%% $Id$
\chapter{Constructive Type Theory}
Martin-L\"of's Constructive Type Theory \cite{martinlof84,nordstrom90} can
be viewed at many different levels. It is a formal system that embodies
the principles of intuitionistic mathematics; it embodies the
interpretation of propositions as types; it is a vehicle for deriving
programs from proofs. The logic is complex and many authors have attempted
to simplify it. Thompson~\cite{thompson91} is a readable and thorough
account of the theory.
Isabelle's original formulation of Type Theory was a kind of sequent
calculus, following Martin-L\"of~\cite{martinlof84}. It included rules for
building the context, namely variable bindings with their types. A typical
judgement was
\[ a(x@1,\ldots,x@n)\in A(x@1,\ldots,x@n) \;
[ x@1\in A@1, x@2\in A@2(x@1), \ldots, x@n\in A@n(x@1,\ldots,x@{n-1}) ]
\]
This sequent calculus was not satisfactory because assumptions like
`suppose $A$ is a type' or `suppose $B(x)$ is a type for all $x$ in $A$'
could not be formalized.
The directory~\ttindexbold{CTT} implements Constructive Type Theory, using
natural deduction. The judgement above is expressed using $\Forall$ and
$\Imp$:
\[ \begin{array}{r@{}l}
\Forall x@1\ldots x@n. &
\List{x@1\in A@1;
x@2\in A@2(x@1); \cdots \;
x@n\in A@n(x@1,\ldots,x@{n-1})} \Imp \\
& \qquad\qquad a(x@1,\ldots,x@n)\in A(x@1,\ldots,x@n)
\end{array}
\]
Assumptions can use all the judgement forms, for instance to express that
$B$ is a family of types over~$A$:
\[ \Forall x . x\in A \Imp B(x)\;{\rm type} \]
To justify the {\CTT} formulation it is probably best to appeal directly
to the semantic explanations of the rules~\cite{martinlof84}, rather than
to the rules themselves. The order of assumptions no longer matters,
unlike in standard Type Theory. Contexts, which are typical of many modern
type theories, are difficult to represent in Isabelle. In particular, it
is difficult to enforce that all the variables in a context are distinct.
The theory has the {\ML} identifier \ttindexbold{CTT.thy}. It does not
use polymorphism. Terms in {\CTT} have type~$i$, the type of individuals.
Types in {\CTT} have type~$t$.
{\CTT} supports all of Type Theory apart from list types, well ordering
types, and universes. Universes could be introduced {\em\`a la Tarski},
adding new constants as names for types. The formulation {\em\`a la
Russell}, where types denote themselves, is only possible if we identify
the meta-types~$i$ and~$o$. Most published formulations of well ordering
types have difficulties involving extensionality of functions; I suggest
that you use some other method for defining recursive types. List types
are easy to introduce by declaring new rules.
{\CTT} uses the 1982 version of Type Theory, with extensional equality.
The computation $a=b\in A$ and the equality $c\in Eq(A,a,b)$ are
interchangeable. Its rewriting tactics prove theorems of the form $a=b\in
A$. It could be modified to have intensional equality, but rewriting
tactics would have to prove theorems of the form $c\in Eq(A,a,b)$ and the
computation rules might require a second simplifier.
\begin{figure} \tabcolsep=1em %wider spacing in tables
\begin{center}
\begin{tabular}{rrr}
\it name & \it meta-type & \it description \\
\idx{Type} & $t \to prop$ & judgement form \\
\idx{Eqtype} & $[t,t]\to prop$ & judgement form\\
\idx{Elem} & $[i, t]\to prop$ & judgement form\\
\idx{Eqelem} & $[i, i, t]\to prop$ & judgement form\\
\idx{Reduce} & $[i, i]\to prop$ & extra judgement form\\[2ex]
\idx{N} & $t$ & natural numbers type\\
\idx{0} & $i$ & constructor\\
\idx{succ} & $i\to i$ & constructor\\
\idx{rec} & $[i,i,[i,i]\to i]\to i$ & eliminator\\[2ex]
\idx{Prod} & $[t,i\to t]\to t$ & general product type\\
\idx{lambda} & $(i\to i)\to i$ & constructor\\[2ex]
\idx{Sum} & $[t, i\to t]\to t$ & general sum type\\
\idx{pair} & $[i,i]\to i$ & constructor\\
\idx{split} & $[i,[i,i]\to i]\to i$ & eliminator\\
\idx{fst} \idx{snd} & $i\to i$ & projections\\[2ex]
\idx{inl} \idx{inr} & $i\to i$ & constructors for $+$\\
\idx{when} & $[i,i\to i, i\to i]\to i$ & eliminator for $+$\\[2ex]
\idx{Eq} & $[t,i,i]\to t$ & equality type\\
\idx{eq} & $i$ & constructor\\[2ex]
\idx{F} & $t$ & empty type\\
\idx{contr} & $i\to i$ & eliminator\\[2ex]
\idx{T} & $t$ & singleton type\\
\idx{tt} & $i$ & constructor
\end{tabular}
\end{center}
\caption{The constants of {\CTT}} \label{ctt-constants}
\end{figure}
\begin{figure} \tabcolsep=1em %wider spacing in tables
\begin{center}
\begin{tabular}{llrrr}
\it symbol &\it name &\it meta-type & \it precedence & \it description \\
\idx{lam} & \idx{lambda} & $(i\To o)\To i$ & 10 & $\lambda$-abstraction
\end{tabular}
\end{center}
\subcaption{Binders}
\begin{center}
\indexbold{*"`}
\indexbold{*"+}
\begin{tabular}{rrrr}
\it symbol & \it meta-type & \it precedence & \it description \\
\tt ` & $[i,i]\to i$ & Left 55 & function application\\
\tt + & $[t,t]\to t$ & Right 30 & sum of two types
\end{tabular}
\end{center}
\subcaption{Infixes}
\indexbold{*"*}
\indexbold{*"-"-">}
\begin{center} \tt\frenchspacing
\begin{tabular}{rrr}
\it external & \it internal & \it standard notation \\
\idx{PROD} $x$:$A$ . $B[x]$ & Prod($A$, $\lambda x.B[x]$) &
\rm product $\prod@{x\in A}B[x]$ \\
\idx{SUM} $x$:$A$ . $B[x]$ & Sum($A$, $\lambda x.B[x]$) &
\rm sum $\sum@{x\in A}B[x]$ \\
$A$ --> $B$ & Prod($A$, $\lambda x.B$) &
\rm function space $A\to B$ \\
$A$ * $B$ & Sum($A$, $\lambda x.B$) &
\rm binary product $A\times B$
\end{tabular}
\end{center}
\subcaption{Translations}
\indexbold{*"=}
\begin{center}
\dquotes
\[ \begin{array}{rcl}
prop & = & type " type" \\
& | & type " = " type \\
& | & term " : " type \\
& | & term " = " term " : " type
\\[2ex]
type & = & \hbox{expression of type~$t$} \\
& | & "PROD~" id " : " type " . " type \\
& | & "SUM~~" id " : " type " . " type
\\[2ex]
term & = & \hbox{expression of type~$i$} \\
& | & "lam " id~id^* " . " term \\
& | & "< " term " , " term " >"
\end{array}
\]
\end{center}
\subcaption{Grammar}
\caption{Syntax of {\CTT}} \label{ctt-syntax}
\end{figure}
%%%%\section{Generic Packages} typedsimp.ML????????????????
