doc-src/TutorialI/Misc/AdvancedInd.thy
author nipkow
Fri Sep 01 19:09:44 2000 +0200 (2000-09-01)
changeset 9792 bbefb6ce5cb2
parent 9723 a977245dfc8a
child 9834 109b11c4e77e
permissions -rw-r--r--
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(*<*)
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theory AdvancedInd = Main:;
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(*>*)
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text{*\noindent
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Now that we have learned about rules and logic, we take another look at the
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finer points of induction. The two questions we answer are: what to do if the
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proposition to be proved is not directly amenable to induction, and how to
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utilize and even derive new induction schemas.
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*};
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subsection{*Massaging the proposition\label{sec:ind-var-in-prems}*};
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text{*
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\noindent
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So far we have assumed that the theorem we want to prove is already in a form
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that is amenable to induction, but this is not always the case:
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*};
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lemma "xs \\<noteq> [] \\<Longrightarrow> hd(rev xs) = last xs";
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apply(induct_tac xs);
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txt{*\noindent
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(where @{term"hd"} and @{term"last"} return the first and last element of a
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non-empty list)
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produces the warning
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\begin{quote}\tt
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Induction variable occurs also among premises!
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\end{quote}
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and leads to the base case
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\begin{isabelle}
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\ 1.\ xs\ {\isasymnoteq}\ []\ {\isasymLongrightarrow}\ hd\ (rev\ [])\ =\ last\ []
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\end{isabelle}
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which, after simplification, becomes
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\begin{isabelle}
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\ 1.\ xs\ {\isasymnoteq}\ []\ {\isasymLongrightarrow}\ hd\ []\ =\ last\ []
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\end{isabelle}
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We cannot prove this equality because we do not know what @{term"hd"} and
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@{term"last"} return when applied to @{term"[]"}.
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The point is that we have violated the above warning. Because the induction
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formula is only the conclusion, the occurrence of @{term"xs"} in the premises is
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not modified by induction. Thus the case that should have been trivial
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becomes unprovable. Fortunately, the solution is easy:
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\begin{quote}
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\emph{Pull all occurrences of the induction variable into the conclusion
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using @{text"\<longrightarrow>"}.}
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\end{quote}
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This means we should prove
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*};
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(*<*)oops;(*>*)
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lemma hd_rev: "xs \\<noteq> [] \\<longrightarrow> hd(rev xs) = last xs";
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(*<*)
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by(induct_tac xs, auto);
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(*>*)
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text{*\noindent
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This time, induction leaves us with the following base case
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\begin{isabelle}
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\ 1.\ []\ {\isasymnoteq}\ []\ {\isasymlongrightarrow}\ hd\ (rev\ [])\ =\ last\ []
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\end{isabelle}
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which is trivial, and @{text"auto"} finishes the whole proof.
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If @{thm[source]hd_rev} is meant to be a simplification rule, you are
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done. But if you really need the @{text"\<Longrightarrow>"}-version of
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@{thm[source]hd_rev}, for example because you want to apply it as an
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introduction rule, you need to derive it separately, by combining it with
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modus ponens:
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*};
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lemmas hd_revI = hd_rev[THEN mp];
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text{*\noindent
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which yields the lemma we originally set out to prove.
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In case there are multiple premises $A@1$, \dots, $A@n$ containing the
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induction variable, you should turn the conclusion $C$ into
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\[ A@1 \longrightarrow \cdots A@n \longrightarrow C \]
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(see the remark?? in \S\ref{??}).
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Additionally, you may also have to universally quantify some other variables,
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which can yield a fairly complex conclusion.
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Here is a simple example (which is proved by @{text"blast"}):
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*};
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lemma simple: "\\<forall>y. A y \\<longrightarrow> B y \<longrightarrow> B y & A y";
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(*<*)by blast;(*>*)
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text{*\noindent
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You can get the desired lemma by explicit
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application of modus ponens and @{thm[source]spec}:
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*};
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lemmas myrule = simple[THEN spec, THEN mp, THEN mp];
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text{*\noindent
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or the wholesale stripping of @{text"\<forall>"} and
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@{text"\<longrightarrow>"} in the conclusion via @{text"rulify"} 
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*};
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lemmas myrule = simple[rulify];
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text{*\noindent
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yielding @{thm"myrule"[no_vars]}.
