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@@ 1,3 +1,56 @@
+[[!toc]]
+
+Substitution and AlphaConversion
+=================================
+
+Intuitively, (a) and (b) express the application of the same function to the argument `y`:
+
+
+(\x. \z. z x) y
+(\x. \y. y x) y
+
+
+One can't just rename variables freely. (a) and (b) are different than what's expressed by:
+
+
+(\z. (\z. z z) y
+
+
+
+Substituting `y` into the body of `(\x. \z. z x)` is unproblematic:
+
+ (\x. \z. z x) y ~~> \z. z y
+
+However, with (b) we have to be more careful. If we just substituted blindly, then we might take the result to be `\y. y y`. But this is the selfapplication function, not the function which accepts an arbitrary argument and applies that argument to the free variable `y`. In fact, the selfapplication function is what (c) reduces to. So if we took (b) to reduce to `\y. y y`, we'd wrongly be counting (b) to be equivalent to (c), instead of (a).
+
+To reduce (b), then, we need to be careful to that no free variables in what we're substituting in get captured by binding λs that they shouldn't be captured by.
+
+In practical terms, you'd just replace (b) with (a) and do the unproblematic substitution into (a).
+
+What attitude should we have to this?
+
+One way to think of it is to identify expressions of the lambda calculus with particular alphabetic sequences. Then (a) and (b) would be distinct expressions, and we'd have to explicitly articulate a rule permitting you to do the kind of variablerenaming that would take you from (a) to (b) (or vice versa). This kind of renaming is called "alphaconversion."
+
+Another way to think of it is to identify expressions not with particular alphabetic sequences, but rather with classes of alphabetic sequences, which stand to each other in the way that (a) and (b) do. That's the way we'll talk. We say that (a) and (b) are just typographically different notations for a *single* lambda formula. As we'll say, the lambda formula written with (a) and the lambda formula written with (b) are literally syntactically identical.
+
+A third way to think is to identify the lambda formula not with classes of alphabetic sequences, but rather with abstract structures that we might draw like this:
+
+
+ λ ... ___ ...
+ ^ 
+ ______
+
+
+Here there are no bound variables, but there are *bound positions*. We can regard formula like (a) and (b) as just helpfully readable ways to designate these abstract structures.
+
+A version of this last approach is known as **de Bruijn notation** for the lambda calculus.
+
+It doesn't matter which of these approaches one takes; the logical properties of the systems are exactly the same. It just affects the particulars of how one states the rules for substitution, and so on. And whether one talks about expressions being literally "syntactically identical," or whether one instead counts them as "equivalent modulu alphaconversion."
+
+(In a bit, we'll discuss other systems that lack variables. Those systems will not just lack variables in the sense that de Bruijn notation does; they will furthermore lack any notion of a bound position.)
+
+
+
Syntactic equality, reduction, convertibility
=============================================
@@ 10,8 +63,12 @@ Define T to be `(\x. x y) z`. Then T and `(\x. x y) z` are syntactically equal,
equivalent to `(\z. z y) z` is that when a lambda binds a set of
occurrences, it doesn't matter which variable serves to carry out the
binding. Either way, the function does the same thing and means the
same thing. Look in the standard treatments for discussions of alpha
equivalence for more detail.]
+same thing.
+Linguistic trivia: some linguistic discussions suppose that alphabetic variance
+has important linguistic consequences (notably Ivan Sag's dissertation).
+Look in the standard treatments for discussions of alpha
+equivalence for more detail. Also, as mentioned below, one of the intriguing
+properties of Combinatory Logic is that alpha equivalence is not an issue.]
This:
@@ 30,16 +87,22 @@ Lambda expressions that have no free variables are known as **combinators**. Her
> **I** is defined to be `\x x`
> **K** is defined to be `\x y. x`, That is, it throws away its
+> **K** is defined to be `\x y. x`. That is, it throws away its
second argument. So `K x` is a constant function from any
(further) argument to `x`. ("K" for "constant".) Compare K
 to our definition of **true**.
+ to our definition of `true`.
+
+> **getfirst** was our function for extracting the first element of an ordered pair: `\fst snd. fst`. Compare this to K and `true` as well.
+
+> **getsecond** was our function for extracting the second element of an ordered pair: `\fst snd. snd`. Compare this to our definition of `false`.
> **getfirst** was our function for extracting the first element of an ordered pair: `\fst snd. fst`. Compare this to **K** and **true** as well.
+> **B** is defined to be: `\f g x. f (g x)`. (So `B f g` is the composition `\x. f (g x)` of `f` and `g`.)
