effect: lock = FREE, in_cs[i] = false
  Goal: Show I(s')
  Proof: After, in_cs[i] = false.
         (false → lock == HELD_BY_i) is true. ✓
 
[Continue for all actions...]
 
// STEP 4: Conclude mutual exclusion
// Show I → MUTEX where MUTEX ≡ ¬(in_cs[i] ∧ in_cs[j]) for i ≠ j
 
Proof:
  Suppose in_cs[i] ∧ in_cs[j] for some i ≠ j.
  By I: lock == HELD_BY_i (since in_cs[i])
  By I: lock == HELD_BY_j (since in_cs[j])
  But HELD_BY_i ≠ HELD_BY_j, contradiction.
  ∴ ¬(in_cs[i] ∧ in_cs[j]). MUTEX holds. ✓

Strengthening the Invariant

For example, if we need to prove P, and P alone isn't inductive, we find Q such that:

P ∧ Q is inductive
P ∧ Q → P (trivially true)

Finding the right strengthening is an art. Common techniques:

Include conditions linking control locations to variable values
Include conditions about what must be true when a process is at certain points
Include global properties that constrain the combined state

The Art of Invariant Discovery

Complete Proof: Peterson's Algorithm Mutual Exclusion

Let us construct a complete, rigorous proof that Peterson's algorithm satisfies mutual exclusion.

The Algorithm (Formalized)

peterson_formal.txt
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// Peterson's Algorithm - Formal Specification
 
// Shared Variables:
//   flag[0], flag[1] : Boolean (initially false)
//   turn : {0, 1}     (initially 0)
 
// Local Variables for process i:
//   pc[i] : {idle, want, wait, cs, exit}  (program counter)
 
// Process i (for i ∈ {0, 1}, let j = 1 - i):
 
// Location transitions:
// idle → want:  flag[i] := true
// want → wait:  turn := j
// wait → cs:    (flag[j] = false) ∨ (turn = i)  [guard]
// cs → exit:    <critical section complete>
// exit → idle:  flag[i] := false
 
// Initial State:
//   pc[0] = idle, pc[1] = idle
//   flag[0] = false, flag[1] = false
//   turn = 0

The Invariant

We propose the following invariant:

I ≡ ∀i ∈ {0,1}: (pc[i] ∈ {wait, cs, exit} → flag[i] = true)
    ∧
    ∀i ∈ {0,1}: (pc[i] = cs → (flag[j] = false ∨ turn = i))

The first conjunct says: if a process is waiting, in CS, or exiting, its flag is set.

The second conjunct says: if a process is in CS, either the other's flag is false or it's this process's turn.

Proof of Initialization

peterson_init_proof.txt
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// Proof: I holds in the initial state
 
Initial state:
  pc[0] = idle, pc[1] = idle
  flag[0] = false, flag[1] = false
  turn = 0
 
First conjunct for i = 0:
  pc[0] ∈ {wait, cs, exit}? No (pc[0] = idle)
  Therefore, (pc[0] ∈ {wait, cs, exit} → flag[0] = true) is vacuously true. ✓
 
First conjunct for i = 1:
  pc[1] ∈ {wait, cs, exit}? No (pc[1] = idle)
  Vacuously true. ✓
 
Second conjunct for i = 0:
  pc[0] = cs? No (pc[0] = idle)
  Vacuously true. ✓
 
Second conjunct for i = 1:
  pc[1] = cs? No (pc[1] = idle)
  Vacuously true. ✓
 
∴ I holds in the initial state. ✓

Proof of Preservation

We must show that each transition preserves I. There are 5 transitions per process; we analyze the critical ones:

peterson_preservation_proof.txt
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// Proof: each transition preserves I
 
// TRANSITION: idle → want (for process i)
// Effect: flag[i] := true, pc[i] := want
 
Before: Assume I holds.
After: 
  - First conjunct for i: pc[i] = want ∉ {wait, cs, exit}, so vacuously true. ✓
  - First conjunct for j: unchanged (we only modified process i's state). ✓ 
  - Second conjuncts: pc[i] ≠ cs, pc[j] unchanged. If pc[j] was not cs,
    still not cs. If pc[j] = cs, then by I before, (flag[i] = false ∨ turn = j).
    But we just set flag[i] = true! Does this break I?
    
    Hmm, potential issue. Let's check:
    If pc[j] = cs before, we had (flag[i] = false ∨ turn = j).
    After, flag[i] = true, so we need turn = j to maintain.
    