\section{Syntax}
The constants are shown in Fig.\ts\ref{ctt-constants}. The infixes include
the function application operator (sometimes called `apply'), and the
2-place type operators. Note that meta-level abstraction and application,
$\lambda x.b$ and $f(a)$, differ from object-level abstraction and
application, \hbox{\tt lam $x$.$b$} and $b{\tt`}a$. A {\CTT}
function~$f$ is simply an individual as far as Isabelle is concerned: its
Isabelle type is~$i$, not say $i\To i$.
\indexbold{*F}\indexbold{*T}\indexbold{*SUM}\indexbold{*PROD}
The empty type is called $F$ and the one-element type is $T$; other finite
sets are built as $T+T+T$, etc. The notation for~{\CTT}
(Fig.\ts\ref{ctt-syntax}) is based on that of Nordstr\"om et
al.~\cite{nordstrom90}. We can write
\begin{ttbox}
SUM y:B. PROD x:A. C(x,y) {\rm for} Sum(B, \%y. Prod(A, \%x. C(x,y)))
\end{ttbox}
The special cases as \hbox{\tt$A$*$B$} and \hbox{\tt$A$-->$B$} abbreviate
general sums and products over a constant family.\footnote{Unlike normal
infix operators, {\tt*} and {\tt-->} merely define abbreviations; there are
no constants~{\tt op~*} and~\hbox{\tt op~-->}.} Isabelle accepts these
abbreviations in parsing and uses them whenever possible for printing.
\begin{figure}
\begin{ttbox}
\idx{refl_type} A type ==> A = A
\idx{refl_elem} a : A ==> a = a : A
\idx{sym_type} A = B ==> B = A
\idx{sym_elem} a = b : A ==> b = a : A
\idx{trans_type} [| A = B; B = C |] ==> A = C
\idx{trans_elem} [| a = b : A; b = c : A |] ==> a = c : A
\idx{equal_types} [| a : A; A = B |] ==> a : B
\idx{equal_typesL} [| a = b : A; A = B |] ==> a = b : B
\idx{subst_type} [| a : A; !!z. z:A ==> B(z) type |] ==> B(a) type
\idx{subst_typeL} [| a = c : A; !!z. z:A ==> B(z) = D(z)
|] ==> B(a) = D(c)
\idx{subst_elem} [| a : A; !!z. z:A ==> b(z):B(z) |] ==> b(a):B(a)
\idx{subst_elemL} [| a = c : A; !!z. z:A ==> b(z) = d(z) : B(z)
|] ==> b(a) = d(c) : B(a)
\idx{refl_red} Reduce(a,a)
\idx{red_if_equal} a = b : A ==> Reduce(a,b)
\idx{trans_red} [| a = b : A; Reduce(b,c) |] ==> a = c : A
\end{ttbox}
\caption{General equality rules} \label{ctt-equality}
\end{figure}
\begin{figure}
\begin{ttbox}
\idx{NF} N type
\idx{NI0} 0 : N
\idx{NI_succ} a : N ==> succ(a) : N
\idx{NI_succL} a = b : N ==> succ(a) = succ(b) : N
\idx{NE} [| p: N; a: C(0);
!!u v. [| u: N; v: C(u) |] ==> b(u,v): C(succ(u))
|] ==> rec(p, a, \%u v.b(u,v)) : C(p)
\idx{NEL} [| p = q : N; a = c : C(0);
!!u v. [| u: N; v: C(u) |] ==> b(u,v)=d(u,v): C(succ(u))
|] ==> rec(p, a, \%u v.b(u,v)) = rec(q,c,d) : C(p)
\idx{NC0} [| a: C(0);
!!u v. [| u: N; v: C(u) |] ==> b(u,v): C(succ(u))
|] ==> rec(0, a, \%u v.b(u,v)) = a : C(0)
\idx{NC_succ} [| p: N; a: C(0);
!!u v. [| u: N; v: C(u) |] ==> b(u,v): C(succ(u))
|] ==> rec(succ(p), a, \%u v.b(u,v)) =
b(p, rec(p, a, \%u v.b(u,v))) : C(succ(p))
\idx{zero_ne_succ} [| a: N; 0 = succ(a) : N |] ==> 0: F
\end{ttbox}
\caption{Rules for type~$N$} \label{ctt-N}
\end{figure}
\begin{figure}
\begin{ttbox}
\idx{ProdF} [| A type; !!x. x:A ==> B(x) type |] ==> PROD x:A.B(x) type
\idx{ProdFL} [| A = C; !!x. x:A ==> B(x) = D(x) |] ==>
PROD x:A.B(x) = PROD x:C.D(x)
\idx{ProdI} [| A type; !!x. x:A ==> b(x):B(x)
|] ==> lam x.b(x) : PROD x:A.B(x)
\idx{ProdIL} [| A type; !!x. x:A ==> b(x) = c(x) : B(x)
|] ==> lam x.b(x) = lam x.c(x) : PROD x:A.B(x)
\idx{ProdE} [| p : PROD x:A.B(x); a : A |] ==> p`a : B(a)
\idx{ProdEL} [| p=q: PROD x:A.B(x); a=b : A |] ==> p`a = q`b : B(a)
\idx{ProdC} [| a : A; !!x. x:A ==> b(x) : B(x)
|] ==> (lam x.b(x)) ` a = b(a) : B(a)
\idx{ProdC2} p : PROD x:A.B(x) ==> (lam x. p`x) = p : PROD x:A.B(x)
\end{ttbox}
\caption{Rules for the product type $\prod@{x\in A}B[x]$} \label{ctt-prod}
\end{figure}
\begin{figure}
\begin{ttbox}
\idx{SumF} [| A type; !!x. x:A ==> B(x) type |] ==> SUM x:A.B(x) type
\idx{SumFL} [| A = C; !!x. x:A ==> B(x) = D(x)
|] ==> SUM x:A.B(x) = SUM x:C.D(x)
\idx{SumI} [| a : A; b : B(a) |] ==> <a,b> : SUM x:A.B(x)
\idx{SumIL} [| a=c:A; b=d:B(a) |] ==> <a,b> = <c,d> : SUM x:A.B(x)
\idx{SumE} [| p: SUM x:A.B(x);
!!x y. [| x:A; y:B(x) |] ==> c(x,y): C(<x,y>)
|] ==> split(p, \%x y.c(x,y)) : C(p)
\idx{SumEL} [| p=q : SUM x:A.B(x);
!!x y. [| x:A; y:B(x) |] ==> c(x,y)=d(x,y): C(<x,y>)
|] ==> split(p, \%x y.c(x,y)) = split(q, \%x y.d(x,y)) : C(p)
\idx{SumC} [| a: A; b: B(a);
!!x y. [| x:A; y:B(x) |] ==> c(x,y): C(<x,y>)
|] ==> split(<a,b>, \%x y.c(x,y)) = c(a,b) : C(<a,b>)
\idx{fst_def} fst(a) == split(a, \%x y.x)
\idx{snd_def} snd(a) == split(a, \%x y.y)
\end{ttbox}
\caption{Rules for the sum type $\sum@{x\in A}B[x]$} \label{ctt-sum}
\end{figure}
\begin{figure}
\begin{ttbox}
\idx{PlusF} [| A type; B type |] ==> A+B type
\idx{PlusFL} [| A = C; B = D |] ==> A+B = C+D
\idx{PlusI_inl} [| a : A; B type |] ==> inl(a) : A+B
\idx{PlusI_inlL} [| a = c : A; B type |] ==> inl(a) = inl(c) : A+B