> **getsecond** was our function for extracting the second element of an ordered pair: `\fst snd. snd`. Compare this to our definition of **false**.
+> **C** is defined to be: `\f x y. f y x`. (So `C f` is a function like `f` except it expects its first two arguments in swapped order.)
> **ω** is defined to be: `\x. x x`
+> **W** is defined to be: `\f x . f x x`. (So `W f` accepts one argument and gives it to `f` twice. What is the meaning of `W multiply`?)
+
+> **ω** (that is, lowercase omega) is defined to be: `\x. x x`
It's possible to build a logical system equally powerful as the lambda calculus (and readily intertranslatable with it) using just combinators, considered as atomic operations. Such a language doesn't have any variables in it: not just no free variables, but no variables at all.
@@ 48,22 +111,153 @@ One can do that with a very spare set of basic combinators. These days the stand
There are some wellknown linguistic applications of Combinatory
Logic, due to Anna Szabolcsi, Mark Steedman, and Pauline Jacobson.
Szabolcsi supposed that the meanings of certain expressions could be
insightfully expressed in the form of combinators. A couple more
combinators:

 **C** is defined to be: `\f x y. f y x` [swap arguments]
+insightfully expressed in the form of combinators.
 **W** is defined to be: `\f x . f x x` [duplicate argument]
For instance, Szabolcsi argues that reflexive pronouns are argument
duplicators.
![test](./szabolcsireflexive.png)
+![reflexive](http://lambda.jimpryor.net/szabolcsireflexive.jpg)
+
+Notice that the semantic value of *himself* is exactly `W`.
+The reflexive pronoun in direct object position combines first with the transitive verb (through compositional magic we won't go into here). The result is an intransitive verb phrase that takes a subject argument, duplicates that argument, and feeds the two copies to the transitive verb meaning.
+
+Note that `W <~~> S(CI)`:
+
+S(CI) ≡
+S((\fxy.fyx)(\x.x)) ~~>
+S(\xy.(\x.x)yx) ~~>
+S(\xy.yx) ≡
+(\fgx.fx(gx))(\xy.yx) ~~>
+\gx.(\xy.yx)x(gx) ~~>
+\gx.(gx)x ≡
+W
+
+Ok, here comes a shift in thinking. Instead of defining combinators as equivalent to certain lambda terms,
+we can define combinators by what they do. If we have the I combinator followed by any expression X,
+I will take that expression as its argument and return that same expression as the result. In pictures,
+
+ IX ~~> X
+
+Thinking of this as a reduction rule, we can perform the following computation
+
+ II(IX) ~~> IIX ~~> IX ~~> X
+
+The reduction rule for K is also straightforward:
+
+ KXY ~~> X
![Szabolcsi's analysis of *himself* as the duplicator combinator](szabolcsireflexive.jpg)
+That is, K throws away its second argument. The reduction rule for S can be constructed by examining
+the defining lambda term:
+ S = \fgx.fx(gx)
These systems are Turing complete. In other words: every computation we know how to describe can be represented in a logical system consisting of only a single primitive operation!
+S takes three arguments, duplicates the third argument, and feeds one copy to the first argument and the second copy to the second argument. So:
+
+ SFGX ~~> FX(GX)
+
+If the meaning of a function is nothing more than how it behaves with respect to its arguments,
+these reduction rules capture the behavior of the combinators S, K, and I completely.
+We can use these rules to compute without resorting to beta reduction. For instance, we can show how the I combinator is equivalent to a certain crafty combination of Ss and Ks:
+
+ SKKX ~~> KX(KX) ~~> X
+
+So the combinator `SKK` is equivalent to the combinator I.
+
+Combinatory Logic is what you have when you choose a set of combinators and regulate their behavior with a set of reduction rules. The most common system uses S, K, and I as defined here.
+
+###The equivalence of the untyped lambda calculus and combinatory logic###
+
+We've claimed that Combinatory Logic is equivalent to the lambda calculus. If that's so, then S, K, and I must be enough to accomplish any computational task imaginable. Actually, S and K must suffice, since we've just seen that we can simulate I using only S and K. In order to get an intuition about what it takes to be Turing complete, imagine what a text editor does:
+it transforms any arbitrary text into any other arbitrary text. The way it does this is by deleting, copying, and reordering characters. We've already seen that K deletes its second argument, so we have deletion covered. S duplicates and reorders, so we have some reason to hope that S and K are enough to define arbitrary functions.