    But wait—we only set OUR flag. Turn was set during j's want→wait.
    Since j is in cs, j already passed the guard, so turn = j at that time.
    Turn only changes in want→wait transitions.
    Since we're in idle→want, we don't change turn.
    ∴ turn = j still holds. ✓
 
// TRANSITION: want → wait (for process i)
// Effect: turn := j, pc[i] := wait
 
Before: Assume I holds.
After:
  - First conjunct for i: pc[i] = wait, flag[i] was set in idle→want,
    and we haven't cleared it. But wait, prove that flag[i] = true.
    
    Issue: We need flag[i] = true when pc[i] = wait.
    Strengthen I: Add "pc[i] ∈ {want, wait, cs, exit} → flag[i] = true"
    (extend the set to include 'want')
    
    OK, assuming strengthened I:
    Before, pc[i] = want, so flag[i] = true.
    We don't modify flag[i] in this transition.
    After, pc[i] = wait, flag[i] = true. ✓
    
  - First conjunct for j: unchanged. ✓
  
  - Second conjunct for i: pc[i] = wait ≠ cs, vacuously true. ✓
  
  - Second conjunct for j: If pc[j] = cs, need (flag[i] = false ∨ turn = i).
    We just set turn = j, so turn ≠ i (since i ≠ j).
    We need flag[i] = false. But flag[i] = true!
    
    This seems like a problem... but wait:
    If pc[j] = cs, then j passed the guard: flag[i] = false ∨ turn = j.
    Since we're now setting turn = j, that's consistent.
    After this transition, turn = j, so (turn = j) is true.
    ∴ (flag[i] = false ∨ turn = j) holds.
    But the invariant for j says (flag[i] = false ∨ turn = j).
    After, turn = j, so this holds. ✓
 
// TRANSITION: wait → cs (for process i)
// Guard: (flag[j] = false) ∨ (turn = i)
// Effect: pc[i] := cs
 
Before: Assume I holds, and guard is true.
After:
  - First conjunct for i: pc[i] = cs, flag[i] was true (strengthened I). ✓
  
  - Second conjunct for i: pc[i] = cs, need (flag[j] = false ∨ turn = i).
    But this was exactly the guard! We only take this transition when
    the guard is true. ∴ The condition holds. ✓
 
// TRANSITION: exit → idle (for process i)  
// Effect: flag[i] := false, pc[i] := idle
 
Before: Assume I holds.
After:
  - First conjunct for i: pc[i] = idle ∉ {want, wait, cs, exit}. Vacuous. ✓
  - We set flag[i] = false. Does this affect second conjunct for j?
    If pc[j] = cs, we need (flag[i] = false ∨ turn = j).
    After, flag[i] = false, so this holds! ✓
 
∴ All transitions preserve I. ✓

Proof that I Implies Mutual Exclusion

peterson_implies_mutex.txt
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// Proof: I → MUTEX
 
// MUTEX ≡ ¬(pc[0] = cs ∧ pc[1] = cs)
 
Proof by contradiction:
  Assume I holds and pc[0] = cs ∧ pc[1] = cs.
  
  By second conjunct of I applied to process 0:
    pc[0] = cs → (flag[1] = false ∨ turn = 0)
    ∴ flag[1] = false ∨ turn = 0   ... (1)
    
  By second conjunct of I applied to process 1:
    pc[1] = cs → (flag[0] = false ∨ turn = 1)
    ∴ flag[0] = false ∨ turn = 1   ... (2)
    
  By first conjunct of I applied to process 0:
    pc[0] = cs → flag[0] = true
    ∴ flag[0] = true   ... (3)
    
  By first conjunct of I applied to process 1:
    pc[1] = cs → flag[1] = true
    ∴ flag[1] = true   ... (4)
    
  From (1) and (4): Since flag[1] = true, we must have turn = 0.
  From (2) and (3): Since flag[0] = true, we must have turn = 1.
  
  But turn cannot be both 0 and 1. Contradiction!
  
∴ ¬(pc[0] = cs ∧ pc[1] = cs). MUTEX holds. ∎

Mutual Exclusion Proven

Variant-Based Proofs for Liveness

Liveness properties assert that something good eventually happens. We cannot prove such properties using invariants alone—invariants only constrain what CAN happen, not what WILL happen.