\idx{PlusI_inr} [| A type; b : B |] ==> inr(b) : A+B
\idx{PlusI_inrL} [| A type; b = d : B |] ==> inr(b) = inr(d) : A+B
\idx{PlusE} [| p: A+B;
!!x. x:A ==> c(x): C(inl(x));
!!y. y:B ==> d(y): C(inr(y))
|] ==> when(p, \%x.c(x), \%y.d(y)) : C(p)
\idx{PlusEL} [| p = q : A+B;
!!x. x: A ==> c(x) = e(x) : C(inl(x));
!!y. y: B ==> d(y) = f(y) : C(inr(y))
|] ==> when(p, \%x.c(x), \%y.d(y)) =
when(q, \%x.e(x), \%y.f(y)) : C(p)
\idx{PlusC_inl} [| a: A;
!!x. x:A ==> c(x): C(inl(x));
!!y. y:B ==> d(y): C(inr(y))
|] ==> when(inl(a), \%x.c(x), \%y.d(y)) = c(a) : C(inl(a))
\idx{PlusC_inr} [| b: B;
!!x. x:A ==> c(x): C(inl(x));
!!y. y:B ==> d(y): C(inr(y))
|] ==> when(inr(b), \%x.c(x), \%y.d(y)) = d(b) : C(inr(b))
\end{ttbox}
\caption{Rules for the binary sum type $A+B$} \label{ctt-plus}
\end{figure}
\begin{figure}
\begin{ttbox}
\idx{EqF} [| A type; a : A; b : A |] ==> Eq(A,a,b) type
\idx{EqFL} [| A=B; a=c: A; b=d : A |] ==> Eq(A,a,b) = Eq(B,c,d)
\idx{EqI} a = b : A ==> eq : Eq(A,a,b)
\idx{EqE} p : Eq(A,a,b) ==> a = b : A
\idx{EqC} p : Eq(A,a,b) ==> p = eq : Eq(A,a,b)
\end{ttbox}
\subcaption{The equality type $Eq(A,a,b)$}
\begin{ttbox}
\idx{FF} F type
\idx{FE} [| p: F; C type |] ==> contr(p) : C
\idx{FEL} [| p = q : F; C type |] ==> contr(p) = contr(q) : C
\end{ttbox}
\subcaption{The empty type $F$}
\begin{ttbox}
\idx{TF} T type
\idx{TI} tt : T
\idx{TE} [| p : T; c : C(tt) |] ==> c : C(p)
\idx{TEL} [| p = q : T; c = d : C(tt) |] ==> c = d : C(p)
\idx{TC} p : T ==> p = tt : T)
\end{ttbox}
\subcaption{The unit type $T$}
\caption{Rules for other {\CTT} types} \label{ctt-others}
\end{figure}
\begin{figure}
\begin{ttbox}
\idx{replace_type} [| B = A; a : A |] ==> a : B
\idx{subst_eqtyparg} [| a=c : A; !!z. z:A ==> B(z) type |] ==> B(a)=B(c)
\idx{subst_prodE} [| p: Prod(A,B); a: A; !!z. z: B(a) ==> c(z): C(z)
|] ==> c(p`a): C(p`a)
\idx{SumIL2} [| c=a : A; d=b : B(a) |] ==> <c,d> = <a,b> : Sum(A,B)
\idx{SumE_fst} p : Sum(A,B) ==> fst(p) : A
\idx{SumE_snd} [| p: Sum(A,B); A type; !!x. x:A ==> B(x) type
|] ==> snd(p) : B(fst(p))
\end{ttbox}
\caption{Derived rules for {\CTT}} \label{ctt-derived}
\end{figure}
\section{Rules of inference}
The rules obey the following naming conventions. Type formation rules have
the suffix~{\tt F}\@. Introduction rules have the suffix~{\tt I}\@.
Elimination rules have the suffix~{\tt E}\@. Computation rules, which
describe the reduction of eliminators, have the suffix~{\tt C}\@. The
equality versions of the rules (which permit reductions on subterms) are
called {\em long} rules; their names have the suffix~{\tt L}\@.
Introduction and computation rules often are further suffixed with
constructor names.
Figures~\ref{ctt-equality}--\ref{ctt-others} shows the rules. Those
for~$N$ include \ttindex{zero_ne_succ}, $0\not=n+1$: the fourth Peano axiom
cannot be derived without universes \cite[page 91]{martinlof84}.
Figure~\ref{ctt-sum} shows the rules for general sums, which include binary
products as a special case, with the projections \ttindex{fst}
and~\ttindex{snd}.
The extra judgement \ttindex{Reduce} is used to implement rewriting. The
judgement ${\tt Reduce}(a,b)$ holds when $a=b:A$ holds. It also holds
when $a$ and $b$ are syntactically identical, even if they are ill-typed,
because rule \ttindex{refl_red} does not verify that $a$ belongs to $A$. These
rules do not give rise to new theorems about the standard judgements ---
note that the only rule that makes use of {\tt Reduce} is \ttindex{trans_red},
whose first premise ensures that $a$ and $b$ (and thus $c$) are well-typed.
Derived rules are shown in Fig.\ts\ref{ctt-derived}. The rule
\ttindex{subst_prodE} is derived from \ttindex{prodE}, and is easier to
use in backwards proof. The rules \ttindex{SumE_fst} and
\ttindex{SumE_snd} express the typing of~\ttindex{fst} and~\ttindex{snd};
together, they are roughly equivalent to~\ttindex{SumE} with the advantage
of creating no parameters. These rules are demonstrated in a proof of the
Axiom of Choice~(\S\ref{ctt-choice}).
All the rules are given in $\eta$-expanded form. For instance, every
occurrence of $\lambda u\,v.b(u,v)$ could be abbreviated to~$b$ in the
rules for~$N$. This permits Isabelle to preserve bound variable names
during backward proof. Names of bound variables in the conclusion (here,
$u$ and~$v$) are matched with corresponding bound variables in the premises.
\section{Rule lists}
The Type Theory tactics provide rewriting, type inference, and logical
reasoning. Many proof procedures work by repeatedly resolving certain Type
Theory rules against a proof state. {\CTT} defines lists --- each with
type
\hbox{\tt thm list} --- of related rules.
\begin{description}
\item[\ttindexbold{form_rls}]
contains formation rules for the types $N$, $\Pi$, $\Sigma$, $+$, $Eq$,
$F$, and $T$.
\item[\ttindexbold{formL_rls}]
contains long formation rules for $\Pi$, $\Sigma$, $+$, and $Eq$. (For
other types use \ttindex{refl_type}.)
\item[\ttindexbold{intr_rls}]
contains introduction rules for the types $N$, $\Pi$, $\Sigma$, $+$, and
$T$.
\item[\ttindexbold{intrL_rls}]
contains long introduction rules for $N$, $\Pi$, $\Sigma$, and $+$. (For
$T$ use \ttindex{refl_elem}.)
\item[\ttindexbold{elim_rls}]
contains elimination rules for the types $N$, $\Pi$, $\Sigma$, $+$, and
$F$. The rules for $Eq$ and $T$ are omitted because they involve no
eliminator.
\item[\ttindexbold{elimL_rls}]
contains long elimination rules for $N$, $\Pi$, $\Sigma$, $+$, and $F$.
\item[\ttindexbold{comp_rls}]
contains computation rules for the types $N$, $\Pi$, $\Sigma$, and $+$.
Those for $Eq$ and $T$ involve no eliminator.
\item[\ttindexbold{basic_defs}]
contains the definitions of \ttindex{fst} and \ttindex{snd}.
\end{description}
\section{Tactics for subgoal reordering}
\begin{ttbox}
test_assume_tac : int -> tactic
typechk_tac : thm list -> tactic
equal_tac : thm list -> tactic
intr_tac : thm list -> tactic
\end{ttbox}
Blind application of {\CTT} rules seldom leads to a proof. The elimination
rules, especially, create subgoals containing new unknowns. These subgoals
unify with anything, causing an undirectional search. The standard tactic
\ttindex{filt_resolve_tac} (see the {\em Reference Manual}) can reject
overly flexible goals; so does the {\CTT} tactic {\tt test_assume_tac}.
Used with the tactical \ttindex{REPEAT_FIRST} they achieve a simple kind of
subgoal reordering: the less flexible subgoals are attempted first. Do
some single step proofs, or study the examples below, to see why this is
necessary.
\begin{description}
\item[\ttindexbold{test_assume_tac} $i$]
uses \ttindex{assume_tac} to solve the subgoal by assumption, but only if
subgoal~$i$ has the form $a\in A$ and the head of $a$ is not an unknown.
Otherwise, it fails.
\item[\ttindexbold{typechk_tac} $thms$]
uses $thms$ with formation, introduction, and elimination rules to check
the typing of constructions. It is designed to solve goals of the form
$a\in \Var{A}$, where $a$ is rigid and $\Var{A}$ is flexible. Thus it
performs Hindley-Milner type inference. The tactic can also solve goals of
the form $A\;\rm type$.
\item[\ttindexbold{equal_tac} $thms$]
uses $thms$ with the long introduction and elimination rules to solve goals
of the form $a=b\in A$, where $a$ is rigid. It is intended for deriving
the long rules for defined constants such as the arithmetic operators. The
tactic can also perform type checking.
\item[\ttindexbold{intr_tac} $thms$]
uses $thms$ with the introduction rules to break down a type. It is
designed for goals like $\Var{a}\in A$ where $\Var{a}$ is flexible and $A$
rigid. These typically arise when trying to prove a proposition~$A$,
expressed as a type.
\end{description}
\section{Rewriting tactics}
\begin{ttbox}
rew_tac : thm list -> tactic
hyp_rew_tac : thm list -> tactic
\end{ttbox}
Object-level simplification is accomplished through proof, using the {\tt
CTT} equality rules and the built-in rewriting functor
\ttindex{TSimpFun}.\footnote{This should not be confused with {\tt
SimpFun}, which is the main rewriting functor; {\tt TSimpFun} is only
useful for {\CTT} and similar logics with type inference rules.}
The rewrites include the computation rules and other equations. The
long versions of the other rules permit rewriting of subterms and subtypes.
Also used are transitivity and the extra judgement form \ttindex{Reduce}.
Meta-level simplification handles only definitional equality.
\begin{description}
\item[\ttindexbold{rew_tac} $thms$]
applies $thms$ and the computation rules as left-to-right rewrites. It
solves the goal $a=b\in A$ by rewriting $a$ to $b$. If $b$ is an unknown
then it is assigned the rewritten form of~$a$. All subgoals are rewritten.
\item[\ttindexbold{hyp_rew_tac} $thms$]
is like {\tt rew_tac}, but includes as rewrites any equations present in
the assumptions.
\end{description}
\section{Tactics for logical reasoning}
Interpreting propositions as types lets {\CTT} express statements of
intuitionistic logic. However, Constructive Type Theory is not just
another syntax for first-order logic. A key question: can assumptions be
deleted after use? Not every occurrence of a type represents a
proposition, and Type Theory assumptions declare variables.
In first-order logic, $\disj$-elimination with the assumption $P\disj Q$
creates one subgoal assuming $P$ and another assuming $Q$, and $P\disj Q$
can be deleted. In Type Theory, $+$-elimination with the assumption $z\in
A+B$ creates one subgoal assuming $x\in A$ and another assuming $y\in B$
(for arbitrary $x$ and $y$). Deleting $z\in A+B$ may render the subgoals
unprovable if other assumptions refer to $z$. Some people might argue that
such subgoals are not even meaningful.
\begin{ttbox}
mp_tac : int -> tactic
add_mp_tac : int -> tactic
safestep_tac : thm list -> int -> tactic
safe_tac : thm list -> int -> tactic
step_tac : thm list -> int -> tactic
pc_tac : thm list -> int -> tactic
\end{ttbox}
These are loosely based on the intuitionistic proof procedures
of~\ttindex{FOL}. For the reasons discussed above, a rule that is safe for
propositional reasoning may be unsafe for type checking; thus, some of the
``safe'' tactics are misnamed.
\begin{description}
\item[\ttindexbold{mp_tac} $i$]
searches in subgoal~$i$ for assumptions of the form $f\in\Pi(A,B)$ and
$a\in A$, where~$A$ may be found by unification. It replaces
$f\in\Pi(A,B)$ by $z\in B(a)$, where~$z$ is a new parameter. The tactic
can produce multiple outcomes for each suitable pair of assumptions. In
short, {\tt mp_tac} performs Modus Ponens among the assumptions.
\item[\ttindexbold{add_mp_tac} $i$]
is like {\tt mp_tac}~$i$ but retains the assumption $f\in\Pi(A,B)$.
\item[\ttindexbold{safestep_tac} $thms$ $i$]
attacks subgoal~$i$ using formation rules and certain other `safe' rules
(\ttindex{FE}, \ttindex{ProdI}, \ttindex{SumE}, \ttindex{PlusE}), calling
{\tt mp_tac} when appropriate. It also uses~$thms$,
which are typically premises of the rule being derived.
\item[\ttindexbold{safe_tac} $thms$ $i$]
tries to solve subgoal~$i$ by backtracking, using {\tt safestep_tac}.
\item[\ttindexbold{step_tac} $thms$ $i$]
tries to reduce subgoal~$i$ using {\tt safestep_tac}, then tries unsafe
rules. It may produce multiple outcomes.