+
+We've already established that the behavior of combinatory terms can be perfectly mimicked by lambda terms: just replace each combinator with its equivalent lambda term, i.e., replace I with `\x.x`, replace K with `\fxy.x`, and replace S with `\fgx.fx(gx)`. How about the other direction? Here is a method for converting an arbitrary lambda term into an equivalent Combinatory Logic term using only S, K, and I. Besides the intrinsic beauty of this mapping, and the importance of what it says about the nature of binding and computation, it is possible to hear an echo of computing with continuations in this conversion strategy (though you wouldn't be able to hear these echos until we've covered a considerable portion of the rest of the course).
+
+Assume that for any lambda term T, [T] is the equivalent combinatory logic term. The we can define the [.] mapping as follows:
+
+ 1. [a] a
+ 2. [(M N)] ([M][N])
+ 3. [\a.a] I
+ 4. [\a.M] KM assumption: a does not occur free in M
+ 5. [\a.(M N)] S[\a.M][\a.N]
+ 6. [\a\b.M] [\a[\b.M]]
+
+It's easy to understand these rules based on what S, K and I do. The first rule says
+that variables are mapped to themselves.
+The second rule says that the way to translate an application is to translate the
+first element and the second element separately.
+The third rule should be obvious.
+The fourth rule should also be fairly selfevident: since what a lambda term such as `\x.y` does it throw away its first argument and return `y`, that's exactly what the combinatory logic translation should do. And indeed, `Ky` is a function that throws away its argument and returns `y`.
+The fifth rule deals with an abstract whose body is an application: the S combinator takes its next argument (which will fill the role of the original variable a) and copies it, feeding one copy to the translation of \a.M, and the other copy to the translation of \a.N. This ensures that any free occurrences of a inside M or N will end up taking on the appropriate value. Finally, the last rule says that if the body of an abstract is itself an abstract, translate the inner abstract first, and then do the outermost. (Since the translation of [\b.M] will not have any lambdas in it, we can be sure that we won't end up applying rule 6 again in an infinite loop.)
+
+[Fussy notes: if the original lambda term has free variables in it, so will the combinatory logic translation. Feel free to worry about this, though you should be confident that it makes sense. You should also convince yourself that if the original lambda term contains no free variablesi.e., is a combinatorthen the translation will consist only of S, K, and I (plus parentheses). One other detail: this translation algorithm builds expressions that combine lambdas with combinators. For instance, the translation of our boolean false `\x.\y.y` is `[\x[\y.y]] = [\x.I] = KI`. In the intermediate stage, we have `\x.I`, which mixes combinators in the body of a lambda abstract. It's possible to avoid this if you want to, but it takes some careful thought. See, e.g., Barendregt 1984, page 156.]
+
+Let's check that the translation of the false boolean behaves as expected by feeding it two arbitrary arguments:
+
+ KIXY ~~> IY ~~> Y
+
+Throws away the first argument, returns the second argumentyep, it works.
+
+Here's a more elaborate example of the translation. The goal is to establish that combinators can reverse order, so we use the T combinator, where `T = \x\y.yx`:
+
+ [\x\y.yx] = [\x[\y.yx]] = [\x.S[\y.y][\y.x]] = [\x.(SI)(Kx)] = S[\x.SI][\x.Kx] = S(K(SI))(S[\x.K][\x.x]) = S(K(SI))(S(KK)I)
+
+We can test this translation by seeing if it behaves like the original lambda term does.
+The orginal lambda term lifts its first argument (think of it as reversing the order of its two arguments):
+
+ S(K(SI))(S(KK)I) X Y =
+ (K(SI))X ((S(KK)I) X) Y =
+ SI ((KK)X (IX)) Y =
+ SI (KX) Y =
+ IY (KX)Y =
+ Y X
+
+Viola: the combinator takes any X and Y as arguments, and returns Y applied to X.
+
+One very nice property of combinatory logic is that there is no need to worry about alphabetic variance, or
+variable collisionsince there are no (bound) variables, there is no possibility of accidental variable capture,
+and so reduction can be performed without any fear of variable collision. We haven't mentioned the intricacies of
+alpha equivalence or safe variable substitution, but they are in fact quite intricate. (The best way to gain
+an appreciation of that intricacy is to write a program that performs lambda reduction.)
+
+Back to linguistic applications: one consequence of the equivalence between the lambda calculus and combinatory
+logic is that anything that can be done by binding variables can just as well be done with combinators.