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You can go one step further and include these derivations already in the
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statement of your original lemma, thus avoiding the intermediate step:
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*};
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lemma myrule[rulify]:  "\\<forall>y. A y \\<longrightarrow> B y \<longrightarrow> B y & A y";
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(*<*)
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by blast;
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(*>*)
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text{*
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\bigskip
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A second reason why your proposition may not be amenable to induction is that
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you want to induct on a whole term, rather than an individual variable. In
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general, when inducting on some term $t$ you must rephrase the conclusion as
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\[ \forall y@1 \dots y@n.~ x = t \longrightarrow C \] where $y@1 \dots y@n$
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are the free variables in $t$ and $x$ is new, and perform induction on $x$
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afterwards. An example appears below.
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*};
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subsection{*Beyond structural and recursion induction*};
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text{*
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So far, inductive proofs where by structural induction for
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primitive recursive functions and recursion induction for total recursive
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functions. But sometimes structural induction is awkward and there is no
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recursive function in sight either that could furnish a more appropriate
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induction schema. In such cases some existing standard induction schema can
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be helpful. We show how to apply such induction schemas by an example.
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Structural induction on @{typ"nat"} is
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usually known as ``mathematical induction''. There is also ``complete
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induction'', where you must prove $P(n)$ under the assumption that $P(m)$
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holds for all $m<n$. In Isabelle, this is the theorem @{thm[source]less_induct}:
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\begin{quote}
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@{thm[display]"less_induct"[no_vars]}
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\end{quote}
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Here is an example of its application.
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*};
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consts f :: "nat => nat";
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axioms f_ax: "f(f(n)) < f(Suc(n))";
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text{*\noindent
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From the above axiom\footnote{In general, the use of axioms is strongly
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discouraged, because of the danger of inconsistencies. The above axiom does
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not introduce an inconsistency because, for example, the identity function
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satisfies it.}
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for @{term"f"} it follows that @{prop"n <= f n"}, which can
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be proved by induction on @{term"f n"}. Following the recipy outlined
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above, we have to phrase the proposition as follows to allow induction:
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*};
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lemma f_incr_lem: "\\<forall>i. k = f i \\<longrightarrow> i \\<le> f i";
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txt{*\noindent
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To perform induction on @{term"k"} using @{thm[source]less_induct}, we use the same
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general induction method as for recursion induction (see
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\S\ref{sec:recdef-induction}):
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*};
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apply(induct_tac k rule:less_induct);
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(*<*)
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apply(rule allI);
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apply(case_tac i);
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 apply(simp);
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(*>*)
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txt{*\noindent
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which leaves us with the following proof state:
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\begin{isabelle}
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\ 1.\ {\isasymAnd}\mbox{n}.\ {\isasymforall}\mbox{m}.\ \mbox{m}\ <\ \mbox{n}\ {\isasymlongrightarrow}\ ({\isasymforall}\mbox{i}.\ \mbox{m}\ =\ f\ \mbox{i}\ {\isasymlongrightarrow}\ \mbox{i}\ {\isasymle}\ f\ \mbox{i})\isanewline
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\ \ \ \ \ \ \ {\isasymLongrightarrow}\ {\isasymforall}\mbox{i}.\ \mbox{n}\ =\ f\ \mbox{i}\ {\isasymlongrightarrow}\ \mbox{i}\ {\isasymle}\ f\ \mbox{i}
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\end{isabelle}
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After stripping the @{text"\<forall>i"}, the proof continues with a case
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distinction on @{term"i"}. The case @{prop"i = 0"} is trivial and we focus on
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the other case:
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\begin{isabelle}
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\ 1.\ {\isasymAnd}\mbox{n}\ \mbox{i}\ \mbox{nat}.\isanewline
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\ \ \ \ \ \ \ {\isasymlbrakk}{\isasymforall}\mbox{m}.\ \mbox{m}\ <\ \mbox{n}\ {\isasymlongrightarrow}\ ({\isasymforall}\mbox{i}.\ \mbox{m}\ =\ f\ \mbox{i}\ {\isasymlongrightarrow}\ \mbox{i}\ {\isasymle}\ f\ \mbox{i});\ \mbox{i}\ =\ Suc\ \mbox{nat}{\isasymrbrakk}\isanewline
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\ \ \ \ \ \ \ {\isasymLongrightarrow}\ \mbox{n}\ =\ f\ \mbox{i}\ {\isasymlongrightarrow}\ \mbox{i}\ {\isasymle}\ f\ \mbox{i}
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\end{isabelle}
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*};
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by(blast intro!: f_ax Suc_leI intro:le_less_trans);
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text{*\noindent
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It is not surprising if you find the last step puzzling.
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The proof goes like this (writing @{term"j"} instead of @{typ"nat"}).
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Since @{prop"i = Suc j"} it suffices to show
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@{prop"j < f(Suc j)"} (by @{thm[source]Suc_leI}: @{thm"Suc_leI"[no_vars]}). This is
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proved as follows. From @{thm[source]f_ax} we have @{prop"f (f j) < f (Suc j)"}
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(1) which implies @{prop"f j <= f (f j)"} (by the induction hypothesis).
