#Typed lambda terms#
-Given a set of types `T`, we define the set of typed lambda terms <code>Λ~T~</code>,
+Given a set of types `T`, we define the set of typed lambda terms <code>Λ_T</code>,
which is the smallest set such that
* each type `t` has an infinite set of distinct variables, {x^t}_1,
σ, then the application `(M N)` has type τ.
* If a variable `a` has type σ, and term `M` has type τ,
- then the abstract `λ a M` has type `σ --> τ`.
+ then the abstract <code>λ a M</code> has type σ --> τ.
The definitions of types and of typed terms should be highly familiar
-to semanticists, except that instead of writing `σ --> τ`,
-linguists (following Montague, who followed Church) write `<σ,
-τ>`. We will use the arrow notation, since it is more iconic.
+to semanticists, except that instead of writing σ --> τ,
+linguists (following Montague, who followed Church) write <σ,
+τ>. We will use the arrow notation, since it is more iconic.
Some examples (assume that `x` has type `o`):
have types `a`, `b`, `c`, and `d`, respectively, then `f` has type `a
--> b --> c --> d == (a --> (b --> (c --> d)))`.
-It is a serious faux pas to associate to the left for types, on a par
-with using your salad fork to stir your tea.
+It is a serious faux pas to associate to the left for types. You may
+as well use your salad fork to stir your tea.
#The simply-typed lambda calculus is strongly normalizing#
-If `M` is a term with type τ in `Λ_T`, then `M` has a
+If `M` is a term with type τ in Λ_T, then `M` has a
normal form. The proof is not particularly complex, but we will not
present it here; see Berendregt or Hankin.
Since Ω does not have a normal form, it follows that Ω
-cannot have a type in `Λ_T`. We can easily see why:
+cannot have a type in Λ_T. We can easily see why:
Ω = (\x.xx)(\x.xx)
Assume Ω has type τ, and `\x.xx` has type σ. Then
because `\x.xx` takes an argument of type σ and returns
-something of type τ, `\x.xx` must also have type `σ -->
-τ`. By repeating this reasoning, `\x.xx` must also have type
-`(σ --> τ) --> τ`; and so on. Since variables have
+something of type τ, `\x.xx` must also have type σ -->
+τ. By repeating this reasoning, `\x.xx` must also have type
+(σ --> τ) --> τ; and so on. Since variables have
finite types, there is no way to choose a type for the variable `x`
that can satisfy all of the requirements imposed on it.
Version 1 type numerals are not a good choice for the simply-typed
lambda calculus. The reason is that each different numberal has a
-different type! For instance, if zero has type σ, and `false`
-has type `τ --> τ --> τ` for some τ, and one is
-represented by the function `\x.x false 0`, then one must have type
-`(τ --> τ --> &tau) --> &sigma --> σ`. But this is a
-different type than zero! Because numbers have different types, it
-becomes impossible to write arithmetic operations that can combine
-zero with one. We would need as many different addition operations as
-we had pairs of numbers that we wanted to add.
+different type! For instance, if zero has type σ, then `false`
+has type τ --> τ --> τ, for some τ. Since one is
+represented by the function `\x.x false 0`, one must have type (τ
+--> τ --> τ) --> σ --> σ. But this is a different
+type than zero! Because each number has a different type, it becomes
+impossible to write arithmetic operations that can combine zero with
+one. We would need as many different addition operations as we had
+pairs of numbers that we wanted to add.
Fortunately, the Church numberals are well behaved with respect to
-types. They can all be given the type `(σ --> σ) -->
-σ --> σ`.
+types. They can all be given the type (σ --> σ) -->
+σ --> σ.
+
+
+
+
+<!--
+
+Mau integrate some mention of this at some point.
+
+http://okmij.org/ftp/Computation/lambda-calc.html#predecessor
+
+
+Predecessor and lists are not representable in simply typed lambda-calculus
+
+ The predecessor of a Church-encoded numeral, or, generally, the encoding of a list with the car and cdr operations are both impossible in the simply typed lambda-calculus. Henk Barendregt's ``The impact of the lambda-calculus in logic and computer science'' (The Bulletin of Symbolic Logic, v3, N2, June 1997) has the following phrase, on p. 186:
+
+ Even for a function as simple as the predecessor lambda definability remained an open problem for a while. From our present knowledge it is tempting to explain this as follows. Although the lambda calculus was conceived as an untyped theory, typeable terms are more intuitive. Now the functions addition and multiplication are defineable by typeable terms, while [101] and [108] have characterized the lambda-defineable functions in the (simply) typed lambda calculus and the predecessor is not among them [the story of the removal of Kleene's four wisdom teeth is skipped...]
+ Ref 108 is R.Statman: The typed lambda calculus is not elementary recursive. Theoretical Comp. Sci., vol 9 (1979), pp. 73-81.
+
+ Since list is a generalization of numeral -- with cons being a successor, append being the addition, tail (aka cdr) being the predecessor -- it follows then the list cannot be encoded in the simply typed lambda-calculus.
+
+ To encode both operations, we need either inductive (generally, recursive) types, or System F with its polymorphism. The first approach is the most common. Indeed, the familiar definition of a list
+
+ data List a = Nil | Cons a (List a)
+
+ gives an (iso-) recursive data type (in Haskell. In ML, it is an inductive data type).
+
+ Lists can also be represented in System F. As a matter of fact, we do not need the full System F (where the type reconstruction is not decidable). We merely need the extension of the Hindley-Milner system with higher-ranked types, which requires a modicum of type annotations and yet is able to infer the types of all other terms. This extension is supported in Haskell and OCaml. With such an extension, we can represent a list by its fold, as shown in the code below. It is less known that this representation is faithful: we can implement all list operations, including tail, drop, and even zip.
+
+
+-->