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Beta equivalence is the smallest equivalence relation containing beta reduction; equivalently, $M{\equiv }_{\beta }N$ when a finite sequence of beta reductions and reversed beta reductions connects $M$ and $N$.

A beta-redex is a subterm of the form $(\lambda x.M)N$, to which beta reduction can be applied.

Two distinct closed lambda terms in normal form for both beta reduction and eta conversion admit a separating context that maps them respectively to the two Church Booleans. Finite Böhm trees and the Böhm-out lemma isolate their first genuine head disagreement. The qualification about eta conversion matters: distinct beta-normal forms that are eta-equivalent cannot be separated this way.

A reduction relation has the Church-Rosser property when its equivalence relation generated by forward and backward steps satisfies $b{\leftrightarrow }^{\ast }c\Rightarrow \exists d(b{\to }^{\ast }d{\leftarrow }^{\ast }c)$. It is equivalent to confluence of ${\to }^{\ast }$. The Church-Rosser theorem establishes it for beta reduction.

A reduction relation $\to$ has the diamond property if $a\to b$ and $a\to c$ imply a common one-step successor $b\to d\leftarrow c$. For relations allowing an identity step, the same definition includes zero-length joins. A finite grid of these diamonds proves confluence of ${\to }^{\ast }$. Ordinary one-step beta reduction fails this property because contracting before duplication can save two separate contractions.

Eta conversion expresses extensionality for a lambda abstraction whose fresh argument is only passed to a function. The side condition prevents removal of a genuine dependency. In a categorical model with $D\cong D^D$, the inverse laws for currying and evaluation validate this conversion as well as beta reduction.

A lambda term is a variable $x$, a lambda application $(MN)$, or a lambda abstraction $\lambda x.M$. The lambda abstraction binds $x$ in $M$ and represents a function whose argument will be substituted for $x$. Computation uses beta reduction:where the substitution is capture-avoiding substitution. For example, $(\lambda x.x)N$ reduces to $N$, while substituting a free $y$ into $\lambda y.x$ first requires renaming its bound variable, giving $\lambda z.y$ rather than $\lambda y.y$.

The untyped lambda calculus allows self-application. If $D=\lambda x.xx$, then $DD$ reduces to itself and has no terminating reduction. Conversely, Church numerals encode natural numbers as repeated application: $\lambda f.\lambda x.f(fx)$ represents two. This small syntax supports general computation while making binding, substitution and termination visible. The simply typed lambda calculus restricts the allowed lambda applications through types.

Represent total machine configuration transitions and tests on Church numerals and Church pairs, then use a fixed-point combinator to iterate until the halting test selects the output. Normal-order beta reduction follows the computation. A nonhalting machine gives an infinite required reduction, so the normal-order normalization theorem rules out a numeral normal form.

If an untyped lambda term has a beta-normal form, leftmost outermost beta reduction reaches a normal form. Thus an infinite normal-order computation certifies the absence of a beta-normal form. This is the normalization result needed to make a lambda simulation respect undefined outputs.

The omega combinator is the divergent untyped lambda calculus termwhose only beta reduction reproduces $\Omega$.

Parallel beta reduction permits simultaneous reduction in the body of a lambda abstraction and both sides of an application, together with $(\lambda x.M)N\Rightarrow M^{\prime }[x:=N^{\prime }]$ when $M\Rightarrow M^{\prime }$ and $N\Rightarrow N^{\prime }$. Variables reduce to themselves. Its capture-avoiding substitution lemma is $M\Rightarrow M^{\prime },N\Rightarrow N^{\prime }\Rightarrow M[x:=N]\Rightarrow M^{\prime }[x:=N^{\prime }]$. Its transitive closure equals ordinary finite beta reduction, while its one-step diamond property yields the Church-Rosser theorem.

Choose all auxiliary bound variables fresh for the free variables of $p,q,r$; alpha equivalence permits this. The operation $\bullet$ is composition of lambda terms, so expand its definition and use beta reduction:The two lambda terms have the same reduct, up to alpha equivalence, hence they are related by beta equivalence:No eta conversion is needed for this associativity calculation.

Write $I\equiv \lambda u.u$ for the I combinator, and choose $x$ outside the free variables of $p$. Both orders of composition of lambda terms first give the same lambda abstraction by beta reduction:The freshness condition is exactly the side condition for eta conversion, which contracts $\lambda x.px$ to $p$. ThereforeHere eta conversion matters: a free variable $p$ is already in beta-normal form, as is $\lambda x.px$, and beta equivalence alone does not equate them.

The argument order in $p$ is a cyclic permutation. Undo it by rotating in the other direction. For $n\ge 1$, choose the lambda termUse fresh bound variables throughout. Substituting $qz$ for the first argument of $p$ and then performing beta reduction givesIndeed, the arguments received by $qz$ are $y_1=x_2,\dots ,y_{n-1}=x_n,y_n=x_1$, so its last argument moves back to the front. Apply eta conversion successively, starting with the innermost lambda abstraction. At each step the bound variable being removed is absent from the preceding expression. Thus the last expression is related to $z$ by beta-eta equivalence, and consequentlyFor the induction needed in the next part, also compute the reverse composition of lambda terms:So this $q$ is a two-sided inverse. The empty argument list at $n=0$ means $p=\lambda z.z$; choose $q=I$. For $n=1$, both displayed rotations give $\lambda z{y}_1.z{y}_1$, which is also related to $I$ by eta conversion. These boundary cases require no exceptional inverse construction.

Use structural induction on the construction of $p\in \mathcal{R}$, with the stronger induction claim that $p$ is an invertible lambda term: there is a lambda term $q$ satisfying bothThe base constructor has exactly the two-sided inverse proved in part (c). Strengthening the claim is useful because the inductive construction needs inverses of its constituent lambda terms in both orders.

For the inductive constructor, let $t_i$ be the two-sided inverse supplied for $s_i$. Thus, for a fresh variable $u$, beta-eta equivalence givesFor $n\ge 1$, defineThis lambda term reverses the cyclic permutation and applies the appropriate inverse to each argument. First use beta reduction to expand one order of composition of lambda terms:For the other order, beta reduction givesThese computations close the stronger induction. For $n=1$, each formula has the single argument involving $s_1$ and $t_1$; for $n=0$, the constructor is $I$ and its inverse is $I$. Every construction is finite, so the inverse is a finite lambda term. There is no requirement that this inverse itself belong to $\mathcal{R}$.

The exact restriction in the ordinary simply typed lambda calculus isThus the unrestricted choice of $A,B$ available in part (c) fails. A reading of the printed sentence as asserting impossibility for every choice of $A,B$ is false. This issue occurs in the original PDF as well as its transcription.^[ref]

To prove necessity, first observe why the ordinary rules cannot hide assumptions. Structural induction on a derivation provesFor a variable, both sides are the singleton containing that variable. For lambda application, both sides are the union of the two premise supports. For lambda abstraction, both sides delete the bound variable. There is no universal typing axiom to interrupt this equality.

Now invert the proposed derivation. Each outer lambda abstraction must have been introduced by the abstraction rule, since the other ordinary rules conclude a variable or a lambda application. The body must therefore have the typing judgmentThe support equality just proved ensures that both displayed assumptions occur; consistency fixes their simple types to those specified by the outer function type. Inverting the lambda application gives a simple type $C$ for its argument and the following premises:There is no flexibility in the supports: $\lambda z.x$ has the single free variable $x$, whereas $yx$ has both free variables $y,x$. Invert the second lambda application. Its two premises must befor some simple type $D$. The variable rule is the only possible final rule for each premise. It forces $D=A$ and $B=D\Rightarrow C$, yielding $B=A\Rightarrow C$.

Conversely, suppose $B=A\Rightarrow C$. The ordinary rules giveThe typing contexts are consistent, so another lambda application gives the body type $A$. Two lambda abstractions discharge $y$ and $x$, establishingThis proves sufficiency and supplies a family of counterexamples to the literal impossibility claim. For a concrete failure of the unrestricted claim, choose $A=\alpha$ and $B=\beta$ as atomic type variables. The type variable $\beta$ is not a function type $\alpha \Rightarrow C$, so no derivation exists for that fixed pair. Here the atomic type variables are fixed type expressions, rather than unknowns that one may change during type inference.

The reason for the restriction is the otherwise discarded lambda application $yx$: the ordinary rules still type-check it. Its subsequent disappearance under beta reduction does not license reversing that beta reduction while preserving every simple type of the result. The beta-normal form $\lambda xy.x$ admits arbitrary $A\Rightarrow (B\Rightarrow A)$, whereas the unreduced lambda term has the additional constraint just derived.

Write ${\to }^{\ast }$ for the reflexive closure of the transitive closure of a relation $\to$, and ${\leftrightarrow }^{\ast }$ for its generated equivalence relation. The diamond property requires every pair $a\to b,a\to c$ to admit a one-step join $b\to d\leftarrow c$. The Church-Rosser property requires a possibly longer join whenever $b{\leftrightarrow }^{\ast }c$. Equivalently, every pair of finite reductions with the same source has a common finite reduct. Identity steps may be included in the diamond convention without affecting the counterexample below.

Let $A=(\lambda z.z)y$ and consider the lambda term $M=(\lambda x.xx)A$. Contracting its outer beta-redex gives $AA$, whereas contracting the inner beta-redex gives $(\lambda x.xx)y$. The only one-step reducts of $AA$ are $yA$ and $Ay$; the other branch has the sole one-step reduct $yy$. These terms are different even under alpha equivalence. Thus one-step beta reduction has no diamond property. Both branches do reach $yy$ by finite beta reduction.

For the positive result, use parallel beta reduction $\Rightarrow$, with rulesStructural induction, renaming binders to avoid capture, proves the substitution lemma: $P\Rightarrow P^{\prime }$ and $Q\Rightarrow Q^{\prime }$ imply $P[x:=Q]\Rightarrow P^{\prime }[x:=Q^{\prime }]$. Define the complete development of a lambda term byand $(PQ{)}^{\bullet }=P^{\bullet }{Q}^{\bullet }$ when the application is not initially a root beta-redex. Induction on a parallel derivation, using the substitution lemma in the contracting case, provesConsequently any two parallel reducts have the common successor $M^{\bullet }$. The ordinary one-step relation is included in the parallel relation, and every parallel step expands into finitely many ordinary steps; their finite closures therefore coincide. A finite grid of parallel diamonds joins any two finite parallel reductions. Finally a finite zigzag of forward and backward reductions can be joined inductively using that confluence. Hence beta reduction has the Church-Rosser property, although its individual steps do not have the diamond property.

A fixed-point combinator is a closed lambda term $Z$ such that $ZF{\equiv }_{\beta }F(ZF)$ for every lambda term $F$. It suffices to check this for a fresh free variable $f$, since capture-avoiding substitution then gives every instance. TakeIf $GZ{\equiv }_{\beta }Z$, applying both sides to $f$ gives $Zf{\equiv }_{\beta }f(Zf)$. Conversely suppose $Z$ is a fixed-point combinator. Its application to a fresh $f$ has a head normal form, because it is beta equivalent to $f(Zf)$. Since $Z$ is closed, head reduction of $Zf$ must first expose a lambda abstraction in $Z$ itself. Thus $Z{\to }_{\beta }^{\ast }\lambda u.P$, andThis establishes the needed abstraction equality by beta reduction, without assuming eta conversion. Abstracting $Zf{\equiv }_{\beta }f(Zf)$ now gives $Z{\equiv }_{\beta }\lambda f.f(Zf){\equiv }_{\beta }GZ$.

Two examples are the Curry fixed-point combinator and the Turing fixed-point combinator:Writing $A_f=\lambda x.f(xx)$, we have $Yf{\to }_{\beta }{A}_f{A}_f{\to }_{\beta }f(A_f{A}_f){\equiv }_{\beta }f(Yf)$, while $\Theta f{\to }_{\beta }^{2}f(\Theta f)$. Different displayed syntax alone would not prove that the two combinators are distinct under beta equivalence. Here is a reduct invariant that does. Every finite reduct of $Y$ is $\lambda f.f^n(A_f{A}_f)$, up to alpha equivalence: $A_f$ is beta-normal and the one remaining beta-redex is its self-application. No such reduct contains the closed subterm $DD$.

Every finite reduct of $\Theta$, in contrast, retains a closed subterm $DD$. The contraction $DD\to \lambda f.f(DDf)$ recreates one. The only other redexes created by these expansions are abstractions applied to variables; contracting them substitutes a variable, so preserves any closed $DD$ inside their bodies. This describes all redexes inductively: expansion inserts only a variable-headed application and a copy of $DD$, and a variable substitution preserves that description. Thus every finite reduct contains $DD$. A common reduct of $Y$ and $\Theta$ is impossible, so the Church-Rosser theorem proves $Y/{\equiv }_{\beta }\Theta$.

Finally put $P=ZK$, with $K=\lambda xy.x$. The fixed-point combinator equation giveswhere $y$ is fresh. If $P$ had a head normal form with $n$ leading abstractions, then $\lambda y.P$ would have one with $n+1$. A reduct of a head normal form preserves its initial abstraction count and variable-headed spine: only its argument subterms can still contract. These two head normal forms therefore cannot have a common reduct, contradicting the Church-Rosser theorem. Hence $ZK$ is unsolvable.

Use Church numeralsExplicit lambda terms for the successor and zero test areDirect beta reduction gives $\mathsf{Succ}{c}_n{\equiv }_{\beta }{c}_{n+1}$, and the Church numeral zero test returns $\mathsf{T}$ exactly at zero. For the predecessor, use the Church pair operationsand setAfter $j$ iterations this pair holds $(c_j,c_{\max (j-1,0)})$: the initial pair gives the case $j=0$, and one step moves the old first coordinate into the second and increments the first. This proves the predecessor equation, including $\mathsf{Pred}{c}_0{\equiv }_{\beta }{c}_0$.

A total function is a lambda-definable function when a closed lambda term sends each tuple of numerals to the numeral of its value. Suppose $G,H$ represent total $g,h$, andFor lambda definition of primitive recursion by pair iteration, putOn numeral inputs, induction gives the pair $(c_j,c_{r(\mathbf{x},j)})$ after $j$ iterations. The base case is the equation for $G$, and the induction step is exactly the equation for $H$. Hence $R$ represents $r$, proving closure of total lambda-definable functions under primitive recursion.

For unbounded minimization, the careful closure statement is partial: if total $f(\mathbf{x},n)$ is lambda-definable, thenis a lambda-definable partial function, defined when a zero exists and otherwise undefined. It need not be total. With $F$ representing $f$, use a fixed-point combinator to defineEach test terminates because $f$ is total. Normal-order beta reduction chooses the return branch at the first zero and does not evaluate the discarded branch. If there is no zero, the required head computation passes through tests for all $n$, never exposing a numeral; standardization theorem for beta reduction rules out a hidden numeral reduct. For partial $f$, the usual mu operator is defined only if all earlier tests terminate and the successful test is zero. This stronger partial closure requires representatives that diverge on undefined inputs, or an explicit strict evaluator; the weaker condition of producing no numeral on undefined inputs alone must not be used as a test for convergence.

The zero function, successor, and projections have immediate numeral representations; function composition in recursion theory is implemented by substitution on total inputs. Thus every primitive recursive function is lambda-definable. The Kleene normal form theorem writes every partial recursive function as a total primitive-recursive decoding after a single minimization of a total primitive-recursive computation-history predicate. The search above consequently represents every partial recursive function, diverging when no successful history exists. Conversely a machine can effectively enumerate the beta reductions of a fixed representing lambda term on numeral inputs until it finds a Church numeral; distinct numerals cannot be beta-equivalent by the Church-Rosser theorem. It computes exactly the represented partial function. Therefore lambda-definable partial functions are precisely partial computable functions, and the total ones are precisely total computable functions.