The Variant Method

To prove liveness, we use variant functions (also called ranking functions or progress measures):

A variant is a function V: States → W where:

W is a well-founded set (no infinite descending chains)
Progress: each step toward the goal decreases V
Bounded: V cannot decrease forever

Common choices for W:

Natural numbers (with usual ordering)
Tuples of naturals (lexicographic ordering)
Finite ordinals

Proving Progress

To prove progress (if CS is empty and processes contend, one enters):

Define a variant measuring 'distance' from entry
Show under fairness, the variant decreases
Since the variant can't decrease forever, entry occurs

variant_proof_structure.txt
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// Structure of a Variant-Based Liveness Proof
 
// Goal: Prove that waiting processes eventually enter CS
 
// STEP 1: Define the progress condition
// C(s) ≡ "processes are waiting but CS is empty"
 
// STEP 2: Define the variant V
// V: States → ℕ  (or some well-founded set)
// V measures "how far" we are from a process entering CS
 
Example for Peterson's:
  V(s) = (2 if both in wait, 1 if one in wait and one not, 0 otherwise)
  
  But this is too coarse. Better:
  
  V(s) = < number of processes ahead in "queue",
           number of steps until guard satisfied >
 
// STEP 3: Show V decreases while C holds
// Under fairness, show that eventually:
//   Either C becomes false (a process enters - goal achieved!)
//   Or V decreases
 
Proof sketch:
  Case: Both processes in wait state, turn = j.
  Then process i with turn = i will satisfy guard (flag[j] ∨ turn = i).
  Since turn = j (for j), process j is blocked.
  But process i's guard (flag[j] ∨ turn = i) evaluates to (true ∨ false) = true.
  Wait, let me reconsider...
  
  If turn = j, then:
  - Process i's guard: flag[j] = true (j is waiting), turn = i = (1-j) ≠ j.
    So flag[j] = true and turn ≠ i. Guard is false for i.
  - Process j's guard: flag[i] = true (i is waiting), turn = j. 
    So turn = j. Guard (flag[i]=false ∨ turn=j) becomes (false ∨ true) = true!
    
  So process j enters. V decreases (one fewer waiting). ✓
 
// STEP 4: Conclude
// Since W is well-founded and V decreases while C holds,
// eventually C becomes false (some process in CS).
// ∴ Progress holds.

Proving Bounded Waiting

Bounded waiting requires a stronger argument: not just that SOME process enters, but that EACH waiting process enters within bounded bypasses.

We use a variant that tracks the specific waiting process:

Vᵢ(s) = "distance for process i to enter CS"

And show that:

Each bypass by other processes decreases Vᵢ (or maintains it)
The bound is computable from the initial Vᵢ

Fairness is Essential

Complete Proof: Peterson's Progress and Bounded Waiting

Let us complete the correctness proof for Peterson's algorithm by proving progress and bounded waiting.

Proof of Progress

peterson_progress_proof.txt
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// Proof: Peterson's algorithm satisfies Progress
 
// Goal: If CS is empty and at least one process is in entry (want or wait),
//       eventually some process enters CS.
 
// Assumption: Weak fairness (continuously enabled → eventually executes)
 
// Case 1: One process (say i) is in entry, the other (j) is in remainder.
 
  Process i is at want or wait.
  j is in remainder (idle state), so flag[j] = false.
  
  If i is at want:
    Under fairness, i eventually executes want → wait (sets turn = j).
    Now i is at wait.
    
  i is at wait:
    Guard: flag[j] = false ∨ turn = i
    Since flag[j] = false (j in remainder), guard is TRUE.
    Under fairness, i eventually executes wait → cs.
    
  ∴ Process i enters CS. Progress satisfied. ✓
 
// Case 2: Both processes are in entry.
 
  Both at want or wait states.
  
  Sub-case 2a: At least one is at want.
    Under fairness, it executes want → wait.
    Eventually, both are at wait.
    
  Sub-case 2b: Both at wait.
    flag[0] = true, flag[1] = true (by invariant).
    turn = 0 or turn = 1.
    
    If turn = 0:
      Process 0's guard: flag[1] = true, turn = 0. 
        Guard = (true ∨ true) = true... wait, that's wrong.
        Guard = (flag[1] = false) ∨ (turn = 0) = false ∨ true = true. ✓
      Process 0 can proceed.
      Under fairness, process 0 enters CS.
      
    If turn = 1:
      By symmetric argument, process 1 enters CS.
    