\item[\ttindexbold{pc_tac} $thms$ $i$]
tries to solve subgoal~$i$ by backtracking, using {\tt step_tac}.
\end{description}
\begin{figure}
\begin{ttbox}
\idx{add_def} a#+b == rec(a, b, \%u v.succ(v))
\idx{diff_def} a-b == rec(b, a, \%u v.rec(v, 0, \%x y.x))
\idx{absdiff_def} a|-|b == (a-b) #+ (b-a)
\idx{mult_def} a#*b == rec(a, 0, \%u v. b #+ v)
\idx{mod_def} a mod b == rec(a, 0,
\%u v. rec(succ(v) |-| b, 0, \%x y.succ(v)))
\idx{div_def} a div b == rec(a, 0,
\%u v. rec(succ(u) mod b, succ(v), \%x y.v))
\end{ttbox}
\subcaption{Definitions of the operators}
\begin{ttbox}
\idx{add_typing} [| a:N; b:N |] ==> a #+ b : N
\idx{addC0} b:N ==> 0 #+ b = b : N
\idx{addC_succ} [| a:N; b:N |] ==> succ(a) #+ b = succ(a #+ b) : N
\idx{add_assoc} [| a:N; b:N; c:N |] ==>
(a #+ b) #+ c = a #+ (b #+ c) : N
\idx{add_commute} [| a:N; b:N |] ==> a #+ b = b #+ a : N
\idx{mult_typing} [| a:N; b:N |] ==> a #* b : N
\idx{multC0} b:N ==> 0 #* b = 0 : N
\idx{multC_succ} [| a:N; b:N |] ==> succ(a) #* b = b #+ (a#*b) : N
\idx{mult_commute} [| a:N; b:N |] ==> a #* b = b #* a : N
\idx{add_mult_dist} [| a:N; b:N; c:N |] ==>
(a #+ b) #* c = (a #* c) #+ (b #* c) : N
\idx{mult_assoc} [| a:N; b:N; c:N |] ==>
(a #* b) #* c = a #* (b #* c) : N
\idx{diff_typing} [| a:N; b:N |] ==> a - b : N
\idx{diffC0} a:N ==> a - 0 = a : N
\idx{diff_0_eq_0} b:N ==> 0 - b = 0 : N
\idx{diff_succ_succ} [| a:N; b:N |] ==> succ(a) - succ(b) = a - b : N
\idx{diff_self_eq_0} a:N ==> a - a = 0 : N
\idx{add_inverse_diff} [| a:N; b:N; b-a=0 : N |] ==> b #+ (a-b) = a : N
\end{ttbox}
\subcaption{Some theorems of arithmetic}
\caption{The theory of arithmetic} \label{ctt-arith}
\end{figure}
\section{A theory of arithmetic}
{\CTT} contains a theory of elementary arithmetic. It proves the
properties of addition, multiplication, subtraction, division, and
remainder, culminating in the theorem
\[ a \bmod b + (a/b)\times b = a. \]
Figure~\ref{ctt-arith} presents the definitions and some of the key
theorems, including commutative, distributive, and associative laws. The
theory has the {\ML} identifier \ttindexbold{arith.thy}. All proofs are on
the file {\tt CTT/arith.ML}.
The operators~\verb'#+', \verb'-', \verb'|-|', \verb'#*', \verb'mod'
and~\verb'div' stand for sum, difference, absolute difference, product,
remainder and quotient, respectively. Since Type Theory has only primitive
recursion, some of their definitions may be obscure.
The difference~$a-b$ is computed by taking $b$ predecessors of~$a$, where
the predecessor function is $\lambda v. {\tt rec}(v, 0, \lambda x\,y.x)$.
The remainder $a\bmod b$ counts up to~$a$ in a cyclic fashion, using 0
as the successor of~$b-1$. Absolute difference is used to test the
equality $succ(v)=b$.
The quotient $a/b$ is computed by adding one for every number $x$
such that $0\leq x \leq a$ and $x\bmod b = 0$.
\section{The examples directory}
This directory contains examples and experimental proofs in {\CTT}.
\begin{description}
\item[{\tt CTT/ex/typechk.ML}]
contains simple examples of type checking and type deduction.
\item[{\tt CTT/ex/elim.ML}]
contains some examples from Martin-L\"of~\cite{martinlof84}, proved using
{\tt pc_tac}.
\item[{\tt CTT/ex/equal.ML}]
contains simple examples of rewriting.
\item[{\tt CTT/ex/synth.ML}]
demonstrates the use of unknowns with some trivial examples of program
synthesis.
\end{description}
\section{Example: type inference}
Type inference involves proving a goal of the form $a\in\Var{A}$, where $a$
is a term and $\Var{A}$ is an unknown standing for its type. The type,
initially
unknown, takes shape in the course of the proof. Our example is the
predecessor function on the natural numbers.
\begin{ttbox}
goal CTT.thy "lam n. rec(n, 0, \%x y.x) : ?A";
{\out Level 0}
{\out lam n. rec(n,0,\%x y. x) : ?A}
{\out 1. lam n. rec(n,0,\%x y. x) : ?A}
\end{ttbox}
Since the term is a Constructive Type Theory $\lambda$-abstraction (not to
be confused with a meta-level abstraction), we apply the rule
\ttindex{ProdI}, for $\Pi$-introduction. This instantiates~$\Var{A}$ to a
product type of unknown domain and range.
\begin{ttbox}
by (resolve_tac [ProdI] 1);
{\out Level 1}
{\out lam n. rec(n,0,\%x y. x) : PROD x:?A1. ?B1(x)}
{\out 1. ?A1 type}
{\out 2. !!n. n : ?A1 ==> rec(n,0,\%x y. x) : ?B1(n)}
\end{ttbox}
Subgoal~1 is too flexible. It can be solved by instantiating~$\Var{A@1}$
to any type, but most instantiations will invalidate subgoal~2. We
therefore tackle the latter subgoal. It asks the type of a term beginning
with {\tt rec}, which can be found by $N$-elimination.\index{*NE}
\begin{ttbox}
by (eresolve_tac [NE] 2);
{\out Level 2}
{\out lam n. rec(n,0,\%x y. x) : PROD x:N. ?C2(x,x)}
{\out 1. N type}
{\out 2. !!n. 0 : ?C2(n,0)}
{\out 3. !!n x y. [| x : N; y : ?C2(n,x) |] ==> x : ?C2(n,succ(x))}
\end{ttbox}
Subgoal~1 is no longer flexible: we now know~$\Var{A@1}$ is the type of
natural numbers. However, let us continue proving nontrivial subgoals.
Subgoal~2 asks, what is the type of~0?\index{*NIO}
\begin{ttbox}
by (resolve_tac [NI0] 2);
{\out Level 3}
{\out lam n. rec(n,0,\%x y. x) : N --> N}
{\out 1. N type}
{\out 2. !!n x y. [| x : N; y : N |] ==> x : N}
\end{ttbox}
The type~$\Var{A}$ is now fully determined. It is the product type
$\prod@{x\in N}N$, which is equivalent to function type $N\to N$ because
there is no dependence on~$x$. But we must prove all the subgoals to show
that the original term is validly typed. Subgoal~2 is provable by
assumption and the remaining subgoal falls by $N$-formation.\index{*NF}
\begin{ttbox}
by (assume_tac 2);
{\out Level 4}
{\out lam n. rec(n,0,\%x y. x) : N --> N}
{\out 1. N type}
\ttbreak
by (resolve_tac [NF] 1);
{\out Level 5}
{\out lam n. rec(n,0,\%x y. x) : N --> N}
{\out No subgoals!}
\end{ttbox}
Calling \ttindex{typechk_tac} can prove this theorem in one step.
Even if the original term is ill-typed, one can infer a type for it, but
unprovable subgoals will be left. As an exercise, try to prove the
following invalid goal:
\begin{ttbox}
goal CTT.thy "lam n. rec(n, 0, \%x y.tt) : ?A";
\end{ttbox}
\section{An example of logical reasoning}
Logical reasoning in Type Theory involves proving a goal of the form
$\Var{a}\in A$, where type $A$ expresses a proposition and $\Var{a}$ is an
unknown standing
for its proof term: a value of type $A$. This term is initially unknown, as
with type inference, and takes shape during the proof. Our example
expresses, by propositions-as-types, a theorem about quantifiers in a
sorted logic:
\[ \infer{(\ex{x\in A}P(x)) \disj (\ex{x\in A}Q(x))}
{\ex{x\in A}P(x)\disj Q(x)}
\]
It it related to a distributive law of Type Theory:
\[ \infer{(A\times B) + (A\times C)}{A\times(B+C)} \]
Generalizing this from $\times$ to $\Sigma$, and making the typing
conditions explicit, yields
\[ \infer{\Var{a} \in (\sum@{x\in A} B(x)) + (\sum@{x\in A} C(x))}
{\hbox{$A$ type} &
\infer*{\hbox{$B(x)$ type}}{[x\in A]} &
\infer*{\hbox{$C(x)$ type}}{[x\in A]} &
p\in \sum@{x\in A} B(x)+C(x)}
\]
To derive this rule, we bind its premises --- returned by~\ttindex{goal}
--- to the {\ML} variable~{\tt prems}.
\begin{ttbox}
val prems = goal CTT.thy
"[| A type; \ttback
\ttback !!x. x:A ==> B(x) type; \ttback
\ttback !!x. x:A ==> C(x) type; \ttback
\ttback p: SUM x:A. B(x) + C(x) \ttback
\ttback |] ==> ?a : (SUM x:A. B(x)) + (SUM x:A. C(x))";
{\out Level 0}
{\out ?a : (SUM x:A. B(x)) + (SUM x:A. C(x))}
{\out 1. ?a : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\ttbreak
{\out val prems = ["A type [A type]",}
{\out "?x : A ==> B(?x) type [!!x. x : A ==> B(x) type]",}
{\out "?x : A ==> C(?x) type [!!x. x : A ==> C(x) type]",}
{\out "p : SUM x:A. B(x) + C(x) [p : SUM x:A. B(x) + C(x)]"]}
{\out : thm list}
\end{ttbox}
One of the premises involves summation ($\Sigma$). Since it is a premise
rather than the assumption of a goal, it cannot be found by
\ttindex{eresolve_tac}. We could insert it by calling
\hbox{\tt \ttindex{cut_facts_tac} prems 1}. Instead, let us resolve the
$\Sigma$-elimination rule with the premises using~{\tt RL}; this forward
inference yields one result, which we supply to
\ttindex{resolve_tac}.\index{*SumE}\index{*RL}
\begin{ttbox}
by (resolve_tac (prems RL [SumE]) 1);
{\out Level 1}
{\out split(p,?c1) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
{\out 1. !!x y.}
{\out [| x : A; y : B(x) + C(x) |] ==>}
{\out ?c1(x,y) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\end{ttbox}
The subgoal has two new parameters, $x$ and~$y$. In the main goal,
$\Var{a}$ has been instantiated with a \ttindex{split} term. The
assumption $y\in B(x) + C(x)$ is eliminated next, causing a case split and
a further parameter,~$xa$. It also inserts~\ttindex{when} into the main goal.
\index{*PlusE}
\begin{ttbox}
by (eresolve_tac [PlusE] 1);
{\out Level 2}
{\out split(p,\%x y. when(y,?c2(x,y),?d2(x,y)))}
{\out : (SUM x:A. B(x)) + (SUM x:A. C(x))}
{\out 1. !!x y xa.}
{\out [| x : A; xa : B(x) |] ==>}
{\out ?c2(x,y,xa) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\ttbreak
{\out 2. !!x y ya.}
{\out [| x : A; ya : C(x) |] ==>}
{\out ?d2(x,y,ya) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\end{ttbox}
To complete the proof object for the main goal, we need to instantiate the
terms $\Var{c@2}(x,y,xa)$ and $\Var{d@2}(x,y,xa)$. We attack subgoal~1 by
introduction of~$+$; since it assumes $xa\in B(x)$, we take the left
injection~(\ttindex{inl}).
\index{*PlusI_inl}
\begin{ttbox}
by (resolve_tac [PlusI_inl] 1);
{\out Level 3}
{\out split(p,\%x y. when(y,\%xa. inl(?a3(x,y,xa)),?d2(x,y)))}
{\out : (SUM x:A. B(x)) + (SUM x:A. C(x))}
{\out 1. !!x y xa. [| x : A; xa : B(x) |] ==> ?a3(x,y,xa) : SUM x:A. B(x)}
{\out 2. !!x y xa. [| x : A; xa : B(x) |] ==> SUM x:A. C(x) type}
\ttbreak
{\out 3. !!x y ya.}
{\out [| x : A; ya : C(x) |] ==>}
{\out ?d2(x,y,ya) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\end{ttbox}
A new subgoal has appeared, to verify that $\sum@{x\in A}C(x)$ is a type.
Continuing with subgoal~1, we apply $\Sigma$-introduction. The main goal
now contains an ordered pair.
\index{*SumI}
\begin{ttbox}
by (resolve_tac [SumI] 1);
{\out Level 4}
{\out split(p,\%x y. when(y,\%xa. inl(<?a4(x,y,xa),?b4(x,y,xa)>),?d2(x,y)))}
{\out : (SUM x:A. B(x)) + (SUM x:A. C(x))}
{\out 1. !!x y xa. [| x : A; xa : B(x) |] ==> ?a4(x,y,xa) : A}
{\out 2. !!x y xa. [| x : A; xa : B(x) |] ==> ?b4(x,y,xa) : B(?a4(x,y,xa))}
{\out 3. !!x y xa. [| x : A; xa : B(x) |] ==> SUM x:A. C(x) type}
{\out 4. !!x y ya.}
{\out [| x : A; ya : C(x) |] ==>}
{\out ?d2(x,y,ya) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\end{ttbox}
The two new subgoals both hold by assumption. Observe how the unknowns
$\Var{a@4}$ and $\Var{b@4}$ are instantiated throughout the proof state.
\begin{ttbox}
by (assume_tac 1);
{\out Level 5}
{\out split(p,\%x y. when(y,\%xa. inl(<x,?b4(x,y,xa)>),?d2(x,y)))}
{\out : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\ttbreak
{\out 1. !!x y xa. [| x : A; xa : B(x) |] ==> ?b4(x,y,xa) : B(x)}
{\out 2. !!x y xa. [| x : A; xa : B(x) |] ==> SUM x:A. C(x) type}
{\out 3. !!x y ya.}
{\out [| x : A; ya : C(x) |] ==>}
{\out ?d2(x,y,ya) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\ttbreak
by (assume_tac 1);
{\out Level 6}
{\out split(p,\%x y. when(y,\%xa. inl(<x,xa>),?d2(x,y)))}
{\out : (SUM x:A. B(x)) + (SUM x:A. C(x))}
{\out 1. !!x y xa. [| x : A; xa : B(x) |] ==> SUM x:A. C(x) type}
{\out 2. !!x y ya.}
{\out [| x : A; ya : C(x) |] ==>}
{\out ?d2(x,y,ya) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\end{ttbox}
Subgoal~1 is just type checking. It yields to \ttindex{typechk_tac},
supplied with the current list of premises.
\begin{ttbox}
by (typechk_tac prems);
{\out Level 7}
{\out split(p,\%x y. when(y,\%xa. inl(<x,xa>),?d2(x,y)))}
{\out : (SUM x:A. B(x)) + (SUM x:A. C(x))}
{\out 1. !!x y ya.}
{\out [| x : A; ya : C(x) |] ==>}
{\out ?d2(x,y,ya) : (SUM x:A. B(x)) + (SUM x:A. C(x))}
\end{ttbox}
The other case is similar. Let us prove it by \ttindex{pc_tac}, and note
the final proof object.
\begin{ttbox}
by (pc_tac prems 1);
{\out Level 8}
{\out split(p,\%x y. when(y,\%xa. inl(<x,xa>),\%y. inr(<x,y>)))}
{\out : (SUM x:A. B(x)) + (SUM x:A. C(x))}
{\out No subgoals!}
\end{ttbox}
Calling \ttindex{pc_tac} after the first $\Sigma$-elimination above also
proves this theorem.
\section{Example: deriving a currying functional}
In simply-typed languages such as {\ML}, a currying functional has the type
\[ (A\times B \to C) \to (A\to (B\to C)). \]
Let us generalize this to~$\Sigma$ and~$\Pi$.
The functional takes a function~$f$ that maps $z:\Sigma(A,B)$
to~$C(z)$; the resulting function maps $x\in A$ and $y\in B(x)$ to
$C(\langle x,y\rangle)$.
Formally, there are three typing premises. $A$ is a type; $B$ is an
$A$-indexed family of types; $C$ is a family of types indexed by
$\Sigma(A,B)$. The goal is expressed using \hbox{\tt PROD f} to ensure
that the parameter corresponding to the functional's argument is really
called~$f$; Isabelle echoes the type using \verb|-->| because there is no
explicit dependence upon~$f$.
\begin{ttbox}
val prems = goal CTT.thy
"[| A type; !!x. x:A ==> B(x) type; \ttback
\ttback !!z. z: (SUM x:A. B(x)) ==> C(z) type \ttback
\ttback |] ==> ?a : PROD f: (PROD z : (SUM x:A . B(x)) . C(z)). \ttback
\ttback (PROD x:A . PROD y:B(x) . C(<x,y>))";
\ttbreak
{\out Level 0}
{\out ?a : (PROD z:SUM x:A. B(x). C(z)) -->}
{\out (PROD x:A. PROD y:B(x). C(<x,y>))}
{\out 1. ?a : (PROD z:SUM x:A. B(x). C(z)) -->}
{\out (PROD x:A. PROD y:B(x). C(<x,y>))}
\ttbreak
{\out val prems = ["A type [A type]",}
{\out "?x : A ==> B(?x) type [!!x. x : A ==> B(x) type]",}
{\out "?z : SUM x:A. B(x) ==> C(?z) type}
{\out [!!z. z : SUM x:A. B(x) ==> C(z) type]"] : thm list}
\end{ttbox}
This is a chance to demonstrate \ttindex{intr_tac}. Here, the tactic
repeatedly applies $\Pi$-introduction, proving the rather
tiresome typing conditions.
Note that $\Var{a}$ becomes instantiated to three nested
$\lambda$-abstractions. It would be easier to read if the bound variable
names agreed with the parameters in the subgoal. Isabelle attempts to give
parameters the same names as corresponding bound variables in the goal, but
this does not always work. In any event, the goal is logically correct.
\begin{ttbox}
by (intr_tac prems);
{\out Level 1}
{\out lam x xa xb. ?b7(x,xa,xb)}
{\out : (PROD z:SUM x:A. B(x). C(z)) --> (PROD x:A. PROD y:B(x). C(<x,y>))}
{\out 1. !!f x y.}
{\out [| f : PROD z:SUM x:A. B(x). C(z); x : A; y : B(x) |] ==>}
{\out ?b7(f,x,y) : C(<x,y>)}
\end{ttbox}
Using $\Pi$-elimination, we solve subgoal~1 by applying the function~$f$.
\index{*ProdE}
\begin{ttbox}
by (eresolve_tac [ProdE] 1);
{\out Level 2}
{\out lam x xa xb. x ` <xa,xb>}
{\out : (PROD z:SUM x:A. B(x). C(z)) --> (PROD x:A. PROD y:B(x). C(<x,y>))}
{\out 1. !!f x y. [| x : A; y : B(x) |] ==> <x,y> : SUM x:A. B(x)}
\end{ttbox}
Finally, we exhibit a suitable argument for the function application. This
is straightforward using introduction rules.
\index{*intr_tac}
\begin{ttbox}
by (intr_tac prems);
{\out Level 3}
{\out lam x xa xb. x ` <xa,xb>}
{\out : (PROD z:SUM x:A. B(x). C(z)) --> (PROD x:A. PROD y:B(x). C(<x,y>))}
{\out No subgoals!}
\end{ttbox}
Calling~\ttindex{pc_tac} would have proved this theorem in one step; it can
also prove an example by Martin-L\"of, related to $\disj$-elimination
\cite[page~58]{martinlof84}.
\section{Example: proving the Axiom of Choice} \label{ctt-choice}
Suppose we have a function $h\in \prod@{x\in A}\sum@{y\in B(x)} C(x,y)$,
which takes $x\in A$ to some $y\in B(x)$ paired with some $z\in C(x,y)$.
Interpreting propositions as types, this asserts that for all $x\in A$
there exists $y\in B(x)$ such that $C(x,y)$. The Axiom of Choice asserts
that we can construct a function $f\in \prod@{x\in A}B(x)$ such that
$C(x,f{\tt`}x)$ for all $x\in A$, where the latter property is witnessed by a
function $g\in \prod@{x\in A}C(x,f{\tt`}x)$.