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Using (1) once more we obtain @{prop"f j < f(Suc j)"} (2) by transitivity
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(@{thm[source]le_less_trans}: @{thm"le_less_trans"[no_vars]}).
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Using the induction hypothesis once more we obtain @{prop"j <= f j"}
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which, together with (2) yields @{prop"j < f (Suc j)"} (again by
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@{thm[source]le_less_trans}).
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This last step shows both the power and the danger of automatic proofs: they
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will usually not tell you how the proof goes, because it can be very hard to
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translate the internal proof into a human-readable format. Therefore
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\S\ref{sec:part2?} introduces a language for writing readable yet concise
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proofs.
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We can now derive the desired @{prop"i <= f i"} from @{text"f_incr"}:
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*};
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lemmas f_incr = f_incr_lem[rulify, OF refl];
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text{*\noindent
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The final @{thm[source]refl} gets rid of the premise @{text"?k = f ?i"}. Again,
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we could have included this derivation in the original statement of the lemma:
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*};
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lemma f_incr[rulify, OF refl]: "\\<forall>i. k = f i \\<longrightarrow> i \\<le> f i";
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(*<*)oops;(*>*)
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text{*
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\begin{exercise}
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From the above axiom and lemma for @{term"f"} show that @{term"f"} is the
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identity.
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\end{exercise}
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In general, @{text"induct_tac"} can be applied with any rule $r$
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whose conclusion is of the form ${?}P~?x@1 \dots ?x@n$, in which case the
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format is
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\begin{quote}
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\isacommand{apply}@{text"(induct_tac"} $y@1 \dots y@n$ @{text"rule:"} $r$@{text")"}
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\end{quote}\index{*induct_tac}%
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where $y@1, \dots, y@n$ are variables in the first subgoal.
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In fact, @{text"induct_tac"} even allows the conclusion of
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$r$ to be an (iterated) conjunction of formulae of the above form, in
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which case the application is
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\begin{quote}
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\isacommand{apply}@{text"(induct_tac"} $y@1 \dots y@n$ @{text"and"} \dots\ @{text"and"} $z@1 \dots z@m$ @{text"rule:"} $r$@{text")"}
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\end{quote}
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*};
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subsection{*Derivation of new induction schemas*};
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text{*\label{sec:derive-ind}
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Induction schemas are ordinary theorems and you can derive new ones
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whenever you wish.  This section shows you how to, using the example
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of @{thm[source]less_induct}. Assume we only have structural induction
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available for @{typ"nat"} and want to derive complete induction. This
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requires us to generalize the statement first:
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*};
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lemma induct_lem: "(\\<And>n::nat. \\<forall>m<n. P m \\<Longrightarrow> P n) \\<Longrightarrow> \\<forall>m<n. P m";
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apply(induct_tac n);
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txt{*\noindent
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The base case is trivially true. For the induction step (@{prop"m <
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Suc n"}) we distinguish two cases: @{prop"m < n"} is true by induction
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hypothesis and @{prop"m = n"} follow from the assumption again using
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the induction hypothesis:
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*};
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apply(blast);
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(* apply(blast elim:less_SucE); *)
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ML"set quick_and_dirty"
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sorry;
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ML"reset quick_and_dirty"
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text{*\noindent
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The elimination rule @{thm[source]less_SucE} expresses the case distinction:
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\begin{quote}
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@{thm[display]"less_SucE"[no_vars]}
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\end{quote}
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Now it is straightforward to derive the original version of
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@{thm[source]less_induct} by manipulting the conclusion of the above lemma:
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instantiate @{term"n"} by @{term"Suc n"} and @{term"m"} by @{term"n"} and
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remove the trivial condition @{prop"n < Sc n"}. Fortunately, this
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happens automatically when we add the lemma as a new premise to the
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desired goal:
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*};
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theorem less_induct: "(\\<And>n::nat. \\<forall>m<n. P m \\<Longrightarrow> P n) \\<Longrightarrow> P n";
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by(insert induct_lem, blast);
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text{*\noindent
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Finally we should mention that HOL already provides the mother of all
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inductions, \emph{wellfounded induction} (@{thm[source]wf_induct}):
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\begin{quote}
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@{thm[display]"wf_induct"[no_vars]}
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\end{quote}
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where @{term"wf r"} means that the relation @{term"r"} is wellfounded.
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For example @{thm[source]less_induct} is the special case where @{term"r"} is
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@{text"<"} on @{typ"nat"}. For details see the library.
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*};
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(*<*)
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end
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(*>*)