+This has given rise to a style of semantic analysis called Variable Free Semantics (in addition to
+Szabolcsi's papers, see, for instance,
+Pauline Jacobson's 1999 *Linguistics and Philosophy* paper, `Towards a variablefree Semantics').
+Somewhat ironically, reading strings of combinators is so difficult that most practitioners of variablefree semantics
+express there meanings using the lambdacalculus rather than combinatory logic; perhaps they should call their
+enterprise Free Variable Free Semantics.
+
+A philosophical application: Quine went through a phase in which he developed a variable free logic.
+
+ Quine, Willard. 1960. Variables explained away. {\it Proceedings of
+ the American Philosophical Society}. Volume 104: 343347. Also in
+ W.~V.~Quine. 1960. {\it Selected Logical Papers}. Random House: New
+ York. 227235.
+
+The reason this was important to Quine is similar to the worries that Jim was talking about
+in the first class in which using nonreferring expressions such as Santa Clause might commit
+one to believing in nonexistant things. Quine's slogan was that `to be is to be the value of a variable'.
+What this was supposed to mean is that if and only if an object could serve as the value of some variable, we
+are committed to recognizing the existence of that object in our ontology.
+Obviously, if there ARE no variables, this slogan has to be rethought.
+
+Quine did not appear to appreciate that Shoenfinkel had already invented combinatory logic, though
+he later wrote an introduction to Shoenfinkel's key paper reprinted in Jean
+van Heijenoort (ed) 1967 *From Frege to Goedel,
+ a source book in mathematical logic, 18791931*.
+Cresswell has also developed a variablefree approach of some philosophical and linguistic interest
+in two books in the 1990's.
+
+A final linguistic application: Steedman's Combinatory Categorial Grammar, where the "Combinatory" is
+from combinatory logic (see especially his 2000 book, *The Syntactic Process*). Steedman attempts to build
+a syntax/semantics interface using a small number of combinators, including T = \xy.yx, B = \fxy.f(xy),
+and our friend S. Steedman used Smullyan's fanciful bird
+names for the combinators, Thrush, Bluebird, and Starling.
+
+Many of these combinatory logics, in particular, the SKI system,
+are Turing complete. In other words: every computation we know how to describe can be represented in a logical system consisting of only a single primitive operation!
Here's more to read about combinatorial logic.
Surely the most entertaining exposition is Smullyan's [[!wikipedia To_Mock_a_Mockingbird]].
@@ 127,16 +321,29 @@ This question highlights that there are different choices to make about how eval
With regard to Q3, it should be intuitively clear that `\x. M x` and `M` will behave the same with respect to any arguments they are given. It can also be proven that no other functions can behave differently with respect to them. However, the logical system you get when etareduction is added to the proof theory is importantly different from the one where only betareduction is permitted.
MORE on extensionality
+If we answer Q2 by permitting reduction inside abstracts, and we also permit etareduction, then where none of y_{1}, ..., y_{n} occur free in M, this:
If we answer Q2 by permitting reduction inside abstracts, and we also permit etareduction, then where neither `y` nor `z` occur in M, this:
+\x y_{1}... y_{n}. M y_{1}... y_{n}
 \x y z. M y z

will etareduce by two steps to:
+will etareduce by n steps to:
\x. M
+The logical system you get when etareduction is added to the proof system has the following property:
+
+> if `M`, `N` are normal forms with no free variables, then M ≡ N
iff `M` and `N` behave the same with respect to every possible sequence of arguments.
+
+That is, when `M` and `N` are (closed normal forms that are) syntactically distinct, there will always be some sequences of arguments L_{1}, ..., L_{n}
such that:
+
+M L_{1} ... L_{n} x y ~~> x
+N L_{1} ... L_{n} x y ~~> y
+
+
+That is, closed normal forms that are not just betareduced but also fully etareduced, will be syntactically different iff they yield different values for some arguments. That is, iff their extensions differ.
+
+So the proof theory with etareduction added is called "extensional," because its notion of normal form makes syntactic identity of closed normal forms coincide with extensional equivalence.
+
+
The evaluation strategy which answers Q1 by saying "reduce arguments first" is known as **callbyvalue**. The evaluation strategy which answers Q1 by saying "substitute arguments in unreduced" is known as **callbyname** or **callbyneed** (the difference between these has to do with efficiency, not semantics).
When one has a callbyvalue strategy that also permits reduction to continue inside unapplied abstracts, that's known as "applicative order" reduction. When one has a callbyname strategy that permits reduction inside abstracts, that's known as "normal order" reduction. Consider an expression of the form: