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\chapter{Zermelo-Fraenkel Set Theory}
\index{set theory|(}

The theory~\thydx{ZF} implements Zermelo-Fraenkel set
theory~\cite{halmos60,suppes72} as an extension of~\texttt{FOL}, classical
first-order logic.  The theory includes a collection of derived natural
deduction rules, for use with Isabelle's classical reasoner.  Some
of it is based on the work of No\"el~\cite{noel}.

A tremendous amount of set theory has been formally developed, including the
basic properties of relations, functions, ordinals and cardinals.  Significant
results have been proved, such as the Schr\"oder-Bernstein Theorem, the
Wellordering Theorem and a version of Ramsey's Theorem.  \texttt{ZF} provides
both the integers and the natural numbers.  General methods have been
developed for solving recursion equations over monotonic functors; these have
been applied to yield constructions of lists, trees, infinite lists, etc.

\texttt{ZF} has a flexible package for handling inductive definitions,
such as inference systems, and datatype definitions, such as lists and
trees.  Moreover it handles coinductive definitions, such as
bisimulation relations, and codatatype definitions, such as streams.  It
provides a streamlined syntax for defining primitive recursive functions over
datatypes.

Published articles~\cite{paulson-set-I,paulson-set-II} describe \texttt{ZF}
less formally than this chapter.  Isabelle employs a novel treatment of
non-well-founded data structures within the standard {\sc zf} axioms including
the Axiom of Foundation~\cite{paulson-mscs}.

\section{Which version of axiomatic set theory?}
The two main axiom systems for set theory are Bernays-G\"odel~({\sc bg})
and Zermelo-Fraenkel~({\sc zf}).  Resolution theorem provers can use {\sc
bg} because it is finite~\cite{boyer86,quaife92}.  {\sc zf} does not
have a finite axiom system because of its Axiom Scheme of Replacement.
This makes it awkward to use with many theorem provers, since instances
of the axiom scheme have to be invoked explicitly.  Since Isabelle has no
difficulty with axiom schemes, we may adopt either axiom system.

These two theories differ in their treatment of {\bf classes}, which are
collections that are too big' to be sets.  The class of all sets,~$V$,
cannot be a set without admitting Russell's Paradox.  In {\sc bg}, both
classes and sets are individuals; $x\in V$ expresses that $x$ is a set.  In
{\sc zf}, all variables denote sets; classes are identified with unary
predicates.  The two systems define essentially the same sets and classes,
with similar properties.  In particular, a class cannot belong to another
class (let alone a set).

Modern set theorists tend to prefer {\sc zf} because they are mainly concerned
with sets, rather than classes.  {\sc bg} requires tiresome proofs that various
collections are sets; for instance, showing $x\in\{x\}$ requires showing that
$x$ is a set.

\begin{figure} \small
\begin{center}
\begin{tabular}{rrr}
\it name      &\it meta-type  & \it description \\
\cdx{Let}     & $[\alpha,\alpha\To\beta]\To\beta$ & let binder\\
\cdx{0}       & $i$           & empty set\\
\cdx{cons}    & $[i,i]\To i$  & finite set constructor\\
\cdx{Upair}   & $[i,i]\To i$  & unordered pairing\\
\cdx{Pair}    & $[i,i]\To i$  & ordered pairing\\
\cdx{Inf}     & $i$   & infinite set\\
\cdx{Pow}     & $i\To i$      & powerset\\
\cdx{Union} \cdx{Inter} & $i\To i$    & set union/intersection \\
\cdx{split}   & $[[i,i]\To i, i] \To i$ & generalized projection\\
\cdx{fst} \cdx{snd}   & $i\To i$      & projections\\
\cdx{converse}& $i\To i$      & converse of a relation\\
\cdx{succ}    & $i\To i$      & successor\\
\cdx{Collect} & $[i,i\To o]\To i$     & separation\\
\cdx{Replace} & $[i, [i,i]\To o] \To i$       & replacement\\
\cdx{PrimReplace} & $[i, [i,i]\To o] \To i$   & primitive replacement\\
\cdx{RepFun}  & $[i, i\To i] \To i$   & functional replacement\\
\cdx{Pi} \cdx{Sigma}  & $[i,i\To i]\To i$     & general product/sum\\
\cdx{domain}  & $i\To i$      & domain of a relation\\
\cdx{range}   & $i\To i$      & range of a relation\\
\cdx{field}   & $i\To i$      & field of a relation\\
\cdx{Lambda}  & $[i, i\To i]\To i$    & $\lambda$-abstraction\\
\cdx{restrict}& $[i, i] \To i$        & restriction of a function\\
\cdx{The}     & $[i\To o]\To i$       & definite description\\
\cdx{if}      & $[o,i,i]\To i$        & conditional\\
\cdx{Ball} \cdx{Bex}  & $[i, i\To o]\To o$    & bounded quantifiers
\end{tabular}
\end{center}
\subcaption{Constants}

\begin{center}
\index{*"" symbol}
\index{*"-"" symbol}
\index{*" symbol}\index{function applications}
\index{*"- symbol}
\index{*": symbol}
\index{*"<"= symbol}
\begin{tabular}{rrrr}
\it symbol  & \it meta-type & \it priority & \it description \\
\tt         & $[i,i]\To i$  &  Left 90      & image \\
\tt -       & $[i,i]\To i$  &  Left 90      & inverse image \\
\tt          & $[i,i]\To i$  &  Left 90      & application \\
\sdx{Int}     & $[i,i]\To i$  &  Left 70      & intersection ($\int$) \\
\sdx{Un}      & $[i,i]\To i$  &  Left 65      & union ($\un$) \\
\tt -         & $[i,i]\To i$  &  Left 65      & set difference ($-$) \$1ex] \tt: & [i,i]\To o & Left 50 & membership (\in) \\ \tt <= & [i,i]\To o & Left 50 & subset (\subseteq) \end{tabular} \end{center} \subcaption{Infixes} \caption{Constants of ZF} \label{zf-constants} \end{figure} \section{The syntax of set theory} The language of set theory, as studied by logicians, has no constants. The traditional axioms merely assert the existence of empty sets, unions, powersets, etc.; this would be intolerable for practical reasoning. The Isabelle theory declares constants for primitive sets. It also extends \texttt{FOL} with additional syntax for finite sets, ordered pairs, comprehension, general union/intersection, general sums/products, and bounded quantifiers. In most other respects, Isabelle implements precisely Zermelo-Fraenkel set theory. Figure~\ref{zf-constants} lists the constants and infixes of~ZF, while Figure~\ref{zf-trans} presents the syntax translations. Finally, Figure~\ref{zf-syntax} presents the full grammar for set theory, including the constructs of FOL. Local abbreviations can be introduced by a \isa{let} construct whose syntax appears in Fig.\ts\ref{zf-syntax}. Internally it is translated into the constant~\cdx{Let}. It can be expanded by rewriting with its definition, \tdx{Let_def}. Apart from \isa{let}, set theory does not use polymorphism. All terms in ZF have type~\tydx{i}, which is the type of individuals and has class~\cldx{term}. The type of first-order formulae, remember, is~\tydx{o}. Infix operators include binary union and intersection (A\un B and A\int B), set difference (A-B), and the subset and membership relations. Note that a\verb|~:|b is translated to \lnot(a\in b), which is equivalent to a\notin b. The union and intersection operators (\bigcup A and \bigcap A) form the union or intersection of a set of sets; \bigcup A means the same as \bigcup@{x\in A}x. Of these operators, only \bigcup A is primitive. The constant \cdx{Upair} constructs unordered pairs; thus \isa{Upair(A,B)} denotes the set~\{A,B\} and \isa{Upair(A,A)} denotes the singleton~\{A\}. General union is used to define binary union. The Isabelle version goes on to define the constant \cdx{cons}: \begin{eqnarray*} A\cup B & \equiv & \bigcup(\isa{Upair}(A,B)) \\ \isa{cons}(a,B) & \equiv & \isa{Upair}(a,a) \un B \end{eqnarray*} The \{a@1, \ldots\} notation abbreviates finite sets constructed in the obvious manner using~\isa{cons} and~\emptyset (the empty set) \isasymin \begin{eqnarray*} \{a,b,c\} & \equiv & \isa{cons}(a,\isa{cons}(b,\isa{cons}(c,\emptyset))) \end{eqnarray*} The constant \cdx{Pair} constructs ordered pairs, as in \isa{Pair(a,b)}. Ordered pairs may also be written within angle brackets, as {\tt<a,b>}. The n-tuple {\tt<a@1,\ldots,a@{n-1},a@n>} abbreviates the nest of pairs\par\nobreak \centerline{\isa{Pair(a@1,\ldots,Pair(a@{n-1},a@n)\ldots).}} In ZF, a function is a set of pairs. A ZF function~f is simply an individual as far as Isabelle is concerned: its Isabelle type is~i, not say i\To i. The infix operator~{\tt} denotes the application of a function set to its argument; we must write~f{\tt}x, not~f(x). The syntax for image is~f{\tt}A and that for inverse image is~f{\tt-}A. \begin{figure} \index{lambda abs@\lambda-abstractions} \index{*"-"> symbol} \index{*"* symbol} \begin{center} \footnotesize\tt\frenchspacing \begin{tabular}{rrr} \it external & \it internal & \it description \\ a \ttilde: b & \ttilde(a : b) & \rm negated membership\\ \ttlbracea@1, \ldots, a@n\ttrbrace & cons(a@1,\ldots,cons(a@n,0)) & \rm finite set \\ <a@1, \ldots, a@{n-1}, a@n> & Pair(a@1,\ldots,Pair(a@{n-1},a@n)\ldots) & \rm ordered n-tuple \\ \ttlbracex:A . P[x]\ttrbrace & Collect(A,\lambda x. P[x]) & \rm separation \\ \ttlbracey . x:A, Q[x,y]\ttrbrace & Replace(A,\lambda x\,y. Q[x,y]) & \rm replacement \\ \ttlbraceb[x] . x:A\ttrbrace & RepFun(A,\lambda x. b[x]) & \rm functional replacement \\ \sdx{INT} x:A . B[x] & Inter(\ttlbraceB[x] . x:A\ttrbrace) & \rm general intersection \\ \sdx{UN} x:A . B[x] & Union(\ttlbraceB[x] . x:A\ttrbrace) & \rm general union \\ \sdx{PROD} x:A . B[x] & Pi(A,\lambda x. B[x]) & \rm general product \\ \sdx{SUM} x:A . B[x] & Sigma(A,\lambda x. B[x]) & \rm general sum \\ A -> B & Pi(A,\lambda x. B) & \rm function space \\ A * B & Sigma(A,\lambda x. B) & \rm binary product \\ \sdx{THE} x . P[x] & The(\lambda x. P[x]) & \rm definite description \\ \sdx{lam} x:A . b[x] & Lambda(A,\lambda x. b[x]) & \rm \lambda-abstraction\\[1ex] \sdx{ALL} x:A . P[x] & Ball(A,\lambda x. P[x]) & \rm bounded \forall \\ \sdx{EX} x:A . P[x] & Bex(A,\lambda x. P[x]) & \rm bounded \exists \end{tabular} \end{center} \caption{Translations for ZF} \label{zf-trans} \end{figure} \begin{figure} \index{*let symbol} \index{*in symbol} \dquotes \[\begin{array}{rcl} term & = & \hbox{expression of type~i} \\ & | & "let"~id~"="~term";"\dots";"~id~"="~term~"in"~term \\ & | & "if"~term~"then"~term~"else"~term \\ & | & "{\ttlbrace} " term\; ("," term)^* " {\ttrbrace}" \\ & | & "< " term\; ("," term)^* " >" \\ & | & "{\ttlbrace} " id ":" term " . " formula " {\ttrbrace}" \\ & | & "{\ttlbrace} " id " . " id ":" term ", " formula " {\ttrbrace}" \\ & | & "{\ttlbrace} " term " . " id ":" term " {\ttrbrace}" \\ & | & term "  " term \\ & | & term " - " term \\ & | & term "  " term \\ & | & term " * " term \\ & | & term " \isasyminter " term \\ & | & term " \isasymunion " term \\ & | & term " - " term \\ & | & term " -> " term \\ & | & "THE~~" id " . " formula\\ & | & "lam~~" id ":" term " . " term \\ & | & "INT~~" id ":" term " . " term \\ & | & "UN~~~" id ":" term " . " term \\ & | & "PROD~" id ":" term " . " term \\ & | & "SUM~~" id ":" term " . " term \\[2ex] formula & = & \hbox{expression of type~o} \\ & | & term " : " term \\ & | & term " \ttilde: " term \\ & | & term " <= " term \\ & | & term " = " term \\ & | & term " \ttilde= " term \\ & | & "\ttilde\ " formula \\ & | & formula " \& " formula \\ & | & formula " | " formula \\ & | & formula " --> " formula \\ & | & formula " <-> " formula \\ & | & "ALL " id ":" term " . " formula \\ & | & "EX~~" id ":" term " . " formula \\ & | & "ALL~" id~id^* " . " formula \\ & | & "EX~~" id~id^* " . " formula \\ & | & "EX!~" id~id^* " . " formula \end{array}$
\caption{Full grammar for ZF} \label{zf-syntax}
\end{figure}

\section{Binding operators}
The constant \cdx{Collect} constructs sets by the principle of {\bf
separation}.  The syntax for separation is
\hbox{\tt\ttlbrace$x$:$A$.\ $P[x]$\ttrbrace}, where $P[x]$ is a formula
that may contain free occurrences of~$x$.  It abbreviates the set \isa{Collect($A$,$\lambda x. P[x]$)}, which consists of all $x\in A$ that
satisfy~$P[x]$.  Note that \isa{Collect} is an unfortunate choice of
name: some set theories adopt a set-formation principle, related to
replacement, called collection.

The constant \cdx{Replace} constructs sets by the principle of {\bf
replacement}.  The syntax
\hbox{\tt\ttlbrace$y$.\ $x$:$A$,$Q[x,y]$\ttrbrace} denotes the set
\isa{Replace($A$,$\lambda x\,y. Q[x,y]$)}, which consists of all~$y$ such
that there exists $x\in A$ satisfying~$Q[x,y]$.  The Replacement Axiom
has the condition that $Q$ must be single-valued over~$A$: for
all~$x\in A$ there exists at most one $y$ satisfying~$Q[x,y]$.  A
single-valued binary predicate is also called a {\bf class function}.

The constant \cdx{RepFun} expresses a special case of replacement,
where $Q[x,y]$ has the form $y=b[x]$.  Such a $Q$ is trivially
single-valued, since it is just the graph of the meta-level
function~$\lambda x. b[x]$.  The resulting set consists of all $b[x]$
for~$x\in A$.  This is analogous to the \ML{} functional \isa{map},
since it applies a function to every element of a set.  The syntax is
\isa{\ttlbrace$b[x]$.\ $x$:$A$\ttrbrace}, which expands to
\isa{RepFun($A$,$\lambda x. b[x]$)}.

\index{*INT symbol}\index{*UN symbol}
General unions and intersections of indexed
families of sets, namely $\bigcup@{x\in A}B[x]$ and $\bigcap@{x\in A}B[x]$,
are written \isa{UN $x$:$A$.\ $B[x]$} and \isa{INT $x$:$A$.\ $B[x]$}.
Their meaning is expressed using \isa{RepFun} as
$\bigcup(\{B[x]. x\in A\}) \qquad\hbox{and}\qquad \bigcap(\{B[x]. x\in A\}).$
General sums $\sum@{x\in A}B[x]$ and products $\prod@{x\in A}B[x]$ can be
constructed in set theory, where $B[x]$ is a family of sets over~$A$.  They
have as special cases $A\times B$ and $A\to B$, where $B$ is simply a set.
This is similar to the situation in Constructive Type Theory (set theory
has dependent sets') and calls for similar syntactic conventions.  The
constants~\cdx{Sigma} and~\cdx{Pi} construct general sums and
products.  Instead of \isa{Sigma($A$,$B$)} and \isa{Pi($A$,$B$)} we may
write
\isa{SUM $x$:$A$.\ $B[x]$} and \isa{PROD $x$:$A$.\ $B[x]$}.
\index{*SUM symbol}\index{*PROD symbol}%
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~\isa{op~*} and~\isa{op~->}.} Isabelle accepts these
abbreviations in parsing and uses them whenever possible for printing.

\index{*THE symbol} As mentioned above, whenever the axioms assert the
existence and uniqueness of a set, Isabelle's set theory declares a constant
for that set.  These constants can express the {\bf definite description}
operator~$\iota x. P[x]$, which stands for the unique~$a$ satisfying~$P[a]$,
if such exists.  Since all terms in ZF denote something, a description is
always meaningful, but we do not know its value unless $P[x]$ defines it
uniquely.  Using the constant~\cdx{The}, we may write descriptions as
\isa{The($\lambda x. P[x]$)} or use the syntax \isa{THE $x$.\ $P[x]$}.

\index{*lam symbol}
Function sets may be written in $\lambda$-notation; $\lambda x\in A. b[x]$
stands for the set of all pairs $\pair{x,b[x]}$ for $x\in A$.  In order for
this to be a set, the function's domain~$A$ must be given.  Using the
constant~\cdx{Lambda}, we may express function sets as \isa{Lambda($A$,$\lambda x. b[x]$)} or use the syntax \isa{lam $x$:$A$.\ $b[x]$}.

Isabelle's set theory defines two {\bf bounded quantifiers}:
\begin{eqnarray*}
\forall x\in A. P[x] &\hbox{abbreviates}& \forall x. x\in A\imp P[x] \\
\exists x\in A. P[x] &\hbox{abbreviates}& \exists x. x\in A\conj P[x]
\end{eqnarray*}
The constants~\cdx{Ball} and~\cdx{Bex} are defined
accordingly.  Instead of \isa{Ball($A$,$P$)} and \isa{Bex($A$,$P$)} we may
write
\isa{ALL $x$:$A$.\ $P[x]$} and \isa{EX $x$:$A$.\ $P[x]$}.

%%%% ZF.thy

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{Let_def}:           Let(s, f) == f(s)

\tdx{Ball_def}:          Ball(A,P) == {\isasymforall}x. x \isasymin A --> P(x)
\tdx{Bex_def}:           Bex(A,P)  == {\isasymexists}x. x \isasymin A & P(x)

\tdx{subset_def}:        A \isasymsubseteq B  == {\isasymforall}x \isasymin A. x \isasymin B
\tdx{extension}:         A = B  <->  A \isasymsubseteq B & B \isasymsubseteq A

\tdx{Union_iff}:         A \isasymin Union(C) <-> ({\isasymexists}B \isasymin C. A \isasymin B)
\tdx{Pow_iff}:           A \isasymin Pow(B) <-> A \isasymsubseteq B
\tdx{foundation}:        A=0 | ({\isasymexists}x \isasymin A. {\isasymforall}y \isasymin x. y \isasymnotin A)

\tdx{replacement}:       ({\isasymforall}x \isasymin A. {\isasymforall}y z. P(x,y) & P(x,z) --> y=z) ==>
b \isasymin PrimReplace(A,P) <-> ({\isasymexists}x{\isasymin}A. P(x,b))
\subcaption{The Zermelo-Fraenkel Axioms}

\tdx{Replace_def}: Replace(A,P) ==
PrimReplace(A, \%x y. (\isasymexists!z. P(x,z)) & P(x,y))
\tdx{RepFun_def}:  RepFun(A,f)  == {\ttlbrace}y . x \isasymin A, y=f(x)\ttrbrace
\tdx{the_def}:     The(P)       == Union({\ttlbrace}y . x \isasymin {\ttlbrace}0{\ttrbrace}, P(y){\ttrbrace})
\tdx{if_def}:      if(P,a,b)    == THE z. P & z=a | ~P & z=b
\tdx{Collect_def}: Collect(A,P) == {\ttlbrace}y . x \isasymin A, x=y & P(x){\ttrbrace}
\tdx{Upair_def}:   Upair(a,b)   ==
{\ttlbrace}y. x\isasymin{}Pow(Pow(0)), x=0 & y=a | x=Pow(0) & y=b{\ttrbrace}
\subcaption{Consequences of replacement}

\tdx{Inter_def}:   Inter(A) == {\ttlbrace}x \isasymin Union(A) . {\isasymforall}y \isasymin A. x \isasymin y{\ttrbrace}
\tdx{Un_def}:      A \isasymunion B  == Union(Upair(A,B))
\tdx{Int_def}:     A \isasyminter B  == Inter(Upair(A,B))
\tdx{Diff_def}:    A - B    == {\ttlbrace}x \isasymin A . x \isasymnotin B{\ttrbrace}
\subcaption{Union, intersection, difference}
\end{alltt*}
\caption{Rules and axioms of ZF} \label{zf-rules}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{cons_def}:    cons(a,A) == Upair(a,a) \isasymunion A
\tdx{succ_def}:    succ(i) == cons(i,i)
\tdx{infinity}:    0 \isasymin Inf & ({\isasymforall}y \isasymin Inf. succ(y) \isasymin Inf)
\subcaption{Finite and infinite sets}

\tdx{Pair_def}:      <a,b>      == {\ttlbrace}{\ttlbrace}a,a{\ttrbrace}, {\ttlbrace}a,b{\ttrbrace}{\ttrbrace}
\tdx{split_def}:     split(c,p) == THE y. {\isasymexists}a b. p=<a,b> & y=c(a,b)
\tdx{fst_def}:       fst(A)     == split(\%x y. x, p)
\tdx{snd_def}:       snd(A)     == split(\%x y. y, p)
\tdx{Sigma_def}:     Sigma(A,B) == {\isasymUnion}x \isasymin A. {\isasymUnion}y \isasymin B(x). {\ttlbrace}<x,y>{\ttrbrace}
\subcaption{Ordered pairs and Cartesian products}

\tdx{converse_def}: converse(r) == {\ttlbrace}z. w\isasymin{}r, {\isasymexists}x y. w=<x,y> & z=<y,x>{\ttrbrace}
\tdx{domain_def}:   domain(r)   == {\ttlbrace}x. w \isasymin r, {\isasymexists}y. w=<x,y>{\ttrbrace}
\tdx{range_def}:    range(r)    == domain(converse(r))
\tdx{field_def}:    field(r)    == domain(r) \isasymunion range(r)
\tdx{image_def}:    r  A      == {\ttlbrace}y\isasymin{}range(r) . {\isasymexists}x \isasymin A. <x,y> \isasymin r{\ttrbrace}
\tdx{vimage_def}:   r - A     == converse(r)A
\subcaption{Operations on relations}

\tdx{lam_def}:   Lambda(A,b) == {\ttlbrace}<x,b(x)> . x \isasymin A{\ttrbrace}
\tdx{apply_def}: fa         == THE y. <a,y> \isasymin f
\tdx{Pi_def}: Pi(A,B) == {\ttlbrace}f\isasymin{}Pow(Sigma(A,B)). {\isasymforall}x\isasymin{}A. \isasymexists!y. <x,y>\isasymin{}f{\ttrbrace}
\tdx{restrict_def}:  restrict(f,A) == lam x \isasymin A. fx
\subcaption{Functions and general product}
\end{alltt*}
\caption{Further definitions of ZF} \label{zf-defs}
\end{figure}

\section{The Zermelo-Fraenkel axioms}
The axioms appear in Fig.\ts \ref{zf-rules}.  They resemble those
presented by Suppes~\cite{suppes72}.  Most of the theory consists of
definitions.  In particular, bounded quantifiers and the subset relation
appear in other axioms.  Object-level quantifiers and implications have
been replaced by meta-level ones wherever possible, to simplify use of the
axioms.

$y \in \isa{PrimReplace}(A,P) \bimp (\exists x\in A. P(x,y))$
subject to the condition that $P(x,y)$ is single-valued for all~$x\in A$.
The Isabelle theory defines \cdx{Replace} to apply
\cdx{PrimReplace} to the single-valued part of~$P$, namely
$(\exists!z. P(x,z)) \conj P(x,y).$
Thus $y\in \isa{Replace}(A,P)$ if and only if there is some~$x$ such that
$P(x,-)$ holds uniquely for~$y$.  Because the equivalence is unconditional,
\isa{Replace} is much easier to use than \isa{PrimReplace}; it defines the
same set, if $P(x,y)$ is single-valued.  The nice syntax for replacement
expands to \isa{Replace}.

Other consequences of replacement include replacement for
meta-level functions
(\cdx{RepFun}) and definite descriptions (\cdx{The}).
Axioms for separation (\cdx{Collect}) and unordered pairs
(\cdx{Upair}) are traditionally assumed, but they actually follow
from replacement~\cite[pages 237--8]{suppes72}.

The definitions of general intersection, etc., are straightforward.  Note
the definition of \isa{cons}, which underlies the finite set notation.
The axiom of infinity gives us a set that contains~0 and is closed under
successor (\cdx{succ}).  Although this set is not uniquely defined,
the theory names it (\cdx{Inf}) in order to simplify the
construction of the natural numbers.

Further definitions appear in Fig.\ts\ref{zf-defs}.  Ordered pairs are
defined in the standard way, $\pair{a,b}\equiv\{\{a\},\{a,b\}\}$.  Recall
that \cdx{Sigma}$(A,B)$ generalizes the Cartesian product of two
sets.  It is defined to be the union of all singleton sets
$\{\pair{x,y}\}$, for $x\in A$ and $y\in B(x)$.  This is a typical usage of
general union.

The projections \cdx{fst} and~\cdx{snd} are defined in terms of the
generalized projection \cdx{split}.  The latter has been borrowed from
Martin-L\"of's Type Theory, and is often easier to use than \cdx{fst}
and~\cdx{snd}.

Operations on relations include converse, domain, range, and image.  The
set $\isa{Pi}(A,B)$ generalizes the space of functions between two sets.
Note the simple definitions of $\lambda$-abstraction (using
\cdx{RepFun}) and application (using a definite description).  The
function \cdx{restrict}$(f,A)$ has the same values as~$f$, but only
over the domain~$A$.

%%%% zf.thy

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{ballI}:     [| !!x. x\isasymin{}A ==> P(x) |] ==> {\isasymforall}x\isasymin{}A. P(x)
\tdx{bspec}:     [| {\isasymforall}x\isasymin{}A. P(x);  x\isasymin{}A |] ==> P(x)
\tdx{ballE}:     [| {\isasymforall}x\isasymin{}A. P(x);  P(x) ==> Q;  x \isasymnotin A ==> Q |] ==> Q

\tdx{ball_cong}:  [| A=A';  !!x. x\isasymin{}A' ==> P(x) <-> P'(x) |] ==>
({\isasymforall}x\isasymin{}A. P(x)) <-> ({\isasymforall}x\isasymin{}A'. P'(x))

\tdx{bexI}:      [| P(x);  x\isasymin{}A |] ==> {\isasymexists}x\isasymin{}A. P(x)
\tdx{bexCI}:     [| {\isasymforall}x\isasymin{}A. ~P(x) ==> P(a);  a\isasymin{}A |] ==> {\isasymexists}x\isasymin{}A. P(x)
\tdx{bexE}:      [| {\isasymexists}x\isasymin{}A. P(x);  !!x. [| x\isasymin{}A; P(x) |] ==> Q |] ==> Q

\tdx{bex_cong}:  [| A=A';  !!x. x\isasymin{}A' ==> P(x) <-> P'(x) |] ==>
({\isasymexists}x\isasymin{}A. P(x)) <-> ({\isasymexists}x\isasymin{}A'. P'(x))
\subcaption{Bounded quantifiers}

\tdx{subsetI}:     (!!x. x \isasymin A ==> x \isasymin B) ==> A \isasymsubseteq B
\tdx{subsetD}:     [| A \isasymsubseteq B;  c \isasymin A |] ==> c \isasymin B
\tdx{subsetCE}:    [| A \isasymsubseteq B;  c \isasymnotin A ==> P;  c \isasymin B ==> P |] ==> P
\tdx{subset_refl}:  A \isasymsubseteq A
\tdx{subset_trans}: [| A \isasymsubseteq B;  B \isasymsubseteq C |] ==> A \isasymsubseteq C

\tdx{equalityI}:   [| A \isasymsubseteq B;  B \isasymsubseteq A |] ==> A = B
\tdx{equalityD1}:  A = B ==> A \isasymsubseteq B
\tdx{equalityD2}:  A = B ==> B \isasymsubseteq A
\tdx{equalityE}:   [| A = B;  [| A \isasymsubseteq B; B \isasymsubseteq A |] ==> P |]  ==>  P
\subcaption{Subsets and extensionality}

\tdx{emptyE}:        a \isasymin 0 ==> P
\tdx{empty_subsetI}:  0 \isasymsubseteq A
\tdx{equals0I}:      [| !!y. y \isasymin A ==> False |] ==> A=0
\tdx{equals0D}:      [| A=0;  a \isasymin A |] ==> P

\tdx{PowI}:          A \isasymsubseteq B ==> A \isasymin Pow(B)
\tdx{PowD}:          A \isasymin Pow(B)  ==>  A \isasymsubseteq B
\subcaption{The empty set; power sets}
\end{alltt*}
\caption{Basic derived rules for ZF} \label{zf-lemmas1}
\end{figure}

\section{From basic lemmas to function spaces}
Faced with so many definitions, it is essential to prove lemmas.  Even
trivial theorems like $A \int B = B \int A$ would be difficult to
prove from the definitions alone.  Isabelle's set theory derives many
rules using a natural deduction style.  Ideally, a natural deduction
rule should introduce or eliminate just one operator, but this is not
always practical.  For most operators, we may forget its definition
and use its derived rules instead.

\subsection{Fundamental lemmas}
Figure~\ref{zf-lemmas1} presents the derived rules for the most basic
operators.  The rules for the bounded quantifiers resemble those for the
ordinary quantifiers, but note that \tdx{ballE} uses a negated assumption
in the style of Isabelle's classical reasoner.  The \rmindex{congruence
rules} \tdx{ball_cong} and \tdx{bex_cong} are required by Isabelle's
simplifier, but have few other uses.  Congruence rules must be specially
derived for all binding operators, and henceforth will not be shown.

Figure~\ref{zf-lemmas1} also shows rules for the subset and equality
relations (proof by extensionality), and rules about the empty set and the
power set operator.

Figure~\ref{zf-lemmas2} presents rules for replacement and separation.
The rules for \cdx{Replace} and \cdx{RepFun} are much simpler than
comparable rules for \isa{PrimReplace} would be.  The principle of
separation is proved explicitly, although most proofs should use the
natural deduction rules for \isa{Collect}.  The elimination rule
\tdx{CollectE} is equivalent to the two destruction rules
\tdx{CollectD1} and \tdx{CollectD2}, but each rule is suited to
particular circumstances.  Although too many rules can be confusing, there
is no reason to aim for a minimal set of rules.

Figure~\ref{zf-lemmas3} presents rules for general union and intersection.
The empty intersection should be undefined.  We cannot have
$\bigcap(\emptyset)=V$ because $V$, the universal class, is not a set.  All
expressions denote something in ZF set theory; the definition of
intersection implies $\bigcap(\emptyset)=\emptyset$, but this value is
arbitrary.  The rule \tdx{InterI} must have a premise to exclude
the empty intersection.  Some of the laws governing intersections require
similar premises.

%the [p] gives better page breaking for the book
\begin{figure}[p]
\begin{alltt*}\isastyleminor
\tdx{ReplaceI}:   [| x\isasymin{}A;  P(x,b);  !!y. P(x,y) ==> y=b |] ==>
b\isasymin{}{\ttlbrace}y. x\isasymin{}A, P(x,y){\ttrbrace}

\tdx{ReplaceE}:   [| b\isasymin{}{\ttlbrace}y. x\isasymin{}A, P(x,y){\ttrbrace};
!!x. [| x\isasymin{}A; P(x,b); {\isasymforall}y. P(x,y)-->y=b |] ==> R
|] ==> R

\tdx{RepFunI}:    [| a\isasymin{}A |] ==> f(a)\isasymin{}{\ttlbrace}f(x). x\isasymin{}A{\ttrbrace}
\tdx{RepFunE}:    [| b\isasymin{}{\ttlbrace}f(x). x\isasymin{}A{\ttrbrace};
!!x.[| x\isasymin{}A;  b=f(x) |] ==> P |] ==> P

\tdx{separation}:  a\isasymin{}{\ttlbrace}x\isasymin{}A. P(x){\ttrbrace} <-> a\isasymin{}A & P(a)
\tdx{CollectI}:    [| a\isasymin{}A;  P(a) |] ==> a\isasymin{}{\ttlbrace}x\isasymin{}A. P(x){\ttrbrace}
\tdx{CollectE}:    [| a\isasymin{}{\ttlbrace}x\isasymin{}A. P(x){\ttrbrace};  [| a\isasymin{}A; P(a) |] ==> R |] ==> R
\tdx{CollectD1}:   a\isasymin{}{\ttlbrace}x\isasymin{}A. P(x){\ttrbrace} ==> a\isasymin{}A
\tdx{CollectD2}:   a\isasymin{}{\ttlbrace}x\isasymin{}A. P(x){\ttrbrace} ==> P(a)
\end{alltt*}
\caption{Replacement and separation} \label{zf-lemmas2}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{UnionI}: [| B\isasymin{}C;  A\isasymin{}B |] ==> A\isasymin{}Union(C)
\tdx{UnionE}: [| A\isasymin{}Union(C);  !!B.[| A\isasymin{}B;  B\isasymin{}C |] ==> R |] ==> R

\tdx{InterI}: [| !!x. x\isasymin{}C ==> A\isasymin{}x;  c\isasymin{}C |] ==> A\isasymin{}Inter(C)
\tdx{InterD}: [| A\isasymin{}Inter(C);  B\isasymin{}C |] ==> A\isasymin{}B
\tdx{InterE}: [| A\isasymin{}Inter(C);  A\isasymin{}B ==> R;  B \isasymnotin C ==> R |] ==> R

\tdx{UN_I}:   [| a\isasymin{}A;  b\isasymin{}B(a) |] ==> b\isasymin{}({\isasymUnion}x\isasymin{}A. B(x))
\tdx{UN_E}:   [| b\isasymin{}({\isasymUnion}x\isasymin{}A. B(x));  !!x.[| x\isasymin{}A;  b\isasymin{}B(x) |] ==> R
|] ==> R

\tdx{INT_I}:  [| !!x. x\isasymin{}A ==> b\isasymin{}B(x);  a\isasymin{}A |] ==> b\isasymin{}({\isasymInter}x\isasymin{}A. B(x))
\tdx{INT_E}:  [| b\isasymin{}({\isasymInter}x\isasymin{}A. B(x));  a\isasymin{}A |] ==> b\isasymin{}B(a)
\end{alltt*}
\caption{General union and intersection} \label{zf-lemmas3}
\end{figure}

%%% upair.thy

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{pairing}:   a\isasymin{}Upair(b,c) <-> (a=b | a=c)
\tdx{UpairI1}:   a\isasymin{}Upair(a,b)
\tdx{UpairI2}:   b\isasymin{}Upair(a,b)
\tdx{UpairE}:    [| a\isasymin{}Upair(b,c);  a=b ==> P;  a=c ==> P |] ==> P
\end{alltt*}
\caption{Unordered pairs} \label{zf-upair1}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{UnI1}:      c\isasymin{}A ==> c\isasymin{}A \isasymunion B
\tdx{UnI2}:      c\isasymin{}B ==> c\isasymin{}A \isasymunion B
\tdx{UnCI}:      (c \isasymnotin B ==> c\isasymin{}A) ==> c\isasymin{}A \isasymunion B
\tdx{UnE}:       [| c\isasymin{}A \isasymunion B;  c\isasymin{}A ==> P;  c\isasymin{}B ==> P |] ==> P

\tdx{IntI}:      [| c\isasymin{}A;  c\isasymin{}B |] ==> c\isasymin{}A \isasyminter B
\tdx{IntD1}:     c\isasymin{}A \isasyminter B ==> c\isasymin{}A
\tdx{IntD2}:     c\isasymin{}A \isasyminter B ==> c\isasymin{}B
\tdx{IntE}:      [| c\isasymin{}A \isasyminter B;  [| c\isasymin{}A; c\isasymin{}B |] ==> P |] ==> P

\tdx{DiffI}:     [| c\isasymin{}A;  c \isasymnotin B |] ==> c\isasymin{}A - B
\tdx{DiffD1}:    c\isasymin{}A - B ==> c\isasymin{}A
\tdx{DiffD2}:    c\isasymin{}A - B ==> c  \isasymnotin  B
\tdx{DiffE}:     [| c\isasymin{}A - B;  [| c\isasymin{}A; c \isasymnotin B |] ==> P |] ==> P
\end{alltt*}
\caption{Union, intersection, difference} \label{zf-Un}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{consI1}:    a\isasymin{}cons(a,B)
\tdx{consI2}:    a\isasymin{}B ==> a\isasymin{}cons(b,B)
\tdx{consCI}:    (a \isasymnotin B ==> a=b) ==> a\isasymin{}cons(b,B)
\tdx{consE}:     [| a\isasymin{}cons(b,A);  a=b ==> P;  a\isasymin{}A ==> P |] ==> P

\tdx{singletonI}:  a\isasymin{}{\ttlbrace}a{\ttrbrace}
\tdx{singletonE}:  [| a\isasymin{}{\ttlbrace}b{\ttrbrace}; a=b ==> P |] ==> P
\end{alltt*}
\caption{Finite and singleton sets} \label{zf-upair2}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{succI1}:    i\isasymin{}succ(i)
\tdx{succI2}:    i\isasymin{}j ==> i\isasymin{}succ(j)
\tdx{succCI}:    (i \isasymnotin j ==> i=j) ==> i\isasymin{}succ(j)
\tdx{succE}:     [| i\isasymin{}succ(j);  i=j ==> P;  i\isasymin{}j ==> P |] ==> P
\tdx{succ_neq_0}:  [| succ(n)=0 |] ==> P
\tdx{succ_inject}: succ(m) = succ(n) ==> m=n
\end{alltt*}
\caption{The successor function} \label{zf-succ}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{the_equality}: [| P(a); !!x. P(x) ==> x=a |] ==> (THE x. P(x))=a
\tdx{theI}:         \isasymexists! x. P(x) ==> P(THE x. P(x))

\tdx{if_P}:          P ==> (if P then a else b) = a
\tdx{if_not_P}:     ~P ==> (if P then a else b) = b

\tdx{mem_asym}:     [| a\isasymin{}b;  b\isasymin{}a |] ==> P
\tdx{mem_irrefl}:   a\isasymin{}a ==> P
\end{alltt*}
\caption{Descriptions; non-circularity} \label{zf-the}
\end{figure}

\subsection{Unordered pairs and finite sets}
Figure~\ref{zf-upair1} presents the principle of unordered pairing, along
with its derived rules.  Binary union and intersection are defined in terms
of ordered pairs (Fig.\ts\ref{zf-Un}).  Set difference is also included.  The
rule \tdx{UnCI} is useful for classical reasoning about unions,
like \isa{disjCI}\@; it supersedes \tdx{UnI1} and
\tdx{UnI2}, but these rules are often easier to work with.  For
intersection and difference we have both elimination and destruction rules.
Again, there is no reason to provide a minimal rule set.

Figure~\ref{zf-upair2} is concerned with finite sets: it presents rules
for~\isa{cons}, the finite set constructor, and rules for singleton
sets.  Figure~\ref{zf-succ} presents derived rules for the successor
function, which is defined in terms of~\isa{cons}.  The proof that
\isa{succ} is injective appears to require the Axiom of Foundation.

Definite descriptions (\sdx{THE}) are defined in terms of the singleton
set~$\{0\}$, but their derived rules fortunately hide this
(Fig.\ts\ref{zf-the}).  The rule~\tdx{theI} is difficult to apply
because of the two occurrences of~$\Var{P}$.  However,
\tdx{the_equality} does not have this problem and the files contain
many examples of its use.

Finally, the impossibility of having both $a\in b$ and $b\in a$
(\tdx{mem_asym}) is proved by applying the Axiom of Foundation to
the set $\{a,b\}$.  The impossibility of $a\in a$ is a trivial consequence.

%%% subset.thy?

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{Union_upper}:    B\isasymin{}A ==> B \isasymsubseteq Union(A)
\tdx{Union_least}:    [| !!x. x\isasymin{}A ==> x \isasymsubseteq C |] ==> Union(A) \isasymsubseteq C

\tdx{Inter_lower}:    B\isasymin{}A ==> Inter(A) \isasymsubseteq B
\tdx{Inter_greatest}: [| a\isasymin{}A; !!x. x\isasymin{}A ==> C \isasymsubseteq x |] ==> C\isasymsubseteq{}Inter(A)

\tdx{Un_upper1}:      A \isasymsubseteq A \isasymunion B
\tdx{Un_upper2}:      B \isasymsubseteq A \isasymunion B
\tdx{Un_least}:       [| A \isasymsubseteq C;  B \isasymsubseteq C |] ==> A \isasymunion B \isasymsubseteq C

\tdx{Int_lower1}:     A \isasyminter B \isasymsubseteq A
\tdx{Int_lower2}:     A \isasyminter B \isasymsubseteq B
\tdx{Int_greatest}:   [| C \isasymsubseteq A;  C \isasymsubseteq B |] ==> C \isasymsubseteq A \isasyminter B

\tdx{Diff_subset}:    A-B \isasymsubseteq A
\tdx{Diff_contains}:  [| C \isasymsubseteq A;  C \isasyminter B = 0 |] ==> C \isasymsubseteq A-B

\tdx{Collect_subset}: Collect(A,P) \isasymsubseteq A
\end{alltt*}
\caption{Subset and lattice properties} \label{zf-subset}
\end{figure}

\subsection{Subset and lattice properties}
The subset relation is a complete lattice.  Unions form least upper bounds;
non-empty intersections form greatest lower bounds.  Figure~\ref{zf-subset}
shows the corresponding rules.  A few other laws involving subsets are
included.
Reasoning directly about subsets often yields clearer proofs than
reasoning about the membership relation.  Section~\ref{sec:ZF-pow-example}
below presents an example of this, proving the equation
${\isa{Pow}(A)\cap \isa{Pow}(B)}= \isa{Pow}(A\cap B)$.

%%% pair.thy

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{Pair_inject1}: <a,b> = <c,d> ==> a=c
\tdx{Pair_inject2}: <a,b> = <c,d> ==> b=d
\tdx{Pair_inject}:  [| <a,b> = <c,d>;  [| a=c; b=d |] ==> P |] ==> P
\tdx{Pair_neq_0}:   <a,b>=0 ==> P

\tdx{fst_conv}:     fst(<a,b>) = a
\tdx{snd_conv}:     snd(<a,b>) = b
\tdx{split}:        split(\%x y. c(x,y), <a,b>) = c(a,b)

\tdx{SigmaI}:     [| a\isasymin{}A;  b\isasymin{}B(a) |] ==> <a,b>\isasymin{}Sigma(A,B)

\tdx{SigmaE}:     [| c\isasymin{}Sigma(A,B);
!!x y.[| x\isasymin{}A; y\isasymin{}B(x); c=<x,y> |] ==> P |] ==> P

\tdx{SigmaE2}:    [| <a,b>\isasymin{}Sigma(A,B);
[| a\isasymin{}A;  b\isasymin{}B(a) |] ==> P   |] ==> P
\end{alltt*}
\caption{Ordered pairs; projections; general sums} \label{zf-pair}
\end{figure}

\subsection{Ordered pairs} \label{sec:pairs}

Figure~\ref{zf-pair} presents the rules governing ordered pairs,
projections and general sums --- in particular, that
$\{\{a\},\{a,b\}\}$ functions as an ordered pair.  This property is
expressed as two destruction rules,
\tdx{Pair_inject1} and \tdx{Pair_inject2}, and equivalently
as the elimination rule \tdx{Pair_inject}.

The rule \tdx{Pair_neq_0} asserts $\pair{a,b}\neq\emptyset$.  This
is a property of $\{\{a\},\{a,b\}\}$, and need not hold for other
encodings of ordered pairs.  The non-standard ordered pairs mentioned below
satisfy $\pair{\emptyset;\emptyset}=\emptyset$.

The natural deduction rules \tdx{SigmaI} and \tdx{SigmaE}
assert that \cdx{Sigma}$(A,B)$ consists of all pairs of the form
$\pair{x,y}$, for $x\in A$ and $y\in B(x)$.  The rule \tdx{SigmaE2}
merely states that $\pair{a,b}\in \isa{Sigma}(A,B)$ implies $a\in A$ and
$b\in B(a)$.

In addition, it is possible to use tuples as patterns in abstractions:
\begin{center}
{\tt\%<$x$,$y$>. $t$} \quad stands for\quad \isa{split(\%$x$ $y$.\ $t$)}
\end{center}
Nested patterns are translated recursively:
{\tt\%<$x$,$y$,$z$>. $t$} $\leadsto$ {\tt\%<$x$,<$y$,$z$>>. $t$} $\leadsto$
\isa{split(\%$x$.\%<$y$,$z$>. $t$)} $\leadsto$ \isa{split(\%$x$. split(\%$y$
$z$.\ $t$))}.  The reverse translation is performed upon printing.
\begin{warn}
The translation between patterns and \isa{split} is performed automatically
by the parser and printer.  Thus the internal and external form of a term
may differ, which affects proofs.  For example the term \isa{(\%<x,y>.<y,x>)<a,b>} requires the theorem \isa{split} to rewrite to
{\tt<b,a>}.
\end{warn}
In addition to explicit $\lambda$-abstractions, patterns can be used in any
variable binding construct which is internally described by a
$\lambda$-abstraction.  Here are some important examples:
\begin{description}
\item[Let:] \isa{let {\it pattern} = $t$ in $u$}
\item[Choice:] \isa{THE~{\it pattern}~.~$P$}
\item[Set operations:] \isa{\isasymUnion~{\it pattern}:$A$.~$B$}
\item[Comprehension:] \isa{{\ttlbrace}~{\it pattern}:$A$~.~$P$~{\ttrbrace}}
\end{description}

%%% domrange.thy?

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{domainI}:     <a,b>\isasymin{}r ==> a\isasymin{}domain(r)
\tdx{domainE}:     [| a\isasymin{}domain(r); !!y. <a,y>\isasymin{}r ==> P |] ==> P
\tdx{domain_subset}: domain(Sigma(A,B)) \isasymsubseteq A

\tdx{rangeI}:      <a,b>\isasymin{}r ==> b\isasymin{}range(r)
\tdx{rangeE}:      [| b\isasymin{}range(r); !!x. <x,b>\isasymin{}r ==> P |] ==> P
\tdx{range_subset}: range(A*B) \isasymsubseteq B

\tdx{fieldI1}:     <a,b>\isasymin{}r ==> a\isasymin{}field(r)
\tdx{fieldI2}:     <a,b>\isasymin{}r ==> b\isasymin{}field(r)
\tdx{fieldCI}:     (<c,a> \isasymnotin r ==> <a,b>\isasymin{}r) ==> a\isasymin{}field(r)

\tdx{fieldE}:      [| a\isasymin{}field(r);
!!x. <a,x>\isasymin{}r ==> P;
!!x. <x,a>\isasymin{}r ==> P
|] ==> P

\tdx{field_subset}:  field(A*A) \isasymsubseteq A
\end{alltt*}
\caption{Domain, range and field of a relation} \label{zf-domrange}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{imageI}:      [| <a,b>\isasymin{}r; a\isasymin{}A |] ==> b\isasymin{}rA
\tdx{imageE}:      [| b\isasymin{}rA; !!x.[| <x,b>\isasymin{}r; x\isasymin{}A |] ==> P |] ==> P

\tdx{vimageI}:     [| <a,b>\isasymin{}r; b\isasymin{}B |] ==> a\isasymin{}r-B
\tdx{vimageE}:     [| a\isasymin{}r-B; !!x.[| <a,x>\isasymin{}r;  x\isasymin{}B |] ==> P |] ==> P
\end{alltt*}
\caption{Image and inverse image} \label{zf-domrange2}
\end{figure}

\subsection{Relations}
Figure~\ref{zf-domrange} presents rules involving relations, which are sets
of ordered pairs.  The converse of a relation~$r$ is the set of all pairs
$\pair{y,x}$ such that $\pair{x,y}\in r$; if $r$ is a function, then
{\cdx{converse}$(r)$} is its inverse.  The rules for the domain
operation, namely \tdx{domainI} and~\tdx{domainE}, assert that
\cdx{domain}$(r)$ consists of all~$x$ such that $r$ contains
some pair of the form~$\pair{x,y}$.  The range operation is similar, and
the field of a relation is merely the union of its domain and range.

Figure~\ref{zf-domrange2} presents rules for images and inverse images.
Note that these operations are generalisations of range and domain,
respectively.

%%% func.thy

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{fun_is_rel}:     f\isasymin{}Pi(A,B) ==> f \isasymsubseteq Sigma(A,B)

\tdx{apply_equality}:  [| <a,b>\isasymin{}f; f\isasymin{}Pi(A,B) |] ==> fa = b
\tdx{apply_equality2}: [| <a,b>\isasymin{}f; <a,c>\isasymin{}f; f\isasymin{}Pi(A,B) |] ==> b=c

\tdx{apply_type}:     [| f\isasymin{}Pi(A,B); a\isasymin{}A |] ==> fa\isasymin{}B(a)
\tdx{apply_Pair}:     [| f\isasymin{}Pi(A,B); a\isasymin{}A |] ==> <a,fa>\isasymin{}f
\tdx{apply_iff}:      f\isasymin{}Pi(A,B) ==> <a,b>\isasymin{}f <-> a\isasymin{}A & fa = b

\tdx{fun_extension}:  [| f\isasymin{}Pi(A,B); g\isasymin{}Pi(A,D);
!!x. x\isasymin{}A ==> fx = gx     |] ==> f=g

\tdx{domain_type}:    [| <a,b>\isasymin{}f; f\isasymin{}Pi(A,B) |] ==> a\isasymin{}A
\tdx{range_type}:     [| <a,b>\isasymin{}f; f\isasymin{}Pi(A,B) |] ==> b\isasymin{}B(a)

\tdx{Pi_type}:        [| f\isasymin{}A->C; !!x. x\isasymin{}A ==> fx\isasymin{}B(x) |] ==> f\isasymin{}Pi(A,B)
\tdx{domain_of_fun}:  f\isasymin{}Pi(A,B) ==> domain(f)=A
\tdx{range_of_fun}:   f\isasymin{}Pi(A,B) ==> f\isasymin{}A->range(f)

\tdx{restrict}:       a\isasymin{}A ==> restrict(f,A)  a = fa
\tdx{restrict_type}:  [| !!x. x\isasymin{}A ==> fx\isasymin{}B(x) |] ==>
restrict(f,A)\isasymin{}Pi(A,B)
\end{alltt*}
\caption{Functions} \label{zf-func1}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{lamI}:     a\isasymin{}A ==> <a,b(a)>\isasymin{}(lam x\isasymin{}A. b(x))
\tdx{lamE}:     [| p\isasymin{}(lam x\isasymin{}A. b(x)); !!x.[| x\isasymin{}A; p=<x,b(x)> |] ==> P
|] ==>  P

\tdx{lam_type}: [| !!x. x\isasymin{}A ==> b(x)\isasymin{}B(x) |] ==> (lam x\isasymin{}A. b(x))\isasymin{}Pi(A,B)

\tdx{beta}:     a\isasymin{}A ==> (lam x\isasymin{}A. b(x))  a = b(a)
\tdx{eta}:      f\isasymin{}Pi(A,B) ==> (lam x\isasymin{}A. fx) = f
\end{alltt*}
\caption{$\lambda$-abstraction} \label{zf-lam}
\end{figure}

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{fun_empty}:           0\isasymin{}0->0
\tdx{fun_single}:          {\ttlbrace}<a,b>{\ttrbrace}\isasymin{}{\ttlbrace}a{\ttrbrace} -> {\ttlbrace}b{\ttrbrace}

\tdx{fun_disjoint_Un}:     [| f\isasymin{}A->B; g\isasymin{}C->D; A \isasyminter C = 0  |] ==>
(f \isasymunion g)\isasymin{}(A \isasymunion C) -> (B \isasymunion D)

\tdx{fun_disjoint_apply1}: [| a\isasymin{}A; f\isasymin{}A->B; g\isasymin{}C->D;  A\isasyminter{}C = 0 |] ==>
(f \isasymunion g)a = fa

\tdx{fun_disjoint_apply2}: [| c\isasymin{}C; f\isasymin{}A->B; g\isasymin{}C->D;  A\isasyminter{}C = 0 |] ==>
(f \isasymunion g)c = gc
\end{alltt*}
\caption{Constructing functions from smaller sets} \label{zf-func2}
\end{figure}

\subsection{Functions}
Functions, represented by graphs, are notoriously difficult to reason
about.  The ZF theory provides many derived rules, which overlap more
than they ought.  This section presents the more important rules.

Figure~\ref{zf-func1} presents the basic properties of \cdx{Pi}$(A,B)$,
the generalized function space.  For example, if $f$ is a function and
$\pair{a,b}\in f$, then $fa=b$ (\tdx{apply_equality}).  Two functions
are equal provided they have equal domains and deliver equals results
(\tdx{fun_extension}).

By \tdx{Pi_type}, a function typing of the form $f\in A\to C$ can be
refined to the dependent typing $f\in\prod@{x\in A}B(x)$, given a suitable
family of sets $\{B(x)\}@{x\in A}$.  Conversely, by \tdx{range_of_fun},
any dependent typing can be flattened to yield a function type of the form
$A\to C$; here, $C=\isa{range}(f)$.

Among the laws for $\lambda$-abstraction, \tdx{lamI} and \tdx{lamE}
describe the graph of the generated function, while \tdx{beta} and
\tdx{eta} are the standard conversions.  We essentially have a
dependently-typed $\lambda$-calculus (Fig.\ts\ref{zf-lam}).

Figure~\ref{zf-func2} presents some rules that can be used to construct
pair, and may form the union of two functions provided their domains are
disjoint.

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{Int_absorb}:        A \isasyminter A = A
\tdx{Int_commute}:       A \isasyminter B = B \isasyminter A
\tdx{Int_assoc}:         (A \isasyminter B) \isasyminter C  =  A \isasyminter (B \isasyminter C)
\tdx{Int_Un_distrib}:    (A \isasymunion B) \isasyminter C  =  (A \isasyminter C) \isasymunion (B \isasyminter C)

\tdx{Un_absorb}:         A \isasymunion A = A
\tdx{Un_commute}:        A \isasymunion B = B \isasymunion A
\tdx{Un_assoc}:          (A \isasymunion B) \isasymunion C  =  A \isasymunion (B \isasymunion C)
\tdx{Un_Int_distrib}:    (A \isasyminter B) \isasymunion C  =  (A \isasymunion C) \isasyminter (B \isasymunion C)

\tdx{Diff_cancel}:       A-A = 0
\tdx{Diff_disjoint}:     A \isasyminter (B-A) = 0
\tdx{Diff_partition}:    A \isasymsubseteq B ==> A \isasymunion (B-A) = B
\tdx{double_complement}: [| A \isasymsubseteq B; B \isasymsubseteq C |] ==> (B - (C-A)) = A
\tdx{Diff_Un}:           A - (B \isasymunion C) = (A-B) \isasyminter (A-C)
\tdx{Diff_Int}:          A - (B \isasyminter C) = (A-B) \isasymunion (A-C)

\tdx{Union_Un_distrib}:  Union(A \isasymunion B) = Union(A) \isasymunion Union(B)
\tdx{Inter_Un_distrib}:  [| a \isasymin A;  b \isasymin B |] ==>
Inter(A \isasymunion B) = Inter(A) \isasyminter Inter(B)

\tdx{Int_Union_RepFun}:  A \isasyminter Union(B) = ({\isasymUnion}C \isasymin B. A \isasyminter C)

\tdx{Un_Inter_RepFun}:   b \isasymin B ==>
A \isasymunion Inter(B) = ({\isasymInter}C \isasymin B. A \isasymunion C)

\tdx{SUM_Un_distrib1}:   (SUM x \isasymin A \isasymunion B. C(x)) =
(SUM x \isasymin A. C(x)) \isasymunion (SUM x \isasymin B. C(x))

\tdx{SUM_Un_distrib2}:   (SUM x \isasymin C. A(x) \isasymunion B(x)) =
(SUM x \isasymin C. A(x)) \isasymunion (SUM x \isasymin C. B(x))

\tdx{SUM_Int_distrib1}:  (SUM x \isasymin A \isasyminter B. C(x)) =
(SUM x \isasymin A. C(x)) \isasyminter (SUM x \isasymin B. C(x))

\tdx{SUM_Int_distrib2}:  (SUM x \isasymin C. A(x) \isasyminter B(x)) =
(SUM x \isasymin C. A(x)) \isasyminter (SUM x \isasymin C. B(x))
\end{alltt*}
\caption{Equalities} \label{zf-equalities}
\end{figure}

\begin{figure}
%\begin{constants}
%  \cdx{1}       & $i$           &       & $\{\emptyset\}$       \\
%  \cdx{bool}    & $i$           &       & the set $\{\emptyset,1\}$     \\
%  \cdx{cond}   & $[i,i,i]\To i$ &       & conditional for \isa{bool}    \\
%  \cdx{not}    & $i\To i$       &       & negation for \isa{bool}       \\
%  \sdx{and}    & $[i,i]\To i$   & Left 70 & conjunction for \isa{bool}  \\
%  \sdx{or}     & $[i,i]\To i$   & Left 65 & disjunction for \isa{bool}  \\
%  \sdx{xor}    & $[i,i]\To i$   & Left 65 & exclusive-or for \isa{bool}
%\end{constants}
%
\begin{alltt*}\isastyleminor
\tdx{bool_def}:      bool == {\ttlbrace}0,1{\ttrbrace}
\tdx{cond_def}:      cond(b,c,d) == if b=1 then c else d
\tdx{not_def}:       not(b)  == cond(b,0,1)
\tdx{and_def}:       a and b == cond(a,b,0)
\tdx{or_def}:        a or b  == cond(a,1,b)
\tdx{xor_def}:       a xor b == cond(a,not(b),b)

\tdx{bool_1I}:       1 \isasymin bool
\tdx{bool_0I}:       0 \isasymin bool
\tdx{boolE}:         [| c \isasymin bool;  c=1 ==> P;  c=0 ==> P |] ==> P
\tdx{cond_1}:        cond(1,c,d) = c
\tdx{cond_0}:        cond(0,c,d) = d
\end{alltt*}
\caption{The booleans} \label{zf-bool}
\end{figure}

\section{Further developments}
The next group of developments is complex and extensive, and only
highlights can be covered here.  It involves many theories and proofs.

Figure~\ref{zf-equalities} presents commutative, associative, distributive,
and idempotency laws of union and intersection, along with other equations.

Theory \thydx{Bool} defines $\{0,1\}$ as a set of booleans, with the usual
operators including a conditional (Fig.\ts\ref{zf-bool}).  Although ZF is a
first-order theory, you can obtain the effect of higher-order logic using
\isa{bool}-valued functions, for example.  The constant~\isa{1} is
translated to \isa{succ(0)}.

\begin{figure}
\index{*"+ symbol}
\begin{constants}
\it symbol    & \it meta-type & \it priority & \it description \\
\tt +         & $[i,i]\To i$  &  Right 65     & disjoint union operator\\
\cdx{Inl}~~\cdx{Inr}  & $i\To i$      &       & injections\\
\cdx{case}    & $[i\To i,i\To i, i]\To i$ &   & conditional for $A+B$
\end{constants}
\begin{alltt*}\isastyleminor
\tdx{sum_def}:   A+B == {\ttlbrace}0{\ttrbrace}*A \isasymunion {\ttlbrace}1{\ttrbrace}*B
\tdx{Inl_def}:   Inl(a) == <0,a>
\tdx{Inr_def}:   Inr(b) == <1,b>
\tdx{case_def}:  case(c,d,u) == split(\%y z. cond(y, d(z), c(z)), u)

\tdx{InlI}:      a \isasymin A ==> Inl(a) \isasymin A+B
\tdx{InrI}:      b \isasymin B ==> Inr(b) \isasymin A+B

\tdx{Inl_inject}:  Inl(a)=Inl(b) ==> a=b
\tdx{Inr_inject}:  Inr(a)=Inr(b) ==> a=b
\tdx{Inl_neq_Inr}: Inl(a)=Inr(b) ==> P

\tdx{sum_iff}:  u \isasymin A+B <-> ({\isasymexists}x\isasymin{}A. u=Inl(x)) | ({\isasymexists}y\isasymin{}B. u=Inr(y))

\tdx{case_Inl}:  case(c,d,Inl(a)) = c(a)
\tdx{case_Inr}:  case(c,d,Inr(b)) = d(b)
\end{alltt*}
\caption{Disjoint unions} \label{zf-sum}
\end{figure}

\subsection{Disjoint unions}

Theory \thydx{Sum} defines the disjoint union of two sets, with
injections and a case analysis operator (Fig.\ts\ref{zf-sum}).  Disjoint
unions play a role in datatype definitions, particularly when there is
mutual recursion~\cite{paulson-set-II}.

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{QPair_def}:      <a;b> == a+b
\tdx{qsplit_def}:     qsplit(c,p)  == THE y. {\isasymexists}a b. p=<a;b> & y=c(a,b)
\tdx{qfsplit_def}:    qfsplit(R,z) == {\isasymexists}x y. z=<x;y> & R(x,y)
\tdx{qconverse_def}:  qconverse(r) == {\ttlbrace}z. w \isasymin r, {\isasymexists}x y. w=<x;y> & z=<y;x>{\ttrbrace}
\tdx{QSigma_def}:     QSigma(A,B)  == {\isasymUnion}x \isasymin A. {\isasymUnion}y \isasymin B(x). {\ttlbrace}<x;y>{\ttrbrace}

\tdx{qsum_def}:       A <+> B      == ({\ttlbrace}0{\ttrbrace} <*> A) \isasymunion ({\ttlbrace}1{\ttrbrace} <*> B)
\tdx{QInl_def}:       QInl(a)      == <0;a>
\tdx{QInr_def}:       QInr(b)      == <1;b>
\tdx{qcase_def}:      qcase(c,d)   == qsplit(\%y z. cond(y, d(z), c(z)))
\end{alltt*}
\caption{Non-standard pairs, products and sums} \label{zf-qpair}
\end{figure}

\subsection{Non-standard ordered pairs}

Theory \thydx{QPair} defines a notion of ordered pair that admits
non-well-founded tupling (Fig.\ts\ref{zf-qpair}).  Such pairs are written
{\tt<$a$;$b$>}.  It also defines the eliminator \cdx{qsplit}, the
converse operator \cdx{qconverse}, and the summation operator
\cdx{QSigma}.  These are completely analogous to the corresponding
versions for standard ordered pairs.  The theory goes on to define a
non-standard notion of disjoint sum using non-standard pairs.  All of these
concepts satisfy the same properties as their standard counterparts; in
addition, {\tt<$a$;$b$>} is continuous.  The theory supports coinductive
definitions, for example of infinite lists~\cite{paulson-mscs}.

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{bnd_mono_def}:  bnd_mono(D,h) ==
h(D)\isasymsubseteq{}D & ({\isasymforall}W X. W\isasymsubseteq{}X --> X\isasymsubseteq{}D --> h(W)\isasymsubseteq{}h(X))

\tdx{lfp_def}:       lfp(D,h) == Inter({\ttlbrace}X \isasymin Pow(D). h(X) \isasymsubseteq X{\ttrbrace})
\tdx{gfp_def}:       gfp(D,h) == Union({\ttlbrace}X \isasymin Pow(D). X \isasymsubseteq h(X){\ttrbrace})

\tdx{lfp_lowerbound}: [| h(A) \isasymsubseteq A;  A \isasymsubseteq D |] ==> lfp(D,h) \isasymsubseteq A

\tdx{lfp_subset}:    lfp(D,h) \isasymsubseteq D

\tdx{lfp_greatest}:  [| bnd_mono(D,h);
!!X. [| h(X) \isasymsubseteq X;  X \isasymsubseteq D |] ==> A \isasymsubseteq X
|] ==> A \isasymsubseteq lfp(D,h)

\tdx{lfp_Tarski}:    bnd_mono(D,h) ==> lfp(D,h) = h(lfp(D,h))

\tdx{induct}:        [| a \isasymin lfp(D,h);  bnd_mono(D,h);
!!x. x \isasymin h(Collect(lfp(D,h),P)) ==> P(x)
|] ==> P(a)

\tdx{lfp_mono}:      [| bnd_mono(D,h);  bnd_mono(E,i);
!!X. X \isasymsubseteq D ==> h(X) \isasymsubseteq i(X)
|] ==> lfp(D,h) \isasymsubseteq lfp(E,i)

\tdx{gfp_upperbound}: [| A \isasymsubseteq h(A);  A \isasymsubseteq D |] ==> A \isasymsubseteq gfp(D,h)

\tdx{gfp_subset}:    gfp(D,h) \isasymsubseteq D

\tdx{gfp_least}:     [| bnd_mono(D,h);
!!X. [| X \isasymsubseteq h(X);  X \isasymsubseteq D |] ==> X \isasymsubseteq A
|] ==> gfp(D,h) \isasymsubseteq A

\tdx{gfp_Tarski}:    bnd_mono(D,h) ==> gfp(D,h) = h(gfp(D,h))

\tdx{coinduct}:      [| bnd_mono(D,h); a \isasymin X; X \isasymsubseteq h(X \isasymunion gfp(D,h)); X \isasymsubseteq D
|] ==> a \isasymin gfp(D,h)

\tdx{gfp_mono}:      [| bnd_mono(D,h);  D \isasymsubseteq E;
!!X. X \isasymsubseteq D ==> h(X) \isasymsubseteq i(X)
|] ==> gfp(D,h) \isasymsubseteq gfp(E,i)
\end{alltt*}
\caption{Least and greatest fixedpoints} \label{zf-fixedpt}
\end{figure}

\subsection{Least and greatest fixedpoints}

The Knaster-Tarski Theorem states that every monotone function over a
complete lattice has a fixedpoint.  Theory \thydx{Fixedpt} proves the
Theorem only for a particular lattice, namely the lattice of subsets of a
set (Fig.\ts\ref{zf-fixedpt}).  The theory defines least and greatest
fixedpoint operators with corresponding induction and coinduction rules.
These are essential to many definitions that follow, including the natural
numbers and the transitive closure operator.  The (co)inductive definition
package also uses the fixedpoint operators~\cite{paulson-CADE}.  See
Davey and Priestley~\cite{davey-priestley} for more on the Knaster-Tarski
Theorem and my paper~\cite{paulson-set-II} for discussion of the Isabelle
proofs.

Monotonicity properties are proved for most of the set-forming operations:
union, intersection, Cartesian product, image, domain, range, etc.  These
are useful for applying the Knaster-Tarski Fixedpoint Theorem.  The proofs
themselves are trivial applications of Isabelle's classical reasoner.

\subsection{Finite sets and lists}

Theory \texttt{Finite} (Figure~\ref{zf-fin}) defines the finite set operator;
$\isa{Fin}(A)$ is the set of all finite sets over~$A$.  The theory employs
Isabelle's inductive definition package, which proves various rules
automatically.  The induction rule shown is stronger than the one proved by
the package.  The theory also defines the set of all finite functions
between two given sets.

\begin{figure}
\begin{alltt*}\isastyleminor
\tdx{Fin.emptyI}      0 \isasymin Fin(A)
\tdx{Fin.consI}       [| a \isasymin A;  b \isasymin Fin(A) |] ==> cons(a,b) \isasymin Fin(A)

\tdx{Fin_induct}
[| b \isasymin Fin(A);
P(0);
!!x y. [| x\isasymin{}A; y\isasymin{}Fin(A); x\isasymnotin{}y; P(y) |] ==> P(cons(x,y))
|] ==> P(b)

\tdx{Fin_mono}:       A \isasymsubseteq B ==> Fin(A) \isasymsubseteq Fin(B)
\tdx{Fin_UnI}:        [| b \isasymin Fin(A);  c \isasymin Fin(A) |] ==> b \isasymunion c \isasymin Fin(A)
\tdx{Fin_UnionI}:     C \isasymin Fin(Fin(A)) ==> Union(C) \isasymin Fin(A)
\tdx{Fin_subset}:     [| c \isasymsubseteq b;  b \isasymin Fin(A) |] ==> c \isasymin Fin(A)
\end{alltt*}
\caption{The finite set operator} \label{zf-fin}
\end{figure}

\begin{figure}
\begin{constants}
\it symbol  & \it meta-type & \it priority & \it description \\
\cdx{list}    & $i\To i$      && lists over some set\\
\cdx{list_case} & $[i, [i,i]\To i, i] \To i$  && conditional for $list(A)$ \\
\cdx{map}     & $[i\To i, i] \To i$   &       & mapping functional\\
\cdx{length}  & $i\To i$              &       & length of a list\\
\cdx{rev}     & $i\To i$              &       & reverse of a list\\
\tt \at       & $[i,i]\To i$  &  Right 60     & append for lists\\
\cdx{flat}    & $i\To i$   &                  & append of list of lists
\end{constants}

\underscoreon %%because @ is used here
\begin{alltt*}\isastyleminor
\tdx{NilI}:       Nil \isasymin list(A)
\tdx{ConsI}:      [| a \isasymin A;  l \isasymin list(A) |] ==> Cons(a,l) \isasymin list(A)

\tdx{List.induct}
[| l \isasymin list(A);
P(Nil);
!!x y. [| x \isasymin A;  y \isasymin list(A);  P(y) |] ==> P(Cons(x,y))
|] ==> P(l)

\tdx{Cons_iff}:       Cons(a,l)=Cons(a',l') <-> a=a' & l=l'
\tdx{Nil_Cons_iff}:    Nil \isasymnoteq Cons(a,l)

\tdx{list_mono}:      A \isasymsubseteq B ==> list(A) \isasymsubseteq list(B)

\tdx{map_ident}:      l\isasymin{}list(A) ==> map(\%u. u, l) = l
\tdx{map_compose}:    l\isasymin{}list(A) ==> map(h, map(j,l)) = map(\%u. h(j(u)), l)
\tdx{map_app_distrib}: xs\isasymin{}list(A) ==> map(h, xs@ys) = map(h,xs)@map(h,ys)
\tdx{map_type}
[| l\isasymin{}list(A); !!x. x\isasymin{}A ==> h(x)\isasymin{}B |] ==> map(h,l)\isasymin{}list(B)
\tdx{map_flat}
ls: list(list(A)) ==> map(h, flat(ls)) = flat(map(map(h),ls))
\end{alltt*}
\caption{Lists} \label{zf-list}
\end{figure}

Figure~\ref{zf-list} presents the set of lists over~$A$, $\isa{list}(A)$.  The
definition employs Isabelle's datatype package, which defines the introduction
and induction rules automatically, as well as the constructors, case operator
(\isa{list\_case}) and recursion operator.  The theory then defines the usual
list functions by primitive recursion.  See theory \texttt{List}.

\subsection{Miscellaneous}

\begin{figure}
\begin{constants}
\it symbol  & \it meta-type & \it priority & \it description \\
\sdx{O}       & $[i,i]\To i$  &  Right 60     & composition ($\circ$) \\
\cdx{id}      & $i\To i$      &       & identity function \\
\cdx{inj}     & $[i,i]\To i$  &       & injective function space\\
\cdx{surj}    & $[i,i]\To i$  &       & surjective function space\\
\cdx{bij}     & $[i,i]\To i$  &       & bijective function space
\end{constants}

\begin{alltt*}\isastyleminor
\tdx{comp_def}: r O s     == {\ttlbrace}xz \isasymin domain(s)*range(r) .
{\isasymexists}x y z. xz=<x,z> & <x,y> \isasymin s & <y,z> \isasymin r{\ttrbrace}
\tdx{id_def}:   id(A)     == (lam x \isasymin A. x)
\tdx{inj_def}:  inj(A,B)  == {\ttlbrace} f\isasymin{}A->B. {\isasymforall}w\isasymin{}A. {\isasymforall}x\isasymin{}A. fw=fx --> w=x {\ttrbrace}
\tdx{surj_def}: surj(A,B) == {\ttlbrace} f\isasymin{}A->B . {\isasymforall}y\isasymin{}B. {\isasymexists}x\isasymin{}A. fx=y {\ttrbrace}
\tdx{bij_def}:  bij(A,B)  == inj(A,B) \isasyminter surj(A,B)

\tdx{left_inverse}:    [| f\isasymin{}inj(A,B);  a\isasymin{}A |] ==> converse(f)(fa) = a
\tdx{right_inverse}:   [| f\isasymin{}inj(A,B);  b\isasymin{}range(f) |] ==>
f(converse(f)b) = b

\tdx{inj_converse_inj}: f\isasymin{}inj(A,B) ==> converse(f) \isasymin inj(range(f),A)
\tdx{bij_converse_bij}: f\isasymin{}bij(A,B) ==> converse(f) \isasymin bij(B,A)

\tdx{comp_type}:     [| s \isasymsubseteq A*B;  r \isasymsubseteq B*C |] ==> (r O s) \isasymsubseteq A*C
\tdx{comp_assoc}:    (r O s) O t = r O (s O t)

\tdx{left_comp_id}:  r \isasymsubseteq A*B ==> id(B) O r = r
\tdx{right_comp_id}: r \isasymsubseteq A*B ==> r O id(A) = r

\tdx{comp_func}:     [| g\isasymin{}A->B; f\isasymin{}B->C |] ==> (f O g) \isasymin A->C
\tdx{comp_func_apply}: [| g\isasymin{}A->B; f\isasymin{}B->C; a\isasymin{}A |] ==> (f O g)a = f(ga)

\tdx{comp_inj}:      [| g\isasymin{}inj(A,B);  f\isasymin{}inj(B,C)  |] ==> (f O g)\isasymin{}inj(A,C)
\tdx{comp_surj}:     [| g\isasymin{}surj(A,B); f\isasymin{}surj(B,C) |] ==> (f O g)\isasymin{}surj(A,C)
\tdx{comp_bij}:      [| g\isasymin{}bij(A,B); f\isasymin{}bij(B,C) |] ==> (f O g)\isasymin{}bij(A,C)

\tdx{left_comp_inverse}:    f\isasymin{}inj(A,B) ==> converse(f) O f = id(A)
\tdx{right_comp_inverse}:   f\isasymin{}surj(A,B) ==> f O converse(f) = id(B)

\tdx{bij_disjoint_Un}:
[| f\isasymin{}bij(A,B);  g\isasymin{}bij(C,D);  A \isasyminter C = 0;  B \isasyminter D = 0 |] ==>
(f \isasymunion g)\isasymin{}bij(A \isasymunion C, B \isasymunion D)

\tdx{restrict_bij}: [| f\isasymin{}inj(A,B); C\isasymsubseteq{}A |] ==> restrict(f,C)\isasymin{}bij(C, fC)
\end{alltt*}
\caption{Permutations} \label{zf-perm}
\end{figure}

The theory \thydx{Perm} is concerned with permutations (bijections) and
related concepts.  These include composition of relations, the identity
relation, and three specialized function spaces: injective, surjective and
bijective.  Figure~\ref{zf-perm} displays many of their properties that
have been proved.  These results are fundamental to a treatment of
equipollence and cardinality.

Theory \thydx{Univ} defines a universe' $\isa{univ}(A)$, which is used by
the datatype package.  This set contains $A$ and the
natural numbers.  Vitally, it is closed under finite products:
$\isa{univ}(A)\times\isa{univ}(A)\subseteq\isa{univ}(A)$.  This theory also
defines the cumulative hierarchy of axiomatic set theory, which
traditionally is written $V@\alpha$ for an ordinal~$\alpha$.  The
universe' is a simple generalization of~$V@\omega$.

Theory \thydx{QUniv} defines a universe' $\isa{quniv}(A)$, which is used by
the datatype package to construct codatatypes such as streams.  It is
analogous to $\isa{univ}(A)$ (and is defined in terms of it) but is closed
under the non-standard product and sum.

\section{Automatic Tools}

ZF provides the simplifier and the classical reasoner.  Moreover it supplies a
specialized tool to infer types' of terms.

\subsection{Simplification and Classical Reasoning}

ZF inherits simplification from FOL but adopts it for set theory.  The
extraction of rewrite rules takes the ZF primitives into account.  It can
strip bounded universal quantifiers from a formula; for example, ${\forall x\in A. f(x)=g(x)}$ yields the conditional rewrite rule $x\in A \Imp f(x)=g(x)$.  Given $a\in\{x\in A. P(x)\}$ it extracts rewrite rules from $a\in A$ and~$P(a)$.  It can also break down $a\in A\int B$ and $a\in A-B$.

The default simpset used by \isa{simp} contains congruence rules for all of ZF's
binding operators.  It contains all the conversion rules, such as
\isa{fst} and
\isa{snd}, as well as the rewrites shown in Fig.\ts\ref{zf-simpdata}.

Classical reasoner methods such as \isa{blast} and \isa{auto} refer to
a rich collection of built-in axioms for all the set-theoretic
primitives.

\begin{figure}
\begin{eqnarray*}
a\in \emptyset        & \bimp &  \bot\\
a \in A \un B      & \bimp &  a\in A \disj a\in B\\
a \in A \int B      & \bimp &  a\in A \conj a\in B\\
a \in A-B             & \bimp &  a\in A \conj \lnot (a\in B)\\
\pair{a,b}\in \isa{Sigma}(A,B)
& \bimp &  a\in A \conj b\in B(a)\\
a \in \isa{Collect}(A,P)      & \bimp &  a\in A \conj P(a)\\
(\forall x \in \emptyset. P(x)) & \bimp &  \top\\
(\forall x \in A. \top)       & \bimp &  \top
\end{eqnarray*}
\caption{Some rewrite rules for set theory} \label{zf-simpdata}
\end{figure}

\subsection{Type-Checking Tactics}
\index{type-checking tactics}

Isabelle/ZF provides simple tactics to help automate those proofs that are
essentially type-checking.  Such proofs are built by applying rules such as
these:
\begin{ttbox}\isastyleminor
[| ?P ==> ?a \isasymin ?A; ~?P ==> ?b \isasymin ?A |]
==> (if ?P then ?a else ?b) \isasymin ?A

[| ?m \isasymin nat; ?n \isasymin nat |] ==> ?m #+ ?n \isasymin nat

?a \isasymin ?A ==> Inl(?a) \isasymin ?A + ?B
\end{ttbox}
In typical applications, the goal has the form $t\in\Var{A}$: in other words,
we have a specific term~$t$ and need to infer its type' by instantiating the
set variable~$\Var{A}$.  Neither the simplifier nor the classical reasoner
does this job well.  The if-then-else rule, and many similar ones, can make
the classical reasoner loop.  The simplifier refuses (on principle) to
instantiate variables during rewriting, so goals such as \isa{i\#+j \isasymin \ ?A}
are left unsolved.

The simplifier calls the type-checker to solve rewritten subgoals: this stage
can indeed instantiate variables.  If you have defined new constants and
proved type-checking rules for them, then declare the rules using
the attribute \isa{TC} and the rest should be automatic.  In
particular, the simplifier will use type-checking to help satisfy
conditional rewrite rules. Call the method \ttindex{typecheck} to
break down all subgoals using type-checking rules. You can add new
type-checking rules temporarily like this:
\begin{isabelle}
\end{isabelle}

%Though the easiest way to invoke the type-checker is via the simplifier,
%specialized applications may require more detailed knowledge of
%the type-checking primitives.  They are modelled on the simplifier's:
%\begin{ttdescription}
%\item[\ttindexbold{tcset}] is the type of tcsets: sets of type-checking rules.
%
%  a tcset.
%
%\item[\ttindexbold{delTCs}] is an infix operator to remove type-checking rules
%  from a tcset.
%
%\item[\ttindexbold{typecheck_tac}] is a tactic for attempting to prove all
%  subgoals using the rules given in its argument, a tcset.
%\end{ttdescription}
%
%Tcsets, like simpsets, are associated with theories and are merged when
%theories are merged.  There are further primitives that use the default tcset.
%\begin{ttdescription}
%\item[\ttindexbold{tcset}] is a function to return the default tcset; use the
%  expression \isa{tcset()}.
%
%
%\item[\ttindexbold{DelTCs}] removes type-checking rules from the default
%  tcset.
%
%\item[\ttindexbold{Typecheck_tac}] calls \isa{typecheck_tac} using the
%  default tcset.
%\end{ttdescription}
%
%To supply some type-checking rules temporarily, using \isa{Addrules} and
%later \isa{Delrules} is the simplest way.  There is also a high-tech
%approach.  Call the simplifier with a new solver expressed using
%\ttindexbold{type_solver_tac} and your temporary type-checking rules.
%\begin{ttbox}\isastyleminor
%by (asm_simp_tac
%     (simpset() setSolver type_solver_tac (tcset() addTCs prems)) 2);
%\end{ttbox}

\section{Natural number and integer arithmetic}

\index{arithmetic|(}

\begin{figure}\small
\index{#*@{\tt\#*} symbol}
\index{*div symbol}
\index{*mod symbol}
\index{#+@{\tt\#+} symbol}
\index{#-@{\tt\#-} symbol}
\begin{constants}
\it symbol  & \it meta-type & \it priority & \it description \\
\cdx{nat}     & $i$                   &       & set of natural numbers \\
\cdx{nat_case}& $[i,i\To i,i]\To i$     &     & conditional for $nat$\\
\tt \#*       & $[i,i]\To i$  &  Left 70      & multiplication \\
\tt div       & $[i,i]\To i$  &  Left 70      & division\\
\tt mod       & $[i,i]\To i$  &  Left 70      & modulus\\
\tt \#+       & $[i,i]\To i$  &  Left 65      & addition\\
\tt \#-       & $[i,i]\To i$  &  Left 65      & subtraction
\end{constants}

\begin{alltt*}\isastyleminor
\tdx{nat_def}: nat == lfp(lam r \isasymin Pow(Inf). {\ttlbrace}0{\ttrbrace} \isasymunion {\ttlbrace}succ(x). x \isasymin r{\ttrbrace}

\tdx{nat_case_def}:  nat_case(a,b,k) ==
THE y. k=0 & y=a | ({\isasymexists}x. k=succ(x) & y=b(x))

\tdx{nat_0I}:           0 \isasymin nat
\tdx{nat_succI}:        n \isasymin nat ==> succ(n) \isasymin nat

\tdx{nat_induct}:
[| n \isasymin nat;  P(0);  !!x. [| x \isasymin nat;  P(x) |] ==> P(succ(x))
|] ==> P(n)

\tdx{nat_case_0}:       nat_case(a,b,0) = a
\tdx{nat_case_succ}:    nat_case(a,b,succ(m)) = b(m)

\tdx{add_0_natify}:     0 #+ n = natify(n)
\tdx{add_succ}:         succ(m) #+ n = succ(m #+ n)

\tdx{mult_type}:        m #* n \isasymin nat
\tdx{mult_0}:           0 #* n = 0
\tdx{mult_succ}:        succ(m) #* n = n #+ (m #* n)
\tdx{mult_commute}:     m #* n = n #* m
\tdx{add_mult_dist}:    (m #+ n) #* k = (m #* k) #+ (n #* k)
\tdx{mult_assoc}:       (m #* n) #* k = m #* (n #* k)
\tdx{mod_div_equality}: m \isasymin nat ==> (m div n)#*n #+ m mod n = m
\end{alltt*}
\caption{The natural numbers} \label{zf-nat}
\end{figure}

\index{natural numbers}

Theory \thydx{Nat} defines the natural numbers and mathematical
induction, along with a case analysis operator.  The set of natural
numbers, here called \isa{nat}, is known in set theory as the ordinal~$\omega$.

Theory \thydx{Arith} develops arithmetic on the natural numbers
(Fig.\ts\ref{zf-nat}).  Addition, multiplication and subtraction are defined
by primitive recursion.  Division and remainder are defined by repeated
subtraction, which requires well-founded recursion; the termination argument
relies on the divisor's being non-zero.  Many properties are proved:
commutative, associative and distributive laws, identity and cancellation
laws, etc.  The most interesting result is perhaps the theorem $a \bmod b + (a/b)\times b = a$.

To minimize the need for tedious proofs of $t\in\isa{nat}$, the arithmetic
operators coerce their arguments to be natural numbers.  The function
\cdx{natify} is defined such that $\isa{natify}(n) = n$ if $n$ is a natural
number, $\isa{natify}(\isa{succ}(x)) = \isa{succ}(\isa{natify}(x))$ for all $x$, and finally
$\isa{natify}(x)=0$ in all other cases.  The benefit is that the addition,
subtraction, multiplication, division and remainder operators always return
natural numbers, regardless of their arguments.  Algebraic laws (commutative,
associative, distributive) are unconditional.  Occurrences of \isa{natify}
as operands of those operators are simplified away.  Any remaining occurrences
can either be tolerated or else eliminated by proving that the argument is a
natural number.

The simplifier automatically cancels common terms on the opposite sides of
subtraction and of relations ($=$, $<$ and $\le$).  Here is an example:
\begin{isabelle}
1. i \#+ j \#+ k \#- j < k \#+ l\isanewline
\isacommand{apply}\ simp\isanewline
1. natify(i) < natify(l)
\end{isabelle}
Given the assumptions \isa{i \isasymin nat} and \isa{l \isasymin nat}, both occurrences of
\cdx{natify} would be simplified away.

\begin{figure}\small
\index{$*@{\tt\$*} symbol}
\index{$+@{\tt\$+} symbol}
\index{$-@{\tt\$-} symbol}
\begin{constants}
\it symbol  & \it meta-type & \it priority & \it description \\
\cdx{int}     & $i$                   &       & set of integers \\
\tt \$* &$[i,i]\To i$& Left 70 & multiplication \\ \tt \$+       & $[i,i]\To i$  &  Left 65      & addition\\
\tt \$- &$[i,i]\To i$& Left 65 & subtraction\\ \tt \$<       & $[i,i]\To o$  &  Left 50      & $<$ on integers\\
\tt \$<= &$[i,i]\To o$& Left 50 &$\le$on integers \end{constants} \begin{alltt*}\isastyleminor \tdx{zadd_0_intify}: 0$+ n = intify(n)

\tdx{zmult_type}:       m $* n \isasymin int \tdx{zmult_0}: 0$* n = 0
\tdx{zmult_commute}:    m $* n = n$* m
\tdx{zadd_zmult_dist}:   (m $+ n)$* k = (m $* k)$+ (n $* k) \tdx{zmult_assoc}: (m$* n) $* k = m$* (n $* k) \end{alltt*} \caption{The integers} \label{zf-int} \end{figure} \index{integers} Theory \thydx{Int} defines the integers, as equivalence classes of natural numbers. Figure~\ref{zf-int} presents a tidy collection of laws. In fact, a large library of facts is proved, including monotonicity laws for addition and multiplication, covering both positive and negative operands. As with the natural numbers, the need for typing proofs is minimized. All the operators defined in Fig.\ts\ref{zf-int} coerce their operands to integers by applying the function \cdx{intify}. This function is the identity on integers and maps other operands to zero. Decimal notation is provided for the integers. Numbers, written as \isa{\#$nnn$} or \isa{\#-$nnn$}, are represented internally in two's-complement binary. Expressions involving addition, subtraction and multiplication of numeral constants are evaluated (with acceptable efficiency) by simplification. The simplifier also collects similar terms, multiplying them by a numerical coefficient. It also cancels occurrences of the same terms on the other side of the relational operators. Example: \begin{isabelle} 1. y \$+ z \$+ \#-3 \$* x \$+ y \$<=  x \$* \#2 \$+
z\isanewline
\isacommand{apply}\ simp\isanewline
1. \#2 \$* y \$<= \#5 \$* x \end{isabelle} For more information on the integers, please see the theories on directory \texttt{ZF/Integ}. \index{arithmetic|)} \section{Datatype definitions} \label{sec:ZF:datatype} \index{*datatype|(} The \ttindex{datatype} definition package of ZF constructs inductive datatypes similar to \ML's. It can also construct coinductive datatypes (codatatypes), which are non-well-founded structures such as streams. It defines the set using a fixed-point construction and proves induction rules, as well as theorems for recursion and case combinators. It supplies mechanisms for reasoning about freeness. The datatype package can handle both mutual and indirect recursion. \subsection{Basics} \label{subsec:datatype:basics} A \isa{datatype} definition has the following form: $\begin{array}{llcl} \mathtt{datatype} & t@1(A@1,\ldots,A@h) & = & constructor^1@1 ~\mid~ \ldots ~\mid~ constructor^1@{k@1} \\ & & \vdots \\ \mathtt{and} & t@n(A@1,\ldots,A@h) & = & constructor^n@1~ ~\mid~ \ldots ~\mid~ constructor^n@{k@n} \end{array}$ Here$t@1$, \ldots,~$t@n$are identifiers and$A@1$, \ldots,~$A@h$are variables: the datatype's parameters. Each constructor specification has the form \dquotesoff $C \hbox{\tt~( } \hbox{\tt"} x@1 \hbox{\tt:} T@1 \hbox{\tt"},\; \ldots,\; \hbox{\tt"} x@m \hbox{\tt:} T@m \hbox{\tt"} \hbox{\tt~)}$ Here$C$is the constructor name, and variables$x@1$, \ldots,~$x@m$are the constructor arguments, belonging to the sets$T@1$, \ldots,$T@m$, respectively. Typically each$T@j$is either a constant set, a datatype parameter (one of$A@1$, \ldots,$A@h$) or a recursive occurrence of one of the datatypes, say$t@i(A@1,\ldots,A@h)$. More complex possibilities exist, but they are much harder to realize. Often, additional information must be supplied in the form of theorems. A datatype can occur recursively as the argument of some function~$F$. This is called a {\em nested} (or \emph{indirect}) occurrence. It is only allowed if the datatype package is given a theorem asserting that$F$is monotonic. If the datatype has indirect occurrences, then Isabelle/ZF does not support recursive function definitions. A simple example of a datatype is \isa{list}, which is built-in, and is defined by \begin{alltt*}\isastyleminor consts list :: "i=>i" datatype "list(A)" = Nil | Cons ("a \isasymin A", "l \isasymin list(A)") \end{alltt*} Note that the datatype operator must be declared as a constant first. However, the package declares the constructors. Here, \isa{Nil} gets type$i$and \isa{Cons} gets type$[i,i]\To i$. Trees and forests can be modelled by the mutually recursive datatype definition \begin{alltt*}\isastyleminor consts tree :: "i=>i" forest :: "i=>i" tree_forest :: "i=>i" datatype "tree(A)" = Tcons ("a{\isasymin}A", "f{\isasymin}forest(A)") and "forest(A)" = Fnil | Fcons ("t{\isasymin}tree(A)", "f{\isasymin}forest(A)") \end{alltt*} Here$\isa{tree}(A)$is the set of trees over$A$,$\isa{forest}(A)$is the set of forests over$A$, and$\isa{tree_forest}(A)$is the union of the previous two sets. All three operators must be declared first. The datatype \isa{term}, which is defined by \begin{alltt*}\isastyleminor consts term :: "i=>i" datatype "term(A)" = Apply ("a \isasymin A", "l \isasymin list(term(A))") monos list_mono type_elims list_univ [THEN subsetD, elim_format] \end{alltt*} is an example of nested recursion. (The theorem \isa{list_mono} is proved in theory \isa{List}, and the \isa{term} example is developed in theory \thydx{Induct/Term}.) \subsubsection{Freeness of the constructors} Constructors satisfy {\em freeness} properties. Constructions are distinct, for example$\isa{Nil}\not=\isa{Cons}(a,l)$, and they are injective, for example$\isa{Cons}(a,l)=\isa{Cons}(a',l') \bimp a=a' \conj l=l'$. Because the number of freeness is quadratic in the number of constructors, the datatype package does not prove them. Instead, it ensures that simplification will prove them dynamically: when the simplifier encounters a formula asserting the equality of two datatype constructors, it performs freeness reasoning. Freeness reasoning can also be done using the classical reasoner, but it is more complicated. You have to add some safe elimination rules rules to the claset. For the \isa{list} datatype, they are called \isa{list.free_elims}. Occasionally this exposes the underlying representation of some constructor, which can be rectified using the command \isa{unfold list.con_defs [symmetric]}. \subsubsection{Structural induction} The datatype package also provides structural induction rules. For datatypes without mutual or nested recursion, the rule has the form exemplified by \isa{list.induct} in Fig.\ts\ref{zf-list}. For mutually recursive datatypes, the induction rule is supplied in two forms. Consider datatype \isa{TF}. The rule \isa{tree_forest.induct} performs induction over a single predicate~\isa{P}, which is presumed to be defined for both trees and forests: \begin{alltt*}\isastyleminor [| x \isasymin tree_forest(A); !!a f. [| a \isasymin A; f \isasymin forest(A); P(f) |] ==> P(Tcons(a, f)); P(Fnil); !!f t. [| t \isasymin tree(A); P(t); f \isasymin forest(A); P(f) |] ==> P(Fcons(t, f)) |] ==> P(x) \end{alltt*} The rule \isa{tree_forest.mutual_induct} performs induction over two distinct predicates, \isa{P_tree} and \isa{P_forest}. \begin{alltt*}\isastyleminor [| !!a f. [| a{\isasymin}A; f{\isasymin}forest(A); P_forest(f) |] ==> P_tree(Tcons(a,f)); P_forest(Fnil); !!f t. [| t{\isasymin}tree(A); P_tree(t); f{\isasymin}forest(A); P_forest(f) |] ==> P_forest(Fcons(t, f)) |] ==> ({\isasymforall}za. za \isasymin tree(A) --> P_tree(za)) & ({\isasymforall}za. za \isasymin forest(A) --> P_forest(za)) \end{alltt*} For datatypes with nested recursion, such as the \isa{term} example from above, things are a bit more complicated. The rule \isa{term.induct} refers to the monotonic operator, \isa{list}: \begin{alltt*}\isastyleminor [| x \isasymin term(A); !!a l. [| a\isasymin{}A; l\isasymin{}list(Collect(term(A), P)) |] ==> P(Apply(a,l)) |] ==> P(x) \end{alltt*} The theory \isa{Induct/Term.thy} derives two higher-level induction rules, one of which is particularly useful for proving equations: \begin{alltt*}\isastyleminor [| t \isasymin term(A); !!x zs. [| x \isasymin A; zs \isasymin list(term(A)); map(f, zs) = map(g, zs) |] ==> f(Apply(x, zs)) = g(Apply(x, zs)) |] ==> f(t) = g(t) \end{alltt*} How this can be generalized to other nested datatypes is a matter for future research. \subsubsection{The \isa{case} operator} The package defines an operator for performing case analysis over the datatype. For \isa{list}, it is called \isa{list_case} and satisfies the equations \begin{ttbox}\isastyleminor list_case(f_Nil, f_Cons, []) = f_Nil list_case(f_Nil, f_Cons, Cons(a, l)) = f_Cons(a, l) \end{ttbox} Here \isa{f_Nil} is the value to return if the argument is \isa{Nil} and \isa{f_Cons} is a function that computes the value to return if the argument has the form$\isa{Cons}(a,l)$. The function can be expressed as an abstraction, over patterns if desired (\S\ref{sec:pairs}). For mutually recursive datatypes, there is a single \isa{case} operator. In the tree/forest example, the constant \isa{tree_forest_case} handles all of the constructors of the two datatypes. \subsection{Defining datatypes} The theory syntax for datatype definitions is shown in the Isabelle/Isar reference manual. In order to be well-formed, a datatype definition has to obey the rules stated in the previous section. As a result the theory is extended with the new types, the constructors, and the theorems listed in the previous section. Codatatypes are declared like datatypes and are identical to them in every respect except that they have a coinduction rule instead of an induction rule. Note that while an induction rule has the effect of limiting the values contained in the set, a coinduction rule gives a way of constructing new values of the set. Most of the theorems about datatypes become part of the default simpset. You never need to see them again because the simplifier applies them automatically. \subsubsection{Specialized methods for datatypes} Induction and case-analysis can be invoked using these special-purpose methods: \begin{ttdescription} \item[\methdx{induct_tac}$x$] applies structural induction on variable$x$to subgoal~1, provided the type of$x$is a datatype. The induction variable should not occur among other assumptions of the subgoal. \end{ttdescription} % % we also have the ind_cases method, but what does it do? In some situations, induction is overkill and a case distinction over all constructors of the datatype suffices. \begin{ttdescription} \item[\methdx{case_tac}$x$] performs a case analysis for the variable~$x$. \end{ttdescription} Both tactics can only be applied to a variable, whose typing must be given in some assumption, for example the assumption \isa{x \isasymin \ list(A)}. The tactics also work for the natural numbers (\isa{nat}) and disjoint sums, although these sets were not defined using the datatype package. (Disjoint sums are not recursive, so only \isa{case_tac} is available.) Structured Isar methods are also available. Below,$t$stands for the name of the datatype. \begin{ttdescription} \item[\methdx{induct} \isa{set:}\$t$] is the Isar induction tactic. \item[\methdx{cases} \isa{set:}\$t$] is the Isar case-analysis tactic. \end{ttdescription} \subsubsection{The theorems proved by a datatype declaration} Here are some more details for the technically minded. Processing the datatype declaration of a set~$t$produces a name space~$t$containing the following theorems: \begin{ttbox}\isastyleminor intros \textrm{the introduction rules} cases \textrm{the case analysis rule} induct \textrm{the standard induction rule} mutual_induct \textrm{the mutual induction rule, if needed} case_eqns \textrm{equations for the case operator} recursor_eqns \textrm{equations for the recursor} simps \textrm{the union of} case_eqns \textrm{and} recursor_eqns con_defs \textrm{definitions of the case operator and constructors} free_iffs \textrm{logical equivalences for proving freeness} free_elims \textrm{elimination rules for proving freeness} defs \textrm{datatype definition(s)} \end{ttbox} Furthermore there is the theorem$C$for every constructor~$C$; for example, the \isa{list} datatype's introduction rules are bound to the identifiers \isa{Nil} and \isa{Cons}. For a codatatype, the component \isa{coinduct} is the coinduction rule, replacing the \isa{induct} component. See the theories \isa{Induct/Ntree} and \isa{Induct/Brouwer} for examples of infinitely branching datatypes. See theory \isa{Induct/LList} for an example of a codatatype. Some of these theories illustrate the use of additional, undocumented features of the datatype package. Datatype definitions are reduced to inductive definitions, and the advanced features should be understood in that light. \subsection{Examples} \subsubsection{The datatype of binary trees} Let us define the set$\isa{bt}(A)$of binary trees over~$A$. The theory must contain these lines: \begin{alltt*}\isastyleminor consts bt :: "i=>i" datatype "bt(A)" = Lf | Br ("a\isasymin{}A", "t1\isasymin{}bt(A)", "t2\isasymin{}bt(A)") \end{alltt*} After loading the theory, we can prove some theorem. We begin by declaring the constructor's typechecking rules as simplification rules: \begin{isabelle} \isacommand{declare}\ bt.intros\ [simp]% \end{isabelle} Our first example is the theorem that no tree equals its left branch. To make the inductive hypothesis strong enough, the proof requires a quantified induction formula, but the \isa{rule\_format} attribute will remove the quantifiers before the theorem is stored. \begin{isabelle} \isacommand{lemma}\ Br\_neq\_left\ [rule\_format]:\ "l\isasymin bt(A)\ ==>\ \isasymforall x\ r.\ Br(x,l,r)\isasymnoteq{}l"\isanewline \ 1.\ l\ \isasymin \ bt(A)\ \isasymLongrightarrow \ \isasymforall x\ r.\ Br(x,\ l,\ r)\ \isasymnoteq \ l% \end{isabelle} This can be proved by the structural induction tactic: \begin{isabelle} \ \ \isacommand{apply}\ (induct\_tac\ l)\isanewline \ 1.\ \isasymforall x\ r.\ Br(x,\ Lf,\ r)\ \isasymnoteq \ Lf\isanewline \ 2.\ \isasymAnd a\ t1\ t2.\isanewline \isaindent{\ 2.\ \ \ \ }\isasymlbrakk a\ \isasymin \ A;\ t1\ \isasymin \ bt(A);\ \isasymforall x\ r.\ Br(x,\ t1,\ r)\ \isasymnoteq \ t1;\ t2\ \isasymin \ bt(A);\isanewline \isaindent{\ 2.\ \ \ \ \ \ \ }\isasymforall x\ r.\ Br(x,\ t2,\ r)\ \isasymnoteq \ t2\isasymrbrakk \isanewline \isaindent{\ 2.\ \ \ \ }\isasymLongrightarrow \ \isasymforall x\ r.\ Br(x,\ Br(a,\ t1,\ t2),\ r)\ \isasymnoteq \ Br(a,\ t1,\ t2) \end{isabelle} Both subgoals are proved using \isa{auto}, which performs the necessary freeness reasoning. \begin{isabelle} \ \ \isacommand{apply}\ auto\isanewline No\ subgoals!\isanewline \isacommand{done} \end{isabelle} An alternative proof uses Isar's fancy \isa{induct} method, which automatically quantifies over all free variables: \begin{isabelle} \isacommand{lemma}\ Br\_neq\_left':\ "l\ \isasymin \ bt(A)\ ==>\ (!!x\ r.\ Br(x,\ l,\ r)\ \isasymnoteq \ l)"\isanewline \ \ \isacommand{apply}\ (induct\ set:\ bt)\isanewline \ 1.\ \isasymAnd x\ r.\ Br(x,\ Lf,\ r)\ \isasymnoteq \ Lf\isanewline \ 2.\ \isasymAnd a\ t1\ t2\ x\ r.\isanewline \isaindent{\ 2.\ \ \ \ }\isasymlbrakk a\ \isasymin \ A;\ t1\ \isasymin \ bt(A);\ \isasymAnd x\ r.\ Br(x,\ t1,\ r)\ \isasymnoteq \ t1;\ t2\ \isasymin \ bt(A);\isanewline \isaindent{\ 2.\ \ \ \ \ \ \ }\isasymAnd x\ r.\ Br(x,\ t2,\ r)\ \isasymnoteq \ t2\isasymrbrakk \isanewline \isaindent{\ 2.\ \ \ \ }\isasymLongrightarrow \ Br(x,\ Br(a,\ t1,\ t2),\ r)\ \isasymnoteq \ Br(a,\ t1,\ t2) \end{isabelle} Compare the form of the induction hypotheses with the corresponding ones in the previous proof. As before, to conclude requires only \isa{auto}. When there are only a few constructors, we might prefer to prove the freenness theorems for each constructor. This is simple: \begin{isabelle} \isacommand{lemma}\ Br\_iff:\ "Br(a,l,r)\ =\ Br(a',l',r')\ <->\ a=a'\ \&\ l=l'\ \&\ r=r'"\isanewline \ \ \isacommand{by}\ (blast\ elim!:\ bt.free\_elims) \end{isabelle} Here we see a demonstration of freeness reasoning using \isa{bt.free\_elims}, but simpler still is just to apply \isa{auto}. An \ttindex{inductive\_cases} declaration generates instances of the case analysis rule that have been simplified using freeness reasoning. \begin{isabelle} \isacommand{inductive\_cases}\ Br\_in\_bt:\ "Br(a,\ l,\ r)\ \isasymin \ bt(A)" \end{isabelle} The theorem just created is \begin{isabelle} \isasymlbrakk Br(a,\ l,\ r)\ \isasymin \ bt(A);\ \isasymlbrakk a\ \isasymin \ A;\ l\ \isasymin \ bt(A);\ r\ \isasymin \ bt(A)\isasymrbrakk \ \isasymLongrightarrow \ Q\isasymrbrakk \ \isasymLongrightarrow \ Q. \end{isabelle} It is an elimination rule that from$\isa{Br}(a,l,r)\in\isa{bt}(A)$lets us infer$a\in A$,$l\in\isa{bt}(A)$and$r\in\isa{bt}(A)$. \subsubsection{Mixfix syntax in datatypes} Mixfix syntax is sometimes convenient. The theory \isa{Induct/PropLog} makes a deep embedding of propositional logic: \begin{alltt*}\isastyleminor consts prop :: i datatype "prop" = Fls | Var ("n \isasymin nat") ("#_"  100) | "=>" ("p \isasymin prop", "q \isasymin prop") (infixr 90) \end{alltt*} The second constructor has a special$\#n$syntax, while the third constructor is an infixed arrow. \subsubsection{A giant enumeration type} This example shows a datatype that consists of 60 constructors: \begin{alltt*}\isastyleminor consts enum :: i datatype "enum" = C00 | C01 | C02 | C03 | C04 | C05 | C06 | C07 | C08 | C09 | C10 | C11 | C12 | C13 | C14 | C15 | C16 | C17 | C18 | C19 | C20 | C21 | C22 | C23 | C24 | C25 | C26 | C27 | C28 | C29 | C30 | C31 | C32 | C33 | C34 | C35 | C36 | C37 | C38 | C39 | C40 | C41 | C42 | C43 | C44 | C45 | C46 | C47 | C48 | C49 | C50 | C51 | C52 | C53 | C54 | C55 | C56 | C57 | C58 | C59 end \end{alltt*} The datatype package scales well. Even though all properties are proved rather than assumed, full processing of this definition takes around two seconds (on a 1.8GHz machine). The constructors have a balanced representation, related to binary notation, so freeness properties can be proved fast. \begin{isabelle} \isacommand{lemma}\ "C00 \isasymnoteq\ C01"\isanewline \ \ \isacommand{by}\ simp \end{isabelle} You need not derive such inequalities explicitly. The simplifier will dispose of them automatically. \index{*datatype|)} \subsection{Recursive function definitions}\label{sec:ZF:recursive} \index{recursive functions|see{recursion}} \index{*primrec|(} \index{recursion!primitive|(} Datatypes come with a uniform way of defining functions, {\bf primitive recursion}. Such definitions rely on the recursion operator defined by the datatype package. Isabelle proves the desired recursion equations as theorems. In principle, one could introduce primitive recursive functions by asserting their reduction rules as axioms. Here is a dangerous way of defining a recursive function over binary trees: \begin{isabelle} \isacommand{consts}\ \ n\_nodes\ ::\ "i\ =>\ i"\isanewline \isacommand{axioms}\isanewline \ \ n\_nodes\_Lf:\ "n\_nodes(Lf)\ =\ 0"\isanewline \ \ n\_nodes\_Br:\ "n\_nodes(Br(a,l,r))\ =\ succ(n\_nodes(l)\ \#+\ n\_nodes(r))" \end{isabelle} Asserting axioms brings the danger of accidentally introducing contradictions. It should be avoided whenever possible. The \ttindex{primrec} declaration is a safe means of defining primitive recursive functions on datatypes: \begin{isabelle} \isacommand{consts}\ \ n\_nodes\ ::\ "i\ =>\ i"\isanewline \isacommand{primrec}\isanewline \ \ "n\_nodes(Lf)\ =\ 0"\isanewline \ \ "n\_nodes(Br(a,\ l,\ r))\ =\ succ(n\_nodes(l)\ \#+\ n\_nodes(r))" \end{isabelle} Isabelle will now derive the two equations from a low-level definition based upon well-founded recursion. If they do not define a legitimate recursion, then Isabelle will reject the declaration. \subsubsection{Syntax of recursive definitions} The general form of a primitive recursive definition is \begin{ttbox}\isastyleminor primrec {\it reduction rules} \end{ttbox} where \textit{reduction rules} specify one or more equations of the form $f \, x@1 \, \dots \, x@m \, (C \, y@1 \, \dots \, y@k) \, z@1 \, \dots \, z@n = r$ such that$C$is a constructor of the datatype,$r$contains only the free variables on the left-hand side, and all recursive calls in$r$are of the form$f \, \dots \, y@i \, \dots$for some$i$. There must be at most one reduction rule for each constructor. The order is immaterial. For missing constructors, the function is defined to return zero. All reduction rules are added to the default simpset. If you would like to refer to some rule by name, then you must prefix the rule with an identifier. These identifiers, like those in the \isa{rules} section of a theory, will be visible in proof scripts. The reduction rules become part of the default simpset, which leads to short proof scripts: \begin{isabelle} \isacommand{lemma}\ n\_nodes\_type\ [simp]:\ "t\ \isasymin \ bt(A)\ ==>\ n\_nodes(t)\ \isasymin \ nat"\isanewline \ \ \isacommand{by}\ (induct\_tac\ t,\ auto) \end{isabelle} You can even use the \isa{primrec} form with non-recursive datatypes and with codatatypes. Recursion is not allowed, but it provides a convenient syntax for defining functions by cases. \subsubsection{Example: varying arguments} All arguments, other than the recursive one, must be the same in each equation and in each recursive call. To get around this restriction, use explict$\lambda$-abstraction and function application. For example, let us define the tail-recursive version of \isa{n\_nodes}, using an accumulating argument for the counter. The second argument,$k$, varies in recursive calls. \begin{isabelle} \isacommand{consts}\ \ n\_nodes\_aux\ ::\ "i\ =>\ i"\isanewline \isacommand{primrec}\isanewline \ \ "n\_nodes\_aux(Lf)\ =\ (\isasymlambda k\ \isasymin \ nat.\ k)"\isanewline \ \ "n\_nodes\_aux(Br(a,l,r))\ =\ \isanewline \ \ \ \ \ \ (\isasymlambda k\ \isasymin \ nat.\ n\_nodes\_aux(r)\ \ \ (n\_nodes\_aux(l)\ \ succ(k)))" \end{isabelle} Now \isa{n\_nodes\_aux(t)\ \ k} is our function in two arguments. We can prove a theorem relating it to \isa{n\_nodes}. Note the quantification over \isa{k\ \isasymin \ nat}: \begin{isabelle} \isacommand{lemma}\ n\_nodes\_aux\_eq\ [rule\_format]:\isanewline \ \ \ \ \ "t\ \isasymin \ bt(A)\ ==>\ \isasymforall k\ \isasymin \ nat.\ n\_nodes\_aux(t)k\ =\ n\_nodes(t)\ \#+\ k"\isanewline \ \ \isacommand{by}\ (induct\_tac\ t,\ simp\_all) \end{isabelle} Now, we can use \isa{n\_nodes\_aux} to define a tail-recursive version of \isa{n\_nodes}: \begin{isabelle} \isacommand{constdefs}\isanewline \ \ n\_nodes\_tail\ ::\ "i\ =>\ i"\isanewline \ \ \ "n\_nodes\_tail(t)\ ==\ n\_nodes\_aux(t)\ \ 0" \end{isabelle} It is easy to prove that \isa{n\_nodes\_tail} is equivalent to \isa{n\_nodes}: \begin{isabelle} \isacommand{lemma}\ "t\ \isasymin \ bt(A)\ ==>\ n\_nodes\_tail(t)\ =\ n\_nodes(t)"\isanewline \ \isacommand{by}\ (simp\ add:\ n\_nodes\_tail\_def\ n\_nodes\_aux\_eq) \end{isabelle} \index{recursion!primitive|)} \index{*primrec|)} \section{Inductive and coinductive definitions} \index{*inductive|(} \index{*coinductive|(} An {\bf inductive definition} specifies the least set~$R$closed under given rules. (Applying a rule to elements of~$R$yields a result within~$R$.) For example, a structural operational semantics is an inductive definition of an evaluation relation. Dually, a {\bf coinductive definition} specifies the greatest set~$R$consistent with given rules. (Every element of~$R$can be seen as arising by applying a rule to elements of~$R$.) An important example is using bisimulation relations to formalise equivalence of processes and infinite data structures. A theory file may contain any number of inductive and coinductive definitions. They may be intermixed with other declarations; in particular, the (co)inductive sets {\bf must} be declared separately as constants, and may have mixfix syntax or be subject to syntax translations. Each (co)inductive definition adds definitions to the theory and also proves some theorems. It behaves identially to the analogous inductive definition except that instead of an induction rule there is a coinduction rule. Its treatment of coinduction is described in detail in a separate paper,% \footnote{It appeared in CADE~\cite{paulson-CADE}; a longer version is distributed with Isabelle as \emph{A Fixedpoint Approach to (Co)Inductive and (Co)Datatype Definitions}.} % which you might refer to for background information. \subsection{The syntax of a (co)inductive definition} An inductive definition has the form \begin{ttbox}\isastyleminor inductive domains {\it domain declarations} intros {\it introduction rules} monos {\it monotonicity theorems} con_defs {\it constructor definitions} type_intros {\it introduction rules for type-checking} type_elims {\it elimination rules for type-checking} \end{ttbox} A coinductive definition is identical, but starts with the keyword \isa{co\-inductive}. The \isa{monos}, \isa{con\_defs}, \isa{type\_intros} and \isa{type\_elims} sections are optional. If present, each is specified as a list of theorems, which may contain Isar attributes as usual. \begin{description} \item[\it domain declarations] are items of the form {\it string\/}~\isa{\isasymsubseteq }~{\it string}, associating each recursive set with its domain. (The domain is some existing set that is large enough to hold the new set being defined.) \item[\it introduction rules] specify one or more introduction rules in the form {\it ident\/}~{\it string}, where the identifier gives the name of the rule in the result structure. \item[\it monotonicity theorems] are required for each operator applied to a recursive set in the introduction rules. There \textbf{must} be a theorem of the form$A\subseteq B\Imp M(A)\subseteq M(B)$, for each premise$t\in M(R_i)$in an introduction rule! \item[\it constructor definitions] contain definitions of constants appearing in the introduction rules. The (co)datatype package supplies the constructors' definitions here. Most (co)inductive definitions omit this section; one exception is the primitive recursive functions example; see theory \isa{Induct/Primrec}. \item[\it type\_intros] consists of introduction rules for type-checking the definition: for demonstrating that the new set is included in its domain. (The proof uses depth-first search.) \item[\it type\_elims] consists of elimination rules for type-checking the definition. They are presumed to be safe and are applied as often as possible prior to the \isa{type\_intros} search. \end{description} The package has a few restrictions: \begin{itemize} \item The theory must separately declare the recursive sets as constants. \item The names of the recursive sets must be identifiers, not infix operators. \item Side-conditions must not be conjunctions. However, an introduction rule may contain any number of side-conditions. \item Side-conditions of the form$x=t$, where the variable~$x$does not occur in~$t$, will be substituted through the rule \isa{mutual\_induct}. \end{itemize} \subsection{Example of an inductive definition} Below, we shall see how Isabelle/ZF defines the finite powerset operator. The first step is to declare the constant~\isa{Fin}. Then we must declare it inductively, with two introduction rules: \begin{isabelle} \isacommand{consts}\ \ Fin\ ::\ "i=>i"\isanewline \isacommand{inductive}\isanewline \ \ \isakeyword{domains}\ \ \ "Fin(A)"\ \isasymsubseteq\ "Pow(A)"\isanewline \ \ \isakeyword{intros}\isanewline \ \ \ \ emptyI:\ \ "0\ \isasymin\ Fin(A)"\isanewline \ \ \ \ consI:\ \ \ "[|\ a\ \isasymin\ A;\ \ b\ \isasymin\ Fin(A)\ |]\ ==>\ cons(a,b)\ \isasymin\ Fin(A)"\isanewline \ \ \isakeyword{type\_intros}\ \ empty\_subsetI\ cons\_subsetI\ PowI\isanewline \ \ \isakeyword{type\_elims}\ \ \ PowD\ [THEN\ revcut\_rl]\end{isabelle} The resulting theory contains a name space, called~\isa{Fin}. The \isa{Fin}$~A$introduction rules can be referred to collectively as \isa{Fin.intros}, and also individually as \isa{Fin.emptyI} and \isa{Fin.consI}. The induction rule is \isa{Fin.induct}. The chief problem with making (co)inductive definitions involves type-checking the rules. Sometimes, additional theorems need to be supplied under \isa{type_intros} or \isa{type_elims}. If the package fails when trying to prove your introduction rules, then set the flag \ttindexbold{trace_induct} to \isa{true} and try again. (See the manual \emph{A Fixedpoint Approach \ldots} for more discussion of type-checking.) In the example above,$\isa{Pow}(A)$is given as the domain of$\isa{Fin}(A)$, for obviously every finite subset of~$A$is a subset of~$A$. However, the inductive definition package can only prove that given a few hints. Here is the output that results (with the flag set) when the \isa{type_intros} and \isa{type_elims} are omitted from the inductive definition above: \begin{alltt*}\isastyleminor Inductive definition Finite.Fin Fin(A) == lfp(Pow(A), \%X. {z\isasymin{}Pow(A) . z = 0 | ({\isasymexists}a b. z = cons(a,b) & a\isasymin{}A & b\isasymin{}X)}) Proving monotonicity... \ttbreak Proving the introduction rules... The type-checking subgoal: 0 \isasymin Fin(A) 1. 0 \isasymin Pow(A) \ttbreak The subgoal after monos, type_elims: 0 \isasymin Fin(A) 1. 0 \isasymin Pow(A) *** prove_goal: tactic failed \end{alltt*} We see the need to supply theorems to let the package prove$\emptyset\in\isa{Pow}(A)$. Restoring the \isa{type_intros} but not the \isa{type_elims}, we again get an error message: \begin{alltt*}\isastyleminor The type-checking subgoal: 0 \isasymin Fin(A) 1. 0 \isasymin Pow(A) \ttbreak The subgoal after monos, type_elims: 0 \isasymin Fin(A) 1. 0 \isasymin Pow(A) \ttbreak The type-checking subgoal: cons(a, b) \isasymin Fin(A) 1. [| a \isasymin A; b \isasymin Fin(A) |] ==> cons(a, b) \isasymin Pow(A) \ttbreak The subgoal after monos, type_elims: cons(a, b) \isasymin Fin(A) 1. [| a \isasymin A; b \isasymin Pow(A) |] ==> cons(a, b) \isasymin Pow(A) *** prove_goal: tactic failed \end{alltt*} The first rule has been type-checked, but the second one has failed. The simplest solution to such problems is to prove the failed subgoal separately and to supply it under \isa{type_intros}. The solution actually used is to supply, under \isa{type_elims}, a rule that changes$b\in\isa{Pow}(A)$to$b\subseteq A$; together with \isa{cons_subsetI} and \isa{PowI}, it is enough to complete the type-checking. \subsection{Further examples} An inductive definition may involve arbitrary monotonic operators. Here is a standard example: the accessible part of a relation. Note the use of~\isa{Pow} in the introduction rule and the corresponding mention of the rule \isa{Pow\_mono} in the \isa{monos} list. If the desired rule has a universally quantified premise, usually the effect can be obtained using \isa{Pow}. \begin{isabelle} \isacommand{consts}\ \ acc\ ::\ "i\ =>\ i"\isanewline \isacommand{inductive}\isanewline \ \ \isakeyword{domains}\ "acc(r)"\ \isasymsubseteq \ "field(r)"\isanewline \ \ \isakeyword{intros}\isanewline \ \ \ \ vimage:\ \ "[|\ r-\isacharbraceleft a\isacharbraceright\ \isasymin\ Pow(acc(r));\ a\ \isasymin \ field(r)\ |] \isanewline \ \ \ \ \ \ \ \ \ \ \ \ \ \ ==>\ a\ \isasymin \ acc(r)"\isanewline \ \ \isakeyword{monos}\ \ Pow\_mono \end{isabelle} Finally, here are some coinductive definitions. We begin by defining lazy (potentially infinite) lists as a codatatype: \begin{isabelle} \isacommand{consts}\ \ llist\ \ ::\ "i=>i"\isanewline \isacommand{codatatype}\isanewline \ \ "llist(A)"\ =\ LNil\ |\ LCons\ ("a\ \isasymin \ A",\ "l\ \isasymin \ llist(A)")\isanewline \end{isabelle} The notion of equality on such lists is modelled as a bisimulation: \begin{isabelle} \isacommand{consts}\ \ lleq\ ::\ "i=>i"\isanewline \isacommand{coinductive}\isanewline \ \ \isakeyword{domains}\ "lleq(A)"\ <=\ "llist(A)\ *\ llist(A)"\isanewline \ \ \isakeyword{intros}\isanewline \ \ \ \ LNil:\ \ "<LNil,\ LNil>\ \isasymin \ lleq(A)"\isanewline \ \ \ \ LCons:\ "[|\ a\ \isasymin \ A;\ <l,l'>\ \isasymin \ lleq(A)\ |]\ \isanewline \ \ \ \ \ \ \ \ \ \ \ \ ==>\ <LCons(a,l),\ LCons(a,l')>\ \isasymin \ lleq(A)"\isanewline \ \ \isakeyword{type\_intros}\ \ llist.intros \end{isabelle} This use of \isa{type_intros} is typical: the relation concerns the codatatype \isa{llist}, so naturally the introduction rules for that codatatype will be required for type-checking the rules. The Isabelle distribution contains many other inductive definitions. Simple examples are collected on subdirectory \isa{ZF/Induct}. The directory \isa{Coind} and the theory \isa{ZF/Induct/LList} contain coinductive definitions. Larger examples may be found on other subdirectories of \isa{ZF}, such as \isa{IMP}, and \isa{Resid}. \subsection{Theorems generated} Each (co)inductive set defined in a theory file generates a name space containing the following elements: \begin{ttbox}\isastyleminor intros \textrm{the introduction rules} elim \textrm{the elimination (case analysis) rule} induct \textrm{the standard induction rule} mutual_induct \textrm{the mutual induction rule, if needed} defs \textrm{definitions of inductive sets} bnd_mono \textrm{monotonicity property} dom_subset \textrm{inclusion in bounding set'} \end{ttbox} Furthermore, each introduction rule is available under its declared name. For a codatatype, the component \isa{coinduct} is the coinduction rule, replacing the \isa{induct} component. Recall that the \ttindex{inductive\_cases} declaration generates simplified instances of the case analysis rule. It is as useful for inductive definitions as it is for datatypes. There are many examples in the theory \isa{Induct/Comb}, which is discussed at length elsewhere~\cite{paulson-generic}. The theory first defines the datatype \isa{comb} of combinators: \begin{alltt*}\isastyleminor consts comb :: i datatype "comb" = K | S | "#" ("p \isasymin comb", "q \isasymin comb") (infixl 90) \end{alltt*} The theory goes on to define contraction and parallel contraction inductively. Then the theory \isa{Induct/Comb.thy} defines special cases of contraction, such as this one: \begin{isabelle} \isacommand{inductive\_cases}\ K\_contractE [elim!]:\ "K -1-> r" \end{isabelle} The theorem just created is \isa{K -1-> r \ \isasymLongrightarrow \ Q}, which expresses that the combinator \isa{K} cannot reduce to anything. (From the assumption \isa{K-1->r}, we can conclude any desired formula \isa{Q}\@.) Similar elimination rules for \isa{S} and application are also generated. The attribute \isa{elim!}\ shown above supplies the generated theorem to the classical reasoner. This mode of working allows effective reasoniung about operational semantics. \index{*coinductive|)} \index{*inductive|)} \section{The outer reaches of set theory} The constructions of the natural numbers and lists use a suite of operators for handling recursive function definitions. I have described the developments in detail elsewhere~\cite{paulson-set-II}. Here is a brief summary: \begin{itemize} \item Theory \isa{Trancl} defines the transitive closure of a relation (as a least fixedpoint). \item Theory \isa{WF} proves the well-founded recursion theorem, using an elegant approach of Tobias Nipkow. This theorem permits general recursive definitions within set theory. \item Theory \isa{Ord} defines the notions of transitive set and ordinal number. It derives transfinite induction. A key definition is {\bf less than}:$i<j$if and only if$i$and$j$are both ordinals and$i\in j$. As a special case, it includes less than on the natural numbers. \item Theory \isa{Epsilon} derives$\varepsilon$-induction and$\varepsilon$-recursion, which are generalisations of transfinite induction and recursion. It also defines \cdx{rank}$(x)$, which is the least ordinal$\alpha$such that$x$is constructed at stage$\alpha$of the cumulative hierarchy (thus$x\in V@{\alpha+1}$). \end{itemize} Other important theories lead to a theory of cardinal numbers. They have not yet been written up anywhere. Here is a summary: \begin{itemize} \item Theory \isa{Rel} defines the basic properties of relations, such as reflexivity, symmetry and transitivity. \item Theory \isa{EquivClass} develops a theory of equivalence classes, not using the Axiom of Choice. \item Theory \isa{Order} defines partial orderings, total orderings and wellorderings. \item Theory \isa{OrderArith} defines orderings on sum and product sets. These can be used to define ordinal arithmetic and have applications to cardinal arithmetic. \item Theory \isa{OrderType} defines order types. Every wellordering is equivalent to a unique ordinal, which is its order type. \item Theory \isa{Cardinal} defines equipollence and cardinal numbers. \item Theory \isa{CardinalArith} defines cardinal addition and multiplication, and proves their elementary laws. It proves that there is no greatest cardinal. It also proves a deep result, namely$\kappa\otimes\kappa=\kappa$for every infinite cardinal~$\kappa$; see Kunen~\cite[page 29]{kunen80}. None of these results assume the Axiom of Choice, which complicates their proofs considerably. \end{itemize} The following developments involve the Axiom of Choice (AC): \begin{itemize} \item Theory \isa{AC} asserts the Axiom of Choice and proves some simple equivalent forms. \item Theory \isa{Zorn} proves Hausdorff's Maximal Principle, Zorn's Lemma and the Wellordering Theorem, following Abrial and Laffitte~\cite{abrial93}. \item Theory \isa{Cardinal\_AC} uses AC to prove simplified theorems about the cardinals. It also proves a theorem needed to justify infinitely branching datatype declarations: if$\kappa$is an infinite cardinal and$|X(\alpha)| \le \kappa$for all$\alpha<\kappa$then$|\union\sb{\alpha<\kappa} X(\alpha)| \le \kappa$. \item Theory \isa{InfDatatype} proves theorems to justify infinitely branching datatypes. Arbitrary index sets are allowed, provided their cardinalities have an upper bound. The theory also justifies some unusual cases of finite branching, involving the finite powerset operator and the finite function space operator. \end{itemize} \section{The examples directories} Directory \isa{HOL/IMP} contains a mechanised version of a semantic equivalence proof taken from Winskel~\cite{winskel93}. It formalises the denotational and operational semantics of a simple while-language, then proves the two equivalent. It contains several datatype and inductive definitions, and demonstrates their use. The directory \isa{ZF/ex} contains further developments in ZF set theory. Here is an overview; see the files themselves for more details. I describe much of this material in other publications~\cite{paulson-set-I,paulson-set-II,paulson-fixedpt-milner}. \begin{itemize} \item File \isa{misc.ML} contains miscellaneous examples such as Cantor's Theorem, the Schr\"oder-Bernstein Theorem and the Composition of homomorphisms' challenge~\cite{boyer86}. \item Theory \isa{Ramsey} proves the finite exponent 2 version of Ramsey's Theorem, following Basin and Kaufmann's presentation~\cite{basin91}. \item Theory \isa{Integ} develops a theory of the integers as equivalence classes of pairs of natural numbers. \item Theory \isa{Primrec} develops some computation theory. It inductively defines the set of primitive recursive functions and presents a proof that Ackermann's function is not primitive recursive. \item Theory \isa{Primes} defines the Greatest Common Divisor of two natural numbers and and the divides'' relation. \item Theory \isa{Bin} defines a datatype for two's complement binary integers, then proves rewrite rules to perform binary arithmetic. For instance,$1359\times {-}2468 = {-}3354012$takes 0.3 seconds. \item Theory \isa{BT} defines the recursive data structure$\isa{bt}(A)$, labelled binary trees. \item Theory \isa{Term} defines a recursive data structure for terms and term lists. These are simply finite branching trees. \item Theory \isa{TF} defines primitives for solving mutually recursive equations over sets. It constructs sets of trees and forests as an example, including induction and recursion rules that handle the mutual recursion. \item Theory \isa{Prop} proves soundness and completeness of propositional logic~\cite{paulson-set-II}. This illustrates datatype definitions, inductive definitions, structural induction and rule induction. \item Theory \isa{ListN} inductively defines the lists of$n$elements~\cite{paulin-tlca}. \item Theory \isa{Acc} inductively defines the accessible part of a relation~\cite{paulin-tlca}. \item Theory \isa{Comb} defines the datatype of combinators and inductively defines contraction and parallel contraction. It goes on to prove the Church-Rosser Theorem. This case study follows Camilleri and Melham~\cite{camilleri92}. \item Theory \isa{LList} defines lazy lists and a coinduction principle for proving equations between them. \end{itemize} \section{A proof about powersets}\label{sec:ZF-pow-example} To demonstrate high-level reasoning about subsets, let us prove the equation${\isa{Pow}(A)\cap \isa{Pow}(B)}= \isa{Pow}(A\cap B)$. Compared with first-order logic, set theory involves a maze of rules, and theorems have many different proofs. Attempting other proofs of the theorem might be instructive. This proof exploits the lattice properties of intersection. It also uses the monotonicity of the powerset operation, from \isa{ZF/mono.ML}: \begin{isabelle} \tdx{Pow_mono}: A \isasymsubseteq B ==> Pow(A) \isasymsubseteq Pow(B) \end{isabelle} We enter the goal and make the first step, which breaks the equation into two inclusions by extensionality:\index{*equalityI theorem} \begin{isabelle} \isacommand{lemma}\ "Pow(A\ Int\ B)\ =\ Pow(A)\ Int\ Pow(B)"\isanewline \ 1.\ Pow(A\ \isasyminter \ B)\ =\ Pow(A)\ \isasyminter \ Pow(B)\isanewline \isacommand{apply}\ (rule\ equalityI)\isanewline \ 1.\ Pow(A\ \isasyminter \ B)\ \isasymsubseteq \ Pow(A)\ \isasyminter \ Pow(B)\isanewline \ 2.\ Pow(A)\ \isasyminter \ Pow(B)\ \isasymsubseteq \ Pow(A\ \isasyminter \ B) \end{isabelle} Both inclusions could be tackled straightforwardly using \isa{subsetI}. A shorter proof results from noting that intersection forms the greatest lower bound:\index{*Int_greatest theorem} \begin{isabelle} \isacommand{apply}\ (rule\ Int\_greatest)\isanewline \ 1.\ Pow(A\ \isasyminter \ B)\ \isasymsubseteq \ Pow(A)\isanewline \ 2.\ Pow(A\ \isasyminter \ B)\ \isasymsubseteq \ Pow(B)\isanewline \ 3.\ Pow(A)\ \isasyminter \ Pow(B)\ \isasymsubseteq \ Pow(A\ \isasyminter \ B) \end{isabelle} Subgoal~1 follows by applying the monotonicity of \isa{Pow} to$A\int
B\subseteq A$; subgoal~2 follows similarly: \index{*Int_lower1 theorem}\index{*Int_lower2 theorem} \begin{isabelle} \isacommand{apply}\ (rule\ Int\_lower1\ [THEN\ Pow\_mono])\isanewline \ 1.\ Pow(A\ \isasyminter \ B)\ \isasymsubseteq \ Pow(B)\isanewline \ 2.\ Pow(A)\ \isasyminter \ Pow(B)\ \isasymsubseteq \ Pow(A\ \isasyminter \ B) \isanewline \isacommand{apply}\ (rule\ Int\_lower2\ [THEN\ Pow\_mono])\isanewline \ 1.\ Pow(A)\ \isasyminter \ Pow(B)\ \isasymsubseteq \ Pow(A\ \isasyminter \ B) \end{isabelle} We are left with the opposite inclusion, which we tackle in the straightforward way:\index{*subsetI theorem} \begin{isabelle} \isacommand{apply}\ (rule\ subsetI)\isanewline \ 1.\ \isasymAnd x.\ x\ \isasymin \ Pow(A)\ \isasyminter \ Pow(B)\ \isasymLongrightarrow \ x\ \isasymin \ Pow(A\ \isasyminter \ B) \end{isabelle} The subgoal is to show$x\in \isa{Pow}(A\cap B)$assuming$x\in\isa{Pow}(A)\cap \isa{Pow}(B)$; eliminating this assumption produces two subgoals. The rule \tdx{IntE} treats the intersection like a conjunction instead of unfolding its definition. \begin{isabelle} \isacommand{apply}\ (erule\ IntE)\isanewline \ 1.\ \isasymAnd x.\ \isasymlbrakk x\ \isasymin \ Pow(A);\ x\ \isasymin \ Pow(B)\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymin \ Pow(A\ \isasyminter \ B) \end{isabelle} The next step replaces the \isa{Pow} by the subset relation~($\subseteq$).\index{*PowI theorem} \begin{isabelle} \isacommand{apply}\ (rule\ PowI)\isanewline \ 1.\ \isasymAnd x.\ \isasymlbrakk x\ \isasymin \ Pow(A);\ x\ \isasymin \ Pow(B)\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymsubseteq \ A\ \isasyminter \ B% \end{isabelle} We perform the same replacement in the assumptions. This is a good demonstration of the tactic \ttindex{drule}:\index{*PowD theorem} \begin{isabelle} \isacommand{apply}\ (drule\ PowD)+\isanewline \ 1.\ \isasymAnd x.\ \isasymlbrakk x\ \isasymsubseteq \ A;\ x\ \isasymsubseteq \ B\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymsubseteq \ A\ \isasyminter \ B% \end{isabelle} The assumptions are that$x$is a lower bound of both$A$and~$B$, but$A\int B$is the greatest lower bound:\index{*Int_greatest theorem} \begin{isabelle} \isacommand{apply}\ (rule\ Int\_greatest)\isanewline \ 1.\ \isasymAnd x.\ \isasymlbrakk x\ \isasymsubseteq \ A;\ x\ \isasymsubseteq \ B\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymsubseteq \ A\isanewline \ 2.\ \isasymAnd x.\ \isasymlbrakk x\ \isasymsubseteq \ A;\ x\ \isasymsubseteq \ B\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymsubseteq \ B% \end{isabelle} To conclude the proof, we clear up the trivial subgoals: \begin{isabelle} \isacommand{apply}\ (assumption+)\isanewline \isacommand{done}% \end{isabelle} We could have performed this proof instantly by calling \ttindex{blast}: \begin{isabelle} \isacommand{lemma}\ "Pow(A\ Int\ B)\ =\ Pow(A)\ Int\ Pow(B)"\isanewline \isacommand{by} \end{isabelle} Past researchers regarded this as a difficult proof, as indeed it is if all the symbols are replaced by their definitions. \goodbreak \section{Monotonicity of the union operator} For another example, we prove that general union is monotonic:${C\subseteq D}$implies$\bigcup(C)\subseteq \bigcup(D)$. To begin, we tackle the inclusion using \tdx{subsetI}: \begin{isabelle} \isacommand{lemma}\ "C\isasymsubseteq D\ ==>\ Union(C)\ \isasymsubseteq \ Union(D)"\isanewline \isacommand{apply}\ (rule\ subsetI)\isanewline \ 1.\ \isasymAnd x.\ \isasymlbrakk C\ \isasymsubseteq \ D;\ x\ \isasymin \ \isasymUnion C\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymin \ \isasymUnion D% \end{isabelle} Big union is like an existential quantifier --- the occurrence in the assumptions must be eliminated early, since it creates parameters. \index{*UnionE theorem} \begin{isabelle} \isacommand{apply}\ (erule\ UnionE)\isanewline \ 1.\ \isasymAnd x\ B.\ \isasymlbrakk C\ \isasymsubseteq \ D;\ x\ \isasymin \ B;\ B\ \isasymin \ C\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymin \ \isasymUnion D% \end{isabelle} Now we may apply \tdx{UnionI}, which creates an unknown involving the parameters. To show \isa{x\ \isasymin \ \isasymUnion D} it suffices to show that~\isa{x} belongs to some element, say~\isa{?B2(x,B)}, of~\isa{D}\@. \begin{isabelle} \isacommand{apply}\ (rule\ UnionI)\isanewline \ 1.\ \isasymAnd x\ B.\ \isasymlbrakk C\ \isasymsubseteq \ D;\ x\ \isasymin \ B;\ B\ \isasymin \ C\isasymrbrakk \ \isasymLongrightarrow \ ?B2(x,\ B)\ \isasymin \ D\isanewline \ 2.\ \isasymAnd x\ B.\ \isasymlbrakk C\ \isasymsubseteq \ D;\ x\ \isasymin \ B;\ B\ \isasymin \ C\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymin \ ?B2(x,\ B) \end{isabelle} Combining the rule \tdx{subsetD} with the assumption \isa{C\ \isasymsubseteq \ D} yields$\Var{a}\in C \Imp \Var{a}\in D$, which reduces subgoal~1. Note that \isa{erule} removes the subset assumption. \begin{isabelle} \isacommand{apply}\ (erule\ subsetD)\isanewline \ 1.\ \isasymAnd x\ B.\ \isasymlbrakk x\ \isasymin \ B;\ B\ \isasymin \ C\isasymrbrakk \ \isasymLongrightarrow \ ?B2(x,\ B)\ \isasymin \ C\isanewline \ 2.\ \isasymAnd x\ B.\ \isasymlbrakk C\ \isasymsubseteq \ D;\ x\ \isasymin \ B;\ B\ \isasymin \ C\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymin \ ?B2(x,\ B) \end{isabelle} The rest is routine. Observe how the first call to \isa{assumption} instantiates \isa{?B2(x,B)} to~\isa{B}\@. \begin{isabelle} \isacommand{apply}\ assumption\ \isanewline \ 1.\ \isasymAnd x\ B.\ \isasymlbrakk C\ \isasymsubseteq \ D;\ x\ \isasymin \ B;\ B\ \isasymin \ C\isasymrbrakk \ \isasymLongrightarrow \ x\ \isasymin \ B% \isanewline \isacommand{apply}\ assumption\ \isanewline No\ subgoals!\isanewline \isacommand{done}% \end{isabelle} Again, \isa{blast} can prove this theorem in one step. The theory \isa{ZF/equalities.thy} has many similar proofs. Reasoning about general intersection can be difficult because of its anomalous behaviour on the empty set. However, \isa{blast} copes well with these. Here is a typical example, borrowed from Devlin~\cite[page 12]{devlin79}: $a\in C \,\Imp\, \inter@{x\in C} \Bigl(A(x) \int B(x)\Bigr) = \Bigl(\inter@{x\in C} A(x)\Bigr) \int \Bigl(\inter@{x\in C} B(x)\Bigr)$ \section{Low-level reasoning about functions} The derived rules \isa{lamI}, \isa{lamE}, \isa{lam_type}, \isa{beta} and \isa{eta} support reasoning about functions in a$\lambda$-calculus style. This is generally easier than regarding functions as sets of ordered pairs. But sometimes we must look at the underlying representation, as in the following proof of~\tdx{fun_disjoint_apply1}. This states that if$f$and~$g$are functions with disjoint domains~$A$and~$C$, and if$a\in A$, then$(f\un g)a = fa$: \begin{isabelle} \isacommand{lemma}\ "[|\ a\ \isasymin \ A;\ \ f\ \isasymin \ A->B;\ \ g\ \isasymin \ C->D;\ \ A\ \isasyminter \ C\ =\ 0\ |] \isanewline \ \ \ \ \ \ \ \ ==>\ (f\ \isasymunion \ g)a\ =\ fa" \end{isabelle} Using \tdx{apply_equality}, we reduce the equality to reasoning about ordered pairs. The second subgoal is to verify that \isa{f\ \isasymunion \ g} is a function, since \isa{Pi(?A,?B)} denotes a dependent function space. \begin{isabelle} \isacommand{apply}\ (rule\ apply\_equality)\isanewline \ 1.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \isanewline \isaindent{\ 1.\ }\isasymLongrightarrow \ \isasymlangle a,\ f\ \ a\isasymrangle \ \isasymin \ f\ \isasymunion \ g\isanewline \ 2.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \isanewline \isaindent{\ 2.\ }\isasymLongrightarrow \ f\ \isasymunion \ g\ \isasymin \ Pi(?A,\ ?B) \end{isabelle} We must show that the pair belongs to~$f$or~$g$; by~\tdx{UnI1} we choose~$f$: \begin{isabelle} \isacommand{apply}\ (rule\ UnI1)\isanewline \ 1.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ \isasymlangle a,\ f\ \ a\isasymrangle \ \isasymin \ f\isanewline \ 2.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \isanewline \isaindent{\ 2.\ }\isasymLongrightarrow \ f\ \isasymunion \ g\ \isasymin \ Pi(?A,\ ?B) \end{isabelle} To show$\pair{a,f`a}\in f$we use \tdx{apply_Pair}, which is essentially the converse of \tdx{apply_equality}: \begin{isabelle} \isacommand{apply}\ (rule\ apply\_Pair)\isanewline \ 1.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ f\ \isasymin \ Pi(?A2,?B2)\isanewline \ 2.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ a\ \isasymin \ ?A2\isanewline \ 3.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \isanewline \isaindent{\ 3.\ }\isasymLongrightarrow \ f\ \isasymunion \ g\ \isasymin \ Pi(?A,\ ?B) \end{isabelle} Using the assumptions$f\in A\to B$and$a\in A$, we solve the two subgoals from \tdx{apply_Pair}. Recall that a$\Pi$-set is merely a generalized function space, and observe that~{\tt?A2} gets instantiated to~\isa{A}. \begin{isabelle} \isacommand{apply}\ assumption\ \isanewline \ 1.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ a\ \isasymin \ A\isanewline \ 2.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \isanewline \isaindent{\ 2.\ }\isasymLongrightarrow \ f\ \isasymunion \ g\ \isasymin \ Pi(?A,\ ?B) \isanewline \isacommand{apply}\ assumption\ \isanewline \ 1.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \isanewline \isaindent{\ 1.\ }\isasymLongrightarrow \ f\ \isasymunion \ g\ \isasymin \ Pi(?A,\ ?B) \end{isabelle} To construct functions of the form$f\un g\$, we apply
\tdx{fun_disjoint_Un}:
\begin{isabelle}
\isacommand{apply}\ (rule\ fun\_disjoint\_Un)\isanewline
\ 1.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ f\ \isasymin \ ?A3\ \isasymrightarrow \ ?B3\isanewline
\ 2.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ g\ \isasymin \ ?C3\ \isasymrightarrow \ ?D3\isanewline
\ 3.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ ?A3\ \isasyminter \ ?C3\ =\ 0
\end{isabelle}
The remaining subgoals are instances of the assumptions.  Again, observe how
unknowns become instantiated:
\begin{isabelle}
\isacommand{apply}\ assumption\ \isanewline
\ 1.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ g\ \isasymin \ ?C3\ \isasymrightarrow \ ?D3\isanewline
\ 2.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ A\ \isasyminter \ ?C3\ =\ 0
\isanewline
\isacommand{apply}\ assumption\ \isanewline
\ 1.\ \isasymlbrakk a\ \isasymin \ A;\ f\ \isasymin \ A\ \isasymrightarrow \ B;\ g\ \isasymin \ C\ \isasymrightarrow \ D;\ A\ \isasyminter \ C\ =\ 0\isasymrbrakk \ \isasymLongrightarrow \ A\ \isasyminter \ C\ =\ 0
\isanewline
\isacommand{apply}\ assumption\ \isanewline
No\ subgoals!\isanewline
\isacommand{done}
\end{isabelle}
See the theories \isa{ZF/func.thy} and \isa{ZF/WF.thy} for more