  ∴ Some process enters CS. Progress satisfied. ✓
 
// Conclusion: Under weak fairness, Peterson's algorithm has Progress. ∎

Proof of Bounded Waiting

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// Proof: Peterson's algorithm satisfies Bounded Waiting with B = 1
 
// Goal: After process i enters its entry section, process j can enter
//       CS at most once before process i enters.
 
// Assume process i enters entry section at time t₀ (sets flag[i] = true).
 
// Case 1: Process j is not in entry/CS at time t₀.
  j is in remainder (idle), so flag[j] = false.
  
  As shown in progress proof, i will enter CS.
  j bypasses i 0 times. ✓
 
// Case 2: Process j is in CS at time t₀.
  j will eventually exit (finite CS execution assumption).
  When j exits:
    - j sets turn = i (in exit section)
    - j sets flag[j] = false
  
  Now, if j wants to enter again:
    j sets flag[j] = true
    j sets turn = i (in want → wait)
    j arrives at wait with turn = i
    
  In the meantime, i is in entry (want or wait).
  
  If i is at wait:
    turn = i (set by j's exit or j's entry!)
    i's guard: flag[j] = true, turn = i
    Guard = (true ∨ true) = true... no wait,
    Guard = (flag[j] = false) ∨ (turn = i) = false ∨ true = TRUE. ✓
    i can enter.
    
  If i is at want:
    Under fairness, i moves to wait.
    i sets turn = j.
    But then j's exit will set turn = i.
    Eventually i sees turn = i and enters.
    
  Key insight: After j exits CS and before i enters CS,
  if j sets turn again (in want→wait), it sets turn = i.
  This guarantees i's guard becomes true.
  
  j entered CS once (was already in at t₀), i then enters.
  j bypasses i at most 1 time. ✓
 
// Case 3: Process j is in entry at time t₀.
  j is at want or wait.
  turn is either 0 or 1.
  
  If turn = j (i.e., it's j's turn):
    j can enter (guard satisfied).
    j enters CS (bypass 1).
    j exits, sets turn = i.
    Now i's guard: (flag[j] = false) ∨ (turn = i).
    After j's exit, flag[j] = false, so guard is true.
    i enters.
    j bypassed i once. ✓
    
  If turn = i:
    j's guard: flag[i] = true, turn = i ≠ j. Guard false.
    j cannot enter.
    i's guard: flag[j] = true, turn = i. Guard true.
    i enters.
    j bypassed i zero times. ✓
 
// Conclusion: In all cases, j enters at most once before i.
// Bounded Waiting holds with B = 1. ∎

Complete Proof Achieved

Summary of Proof Techniques

Let us consolidate the proof techniques we have developed:

Technique Taxonomy

Proof Techniques by Property
Property	Type	Technique	Key Elements
Mutual Exclusion	Safety	Invariant	Find I such that I is inductive and I → MUTEX
Absence of Deadlock	Safety	Invariant	Show ¬(all waiting ∧ none can proceed)
Progress	Liveness	Variant + Fairness	Show variant decreases until entry
Bounded Waiting	Liveness	Bypass Counting	Count max bypasses per waiting process
Termination	Liveness	Variant	Show execution eventually completes

Common Invariant Patterns

Certain invariant patterns recur across synchronization algorithms:

Reusable Invariant Patterns

•Location → Variable: at(L) → var = val (e.g., in CS implies lock held)
•Gate condition: in_cs[i] → gate_condition[i] where gate prevents simultaneous entry
•Turn exclusivity: turn has one value at a time, so two different turn values cannot hold
•Flag semantics: flag[i] = true ↔ process i is interested
•Conservation: Σ in_cs[i] + Σ waiting[i] + Σ idle[i] = n (all processes accounted for)

Common Proof Pitfalls

Common Errors in Correctness Proofs
Pitfall	Description	Remedy
Weak invariant	Invariant doesn't imply property	Strengthen with additional conjuncts
Non-inductive	Invariant not preserved by all transitions	Strengthen or relax invariant
Missed transition	Forgot to check some program action	Systematically enumerate all actions
Conflated atomicity	Treated non-atomic as atomic	Decompose into true atomic steps
Omitted fairness	Proved liveness without fairness	State fairness assumption explicitly
Case explosion	Didn't use structure to simplify	Use symmetry and case merging

Proof Development Process

Start with the property to prove. 2. Guess an invariant (for safety) or variant (for liveness). 3. Try to prove initialization and preservation. 4. When a step fails, analyze why and strengthen. 5. Iterate until complete. 6. Verify the invariant implies the desired property.