In principle, the Axiom of Choice is simple to derive in Constructive Type
Theory \cite[page~50]{martinlof84}. The following definitions work:
\begin{eqnarray*}
f & \equiv & {\tt fst} \circ h \\
g & \equiv & {\tt snd} \circ h
\end{eqnarray*}
But a completely formal proof is hard to find. Many of the rules can be
applied in a multiplicity of ways, yielding a large number of higher-order
unifiers. The proof can get bogged down in the details. But with a
careful selection of derived rules (recall Fig.\ts\ref{ctt-derived}) and
the type checking tactics, we can prove the theorem in nine steps.
\begin{ttbox}
val prems = goal CTT.thy
"[| A type; !!x. x:A ==> B(x) type; \ttback
\ttback !!x y.[| x:A; y:B(x) |] ==> C(x,y) type \ttback
\ttback |] ==> ?a : PROD h: (PROD x:A. SUM y:B(x). C(x,y)). \ttback
\ttback (SUM f: (PROD x:A. B(x)). PROD x:A. C(x, f`x))";
{\out Level 0}
{\out ?a : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
{\out 1. ?a : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
\ttbreak
{\out val prems = ["A type [A type]",}
{\out "?x : A ==> B(?x) type [!!x. x : A ==> B(x) type]",}
{\out "[| ?x : A; ?y : B(?x) |] ==> C(?x, ?y) type}
{\out [!!x y. [| x : A; y : B(x) |] ==> C(x, y) type]"]}
{\out : thm list}
\end{ttbox}
First, \ttindex{intr_tac} applies introduction rules and performs routine
type checking. This instantiates~$\Var{a}$ to a construction involving
three $\lambda$-abstractions and an ordered pair.
\begin{ttbox}
by (intr_tac prems);
{\out Level 1}
{\out lam x. <lam xa. ?b7(x,xa),lam xa. ?b8(x,xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
\ttbreak
{\out 1. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out ?b7(h,x) : B(x)}
\ttbreak
{\out 2. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out ?b8(h,x) : C(x,(lam x. ?b7(h,x)) ` x)}
\end{ttbox}
Subgoal~1 asks to find the choice function itself, taking $x\in A$ to some
$\Var{b@7}(h,x)\in B(x)$. Subgoal~2 asks, given $x\in A$, for a proof
object $\Var{b@8}(h,x)$ to witness that the choice function's argument and
result lie in the relation~$C$. This latter task will take up most of the
proof.
\index{*ProdE}\index{*SumE_fst}\index{*RS}
\begin{ttbox}
by (eresolve_tac [ProdE RS SumE_fst] 1);
{\out Level 2}
{\out lam x. <lam xa. fst(x ` xa),lam xa. ?b8(x,xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
\ttbreak
{\out 1. !!h x. x : A ==> x : A}
{\out 2. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out ?b8(h,x) : C(x,(lam x. fst(h ` x)) ` x)}
\end{ttbox}
Above, we have composed \ttindex{fst} with the function~$h$. Unification
has deduced that the function must be applied to $x\in A$.
\begin{ttbox}
by (assume_tac 1);
{\out Level 3}
{\out lam x. <lam xa. fst(x ` xa),lam xa. ?b8(x,xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
{\out 1. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out ?b8(h,x) : C(x,(lam x. fst(h ` x)) ` x)}
\end{ttbox}
Before we can compose \ttindex{snd} with~$h$, the arguments of $C$ must be
simplified. The derived rule \ttindex{replace_type} lets us replace a type
by any equivalent type, shown below as the schematic term $\Var{A@{13}}(h,x)$:
\begin{ttbox}
by (resolve_tac [replace_type] 1);
{\out Level 4}
{\out lam x. <lam xa. fst(x ` xa),lam xa. ?b8(x,xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
\ttbreak
{\out 1. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out C(x,(lam x. fst(h ` x)) ` x) = ?A13(h,x)}
\ttbreak
{\out 2. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out ?b8(h,x) : ?A13(h,x)}
\end{ttbox}
The derived rule \ttindex{subst_eqtyparg} lets us simplify a type's
argument (by currying, $C(x)$ is a unary type operator):
\begin{ttbox}
by (resolve_tac [subst_eqtyparg] 1);
{\out Level 5}
{\out lam x. <lam xa. fst(x ` xa),lam xa. ?b8(x,xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
\ttbreak
{\out 1. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out (lam x. fst(h ` x)) ` x = ?c14(h,x) : ?A14(h,x)}
\ttbreak
{\out 2. !!h x z.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A;}
{\out z : ?A14(h,x) |] ==>}
{\out C(x,z) type}
\ttbreak
{\out 3. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out ?b8(h,x) : C(x,?c14(h,x))}
\end{ttbox}
Subgoal~1 requires simply $\beta$-contraction, which is the rule
\ttindex{ProdC}. The term $\Var{c@{14}}(h,x)$ in the last subgoal
receives the contracted result.
\begin{ttbox}
by (resolve_tac [ProdC] 1);
{\out Level 6}
{\out lam x. <lam xa. fst(x ` xa),lam xa. ?b8(x,xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
\ttbreak
{\out 1. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out x : ?A15(h,x)}
\ttbreak
{\out 2. !!h x xa.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A;}
{\out xa : ?A15(h,x) |] ==>}
{\out fst(h ` xa) : ?B15(h,x,xa)}
\ttbreak
{\out 3. !!h x z.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A;}
{\out z : ?B15(h,x,x) |] ==>}
{\out C(x,z) type}
\ttbreak
{\out 4. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out ?b8(h,x) : C(x,fst(h ` x))}
\end{ttbox}
Routine type checking goals proliferate in Constructive Type Theory, but
\ttindex{typechk_tac} quickly solves them. Note the inclusion of
\ttindex{SumE_fst} along with the premises.
\begin{ttbox}
by (typechk_tac (SumE_fst::prems));
{\out Level 7}
{\out lam x. <lam xa. fst(x ` xa),lam xa. ?b8(x,xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
\ttbreak
{\out 1. !!h x.}
{\out [| h : PROD x:A. SUM y:B(x). C(x,y); x : A |] ==>}
{\out ?b8(h,x) : C(x,fst(h ` x))}
\end{ttbox}
We are finally ready to compose \ttindex{snd} with~$h$.
\index{*ProdE}\index{*SumE_snd}\index{*RS}
\begin{ttbox}
by (eresolve_tac [ProdE RS SumE_snd] 1);
{\out Level 8}
{\out lam x. <lam xa. fst(x ` xa),lam xa. snd(x ` xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
\ttbreak
{\out 1. !!h x. x : A ==> x : A}
{\out 2. !!h x. x : A ==> B(x) type}
{\out 3. !!h x xa. [| x : A; xa : B(x) |] ==> C(x,xa) type}
\end{ttbox}
The proof object has reached its final form. We call \ttindex{typechk_tac}
to finish the type checking.
\begin{ttbox}
by (typechk_tac prems);
{\out Level 9}
{\out lam x. <lam xa. fst(x ` xa),lam xa. snd(x ` xa)>}
{\out : (PROD x:A. SUM y:B(x). C(x,y)) -->}
{\out (SUM f:PROD x:A. B(x). PROD x:A. C(x,f ` x))}
{\out No subgoals!}
\end{ttbox}