This also gives the full partial minimization closure without assuming that an arbitrary weak representative diverges on every undefined input. Run its effective numeral evaluator successively at $n=0,1,2,\dots$; stop at the first zero, and run forever if an earlier evaluation fails to terminate or no zero is found. This is exactly the sequential-domain convention for the mu operator. The resulting partial computable function has a lambda representation by the construction just proved. Thus partial lambda-definable functions are closed under partial minimization; a minimized function belongs to the total family only when the required successful search exists for every input.

The Böhm tree $BT(M)$ is $\perp$ if $M$ is an unsolvable lambda term. Otherwise choose a head normal formIts root is labeled by the binder prefix and head variable, and its ordered children are $BT(M_1),\dots ,BT(M_m)$. Repeat recursively, possibly forever. Bound names are identified by alpha equivalence; the Church-Rosser theorem makes this tree independent of the selected head reduct. An address $\alpha$ is a finite sequence of child numbers; $M_{\alpha }$ denotes the corresponding argument term encountered after these head reductions.

A Böhm transformation is a finite composition of capture-avoiding substitutions and applications to chosen auxiliary terms. Here is a constructive Böhm-out lemma. Along the finite path to $\alpha$, choose $q$ at least the number of children of each preceding head node, and defineInitially replace all relevant free variables by $U_q$. At a head node with $n$ initial binders and $m$ children, supply its $n$ binders with $U_q$. Its head is now $U_q$ and its arguments are the substituted children $N_1,\dots ,N_m$. Apply $q-m$ extra copies of $U_q$ as padding and then $P_i^q$. Direct beta reduction givesRepeat with the next child number. The term reached at each stage is the selected original child with the accumulated substitutions for its ancestors' binders. Thus the finite composite $\pi$ satisfies $M\pi {\equiv }_{\beta }{M}_{\alpha }\sigma$ for a suitable substitution $\sigma$. The empty path uses the identity transformation. Tupling, rather than immediately replacing a head variable by a fixed projection, is important: the same variable can occur later where a different child must be selected.

For the Böhm separation theorem, two distinct closed normal forms for both beta reduction and eta conversion have finite Böhm trees with a genuine disagreement in a head label or argument structure. A Böhm transformation first extracts their common-prefix node. Projections at a differing head variable select different supplied Church Booleans; differing arities are handled by padding, with the exclusion of mere eta expansions ensuring a genuine disagreement. This yields a context separating the two normal forms. Distinct beta-normal forms that differ only by eta conversion are not covered by this separation statement.

The PDF prints a free $\alpha$ in both the definitions of $K$ and $\mathsf{T}$. Taken literally, this makes the last request impossible. To see why, suppose $M{K}_{\alpha }{\equiv }_{\beta }{K}_{\alpha }$, where $K_{\alpha }=\lambda xy.\alpha$, and choose a fresh $f$. By solvability and its stability under substitution, $Mf$ has a head normal form $\lambda x_1\cdots x_n.h{P}_1\cdots P_r$. If $h\ne f$, substituting either $K_{\alpha }$ or $S$ leaves the same variable-headed spine; it cannot be $K_{\alpha }$ in one case and $\mathsf{F}$ in the other. If $h=f$, obtaining $K_{\alpha }$ requires $r\le 2$ and $n+2-r=2$, hence $n=r$. Substituting $S=\lambda xyz.xz(yz)$ then leaves at least $n+3-r=3$ leading abstractions, whereas $\mathsf{F}$ has two. A variable-headed normal form cannot lose initial abstractions under beta reduction. The Church-Rosser theorem rules out the required equality.

For the intended closed first projection $K=\mathsf{T}=\lambda xy.x$ and $\mathsf{F}=\lambda xy.y=KI$, an explicit answer isIndeed $MK{\to }_{\beta }^{\ast }(KK)I{\to }_{\beta }K=\mathsf{T}$, whereas $MS{\to }_{\beta }^{\ast }(KK)I((KI)I){\to }_{\beta }KI=\mathsf{F}$. These computations solve the corrected closed-combinator version while identifying the obstruction in the literal printed one.