Mechanized Proofs and Verification Tools

While hand-written proofs are valuable for understanding, they are error-prone. Mechanized proofs use computer assistance to verify correctness with mathematical certainty.

Interactive Theorem Provers

Provers like Coq, Isabelle/HOL, and Lean allow humans to construct proofs interactively, with the computer verifying each step:

Human provides proof strategy and key insights
Computer checks all logical steps rigorously
Final proof is machine-verified

Model Checkers (State-Space Exploration)

Tools like SPIN, TLA+, and UPPAAL exhaustively explore the state space:

Automatically check all reachable states
Report counterexamples if property fails
Limited by state explosion, but highly automated

tlaplus_peterson.tla
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--------------------------- MODULE Peterson ---------------------------
EXTENDS Integers
 
VARIABLES flag, turn, pc
 
vars == <<flag, turn, pc>>
 
Init == /\ flag = [i \in {0, 1} |-> FALSE]
        /\ turn \in {0, 1}
        /\ pc = [i \in {0, 1} |-> "idle"]
 
(* Process i transitions *)
Enter(i) == /\ pc[i] = "idle"
            /\ pc' = [pc EXCEPT ![i] = "want"]
            /\ flag' = [flag EXCEPT ![i] = TRUE]
            /\ UNCHANGED turn
 
SetTurn(i) == LET j == 1 - i IN
              /\ pc[i] = "want"
              /\ pc' = [pc EXCEPT ![i] = "wait"]
              /\ turn' = j
              /\ UNCHANGED flag
 
EnterCS(i) == LET j == 1 - i IN
              /\ pc[i] = "wait"
              /\ (flag[j] = FALSE \/ turn = i)
              /\ pc' = [pc EXCEPT ![i] = "cs"]
              /\ UNCHANGED <<flag, turn>>
 
ExitCS(i) == LET j == 1 - i IN
             /\ pc[i] = "cs"
             /\ pc' = [pc EXCEPT ![i] = "idle"]
             /\ flag' = [flag EXCEPT ![i] = FALSE]
             /\ turn' = j
 
Next == \E i \in {0, 1}: Enter(i) \/ SetTurn(i) \/ EnterCS(i) \/ ExitCS(i)
 
(* Mutual exclusion property *)
MutualExclusion == ~(pc[0] = "cs" /\ pc[1] = "cs")
 
(* Specification *)
Spec == Init /\ [][Next]_vars /\ WF_vars(Next)
 
THEOREM Spec => []MutualExclusion
 
=========================================================================

Industrial Use of Verification

Formal verification is used in high-stakes domains:

Hardware verification: CPU designs are now routinely model-checked
Aerospace: Flight control systems require formal proofs
Security protocols: TLS, authentication protocols are verified
Distributed systems: Amazon uses TLA+ for DynamoDB, S3
Compilers: CompCert is a formally verified C compiler

For synchronization specifically, libraries like liblfds (lock-free data structures) have undergone formal verification.

When to Use Formal Verification

Summary: Correctness Proofs

We have explored the rigorous mathematical foundations of correctness proofs for synchronization algorithms. Let us consolidate the key concepts.

Key Takeaways

•Formal proofs provide mathematical certainty that algorithms meet their requirements, transcending what testing can achieve.
•Safety properties (like mutual exclusion) are proved using inductive invariants—predicates true initially and preserved by all transitions.
•Liveness properties (like progress) require fairness assumptions and variant functions that decrease toward the goal.
•Proving mutual exclusion involves finding an invariant that implies no two processes are in CS simultaneously, then proving the invariant is inductive.
•Proving progress involves showing that from any waiting state, under fairness, some process eventually enters CS.
•Proving bounded waiting involves counting the maximum bypasses a process can experience and showing it's finite.
•Mechanized verification tools (TLA+, Coq, SPIN) provide machine-checked certainty and are used in industry for critical systems.

Module Complete

With this page, we have completed Module 3: Requirements for Solutions. You now possess:

Deep understanding of the three fundamental requirements: mutual exclusion, progress, and bounded waiting
Analytical skills to evaluate any synchronization solution systematically
Proof techniques to rigorously verify that algorithms are correct

Next chapters will explore:

Peterson's and Dekker's algorithms in detail
Hardware atomic operations (test-and-set, compare-and-swap)
OS synchronization primitives (mutexes, semaphores, monitors)
Classic synchronization problems (producer-consumer, readers-writers, dining philosophers)

Module Complete